| TOC |
|
By submitting this Internet-Draft, each author represents that any applicable patent or other IPR claims of which he or she is aware have been or will be disclosed, and any of which he or she becomes aware will be disclosed, in accordance with Section 6 of BCP 79.
Internet-Drafts are working documents of the Internet Engineering Task Force (IETF), its areas, and its working groups. Note that other groups may also distribute working documents as Internet-Drafts.
Internet-Drafts are draft documents valid for a maximum of six months and may be updated, replaced, or obsoleted by other documents at any time. It is inappropriate to use Internet-Drafts as reference material or to cite them other than as “work in progress.”
The list of current Internet-Drafts can be accessed at http://www.ietf.org/ietf/1id-abstracts.txt.
The list of Internet-Draft Shadow Directories can be accessed at http://www.ietf.org/shadow.html.
This Internet-Draft will expire on November 8, 2008.
Copyright © The IETF Trust (2008).
This Internet-Draft describes NFS version 4 minor version one, including features retained from the base protocol and protocol extensions made subsequently. Major extensions introduced in NFS version 4 minor version one include: Sessions, Directory Delegations, and parallel NFS (pNFS).
The key words "MUST", "MUST NOT", "REQUIRED", "SHALL", "SHALL NOT", "SHOULD", "SHOULD NOT", "RECOMMENDED", "MAY", and "OPTIONAL" in this document are to be interpreted as described in RFC 2119 (Bradner, S., “Key words for use in RFCs to Indicate Requirement Levels,” March 1997.) [1].
1.
Introduction
1.1.
The NFS Version 4 Minor Version 1 Protocol
1.2.
Scope of this Document
1.3.
NFSv4 Goals
1.4.
NFSv4.1 Goals
1.5.
General Definitions
1.6.
Overview of NFSv4.1 Features
1.6.1.
RPC and Security
1.6.2.
Protocol Structure
1.6.3.
File System Model
1.6.4.
Locking Facilities
1.7.
Differences from NFSv4.0
2.
Core Infrastructure
2.1.
Introduction
2.2.
RPC and XDR
2.2.1.
RPC-based Security
2.3.
COMPOUND and CB_COMPOUND
2.4.
Client Identifiers and Client Owners
2.4.1.
Upgrade from NFSv4.0 to NFSv4.1
2.4.2.
Server Release of Client ID
2.4.3.
Resolving Client Owner Conflicts
2.5.
Server Owners
2.6.
Security Service Negotiation
2.6.1.
NFSv4.1 Security Tuples
2.6.2.
SECINFO and SECINFO_NO_NAME
2.6.3.
Security Error
2.7.
Minor Versioning
2.8.
Non-RPC-based Security Services
2.8.1.
Authorization
2.8.2.
Auditing
2.8.3.
Intrusion Detection
2.9.
Transport Layers
2.9.1.
REQUIRED and RECOMMENDED Properties of Transports
2.9.2.
Client and Server Transport Behavior
2.9.3.
Ports
2.10.
Session
2.10.1.
Motivation and Overview
2.10.2.
NFSv4 Integration
2.10.3.
Channels
2.10.4.
Trunking
2.10.5.
Exactly Once Semantics
2.10.6.
RDMA Considerations
2.10.7.
Sessions Security
2.10.8.
The SSV GSS Mechanism
2.10.9.
Session Mechanics - Steady State
2.10.10.
Session Inactivity Timer
2.10.11.
Session Mechanics - Recovery
2.10.12.
Parallel NFS and Sessions
3.
Protocol Constants and Data Types
3.1.
Basic Constants
3.2.
Basic Data Types
3.3.
Structured Data Types
4.
Filehandles
4.1.
Obtaining the First Filehandle
4.1.1.
Root Filehandle
4.1.2.
Public Filehandle
4.2.
Filehandle Types
4.2.1.
General Properties of a Filehandle
4.2.2.
Persistent Filehandle
4.2.3.
Volatile Filehandle
4.3.
One Method of Constructing a Volatile Filehandle
4.4.
Client Recovery from Filehandle Expiration
5.
File Attributes
5.1.
REQUIRED Attributes
5.2.
RECOMMENDED Attributes
5.3.
Named Attributes
5.4.
Classification of Attributes
5.5.
Set-Only and Get-Only Attributes
5.6.
REQUIRED Attributes - List and Definition References
5.7.
RECOMMENDED Attributes - List and Definition References
5.8.
Attribute
Definitions
5.8.1.
Definitions of REQUIRED Attributes
5.8.2.
Definitions of Uncategorized RECOMMENDED Attributes
5.9.
Interpreting owner and owner_group
5.10.
Character Case Attributes
5.11.
Directory Notification Attributes
5.12.
pNFS Attribute Definitions
5.13.
Retention Attributes
6.
Access Control Attributes
6.1.
Goals
6.2.
File Attributes Discussion
6.2.1.
Attribute 12: acl
6.2.2.
Attribute 58: dacl
6.2.3.
Attribute 59: sacl
6.2.4.
Attribute 33: mode
6.2.5.
Attribute 74: mode_set_masked
6.3.
Common Methods
6.3.1.
Interpreting an ACL
6.3.2.
Computing a Mode Attribute from an ACL
6.4.
Requirements
6.4.1.
Setting the mode and/or ACL Attributes
6.4.2.
Retrieving the mode and/or ACL Attributes
6.4.3.
Creating New Objects
7.
Single-server Namespace
7.1.
Server Exports
7.2.
Browsing Exports
7.3.
Server Pseudo File System
7.4.
Multiple Roots
7.5.
Filehandle Volatility
7.6.
Exported Root
7.7.
Mount Point Crossing
7.8.
Security Policy and Namespace Presentation
8.
State Management
8.1.
Client and Session ID
8.2.
Stateid Definition
8.2.1.
Stateid Types
8.2.2.
Stateid Structure
8.2.3.
Special Stateids
8.2.4.
Stateid Lifetime and Validation
8.2.5.
Stateid Use for I/O Operations
8.2.6.
Stateid Use for SETATTR Operations
8.3.
Lease Renewal
8.4.
Crash Recovery
8.4.1.
Client Failure and Recovery
8.4.2.
Server Failure and Recovery
8.4.3.
Network Partitions and Recovery
8.5.
Server Revocation of Locks
8.6.
Short and Long Leases
8.7.
Clocks, Propagation Delay, and Calculating Lease Expiration
8.8.
Obsolete Locking Infrastructure From NFSv4.0
9.
File Locking and Share Reservations
9.1.
Opens and Byte-Range Locks
9.1.1.
State-owner Definition
9.1.2.
Use of the Stateid and Locking
9.2.
Lock Ranges
9.3.
Upgrading and Downgrading Locks
9.4.
Stateid Seqid Values and Byte-Range Locks
9.5.
Issues with Multiple Open-Owners
9.6.
Blocking Locks
9.7.
Share Reservations
9.8.
OPEN/CLOSE Operations
9.9.
Open Upgrade and Downgrade
9.10.
Parallel OPENs
9.11.
Reclaim of Open and Byte-Range Locks
10.
Client-Side Caching
10.1.
Performance Challenges for Client-Side Caching
10.2.
Delegation and Callbacks
10.2.1.
Delegation Recovery
10.3.
Data Caching
10.3.1.
Data Caching and OPENs
10.3.2.
Data Caching and File Locking
10.3.3.
Data Caching and Mandatory File Locking
10.3.4.
Data Caching and File Identity
10.4.
Open Delegation
10.4.1.
Open Delegation and Data Caching
10.4.2.
Open Delegation and File Locks
10.4.3.
Handling of CB_GETATTR
10.4.4.
Recall of Open Delegation
10.4.5.
Clients that Fail to Honor Delegation Recalls
10.4.6.
Delegation Revocation
10.4.7.
Delegations via WANT_DELEGATION
10.5.
Data Caching and Revocation
10.5.1.
Revocation Recovery for Write Open Delegation
10.6.
Attribute Caching
10.7.
Data and Metadata Caching and Memory Mapped Files
10.8.
Name and Directory Caching without Directory Delegations
10.8.1.
Name Caching
10.8.2.
Directory Caching
10.9.
Directory Delegations
10.9.1.
Introduction to Directory Delegations
10.9.2.
Directory Delegation Design
10.9.3.
Attributes in Support of Directory Notifications
10.9.4.
Directory Delegation Recall
10.9.5.
Directory Delegation Recovery
11.
Multi-Server Namespace
11.1.
Location Attributes
11.2.
File System Presence or Absence
11.3.
Getting Attributes for an Absent File System
11.3.1.
GETATTR Within an Absent File System
11.3.2.
READDIR and Absent File Systems
11.4.
Uses of Location Information
11.4.1.
File System Replication
11.4.2.
File System Migration
11.4.3.
Referrals
11.5.
Location Entries and Server Identity
11.6.
Additional Client-side Considerations
11.7.
Effecting File System Transitions
11.7.1.
File System Transitions and Simultaneous Access
11.7.2.
Simultaneous Use and Transparent Transitions
11.7.3.
Filehandles and File System Transitions
11.7.4.
Fileids and File System Transitions
11.7.5.
Fsids and File System Transitions
11.7.6.
The Change Attribute and File System Transitions
11.7.7.
Lock State and File System Transitions
11.7.8.
Write Verifiers and File System Transitions
11.7.9.
Readdir Cookies and Verifiers and File System Transitions
11.7.10.
File System Data and File System Transitions
11.8.
Effecting File System Referrals
11.8.1.
Referral Example (LOOKUP)
11.8.2.
Referral Example (READDIR)
11.9.
The Attribute fs_locations
11.10.
The Attribute fs_locations_info
11.10.1.
The fs_locations_server4 Structure
11.10.2.
The fs_locations_info4 Structure
11.10.3.
The fs_locations_item4 Structure
11.11.
The Attribute fs_status
12.
Parallel NFS (pNFS)
12.1.
Introduction
12.2.
pNFS Definitions
12.2.1.
Metadata
12.2.2.
Metadata Server
12.2.3.
pNFS Client
12.2.4.
Storage Device
12.2.5.
Storage Protocol
12.2.6.
Control Protocol
12.2.7.
Layout Types
12.2.8.
Layout
12.2.9.
Layout Iomode
12.2.10.
Device IDs
12.3.
pNFS Operations
12.4.
pNFS Attributes
12.5.
Layout Semantics
12.5.1.
Guarantees Provided by Layouts
12.5.2.
Getting a Layout
12.5.3.
Layout Stateid
12.5.4.
Committing a Layout
12.5.5.
Recalling a Layout
12.5.6.
Revoking Layouts
12.5.7.
Metadata Server Write Propagation
12.6.
pNFS Mechanics
12.7.
Recovery
12.7.1.
Recovery from Client Restart
12.7.2.
Dealing with Lease Expiration on the Client
12.7.3.
Dealing with Loss of Layout State on the Metadata Server
12.7.4.
Recovery from Metadata Server Restart
12.7.5.
Operations During Metadata Server Grace Period
12.7.6.
Storage Device Recovery
12.8.
Metadata and Storage Device Roles
12.9.
Security Considerations for pNFS
13.
PNFS: NFSv4.1 File Layout Type
13.1.
Client ID and Session Considerations
13.1.1.
Sessions Considerations for Data Servers
13.2.
File Layout Definitions
13.3.
File Layout Data Types
13.4.
Interpreting the File Layout
13.4.1.
Determining the Stripe Unit Number
13.4.2.
Interpreting the File Layout Using Sparse Packing
13.4.3.
Interpreting the File Layout Using Dense Packing
13.4.4.
Sparse and Dense Stripe Unit Packing
13.5.
Data Server Multipathing
13.6.
Operations Sent to NFSv4.1 Data Servers
13.7.
COMMIT Through Metadata Server
13.8.
The Layout Iomode
13.9.
Metadata and Data Server State Coordination
13.9.1.
Global Stateid Requirements
13.9.2.
Data Server State Propagation
13.10.
Data Server Component File Size
13.11.
Layout Revocation and Fencing
13.12.
Security Considerations for the File Layout Type
14.
Internationalization
14.1.
Stringprep profile for the utf8str_cs type
14.2.
Stringprep profile for the utf8str_cis type
14.3.
Stringprep profile for the utf8str_mixed type
14.4.
UTF-8 Capabilities
14.5.
UTF-8 Related Errors
15.
Error Values
15.1.
Error Definitions
15.1.1.
General Errors
15.1.2.
Filehandle Errors
15.1.3.
Compound Structure Errors
15.1.4.
File System Errors
15.1.5.
State Management Errors
15.1.6.
Security Errors
15.1.7.
Name Errors
15.1.8.
Locking Errors
15.1.9.
Reclaim Errors
15.1.10.
pNFS Errors
15.1.11.
Session Use Errors
15.1.12.
Session Management Errors
15.1.13.
Client Management Errors
15.1.14.
Delegation Errors
15.1.15.
Attribute Handling Errors
15.1.16.
Obsoleted Errors
15.2.
Operations and their valid errors
15.3.
Callback operations and their valid errors
15.4.
Errors and the operations that use them
16.
NFSv4.1 Procedures
16.1.
Procedure 0: NULL - No Operation
16.2.
Procedure 1: COMPOUND - Compound Operations
17.
Operations: REQUIRED, RECOMMENDED, or OPTIONAL
18.
NFSv4.1 Operations
18.1.
Operation 3: ACCESS - Check Access Rights
18.2.
Operation 4: CLOSE - Close File
18.3.
Operation 5: COMMIT - Commit Cached Data
18.4.
Operation 6: CREATE - Create a Non-Regular File Object
18.5.
Operation 7: DELEGPURGE - Purge Delegations Awaiting Recovery
18.6.
Operation 8: DELEGRETURN - Return Delegation
18.7.
Operation 9: GETATTR - Get Attributes
18.8.
Operation 10: GETFH - Get Current Filehandle
18.9.
Operation 11: LINK - Create Link to a File
18.10.
Operation 12: LOCK - Create Lock
18.11.
Operation 13: LOCKT - Test For Lock
18.12.
Operation 14: LOCKU - Unlock File
18.13.
Operation 15: LOOKUP - Lookup Filename
18.14.
Operation 16: LOOKUPP - Lookup Parent Directory
18.15.
Operation 17: NVERIFY - Verify Difference in Attributes
18.16.
Operation 18: OPEN - Open a Regular File
18.17.
Operation 19: OPENATTR - Open Named Attribute Directory
18.18.
Operation 21: OPEN_DOWNGRADE - Reduce Open File Access
18.19.
Operation 22: PUTFH - Set Current Filehandle
18.20.
Operation 23: PUTPUBFH - Set
Public Filehandle
18.21.
Operation 24: PUTROOTFH - Set Root Filehandle
18.22.
Operation 25: READ - Read from File
18.23.
Operation 26: READDIR - Read Directory
18.24.
Operation 27: READLINK - Read Symbolic Link
18.25.
Operation 28: REMOVE - Remove File System Object
18.26.
Operation 29: RENAME - Rename Directory Entry
18.27.
Operation 31: RESTOREFH - Restore Saved Filehandle
18.28.
Operation 32: SAVEFH - Save Current Filehandle
18.29.
Operation 33: SECINFO - Obtain Available Security
18.30.
Operation 34: SETATTR - Set Attributes
18.31.
Operation 37: VERIFY - Verify Same Attributes
18.32.
Operation 38: WRITE - Write to File
18.33.
Operation 40: BACKCHANNEL_CTL - Backchannel control
18.34.
Operation 41: BIND_CONN_TO_SESSION
18.35.
Operation 42: EXCHANGE_ID - Instantiate Client ID
18.36.
Operation 43: CREATE_SESSION - Create New Session and Confirm Client ID
18.37.
Operation 44: DESTROY_SESSION - Destroy existing session
18.38.
Operation 45: FREE_STATEID - Free stateid with no locks
18.39.
Operation 46: GET_DIR_DELEGATION - Get a directory delegation
18.40.
Operation 47: GETDEVICEINFO - Get Device Information
18.41.
Operation 48: GETDEVICELIST - Get All Device Mappings for a File System
18.42.
Operation 49: LAYOUTCOMMIT - Commit writes made using a layout
18.43.
Operation 50: LAYOUTGET - Get Layout Information
18.44.
Operation 51: LAYOUTRETURN - Release Layout Information
18.45.
Operation 52: SECINFO_NO_NAME - Get Security on Unnamed Object
18.46.
Operation 53: SEQUENCE - Supply per-procedure sequencing and control
18.47.
Operation 54: SET_SSV - Update SSV for a Client ID
18.48.
Operation 55: TEST_STATEID - Test stateids for validity
18.49.
Operation 56: WANT_DELEGATION - Request Delegation
18.50.
Operation 57: DESTROY_CLIENTID - Destroy existing client ID
18.51.
Operation 58: RECLAIM_COMPLETE - Indicates Reclaims Finished
18.52.
Operation 10044: ILLEGAL - Illegal operation
19.
NFSv4.1 Callback Procedures
19.1.
Procedure 0: CB_NULL - No Operation
19.2.
Procedure 1: CB_COMPOUND - Compound Operations
20.
NFSv4.1 Callback Operations
20.1.
Operation 3: CB_GETATTR - Get Attributes
20.2.
Operation 4: CB_RECALL - Recall a Delegation
20.3.
Operation 5: CB_LAYOUTRECALL - Recall Layout from Client
20.4.
Operation 6: CB_NOTIFY - Notify directory changes
20.5.
Operation 7: CB_PUSH_DELEG - Offer Delegation to Client
20.6.
Operation 8: CB_RECALL_ANY - Keep any N recallable objects
20.7.
Operation 9: CB_RECALLABLE_OBJ_AVAIL - Signal Resources for Recallable Objects
20.8.
Operation 10: CB_RECALL_SLOT - change flow control limits
20.9.
Operation 11: CB_SEQUENCE - Supply backchannel sequencing and control
20.10.
Operation 12: CB_WANTS_CANCELLED - Cancel Pending Delegation Wants
20.11.
Operation 13: CB_NOTIFY_LOCK - Notify of possible lock availability
20.12.
Operation 14: CB_NOTIFY_DEVICEID - Notify device ID changes
20.13.
Operation 10044: CB_ILLEGAL - Illegal Callback Operation
21.
Security Considerations
22.
IANA Considerations
22.1.
Named Attribute Definitions
22.2.
ONC RPC Network Identifiers (netids)
22.3.
Defining New Notifications
22.4.
Defining New Layout Types
22.5.
Path Variable Definitions
22.5.1.
Path Variable Values
22.5.2.
Path Variable Names
23.
References
23.1.
Normative References
23.2.
Informative References
Appendix A.
Acknowledgments
§
Authors' Addresses
§
Intellectual Property and Copyright Statements
| TOC |
| TOC |
The NFS version 4 minor version 1 (NFSv4.1) protocol is the second minor version of the NFS version 4 (NFSv4) protocol. The first minor version, NFSv4.0 is described in [21] (Shepler, S., Callaghan, B., Robinson, D., Thurlow, R., Beame, C., Eisler, M., and D. Noveck, “Network File System (NFS) version 4 Protocol,” April 2003.). It generally follows the guidelines for minor versioning model listed in Section 10 of RFC 3530. However, it diverges from guidelines 11 ("a client and server that supports minor version X must support minor versions 0 through X-1"), and 12 ("no features may be introduced as mandatory in a minor version"). These divergences are due to the introduction of the sessions model for managing non-idempotent operations and the RECLAIM_COMPLETE operation. These two new features are infrastructural in nature and simplify implementation of existing and other new features. Making them anything but REQUIRED would add undue complexity to protocol definition and implementation. NFSv4.1 accordingly updates the Minor Versioning guidelines (Minor Versioning).
As a minor version, NFSv4.1 is consistent with the overall goals for NFSv4, but extends the protocol so as to better meet those goals, based on experiences with NFSv4.0. In addition, NFSv4.1 has adopted some additional goals, which motivate some of the major extensions in NFSv4.1.
| TOC |
This document describes the NFSv4.1 protocol. With respect to NFSv4.0, this document does not:
| TOC |
The NFSv4 protocol is a further revision of the NFS protocol defined already by NFSv3 [22] (Callaghan, B., Pawlowski, B., and P. Staubach, “NFS Version 3 Protocol Specification,” June 1995.). It retains the essential characteristics of previous versions: easy recovery; independence of transport protocols, operating systems and file systems; simplicity; and good performance. NFSv4 has the following goals:
| TOC |
NFSv4.1 has the following goals, within the framework established by the overall NFSv4 goals.
| TOC |
The following definitions are provided for the purpose of providing an appropriate context for the reader.
- Byte
- This document defines a byte as an octet, i.e. a datum exactly 8 bits in length.
- Client
- The "client" is the entity that accesses the NFS server's resources. The client may be an application which contains the logic to access the NFS server directly. The client may also be the traditional operating system client that provides remote file system services for a set of applications.
A client is uniquely identified by a Client Owner.
With reference to file locking, the client is also the entity that maintains a set of locks on behalf of one or more applications. This client is responsible for crash or failure recovery for those locks it manages.
Note that multiple clients may share the same transport and connection and multiple clients may exist on the same network node.- Client ID
- A 64-bit quantity used as a unique, short-hand reference to a client supplied Verifier and client owner. The server is responsible for supplying the client ID.
- Client Owner
- The client owner is a unique string, opaque to the server, which identifies a client. Multiple network connections and source network addresses originating from those connections may share a client owner. The server is expected to treat requests from connnections with the same client owner as coming from the same client.
- File System
- The collection of objects on a server (as identified by the major identifier of a Server Owner, which is defined later in this section), that share the same fsid attribute (see Section 5.8.1.9 (Attribute 8: fsid)).
- Lease
- An interval of time defined by the server for which the client is irrevocably granted a lock. At the end of a lease period the lock may be revoked if the lease has not been extended. The lock must be revoked if a conflicting lock has been granted after the lease interval.
All leases granted by a server have the same fixed interval. Note that the fixed interval was chosen to alleviate the expense a server would have in maintaining state about variable length leases across server failures.- Lock
- The term "lock" is used to refer to byte-range (in UNIX environments, also known as record) locks, share reservations, delegations, or layouts unless specifically stated otherwise.
- Server
- The "Server" is the entity responsible for coordinating client access to a set of file systems and is identified by a Server owner. A server can span multiple network addresses.
- Server Owner
- The "Server Owner" identifies the server to the client. The server owner consists of a major and minor identifier. When the client has two connections each to a peer with the same major identifier, the client assumes both peers are the same server (the server namespace is the same via each connection), and assumes and lock state is sharable across both connections. When each peer has both the same major and minor identifier, the client assumes each connection might be associatable with the same session.
- Stable Storage
- NFSv4.1 servers must be able to recover without data loss from multiple power failures (including cascading power failures, that is, several power failures in quick succession), operating system failures, and hardware failure of components other than the storage medium itself (for example, disk, nonvolatile RAM).
Some examples of stable storage that are allowable for an NFS server include:
- Media commit of data, that is, the modified data has been successfully written to the disk media, for example, the disk platter.
- An immediate reply disk drive with battery-backed on- drive intermediate storage or uninterruptible power system (UPS).
- Server commit of data with battery-backed intermediate storage and recovery software.
- Cache commit with uninterruptible power system (UPS) and recovery software.
- Stateid
- A 128-bit quantity returned by a server that uniquely defines the open and locking state provided by the server for a specific open-owner or lock-owner/open-owner pair for a specific file and type of lock.
- Verifier
- A 64-bit quantity generated by the client that the server can use to determine if the client has restarted and lost all previous lock state.
| TOC |
To provide a reasonable context for the reader, the major features of the NFSv4.1 protocol will be reviewed in brief. This will be done to provide an appropriate context for both the reader who is familiar with the previous versions of the NFS protocol and the reader that is new to the NFS protocols. For the reader new to the NFS protocols, there is still a set of fundamental knowledge that is expected. The reader should be familiar with the XDR and RPC protocols as described in [2] (Eisler, M., “XDR: External Data Representation Standard,” May 2006.) and [3] (Srinivasan, R., “RPC: Remote Procedure Call Protocol Specification Version 2,” August 1995.). A basic knowledge of file systems and distributed file systems is expected as well.
In general this specification of NFSv4.1 will not distinguish those added in minor version one from those present in the base protocol but will treat NFSv4.1 as a unified whole. See Section 1.7 (Differences from NFSv4.0) for a summary of the differences between NFSv4.0 and NFSv4.1.
| TOC |
As with previous versions of NFS, the External Data Representation (XDR) and Remote Procedure Call (RPC) mechanisms used for the NFSv4.1 protocol are those defined in [2] (Eisler, M., “XDR: External Data Representation Standard,” May 2006.) and [3] (Srinivasan, R., “RPC: Remote Procedure Call Protocol Specification Version 2,” August 1995.). To meet end-to-end security requirements, the RPCSEC_GSS framework [4] (Eisler, M., Chiu, A., and L. Ling, “RPCSEC_GSS Protocol Specification,” September 1997.) will be used to extend the basic RPC security. With the use of RPCSEC_GSS, various mechanisms can be provided to offer authentication, integrity, and privacy to the NFSv4 protocol. Kerberos V5 will be used as described in [5] (Zhu, L., Jaganathan, K., and S. Hartman, “The Kerberos Version 5 Generic Security Service Application Program Interface (GSS-API) Mechanism Version 2,” July 2005.) to provide one security framework. The LIPKEY and SPKM-3 GSS-API mechanisms described in [6] (Eisler, M., “LIPKEY - A Low Infrastructure Public Key Mechanism Using SPKM,” June 2000.) will be used to provide for the use of user password and client/server public key certificates by the NFSv4 protocol. With the use of RPCSEC_GSS, other mechanisms may also be specified and used for NFSv4.1 security.
To enable in-band security negotiation, the NFSv4.1 protocol has operations which provide the client a method of querying the server about its policies regarding which security mechanisms must be used for access to the server's file system resources. With this, the client can securely match the security mechanism that meets the policies specified at both the client and server.
| TOC |
| TOC |
Unlike NFSv3, which used a series of ancillary protocols (e.g. NLM, NSM, MOUNT), within all minor versions of NFSv4 a single RPC protocol is used to make requests to the server. Facilities that had been separate protocols, such as locking, are now integrated within a single unified protocol.
| TOC |
Minor version one supports high-performance data access to a clustered server implementation by enabling a separation of metadata access and data access, with the latter done to multiple servers in parallel.
Such parallel data access is controlled by recallable objects known as "layouts", which are integrated into the protocol locking model. Clients direct requests for data access to a set of data servers specified by the layout via a data storage protocol which may be NFSv4.1 or may be another protocol.
| TOC |
The general file system model used for the NFSv4.1 protocol is the same as previous versions. The server file system is hierarchical with the regular files contained within being treated as opaque byte streams. In a slight departure, file and directory names are encoded with UTF-8 to deal with the basics of internationalization.
The NFSv4.1 protocol does not require a separate protocol to provide for the initial mapping between path name and filehandle. All file systems exported by a server are presented as a tree so that all file systems are reachable from a special per-server global root filehandle. This allows LOOKUP operations to be used to perform functions previously provided by the MOUNT protocol. The server provides any necessary pseudo file systems to bridge any gaps that arise due to unexported gaps between exported file systems.
| TOC |
As in previous versions of the NFS protocol, opaque filehandles are used to identify individual files and directories. Lookup-type and create operations translate file and directory names to filehandles which are then used to identify objects in subsequent operations.
The NFSv4.1 protocol provides support for persistent filehandles, guaranteed to be valid for the lifetime of the file system object designated. In addition it provides support to servers to provide filehandles with more limited validity guarantees, called volatile filehandles.
| TOC |
The NFSv4.1 protocol has a rich and extensible attribute structure, which is divided into REQUIRED, RECOMMENDED, and named attributes.
The acl, sacl, and dacl attributes compose a set of RECOMMENDED file attributes that make up the Access Control List (ACL) of a file (Section 6 (Access Control Attributes)). These attributes provide for directory and file access control beyond the model used in NFSv3. The ACL definition allows for specification of specific sets of permissions for individual users and groups. In addition, ACL inheritance allows propagation of access permissions and restriction down a directory tree as file system objects are created.
A named attribute is an opaque byte stream that is associated with a directory or file and referred to by a string name. Named attributes are meant to be used by client applications as a method to associate application-specific data with a regular file or directory. NFSv4.1 modifies named attributes relative to NFSv4.0 by tightening the allowed operations in order to prevent the development of non-interoperable implementation. See Section 5.3 (Named Attributes) for details.
| TOC |
NFSv4.1 contains a number of features to allow implementation of namespaces that cross server boundaries and that allow and facilitate a non-disruptive transfer of support for individual file systems between servers. They are all based upon attributes that allow one file system to specify alternate or new locations for that file system.
These attributes may be used together with the concept of absent file systems, which provide specifications for additional locations but no actual file system content. This allows a number of important facilities:
| TOC |
As mentioned previously, NFS v4.1 is a single protocol which includes locking facilities. These locking facilities include support for many types of locks including a number of sorts of recallable locks. Recallable locks such as delegations allow the client to be assured that certain events will not occur so long as that lock is held. When circumstances change, the lock is recalled via a callback request. The assurances provided by delegations allow more extensive caching to be done safely when circumstances allow it.
The types of locks are:
All locks for a given client are tied together under a single client-wide lease. All requests made on sessions associated with the client renew that lease. When leases are not promptly renewed locks are subject to revocation. In the event of server restart, clients have the opportunity to safely reclaim their locks within a special grace period.
| TOC |
The following summarizes the major differences between minor version one and the base protocol:
| TOC |
| TOC |
NFSv4.1 relies on core infrastructure common to nearly every operation. This core infrastructure is described in the remainder of this section.
| TOC |
The NFSv4.1 protocol is a Remote Procedure Call (RPC) application that uses RPC version 2 and the corresponding eXternal Data Representation (XDR) as defined in [3] (Srinivasan, R., “RPC: Remote Procedure Call Protocol Specification Version 2,” August 1995.) and [2] (Eisler, M., “XDR: External Data Representation Standard,” May 2006.).
| TOC |
Previous NFS versions have been thought of as having a host-based authentication model, where the NFS server authenticates the NFS client, and trusts the client to authenticate all users. Actually, NFS has always depended on RPC for authentication. One of the first forms of RPC authentication, AUTH_SYS, had no strong authentication, and required a host-based authentication approach. NFSv4.1 also depends on RPC for basic security services, and mandates RPC support for a user-based authentication model. The user-based authentication model has user principals authenticated by a server, and in turn the server authenticated by user principals. RPC provides some basic security services which are used by NFSv4.1.
| TOC |
As described in section 7.2 "Authentication" of [3] (Srinivasan, R., “RPC: Remote Procedure Call Protocol Specification Version 2,” August 1995.), RPC security is encapsulated in the RPC header, via a security or authentication flavor, and information specific to the specified security flavor. Every RPC header conveys information used to identify and authenticate a client and server. As discussed in Section 2.2.1.1.1 (RPCSEC_GSS and Security Services), some security flavors provide additional security services.
NFSv4.1 clients and servers MUST implement RPCSEC_GSS. (This requirement to implement is not a requirement to use.) Other flavors, such as AUTH_NONE, and AUTH_SYS, MAY be implemented as well.
| TOC |
RPCSEC_GSS ([4] (Eisler, M., Chiu, A., and L. Ling, “RPCSEC_GSS Protocol Specification,” September 1997.)) uses the functionality of GSS-API [7] (Linn, J., “Generic Security Service Application Program Interface Version 2, Update 1,” January 2000.). This allows for the use of various security mechanisms by the RPC layer without the additional implementation overhead of adding RPC security flavors.
| TOC |
Via the GSS-API, RPCSEC_GSS can be used to identify and authenticate users on clients to servers, and servers to users. It can also perform integrity checking on the entire RPC message, including the RPC header, and the arguments or results. Finally, privacy, usually via encryption, is a service available with RPCSEC_GSS. Privacy is performed on the arguments and results. Note that if privacy is selected, integrity, authentication, and identification are enabled. If privacy is not selected, but integrity is selected, authentication and identification are enabled. If integrity and privacy are not selected, but authentication is enabled, identification is enabled. RPCSEC_GSS does not provide identification as a separate service.
Although GSS-API has an authentication service distinct from its privacy and integrity services, GSS-API's authentication service is not used for RPCSEC_GSS's authentication service. Instead, each RPC request and response header is integrity protected with the GSS-API integrity service, and this allows RPCSEC_GSS to offer per-RPC authentication and identity. See [4] (Eisler, M., Chiu, A., and L. Ling, “RPCSEC_GSS Protocol Specification,” September 1997.) for more information.
NFSv4.1 client and servers MUST support RPCSEC_GSS's integrity and authentication service. NFSv4.1 servers MUST support RPCSEC_GSS's privacy service.
| TOC |
RPCSEC_GSS, via GSS-API, normalizes access to mechanisms that provide security services. Therefore NFSv4.1 clients and servers MUST support three security mechanisms: Kerberos V5, SPKM-3, and LIPKEY.
The use of RPCSEC_GSS requires selection of: mechanism, quality of protection (QOP), and service (authentication, integrity, privacy). For the mandated security mechanisms, NFSv4.1 specifies that a QOP of zero (0) is used, leaving it up to the mechanism or the mechanism's configuration to use an appropriate level of protection that QOP zero maps to. Each mandated mechanism specifies minimum set of cryptographic algorithms for implementing integrity and privacy. NFSv4.1 clients and servers MUST be implemented on operating environments that comply with the REQUIRED cryptographic algorithms of each REQUIRED mechanism.
| TOC |
The Kerberos V5 GSS-API mechanism as described in [5] (Zhu, L., Jaganathan, K., and S. Hartman, “The Kerberos Version 5 Generic Security Service Application Program Interface (GSS-API) Mechanism Version 2,” July 2005.) MUST be implemented with the RPCSEC_GSS services as specified in the following table:
column descriptions: 1 == number of pseudo flavor 2 == name of pseudo flavor 3 == mechanism's OID 4 == RPCSEC_GSS service 5 == NFSv4.1 clients MUST support 6 == NFSv4.1 servers MUST support 1 2 3 4 5 6 ------------------------------------------------------------------ 390003 krb5 1.2.840.113554.1.2.2 rpc_gss_svc_none yes yes 390004 krb5i 1.2.840.113554.1.2.2 rpc_gss_svc_integrity yes yes 390005 krb5p 1.2.840.113554.1.2.2 rpc_gss_svc_privacy no yes
Note that the number and name of the pseudo flavor is presented here as a mapping aid to the implementor. Because the NFSv4.1 protocol includes a method to negotiate security and it understands the GSS-API mechanism, the pseudo flavor is not needed. The pseudo flavor is needed for the NFSv3 since the security negotiation is done via the MOUNT protocol as described in [23] (Eisler, M., “NFS Version 2 and Version 3 Security Issues and the NFS Protocol's Use of RPCSEC_GSS and Kerberos V5,” June 1999.).
| TOC |
The LIPKEY V5 GSS-API mechanism as described in [6] (Eisler, M., “LIPKEY - A Low Infrastructure Public Key Mechanism Using SPKM,” June 2000.) MUST be implemented with the RPCSEC_GSS services as specified in the following table:
1 2 3 4 5 6 ------------------------------------------------------------------ 390006 lipkey 1.3.6.1.5.5.9 rpc_gss_svc_none yes yes 390007 lipkey-i 1.3.6.1.5.5.9 rpc_gss_svc_integrity yes yes 390008 lipkey-p 1.3.6.1.5.5.9 rpc_gss_svc_privacy no yes
| TOC |
The SPKM-3 GSS-API mechanism as described in [6] (Eisler, M., “LIPKEY - A Low Infrastructure Public Key Mechanism Using SPKM,” June 2000.) MUST be implemented with the RPCSEC_GSS services as specified in the following table:
1 2 3 4 5 6 ------------------------------------------------------------------ 390009 spkm3 1.3.6.1.5.5.1.3 rpc_gss_svc_none yes yes 390010 spkm3i 1.3.6.1.5.5.1.3 rpc_gss_svc_integrity yes yes 390011 spkm3p 1.3.6.1.5.5.1.3 rpc_gss_svc_privacy no yes
| TOC |
Regardless of what security mechanism under RPCSEC_GSS is being used, the NFS server, MUST identify itself in GSS-API via a GSS_C_NT_HOSTBASED_SERVICE name type. GSS_C_NT_HOSTBASED_SERVICE names are of the form:
service@hostname
For NFS, the "service" element is
nfs
Implementations of security mechanisms will convert nfs@hostname to various different forms. For Kerberos V5, LIPKEY, and SPKM-3, the following form is RECOMMENDED:
nfs/hostname
| TOC |
A significant departure from the versions of the NFS protocol before NFSv4 is the introduction of the COMPOUND procedure. For the NFSv4 protocol, in all minor versions, there are exactly two RPC procedures, NULL and COMPOUND. The COMPOUND procedure is defined as a series of individual operations and these operations perform the sorts of functions performed by traditional NFS procedures.
The operations combined within a COMPOUND request are evaluated in order by the server, without any atomicity guarantees. A limited set of facilities exist to pass results from one operation to another. Once an operation returns a failing result, the evaluation ends and the results of all evaluated operations are returned to the client.
With the use of the COMPOUND procedure, the client is able to build simple or complex requests. These COMPOUND requests allow for a reduction in the number of RPCs needed for logical file system operations. For example, multi-component lookup requests can be constructed by combining multiple LOOKUP operations. Those can be further combined with operations such as GETATTR, READDIR, or OPEN plus READ to do more complicated sets of operation without incurring additional latency.
NFSv4.1 also contains a considerable set of callback operations in which the server makes an RPC directed at the client. Callback RPC's have a similar structure to that of the normal server requests. In all minor versions of the NFSv4 protocol there are two callback RPC procedures, CB_NULL and CB_COMPOUND. The CB_COMPOUND procedure is defined in an analogous fashion to that of COMPOUND with its own set of callback operations.
The addition of new server and callback operations within the COMPOUND and CB_COMPOUND request framework provides a means of extending the protocol in subsequent minor versions.
Except for a small number of operations needed for session creation, server requests and callback requests are performed within the context of a session. Sessions provide a client context for every request and support robust reply protection for non-idempotent requests.
| TOC |
For each operation that obtains or depends on locking state, the specific client must be identifiable by the server.
Each distinct client instance is represented by a client ID. A client ID is a 64-bit identifier representing a specific client at a given time. The client ID is changed whenever the client re-initializes, and may change when the server re-initializes. Client IDs are used to support lock identification and crash recovery.
During steady state operation, the client ID associated with each operation is derived from the session (see Section 2.10 (Session)) on which the operation is sent. A session is associated with a client ID when the session is created.
Unlike NFSv4.0, the only NFSv4.1 operations possible before a client ID is established are those needed to establish the client ID.
A sequence of an EXCHANGE_ID operation followed by a CREATE_SESSION operation using that client ID (eir_clientid as returned from EXCHANGE_ID) is required to establish and confirm the client ID on the server. Establishment of identification by a new incarnation of the client also has the effect of immediately releasing any locking state that a previous incarnation of that same client might have had on the server. Such released state would include all lock, share reservation, layout state, and where the server is not supporting the CLAIM_DELEGATE_PREV claim type, all delegation state associated with the same client with the same identity. For discussion of delegation state recovery, see Section 10.2.1 (Delegation Recovery). For discussion of layout state recovery see Section 12.7.1 (Recovery from Client Restart).
Releasing such state requires that the server be able to determine that one client instance is the successor of another. Where this cannot be done, for any of a number of reasons, the locking state will remain for a time subject to lease expiration (see Section 8.3 (Lease Renewal)) and the new client will need to wait for such state to be removed, if it makes conflicting lock requests.
Client identification is encapsulated in the following Client Owner data type:
struct client_owner4 {
verifier4 co_verifier;
opaque co_ownerid<NFS4_OPAQUE_LIMIT>;
};
The first field, co_verifier, is a client incarnation verifier. The server will start the process of canceling the client's leased state if co_verifier is different than what the server has previously recorded for the identified client (as specified in the co_ownerid field).
The second field, co_ownerid is a variable length string that uniquely defines the client so that subsequent instances of the same client bear the same co_ownerid with a different verifier.
There are several considerations for how the client generates the co_ownerid string:
Given the above considerations, an example of a well generated co_ownerid string is one that includes:
The client ID is assigned by the server (the eir_clientid result from EXCHANGE_ID) and should be chosen so that it will not conflict with a client ID previously assigned by the server. This applies across server restarts.
In the event of a server restart, a client may find out that its current client ID is no longer valid when it receives a NFS4ERR_STALE_CLIENTID error. The precise circumstances depend on the characteristics of the sessions involved, specifically whether the session is persistent (see Section 2.10.5.5 (Persistence)), but in each case the client will receive this error when it attempts to establish a new session with the existing client ID and receives the error NFS4ERR_STALE_CLIENTID, indicating that a new client ID must be obtained via EXCHANGE_ID and the new session established with that client ID.
When a session is not persistent, the client will find out that it needs to create a new session as a result of getting an NFS4ERR_BADSESSION, since the session in question was lost as part of a server restart. When the existing client ID is presented to a server as part of creating a session and that client ID is not recognized, as would happen after a server restart, the server will reject the request with the error NFS4ERR_STALE_CLIENTID.
In the case of the session being persistent, the client will re-establish communication using the existing session after the restart. This session will be associated with the existing client ID but may only be used to retransmit operations that the client previously transmitted and did not see replies to. Replies to operations that the server previously performed will come from the reply cache, otherwise NFS4ERR_DEADSESSION will be returned. Hence, such a session is referred to as "dead". In this situation, in order to perform new operations, the client must establish a new session. If an attempt is made to establish this new session with the existing client ID, the server will reject the request with NFS4ERR_STALE_CLIENTID.
When NFS4ERR_STALE_CLIENTID is received in either of these situations, the client must obtain a new client ID by use of the EXCHANGE_ID operation, then use that client ID as the basis of a new session, and then proceed to any other necessary recovery for the server restart case (See Section 8.4.2 (Server Failure and Recovery)).
See the descriptions of EXCHANGE_ID (Section 18.35 (Operation 42: EXCHANGE_ID - Instantiate Client ID)) and CREATE_SESSION (Section 18.36 (Operation 43: CREATE_SESSION - Create New Session and Confirm Client ID)) for a complete specification of these operations.
| TOC |
To facilitate upgrade from NFSv4.0 to NFSv4.1, a server may compare a client_owner4 in an EXCHANGE_ID with an nfs_client_id4 established using the SETCLIENTID operation of NFSv4.0. A server that does so will allow an upgraded client to avoid waiting until the lease (i.e. the lease established by the NFSv4.0 instance client) expires. This requires the client_owner4 be constructed the same way as the nfs_client_id4. If the latter's contents included the server's network address (per the recommendations of the NFSv4.0 specification [21] (Shepler, S., Callaghan, B., Robinson, D., Thurlow, R., Beame, C., Eisler, M., and D. Noveck, “Network File System (NFS) version 4 Protocol,” April 2003.)), and the NFSv4.1 client does not wish to use a client ID that prevents trunking, it should send two EXCHANGE_ID operations. The first EXCHANGE_ID will have a client_owner4 equal to the nfs_client_id4. This will clear the state created by the NFSv4.0 client. The second EXCHANGE_ID will not have the server's network address. The state created for the second EXCHANGE_ID will not have to wait for lease expiration, because there will be no state to expire.
| TOC |
NFSv4.1 introduces a new operation called DESTROY_CLIENTID (Section 18.50 (Operation 57: DESTROY_CLIENTID - Destroy existing client ID)) which the client SHOULD use to destroy a client ID it no longer needs. This permits graceful, bilateral release of a client ID. The operation cannot be used if there are sessions associated with the client ID, or state with an unexpired lease.
If the server determines that the client holds no associated state for its client ID (including sessions, opens, locks, delegations, layouts, and wants), the server may choose to unilaterally release the client ID in order to conserve resources. If the client contacts the server after this release, the server must ensure the client receives the appropriate error so that it will use the EXCHANGE_ID/CREATE_SESSION sequence to establish a new client ID. The server ought to be very hesitant to release a client ID since the resulting work on the client to recover from such an event will be the same burden as if the server had failed and restarted. Typically a server would not release a client ID unless there had been no activity from that client for many minutes. As long as there are sessions, opens, locks, delegations, layouts, or wants, the server MUST NOT release the client ID. See Section 2.10.11.1.4 (Loss of Session) for discussion on releasing inactive sessions.
| TOC |
When the server gets an EXCHANGE_ID for a client owner that currently has no state, or that has state, but the lease has expired, the server MUST allow the EXCHANGE_ID, and confirm the new client ID if followed by the appropriate CREATE_SESSION.
When the server gets an EXCHANGE_ID for a new incarnation of a client owner that currently has an old incarnation with state and an unexpired lease, the server is allowed to dispose of the state of the previous incarnation of the client owner if one of the following are true:
If none of the above situations apply, the server MUST return NFS4ERR_CLID_INUSE.
If the server accepts the principal and co_ownerid as matching that which created the client ID, and the co_verifier in the EXCHANGE_ID differs from the co_verifier used when the client ID was created, then after the server receives a CREATE_SESSION that confirms the client ID, the server deletes state. If the co_verifier values are the same, (e.g. the client is either updating properties of the client ID (Section 18.35 (Operation 42: EXCHANGE_ID - Instantiate Client ID)), or the client is attempting trunking (Section 2.10.4 (Trunking)) the server MUST NOT delete state.
| TOC |
The Server Owner is similar to a Client Owner (Section 2.4 (Client Identifiers and Client Owners)), but unlike the Client Owner, there is no shorthand server ID. The Server Owner is defined in the following data type:
struct server_owner4 {
uint64_t so_minor_id;
opaque so_major_id<NFS4_OPAQUE_LIMIT>;
};
The Server Owner is returned from EXCHANGE_ID. When the so_major_id fields are the same in two EXCHANGE_ID results, the connections each EXCHANGE_ID were sent over can be assumed to address the same Server (as defined in Section 1.5 (General Definitions)). If the so_minor_id fields are also the same, then not only do both connections connect to the same server, but the session can be shared across both connections. The reader is cautioned that multiple servers may deliberately or accidentally claim to have the same so_major_id or so_major_id/so_minor_id; the reader should examine Section 2.10.4 (Trunking) and Section 18.35 (Operation 42: EXCHANGE_ID - Instantiate Client ID) in order to avoid acting on falsely matching Server Owner values.
The considerations for generating a so_major_id are similar to that for generating a co_ownerid string (see Section 2.4 (Client Identifiers and Client Owners)). The consequences of two servers generating conflicting so_major_id values are less dire than they are for co_ownerid conflicts because the client can use RPCSEC_GSS to compare the authenticity of each server (see Section 2.10.4 (Trunking)).
| TOC |
With the NFSv4.1 server potentially offering multiple security mechanisms, the client needs a method to determine or negotiate which mechanism is to be used for its communication with the server. The NFS server may have multiple points within its file system namespace that are available for use by NFS clients. These points can be considered security policy boundaries, and in some NFS implementations are tied to NFS export points. In turn the NFS server may be configured such that each of these security policy boundaries may have different or multiple security mechanisms in use.
The security negotiation between client and server must be done with a secure channel to eliminate the possibility of a third party intercepting the negotiation sequence and forcing the client and server to choose a lower level of security than required or desired. See Section 21 (Security Considerations) for further discussion.
| TOC |
An NFS server can assign one or more "security tuples" to each security policy boundary in its namespace. Each security tuple consists of a security flavor (see Section 2.2.1.1 (RPC Security Flavors)), and if the flavor is RPCSEC_GSS, a GSS-API mechanism OID, a GSS-API quality of protection, and an RPCSEC_GSS service.
| TOC |
The SECINFO and SECINFO_NO_NAME operations allow the client to determine, on a per filehandle basis, what security tuple is to be used for server access. In general, the client will not have to use either operation except during initial communication with the server or when the client crosses security policy boundaries at the server. However, the server's policies may also change at any time and force the client to negotiate a new security tuple.
Where the use of different security tuples would affect the type of access that would be allowed if a request was sent over the same connection used for the SECINFO or SECINFO_NO_NAME operation (e.g. read-only vs. read-write) access, security tuples that allow greater access should be presented first. Where the general level of access is the same and different security flavors limit the range of principals whose privileges are recognized (e.g. allowing or disallowing root access), flavors supporting the greatest range of principals should be listed first.
| TOC |
Based on the assumption that each NFSv4.1 client and server must support a minimum set of security (i.e., LIPKEY, SPKM-3, and Kerberos-V5 all under RPCSEC_GSS), the NFS client will initiate file access to the server with one of the minimal security tuples. During communication with the server, the client may receive an NFS error of NFS4ERR_WRONGSEC. This error allows the server to notify the client that the security tuple currently being used contravenes the server's security policy. The client is then responsible for determining (see Section 2.6.3.1 (Using NFS4ERR_WRONGSEC, SECINFO, and SECINFO_NO_NAME)) what security tuples are available at the server and choosing one which is appropriate for the client.
| TOC |
This section explains of the mechanics of NFSv4.1 security negotiation.
| TOC |
The term "put filehandle operation" refers to PUTROOTFH, PUTPUBFH, PUTFH, and RESTOREFH. Each of the subsections herein describes how the server handles a subseries of operations that starts with a put filehandle operation.
| TOC |
The client is saving a filehandle for a future RESTOREFH, LINK, or RENAME. SAVEFH MUST NOT return NFS4ERR_WRONGSEC. To determine whether the put filehandle operation returns NFS4ERR_WRONGSEC or not, the server implementation pretends SAVEFH is not in the series of operations and examines which of the situations described in the other subsections of Section 2.6.3.1.1 (Put Filehandle Operations) apply.
| TOC |
For a series of N put filehandle operations, the server MUST NOT return NFS4ERR_WRONGSEC to the first N-1 put filehandle operations. The N'th put filehandle operation is handled as if it is the first in a subseries of operations. For example if the server received PUTFH, PUTROOTFH, LOOKUP, then the PUTFH is ignored for NFS4ERR_WRONGSEC purposes, and the PUTROOTFH, LOOKUP subseries is processed as according to Section 2.6.3.1.1.3 (Put Filehandle Operation + LOOKUP (or OPEN of an Existing Name)).
| TOC |
This situation also applies to a put filehandle operation followed by a LOOKUP or an OPEN operation that specifies an existing component name.
In this situation, the client is potentially crossing a security policy boundary, and the set of security tuples the parent directory supports may differ from those of the child. The server implementation may decide whether to impose any restrictions on security policy administration. There are at least three approaches (sec_policy_child is the tuple set of the child export, sec_policy_parent is that of the parent).
- a)
- sec_policy_child <= sec_policy_parent (<= for subset). This means that the set of security tuples specified on the security policy of a child directory is always a subset of that of its parent directory.
- b)
- sec_policy_child ^ sec_policy_parent != {} (^ for intersection, {} for the empty set). This means that the security tuples specified on the security policy of a child directory always has a non empty intersection with that of the parent.
- c)
- sec_policy_child ^ sec_policy_parent == {}. This means that the set of tuples specified on the security policy of a child directory may not intersect with that of the parent. In other words, there are no restrictions on how the system administrator may set up these tuples.
In order for a server to support approaches (b) (for the case when a client chooses a flavor that is not a member of sec_policy_parent) and (c), the put filehandle operation cannot return NFS4ERR_WRONGSEC when there is a security tuple mismatch. Instead, it should be returned from the LOOKUP (or OPEN by existing component name) that follows.
Since the above guideline does not contradict approach (a), it should be followed in general. Even if approach (a) is implemented, it is possible for the security tuple used to be acceptable for the target of LOOKUP but not for the filehandles used in the put filehandle operation. The put filehandle operation could be a PUTROOTFH or PUTPUBFH, where the client cannot know the security tuples for the root or public filehandle. Or the security policy for the filehandle used by the put filehandle operation could have changed since the time the filehandle was obtained.
Therefore, an NFSv4.1 server MUST NOT return NFS4ERR_WRONGSEC in response to the put filehandle operation if the operation is immediately followed by a LOOKUP or an OPEN by component name.
| TOC |
Since SECINFO only works its way down, there is no way LOOKUPP can return NFS4ERR_WRONGSEC without SECINFO_NO_NAME. SECINFO_NO_NAME solves this issue via style SECINFO_STYLE4_PARENT, which works in the opposite direction as SECINFO. As with Section 2.6.3.1.1.3 (Put Filehandle Operation + LOOKUP (or OPEN of an Existing Name)), a put filehandle operation that is followed by a LOOKUPP MUST NOT return NFS4ERR_WRONGSEC. If the server does not support SECINFO_NO_NAME, the client's only recourse is to send the put filehandle operation, LOOKUPP, GETFH sequence of operations with every security tuple it supports.
Regardless of whether SECINFO_NO_NAME is supported, an NFSv4.1 server MUST NOT return NFS4ERR_WRONGSEC in response to a put filehandle operation if the operation is immediately followed by a LOOKUPP.
| TOC |
A security sensitive client is allowed to choose a strong security tuple when querying a server to determine a file object's permitted security tuples. The security tuple chosen by the client does not have to be included in the tuple list of the security policy of the either parent directory indicated in the put filehandle operation, or the child file object indicated in SECINFO (or any parent directory indicated in SECINFO_NO_NAME). Of course the server has to be configured for whatever security tuple the client selects, otherwise the request will fail at RPC layer with an appropriate authentication error.
In theory, there is no connection between the security flavor used by SECINFO or SECINFO_NO_NAME and those supported by the security policy. But in practice, the client may start looking for strong flavors from those supported by the security policy, followed by those in the REQUIRED set.
The NFSv4.1 server MUST NOT return NFS4ERR_WRONGSEC to a put filehandle operation that is immediately followed by SECINFO or SECINFO_NO_NAME. The NFSv4.1 server MUST NOT return NFS4ERR_WRONGSEC from SECINFO or SECINFO_NO_NAME.
| TOC |
The NFSv4.1 server MUST NOT return NFS4ERR_WRONGSEC.
| TOC |
"Anything Else" includes OPEN by filehandle.
The security policy enforcement applies to the filehandle specified in the put filehandle operation. Therefore the put filehandle operation must return NFS4ERR_WRONGSEC when there is a security tuple mismatch. This avoids the complexity adding NFS4ERR_WRONGSEC as an allowable error to every other operation.
A COMPOUND containing the series put filehandle operation + SECINFO_NO_NAME (style SECINFO_STYLE4_CURRENT_FH) is an efficient way for the client to recover from NFS4ERR_WRONGSEC.
The NFSv4.1 server MUST NOT return NFS4ERR_WRONGSEC to any operation other than a put filehandle operation, LOOKUP, LOOKUPP, and OPEN (by component name).
| TOC |
Suppose a client sends a COMPOUND procedure containing the series SEQUENCE, PUTFH, SECINFO_NONAME, READ, and suppose the security tuple used does not match that required for the target file. By rule (see Section 2.6.3.1.1.5 (Put Filehandle Operation + SECINFO/SECINFO_NO_NAME)), neither PUTFH nor SECINFO_NO_NAME can return NFS4ERR_WRONGSEC. By rule (see Section 2.6.3.1.1.7 (Put Filehandle Operation + Anything Else)), READ cannot return NFS4ERR_WRONGSEC. The issue is resolved by the fact that SECINFO and SECINFO_NO_NAME consume the current filehandle (note that this is a change from NFSv4.0). This leaves no current filehandle for READ to use, and READ returns NFS4ERR_NOFILEHANDLE.
| TOC |
The LINK and RENAME operations use both the current and saved filehandles. When the current filehandle is injected into a series of operations via a put filehandle operation, the server MUST return NFS4ERR_WRONGSEC, per Section 2.6.3.1.1 (Put Filehandle Operations). LINK and RENAME MAY return NFS4ERR_WRONGSEC if the security policy of the saved filehandle rejects the security flavor used in the COMPOUND request's credentials. If the server does so, then if there is no intersection between the security policies of saved and current filehandles, this means it will be impossible for client to perform the intended LINK or RENAME operation.
For example, suppose the client sends this COMPOUND request: SEQUENCE, PUTFH bFH, SAVEFH, PUTFH aFH, RENAME "c" "d", where filehandles bFH and aFH refer to different directories. Suppose no common security tuple exists between the security policies of aFH and bFH. If the client sends the request using credentials acceptable to bFH's security policy but not aFH's policy, then the PUTFH aFH operation will fail with NFS4ERR_WRONGSEC. After a SECINFO_NO_NAME request, the client sends SEQUENCE, PUTFH bFH, SAVEFH, PUTFH aFH, RENAME "c" "d", using credentials acceptable to aFH's security policy, but not bFH's policy. The server returns NFS4ERR_WRONGSEC on the RENAME operation.
To prevent a client from an endless sequence of a request containing LINK or RENAME, followed by a request containing SECINFO_NO_NAME, the server MUST detect when the security policies of the current and saved filehandles have no mutually acceptable security tuple, and MUST NOT NFS4ERR_WRONGSEC in that situation. Instead the server MUST return NFS4ERR_XDEV.
Thus while a server MAY return NFS4ERR_WRONGSEC from LINK and RENAME, the server implementor may reasonably decide the consequences are not worth the security benefits, and so allow the security policy of the current filehandle to override that of the saved filehandle.
| TOC |
To address the requirement of an NFS protocol that can evolve as the need arises, the NFSv4.1 protocol contains the rules and framework to allow for future minor changes or versioning.
The base assumption with respect to minor versioning is that any future accepted minor version must follow the IETF process and be documented in a standards track RFC. Therefore, each minor version number will correspond to an RFC. Minor version zero of the NFSv4 protocol is represented by [21] (Shepler, S., Callaghan, B., Robinson, D., Thurlow, R., Beame, C., Eisler, M., and D. Noveck, “Network File System (NFS) version 4 Protocol,” April 2003.), and minor version one is represented by this document [Comment.1] (RFC Editor: change "document" to "RFC" when we publish). The COMPOUND and CB_COMPOUND procedures support the encoding of the minor version being requested by the client.
The following items represent the basic rules for the development of minor versions. Note that a future minor version may decide to modify or add to the following rules as part of the minor version definition.
| TOC |
As described in Section 2.2.1.1.1.1 (Identification, Authentication, Integrity, Privacy), NFSv4.1 relies on RPC for identification, authentication, integrity, and privacy. NFSv4.1 itself provides or enables additional security services as described in the next several subsections.
| TOC |
Authorization to access a file object via an NFSv4.1 operation is ultimately determined by the NFSv4.1 server. A client can predetermine its access to a file object via the OPEN (Section 18.16 (Operation 18: OPEN - Open a Regular File)) and the ACCESS (Section 18.1 (Operation 3: ACCESS - Check Access Rights)) operations.
Principals with appropriate access rights can modify the authorization on a file object via the SETATTR (Section 18.30 (Operation 34: SETATTR - Set Attributes)) operation. Attributes that affect access rights include: mode, owner, owner_group, acl, dacl, and sacl. See Section 5 (File Attributes).
| TOC |
NFSv4.1 provides auditing on a per file object basis, via the acl and sacl attributes as described in Section 6 (Access Control Attributes). It is outside the scope of this specification to specify audit log formats or management policies.
| TOC |
NFSv4.1 provides alarm control on a per file object basis, via the acl and sacl attributes as described in Section 6 (Access Control Attributes). Alarms may serve as the basis for intrusion detection. It is outside the scope of this specification to specify heuristics for detecting intrusion via alarms.
| TOC |
| TOC |
NFSv4.1 works over RDMA and non-RDMA_based transports with the following attributes:
Where an NFSv4.1 implementation supports operation over the IP network protocol, any transport used between NFS and IP MUST be among the IETF-approved congestion control transport protocols. At the time this document was written, the only two transports that had the above attributes were TCP and SCTP. To enhance the possibilities for interoperability, an NFSv4.1 implementation MUST support operation over the TCP transport protocol.
Even if NFSv4.1 is used over a non-IP network protocol, it is RECOMMENDED that the transport support congestion control.
It is permissible for a connectionless transport to be used under NFSv4.1, however reliable and in-order delivery of data by the connectionless transport is REQUIRED. NFSv4.1 assumes that a client transport address and server transport address used to send data over a transport together constitute a connection, even if the underlying transport eschews the concept of a connection.
| TOC |
If a connection-oriented transport (e.g. TCP) is used, the client and server SHOULD use long lived connections for at least three reasons:
In order to reduce congestion, if a connection-oriented transport is used, and the request is not the NULL procedure,
When sending a reply, the replier MUST send the reply to the same full network address (e.g. if using an IP-based transport, the source port of the requester is part of the full network address) that the requester sent the request from. If using a connection-oriented transport, replies MUST be sent on the same connection the request was received from.
If a connection is dropped after the replier receives the request but before the replier sends the reply, the replier might have an pending reply. If a connection is established with the same source and destination full network address as the dropped connection, then the replier MUST NOT send the reply until the client retries the request. The reason for this prohibition is that the client MAY retry a request over a different connection than is associated with the session.
When using RDMA transports there are other reasons for not tolerating retries over the same connection:
In addition, as described in Section 2.10.5.2 (Retry and Replay of Reply), while a session is active, the NFSv4.1 requester MUST NOT stop waiting for a reply.
| TOC |
Historically, NFSv3 servers have listened over TCP port 2049. The registered port 2049 [25] (Reynolds, J., “Assigned Numbers: RFC 1700 is Replaced by an On-line Database,” January 2002.) for the NFS protocol should be the default configuration. NFSv4.1 clients SHOULD NOT use the RPC binding protocols as described in [26] (Srinivasan, R., “Binding Protocols for ONC RPC Version 2,” August 1995.).
| TOC |
| TOC |
Previous versions and minor versions of NFS have suffered from the following:
Through the introduction of a session, NFSv4.1 addresses the above shortfalls with practical solutions:
A session is a dynamically created, long-lived server object created by a client, used over time from one or more transport connections. Its function is to maintain the server's state relative to the connection(s) belonging to a client instance. This state is entirely independent of the connection itself, and indeed the state exists whether the connection exists or not. A client may have one or more sessions associated with it so that client-associated state may be accessed using any of the sessions associated with that client's client ID, when connections are associated with those sessions. When no connections are associated with any of a client ID's sessions for an extended time, such objects as locks, opens, delegations, layouts, etc. are subject to expiration. The session serves as an object representing a means of access by a client to the associated client state on the server, independent of the physical means of access to that state.
A single client may create multiple sessions. A single session MUST NOT serve multiple clients.
| TOC |
Sessions are part of NFSv4.1 and not NFSv4.0. Normally, a major infrastructure change such as sessions would require a new major version number to an ONC RPC program like NFS. However, because NFSv4 encapsulates its functionality in a single procedure, COMPOUND, and because COMPOUND can support an arbitrary number of operations, sessions have been added to NFSv4.1 with little difficulty. COMPOUND includes a minor version number field, and for NFSv4.1 this minor version is set to 1. When the NFSv4 server processes a COMPOUND with the minor version set to 1, it expects a different set of operations than it does for NFSv4.0. NFSv4.1 defines the SEQUENCE operation, which is required for every COMPOUND that operates over an established session, with the exception of some session administration operations, such as DESTROY_SESSION (Section 18.37 (Operation 44: DESTROY_SESSION - Destroy existing session)).
| TOC |
In NFSv4.1, when the SEQUENCE operation is present, it MUST be the first operation in the COMPOUND procedure. The primary purpose of SEQUENCE is to carry the session identifier. The session identifier associates all other operations in the COMPOUND procedure with a particular session. SEQUENCE also contains required information for maintaining EOS (see Section 2.10.5 (Exactly Once Semantics)). Session-enabled NFSv4.1 COMPOUND requests thus have the form:
+-----+--------------+-----------+------------+-----------+----
| tag | minorversion | numops |SEQUENCE op | op + args | ...
| | (== 1) | (limited) | + args | |
+-----+--------------+-----------+------------+-----------+----
and the reply's structure is:
+------------+-----+--------+-------------------------------+--//
|last status | tag | numres |status + SEQUENCE op + results | //
+------------+-----+--------+-------------------------------+--//
//-----------------------+----
// status + op + results | ...
//-----------------------+----
A CB_COMPOUND procedure request and reply has a similar form to COMPOUND, but instead of a SEQUENCE operation, there is a CB_SEQUENCE operation. CB_COMPOUND also has an additional field called "callback_ident", which is superfluous in NFSv4.1 and MUST be ignored by the client. CB_SEQUENCE has the same information as SEQUENCE, and also includes other information needed to resolve callback races (Section 2.10.5.3 (Resolving Server Callback Races)).
| TOC |
Each client ID (Section 2.4 (Client Identifiers and Client Owners)) can have zero or more active sessions. A client ID and associated session are required to perform file access in NFSv4.1. Each time a session is used (whether by a client sending a request to the server, or the client replying to a callback request from the server), the state leased to its associated client ID is automatically renewed.
State such as share reservations, locks, delegations, and layouts (Section 1.6.4 (Locking Facilities)) is tied to the client ID. Client state is not tied to any individual session. Successive state changing operations from a given state owner MAY go over different sessions, provided the session is associated with the same client ID. A callback MAY arrive over a different session than from the session that originally acquired the state pertaining to the callback. For example, if session A is used to acquire a delegation, a request to recall the delegation MAY arrive over session B if both sessions are associated with the same client ID. Section 2.10.7.1 (Session Callback Security) and Section 2.10.7.2 (Backchannel RPC Security) discuss the security considerations around callbacks.
| TOC |
A channel is not a connection. A channel represents the direction ONC RPC requests are sent.
Each session has one or two channels: the fore channel and the backchannel. Because there are at most two channels per session, and because each channel has a distinct purpose, channels are not assigned identifiers.
The fore channel is used for ordinary requests from the client to the server, and carries COMPOUND requests and responses. A session always has a fore channel.
The backchannel used for callback requests from server to client, and carries CB_COMPOUND requests and responses. Whether there is a backchannel or not is a decision by the client, however many features of NFSv4.1 require a backchannel. NFSv4.1 servers MUST support backchannels.
Each session has resources for each channel, including separate reply caches (see Section 2.10.5.1 (Slot Identifiers and Reply Cache)). Note that even the backchannel requires a reply cache because some callback operations are nonidempotent.
| TOC |
Each channel is associated with zero or more transport connections. A connection can be associated with one channel or both channels of a session; the client and server negotiate whether a connection will carry traffic for one channel or both channels via the CREATE_SESSION (Section 18.36 (Operation 43: CREATE_SESSION - Create New Session and Confirm Client ID)) and the BIND_CONN_TO_SESSION (Section 18.34 (Operation 41: BIND_CONN_TO_SESSION)) operations. When a session is created via CREATE_SESSION, the connection that transported the CREATE_SESSION request is automatically associated with the fore channel, and optionally the backchannel. If the client specifies no state protection (Section 18.35 (Operation 42: EXCHANGE_ID - Instantiate Client ID)) when the session is created, then when SEQUENCE is transmitted on a different connection, the connection is automatically associated with the fore channel of the session specified in the SEQUENCE operation.
A connection's association with a session is not exclusive. A connection associated with the channel(s) of one session may be simultaneously associated with the channel(s) of other sessions including sessions associated with other client IDs.
It is permissible for connections of multiple transport types to be associated with the same channel. For example both a TCP and RDMA connection can be associated with the fore channel. In the event an RDMA and non-RDMA connection are associated with the same channel, the maximum number of slots SHOULD be at least one more than the total number of RDMA credits (Section 2.10.5.1 (Slot Identifiers and Reply Cache). This way if all RDMA credits are used, the non-RDMA connection can have at least one outstanding request. If a server supports multiple transport types, it MUST allow a client to associate connections from each transport to a channel.
It is permissible for a connection of one type of transport to be associated with the fore channel, and a connection of a different type to be associated with the backchannel.
| TOC |
Trunking is the use of multiple connections between a client and server in order to increase the speed of data transfer. NFSv4.1 supports two types of trunking: session trunking and client ID trunking. NFSv4.1 servers MUST support trunking.
Session trunking is essentially the association of multiple connections, each with potentially different target and/or source network addresses, to the same session.
Client ID trunking is the association of multiple sessions to the same client ID, major server owner ID (Section 2.5 (Server Owners)), and server scope (Section 11.7.7 (Lock State and File System Transitions)). When two servers return the same major server owner and server scope it means the two servers are cooperating on locking state management which is a prerequisite for client ID trunking.
Understanding and distinguishing session and client ID trunking requires understanding how the results of the EXCHANGE_ID (Section 18.35 (Operation 42: EXCHANGE_ID - Instantiate Client ID)) operation identify a server. Suppose a client sends EXCHANGE_ID over two different connections each with a possibly different target network address but each EXCHANGE_ID with the same value in the eia_clientowner field. If the same NFSv4.1 server is listening over each connection, then each EXCHANGE_ID result MUST return the same values of eir_clientid, eir_server_owner.so_major_id and eir_server_scope. The client can then treat each connection as referring to the same server (subject to verification, see Paragraph 5 later in this section), and it can use each connection to trunk requests and replies. The question is whether session trunking and/or client ID trunking applies.
- Session Trunking
- If the eia_clientowner argument is the same in two different EXCHANGE_ID requests, and the eir_clientid, eir_server_owner.so_major_id, eir_server_owner.so_minor_id, and eir_server_scope results match in both EXCHANGE_ID results, then the client is permitted to perform session trunking. If the client has no session mapping to the tuple of eir_clientid, eir_server_owner.so_major_id, eir_server_scope, eir_server_owner.so_minor_id, then it creates the session via a CREATE_SESSION operation over one of the connections, which associates the connection to the session. If there is a session for the tuple, the client can send BIND_CONN_TO_SESSION to associate the connection to the session. (Of course, if the client does not want to use session trunking, it can invoke CREATE_SESSION on the connection. This will result in client ID trunking as described below.)
- Client ID Trunking
- If the eia_clientowner argument is the same in two different EXCHANGE_ID requests, and the eir_clientid, eir_server_owner.so_major_id, and eir_server_scope results match in both EXCHANGE_ID results, but the eir_server_owner.so_minor_id results do not match then the client is permitted to perform client ID trunking. The client can associate each connection with different sessions, where each session is associated with the same server.
Of course, even if the eir_server_owner.so_minor_id fields do match, the client is free to employ client ID trunking instead of session trunking.
The client completes the act of client ID trunking by invoking CREATE_SESSION on each connection, using the same client ID that was returned in eir_clientid. These invocations create two sessions and also associate each connection with each session.
When doing client ID trunking, locking state is shared across sessions associated with the same client ID. This requires the server to coordinate state across sessions.
When two servers over two connections claim matching or partially matching eir_server_owner, eir_server_scope, and eir_clientid values, the client does not have to trust the servers' claims. The client may verify these claims before trunking traffic in the following ways:
| TOC |
Via the session, NFSv4.1 offers Exactly Once Semantics (EOS) for requests sent over a channel. EOS is supported on both the fore and back channels.
Each COMPOUND or CB_COMPOUND request that is sent with a leading SEQUENCE or CB_SEQUENCE operation MUST be executed by the receiver exactly once. This requirement holds regardless of whether the request is sent with reply caching specified (see Section 2.10.5.1.3 (Optional Reply Caching)). The requirement holds even if the requester is issuing the request over a session created between a pNFS data client and pNFS data server. To understand the rationale for this requirement, divide the requests into three classifications:
An example of a non-idempotent request is RENAME. If is obvious that if a replier executes the same RENAME request twice, and the first execution succeeds, the re-execution will fail. If the replier returns the result from the re-execution, this result is incorrect. Therefore, EOS is required for nonidempotent requests.
An example of an idempotent modifying request is a COMPOUND request containing a WRITE operation. Repeated execution of the same WRITE has the same effect as execution of that write a single time. Nevertheless, enforcing EOS for WRITEs and other idempotent modifying requests is necessary to avoid data corruption.
Suppose a client sends WRITE A to a noncompliant server that does not enforce EOS, and receives no response, perhaps due to a network partition. The client reconnects to the server and re-sends WRITE A. Now, the server has outstanding two instances of A. The server can be in a situation in which it executes and replies to the retry of A, while the first A is still waiting in the server's internal I/O system for some resource. Upon receiving the reply to the second attempt of WRITE A, the client believes its write is done so it is free to send WRITE B which overlaps the range of A. When the original A is dispatched from the server's I/O system, and executed (thus the second time A will have been written), then what has been written by B can be overwritten and thus corrupted.
An example of an idempotent non-modifying request is a COMPOUND containing SEQUENCE, PUTFH, READLINK and nothing else. The re-execution of a such a request will not cause data corruption, or produce an incorrect result. Nonetheless, to keep the implementation simple, the replier MUST enforce EOS for all requests whether idempotent and non-modifying or not.
Note that true and complete EOS is not possible unless the server persists the reply cache in stable storage, unless the server is somehow implemented to never require a restart (indeed if such a server exists, the distinction between a reply cache kept in stable storage versus one that is not is one without meaning). See Section 2.10.5.5 (Persistence) for a discussion of persistence in the reply cache. Regardless, even if the server does not persist the reply cache, EOS improves robustness and correctness over previous versions of NFS because the legacy duplicate request/reply caches were based on the ONC RPC transaction identifier (XID). Section 2.10.5.1 (Slot Identifiers and Reply Cache) explains the shortcomings of the XID as a basis for a reply cache and describes how NFSv4.1 sessions improve upon the XID.
| TOC |
The RPC layer provides a transaction ID (XID), which, while required to be unique, is not convenient for tracking requests for two reasons. First, the XID is only meaningful to the requester; it cannot be interpreted by the replier except to test for equality with previously sent requests. When consulting an RPC-based duplicate request cache, the opaqueness of the XID requires a computationally expensive lookup (often via a hash that includes XID and source address). NFSv4.1 requests use a non-opaque slot id which is an index into a slot table, which is far more efficient. Second, because RPC requests can be executed by the replier in any order, there is no bound on the number of requests that may be outstanding at any time. To achieve perfect EOS using ONC RPC would require storing all replies in the reply cache. XIDs are 32 bits; storing over four billion (2^32) replies in the reply cache is not practical. In practice, previous versions of NFS have chosen to store a fixed number of replies in the cache, and use a least recently used (LRU) approach to replacing cache entries with new entries when the cache is full. In NFSv4.1, the number of outstanding requests is bounded by the size of the slot table, and a sequence id per slot is used to tell the replier when it is safe to delete a cached reply.
In the NFSv4.1 reply cache, when the requester sends a new request, it selects a slot id in the range 0..N, where N is the replier's current maximum slot id granted to the requester on the session over which the request is to be sent. The value of N starts out as equal to ca_maxrequests - 1 (Section 18.36 (Operation 43: CREATE_SESSION - Create New Session and Confirm Client ID)), but can be adjusted by the response to SEQUENCE or CB_SEQUENCE as described later in this section. The slot id must be unused by any of the requests which the requester has already active on the session. "Unused" here means the requester has no outstanding request for that slot id.
A slot contains a sequence id and the cached reply corresponding to the request sent with that sequence id. The sequence id is a 32 bit unsigned value, and is therefore in the range 0..0xFFFFFFFF (2^32 - 1). The first time a slot is used, the requester MUST specify a sequence id of one (1) (Section 18.36 (Operation 43: CREATE_SESSION - Create New Session and Confirm Client ID)). Each time a slot is reused, the request MUST specify a sequence id that is one greater than that of the previous request on the slot. If the previous sequence id was 0xFFFFFFFF, then the next request for the slot MUST have the sequence id set to zero (i.e. (2^32 - 1) + 1 mod 2^32).
The sequence id accompanies the slot id in each request. It is for the critical check at the server: it used to efficiently determine whether a request using a certain slot id is a retransmit or a new, never-before-seen request. It is not feasible for the client to assert that it is retransmitting to implement this, because for any given request the client cannot know whether the server has seen it unless the server actually replies. Of course, if the client has seen the server's reply, the client would not retransmit.
The replier compares each received request's sequence id with the last one previously received for that slot id, to see if the new request is:
Unlike the XID, the slot id is always within a specific range; this has two implications. The first implication is that for a given session, the replier need only cache the results of a limited number of COMPOUND requests . The second implication derives from the first, which is that unlike XID-indexed reply caches (also known as duplicate request caches - DRCs), the slot id-based reply cache cannot be overflowed. Through use of the sequence id to identify retransmitted requests, the replier does not need to actually cache the request itself, reducing the storage requirements of the reply cache further. These facilities make it practical to maintain all the required entries for an effective reply cache.
The slot id, sequence id, and session id therefore take over the traditional role of the XID and source network address in the replier's reply cache implementation. This approach is considerably more portable and completely robust - it is not subject to the reassignment of ports as clients reconnect over IP networks. In addition, the RPC XID is not used in the reply cache, enhancing robustness of the cache in the face of any rapid reuse of XIDs by the requester. While the replier does not care about the XID for the purposes of reply cache management (but the replier MUST return the same XID that was in the request), nonetheless there are considerations for the XID in NFSv4.1 that are the same as all other previous versions of NFS. The RPC XID remains in each message and must be formulated in NFSv4.1 requests as in any other ONC RPC request. The reasons include:
Given that well formulated XIDs continue to be required, this begs the question why SEQUENCE and CB_SEQUENCE replies have a session id, slot id and sequence id? Having the session id in the reply means the requester does not have to use the XID to lookup the session id, which would be necessary if the connection were associated with multiple sessions. Having the slot id and sequence id in the reply means requester does not have to use the XID to lookup the slot id and sequence id. Furhermore, since the XID is only 32 bits, it is too small to guarantee the re-association of a reply with its request ([27] (Werme, R., “RPC XID Issues,” February 1996.)); having session id, slot id, and sequence id in the reply allows the client to validate that the reply in fact belongs to the matched request.
The SEQUENCE (and CB_SEQUENCE) operation also carries a "highest_slotid" value which carries additional requester slot usage information. The requester must always indicate the slot id representing the outstanding request with the highest-numbered slot value. The requester should in all cases provide the most conservative value possible, although it can be increased somewhat above the actual instantaneous usage to maintain some minimum or optimal level. This provides a way for the requester to yield unused request slots back to the replier, which in turn can use the information to reallocate resources.
The replier responds with both a new target highest_slotid, and an enforced highest_slotid, described as follows:
| TOC |
When a SEQUENCE or CB_SEQUENCE operation is successfully executed, its reply MUST always be cached. Specifically, session id, sequence id, and slot id MUST be cached in the reply cache. The reply from SEQUENCE also includes the highest slot id, target highest slot id, and status flags. Instead of caching these values, the server MAY re-compute the values from the current state of the fore channel, session and/or client ID as appropriate. Similarly, the reply from CB_SEQUENCE includes a highest slot id and target highest slot id. The client MAY re-compute the values from the current state of the session as appropriate.
Regardless of whether a replier is re-computing highest slot id, target slot id, and status on replies to retries or not, the requester MUST NOT assume the values are being re-computed whenever it receives a reply after a retry is sent, since it has no way of knowing whether the reply it has received was sent by the server in response to the retry, or is a delayed response to the original request. Therefore, it may be the case that highest slot id, target slot id, or status bits may reflect the state of affairs when the request was first executed. Although acting based on such delayed information is valid, it may cause the receiver to do unneeded work. Requesters MAY choose to send additional requests to get the current state of affairs or use the state of affairs reported by subsequent requests, in preference to acting immediately on data which may be out of date.
| TOC |
Any time SEQUENCE or CB_SEQUENCE return an error, the sequence id of the slot MUST NOT change. The replier MUST NOT modify the reply cache entry for the slot whenever an error is returned from SEQUENCE or CB_SEQUENCE.
| TOC |
On a per-request basis the requester can choose to direct the replier to cache the reply to all operations after the first operation (SEQUENCE or CB_SEQUENCE) via the sa_cachethis or csa_cachethis fields of the arguments to SEQUENCE or CB_SEQUENCE. The reason it would not direct the replier to cache the entire reply is that the request is composed of all idempotent operations [24] (Juszczak, C., “Improving the Performance and Correctness of an NFS Server,” June 1990.). Caching the reply may offer little benefit. If the reply is too large (see Section 2.10.5.4 (COMPOUND and CB_COMPOUND Construction Issues)), it may not be cacheable anyway. Even if the reply to idempotent request is small enough to cache, unnecessarily caching the reply slows down the server and increases RPC latency.
Whether the requester requests the reply to be cached or not has no effect on the slot processing. If the results of SEQUENCE or CB_SEQUENCE are NFS4_OK, then the slot's sequence id MUST be incremented by one. If a requester does not direct the replier to cache the reply, the replier MUST do one of following:
| TOC |
A requester MUST NOT retry a request, unless the connection it used to send the request disconnects. The requester can then reconnect and re-send the request, or it can re-send the request over a different connection that is associated with the same session.
If the requester is a server wanting to re-send a callback operation over the backchannel of session, the requester of course cannot reconnect because only the client can associate connections with the backchannel. The server can re-send the request over another connection that is bound to the same session's backchannel. If there is no such connection, the server MUST indicate that the session has no backchannel by setting the SEQ4_STATUS_CB_PATH_DOWN_SESSION flag bit in the response to the next SEQUENCE operation from the client. The client MUST then associate a connection with the session (or destroy the session).
Note that it is not fatal for a client to retry without a disconnect between the request and retry. However the retry does consume resources, especially with RDMA, where each request, retry or not, consumes a credit. Retries for no reason, especially retries sent shortly after the previous attempt, are a poor use of network bandwidth and defeat the purpose of a transport's inherent congestion control system.
A requester MUST wait for a reply to a request before using the slot for another request. If it does not wait for a reply, then the requester does not know what sequence id to use for the slot on its next request. For example, suppose a requester sends a request with sequence id 1, and does not wait for the response. The next time it uses the slot, it sends the new request with sequence id 2. If the replier has not seen the request with sequence id 1, then the replier is not expecting sequence id 2, and rejects the requester's new request with NFS4ERR_SEQ_MISORDERED (as the result from SEQUENCE or CB_SEQUENCE).
RDMA fabrics do not guarantee that the memory handles (Steering Tags) within each RPC/RDMA "chunk" ([8] (Talpey, T. and B. Callaghan, “RDMA Transport for ONC RPC - A Work in Progress,” May 2007.)) are valid on a scope outside that of a single connection. Therefore, handles used by the direct operations become invalid after connection loss. The server must ensure that any RDMA operations which must be replayed from the reply cache use the newly provided handle(s) from the most recent request.
A retry might be sent while the original request is still in progress on the replier. The replier SHOULD deal with the issue by returning NFS4ERR_DELAY as the reply to SEQUENCE or CB_SEQUENCE operation, but implementations MAY return NFS4ERR_MISORDERED. Since errors from SEQUENCE and CB_SEQUENCE are never recorded in the reply cache, this approach allows the results of the execution of the original request to be properly recorded in the reply cache (assuming the requester specified the reply to be cached).
| TOC |
It is possible for server callbacks to arrive at the client before the reply from related fore channel operations. For example, a client may have been granted a delegation to a file it has opened, but the reply to the OPEN (informing the client of the granting of the delegation) may be delayed in the network. If a conflicting operation arrives at the server, it will recall the delegation using the backchannel, which may be on a different transport connection, perhaps even a different network, or even a different session associated with the same client ID
The presence of a session between client and server alleviates this issue. When a session is in place, each client request is uniquely identified by its { session id, slot id, sequence id } triple. By the rules under which slot entries (reply cache entries) are retired, the server has knowledge whether the client has "seen" each of the server's replies. The server can therefore provide sufficient information to the client to allow it to disambiguate between an erroneous or conflicting callback race condition.
For each client operation which might result in some sort of server callback, the server SHOULD "remember" the { session id, slot id, sequence id } triple of the client request until the slot id retirement rules allow the server to determine that the client has, in fact, seen the server's reply. Until the time the { session id, slot id, sequence id } request triple can be retired, any recalls of the associated object MUST carry an array of these referring identifiers (in the CB_SEQUENCE operation's arguments), for the benefit of the client. After this time, it is not necessary for the server to provide this information in related callbacks, since it is certain that a race condition can no longer occur.
The CB_SEQUENCE operation which begins each server callback carries a list of "referring" { session id, slot id, sequence id } triples. If the client finds the request corresponding to the referring session id, slot id and sequence id to be currently outstanding (i.e. the server's reply has not been seen by the client), it can determine that the callback has raced the reply, and act accordingly. If the client does not find the request corresponding the referring triple to be outstanding (including the case of a session id referring to a destroyed session), then there is no race with respect to this triple. The server SHOULD limit the referring triples to requests that refer to just those that apply to the objects referred to in the CB_COMPOUND procedure.
The client must not simply wait forever for the expected server reply to arrive before responding to the CB_COMPOUND that won the race, because it is possible that it will be delayed indefinitely. The client should assume the likely case that the reply will arrive within the average round trip time for COMPOUND requests to the server, and wait that period of time. If that period of time expires it can respond to the CB_COMPOUND with NFS4ERR_DELAY.
There are other scenarios under which callbacks may race replies. Among them are pNFS layout recalls as described in Section 12.5.5.2 (Sequencing of Layout Operations).
| TOC |
Very large requests and replies may pose both buffer management issues (especially with RDMA) and reply cache issues. When the session is created, (Section 18.36 (Operation 43: CREATE_SESSION - Create New Session and Confirm Client ID)), for each channel (fore and back), the client and server negotiate the maximum sized request they will send or process (ca_maxrequestsize), the maximum sized reply they will return or process (ca_maxresponsesize), and the maximum sized reply they will store in the reply cache (ca_maxresponsesize_cached).
If a request exceeds ca_maxrequestsize, the reply will have the status NFS4ERR_REQ_TOO_BIG. A replier MAY return NFS4ERR_REQ_TOO_BIG as the status for first operation (SEQUENCE or CB_SEQUENCE) in the request (which means no operations in the request executed, and the state of the slot in the reply cache is unchanged), or it MAY opt to return it on a subsequent operation in the same COMPOUND or CB_COMPOUND request (which means at least one operation did execute and the state of the slot in reply cache does change). The replier SHOULD set NFS4ERR_REQ_TOO_BIG on the operation that exceeds ca_maxrequestsize.
If a reply exceeds ca_maxresponsesize, the reply will have the status NFS4ERR_REP_TOO_BIG. A replier MAY return NFS4ERR_REP_TOO_BIG as the status for first operation (SEQUENCE or CB_SEQUENCE) in the request, or it MAY opt to return it on a subsequent operation (in the same COMPOUND or CB_COMPOUND reply). A replier MAY return NFS4ERR_REP_TOO_BIG in the reply to SEQUENCE or CB_SEQUENCE, even if the response would still exceed ca_maxresponsesize.
If sa_cachethis or csa_cachethis are TRUE, then the replier MUST cache a reply except if an error is returned by the SEQUENCE or CB_SEQUENCE operation (see Section 2.10.5.1.2 (Errors from SEQUENCE and CB_SEQUENCE)). If the reply exceeds ca_maxresponsesize_cached, (and sa_cachethis or csa_cachethis are TRUE) then the server MUST return NFS4ERR_REP_TOO_BIG_TO_CACHE. Even if NFS4ERR_REP_TOO_BIG_TO_CACHE (or any other error for that matter) is returned on a operation other than first operation (SEQUENCE or CB_SEQUENCE), then the reply MUST be cached if sa_cachethis or csa_cachethis are TRUE. For example, if a COMPOUND has eleven operations, including SEQUENCE, the fifth operation is a RENAME, and the tenth operation is a READ for one million bytes, the server may return NFS4ERR_REP_TOO_BIG_TO_CACHE on the tenth operation. Since the server executed several operations, especially the non-idempotent RENAME, the client's request to cache the reply needs to be honored in order for correct operation of exactly once semantics. If the client retries the request, the server will have cached a reply that contains results for ten of the eleven requested operations, with the tenth operation having a status of NFS4ERR_REP_TOO_BIG_TO_CACHE.
A client needs to take care that when sending operations that change the current filehandle (except for PUTFH, PUTPUBFH, PUTROOTFH and RESTOREFH) that it not exceed the maximum reply buffer before the GETFH operation. Otherwise the client will have to retry the operation that changed the current filehandle, in order to obtain the desired filehandle. For the OPEN operation (see Section 18.16 (Operation 18: OPEN - Open a Regular File)), retry is not always available as an option. The following guidelines for the handling of filehandle changing operations are advised:
| TOC |
Since the reply cache is bounded, it is practical for the reply cache to persist across server restarts. The replier MUST persist the following information if it agreed to persist the session (when the session was created; see Section 18.36 (Operation 43: CREATE_SESSION - Create New Session and Confirm Client ID)):
The above are sufficient for a replier to provide EOS semantics for any requests that were sent and executed before the server restarted. If the replier is a client then there is no need for it to persist any more information, unless the client will be persisting all other state across client restart. In which case, the server will never see any NFSv4.1-level protocol manifestation of a client restart. If the replier is a server, with just the slot table and session id persisting, any requests the client retries after the server restart will return the results that are cached in reply cache. and any new requests (i.e. the sequence id is one (1) greater than the slot's sequence id) MUST be rejected with NFS4ERR_DEADSESSION (returned by SEQUENCE). Such a session is considered dead. A server MAY re-animate a session after a server restart so that the session will accept new requests as well as retries. To re-animate a session the server needs to persist additional information through server restart:
A persistent reply cache places certain demands on the server. The execution of the sequence of operations (starting with SEQUENCE) and placement of its results in the persistent cache MUST be atomic. If a client retries an sequence of operations that was previously executed on the server the only acceptable outcomes are either the original cached reply or an indication that client ID or session has been lost (indicating a catastrophic loss of the reply cache or a session that has been deleted because the client failed to use the session for an extended period of time).
A server could fail and restart in the middle of a COMPOUND procedure that contains one or more non-idempotent or idempotent-but-modifying operations. This creates an even higher challenge for atomic execution and placement of results in the reply cache. One way to view the problem is as a single transaction consisting of each operation in the COMPOUND followed by storing the result in persistent storage, then finally a transaction commit. If there is a failure before the transaction is committed, then the server rolls back the transaction. If server itself fails, then when it restarts, its recovery logic could roll back the transaction before starting the NFSv4.1 server.
While the description of the implementation for atomic execution of the request and caching of the reply is beyond the scope of this document, an example implementation for NFSv2 [28] (Nowicki, B., “NFS: Network File System Protocol specification,” March 1989.) is described in [29] (Bhide, A., Elnozahy, E., and S. Morgan, “A Highly Available Network Server,” January 1991.).
| TOC |
A complete discussion of the operation of RPC-based protocols over RDMA transports is in [8] (Talpey, T. and B. Callaghan, “RDMA Transport for ONC RPC - A Work in Progress,” May 2007.). A discussion of the operation of NFSv4, including NFSv4.1, over RDMA is in [9] (Talpey, T. and B. Callaghan, “NFS Direct Data Placement - A Work in Progress,” May 2007.). Where RDMA is considered, this specification assumes the use of such a layering; it addresses only the upper layer issues relevant to making best use of RPC/RDMA.
| TOC |
RDMA requires its consumers to register memory and post buffers of a specific size and number for receive operations.
Registration of memory can be a relatively high-overhead operation, since it requires pinning of buffers, assignment of attributes (e.g. readable/writable), and initialization of hardware translation. Preregistration is desirable to reduce overhead. These registrations are specific to hardware interfaces and even to RDMA connection endpoints, therefore negotiation of their limits is desirable to manage resources effectively.
Following basic registration, these buffers must be posted by the RPC layer to handle receives. These buffers remain in use by the RPC/NFSv4.1 implementation; the size and number of them must be known to the remote peer in order to avoid RDMA errors which would cause a fatal error on the RDMA connection.
NFSv4.1 manages slots as resources on a per session basis (see Section 2.10 (Session)), while RDMA connections manage credits on a per connection basis. This means that in order for a peer to send data over RDMA to a remote buffer, it has to have both an NFSv4.1 slot, and an RDMA credit. If multiple RDMA connections are associated with a session, then if the total number of credits across all RDMA connections associated with the session is X, and the number slots in the session is Y, then the maximum number of outstanding requests is lesser of X and Y.
| TOC |
Previous versions of NFS do not provide flow control; instead they rely on the windowing provided by transports like TCP to throttle requests. This does not work with RDMA, which provides no operation flow control and will terminate a connection in error when limits are exceeded. Limits such as maximum number of requests outstanding are therefore negotiated when a session is created (see the ca_maxrequests field in Section 18.36 (Operation 43: CREATE_SESSION - Create New Session and Confirm Client ID)). These limits then provide the maxima which each connection associated with the session's channel(s) must remain within. RDMA connections are managed within these limits as described in section 3.3 ("Flow Control"[Comment.2] (RFC Editor: please verify section and title of the RPCRDMA document)) of [8] (Talpey, T. and B. Callaghan, “RDMA Transport for ONC RPC - A Work in Progress,” May 2007.); if there are multiple RDMA connections, then the maximum number of requests for a channel will be divided among the RDMA connections. Put a different way, the onus is on the replier to ensure that total number of RDMA credits across all connections associated with the replier's channel does exceed the channel's maximum number of outstanding requests.
The limits may also be modified dynamically at the replier's choosing by manipulating certain parameters present in each NFSv4.1 reply. In addition, the CB_RECALL_SLOT callback operation (see Section 20.8 (Operation 10: CB_RECALL_SLOT - change flow control limits)) can be sent by a server to a client to return RDMA credits to the server, thereby lowering the maximum number of requests a client can have outstanding to the server.
| TOC |
Header padding is requested by each peer at session initiation (see the ca_headerpadsize argument to CREATE_SESSION in Section 18.36 (Operation 43: CREATE_SESSION - Create New Session and Confirm Client ID)), and subsequently used by the RPC RDMA layer, as described in [8] (Talpey, T. and B. Callaghan, “RDMA Transport for ONC RPC - A Work in Progress,” May 2007.). Zero padding is permitted.
Padding leverages the useful property that RDMA preserve alignment of data, even when they are placed into anonymous (untagged) buffers. If requested, client inline writes will insert appropriate pad bytes within the request header to align the data payload on the specified boundary. The client is encouraged to add sufficient padding (up to the negotiated size) so that the "data" field of the NFSv4.1 WRITE operation is aligned. Most servers can make good use of such padding, which allows them to chain receive buffers in such a way that any data carried by client requests will be placed into appropriate buffers at the server, ready for file system processing. The receiver's RPC layer encounters no overhead from skipping over pad bytes, and the RDMA layer's high performance makes the insertion and transmission of padding on the sender a significant optimization. In this way, the need for servers to perform RDMA Read to satisfy all but the largest client writes is obviated. An added benefit is the reduction of message round trips on the network - a potentially good trade, where latency is present.
The value to choose for padding is subject to a number of criteria. A primary source of variable-length data in the RPC header is the authentication information, the form of which is client-determined, possibly in response to server specification. The contents of COMPOUNDs, sizes of strings such as those passed to RENAME, etc. all go into the determination of a maximal NFSv4.1 request size and therefore minimal buffer size. The client must select its offered value carefully, so as not to overburden the server, and vice- versa. The payoff of an appropriate padding value is higher performance. [Comment.3] (RFC editor please keep this diagram on one page.)
Sender gather:
|RPC Request|Pad bytes|Length| -> |User data...|
\------+----------------------/ \
\ \
\ Receiver scatter: \-----------+- ...
/-----+----------------\ \ \
|RPC Request|Pad|Length| -> |FS buffer|->|FS buffer|->...
In the above case, the server may recycle unused buffers to the next posted receive if unused by the actual received request, or may pass the now-complete buffers by reference for normal write processing. For a server which can make use of it, this removes any need for data copies of incoming data, without resorting to complicated end-to-end buffer advertisement and management. This includes most kernel-based and integrated server designs, among many others. The client may perform similar optimizations, if desired.
| TOC |
Some RDMA transports (for example [10] (Recio, P., Metzler, B., Culley, P., Hilland, J., and D. Garcia, “A Remote Direct Memory Access Protocol Specification,” October 2007.)), permit a "streaming" (non-RDMA) phase, where ordinary traffic might flow before "stepping up" to RDMA mode, commencing RDMA traffic. Some RDMA transports start connections always in RDMA mode. NFSv4.1 allows, but does not assume, a streaming phase before RDMA mode. When a connection is associated with a session, the client and server negotiate whether the connection is used in RDMA or non-RDMA mode (see Section 18.36 (Operation 43: CREATE_SESSION - Create New Session and Confirm Client ID) and Section 18.34 (Operation 41: BIND_CONN_TO_SESSION)).
| TOC |
| TOC |
Via session / connection association, NFSv4.1 improves security over that provided by NFSv4.0 for the backchannel. The connection is client-initiated (see Section 18.34 (Operation 41: BIND_CONN_TO_SESSION)), and subject to the same firewall and routing checks as the fore channel. The connection cannot be hijacked by an attacker who connects to the client port prior to the intended server as is possible with NFSv4.0. At the client's option (see Section 18.35 (Operation 42: EXCHANGE_ID - Instantiate Client ID)), connection association is fully authenticated before being activated (see Section 18.34 (Operation 41: BIND_CONN_TO_SESSION)). Traffic from the server over the backchannel is authenticated exactly as the client specifies (see Section 2.10.7.2 (Backchannel RPC Security)).
| TOC |
When the NFSv4.1 client establishes the backchannel, it informs the server of the security flavors and principals to use when sending requests. If the security flavor is RPCSEC_GSS, the client expresses the principal in the form of an established RPCSEC_GSS context. The server is free to use any of the flavor/principal combinations the client offers, but it MUST NOT use unoffered combinations. This way, the client need not provide a target GSS principal for the backchannel as it did with NFSv4.0, nor the server have to implement an RPCSEC_GSS initiator as it did with NFSv4.0 [21] (Shepler, S., Callaghan, B., Robinson, D., Thurlow, R., Beame, C., Eisler, M., and D. Noveck, “Network File System (NFS) version 4 Protocol,” April 2003.).
The CREATE_SESSION (Section 18.36 (Operation 43: CREATE_SESSION - Create New Session and Confirm Client ID)) and BACKCHANNEL_CTL (Section 18.33 (Operation 40: BACKCHANNEL_CTL - Backchannel control)) operations allow the client to specify flavor/principal combinations.
Also note that the SP4_SSV state protection mode (see Section 18.35 (Operation 42: EXCHANGE_ID - Instantiate Client ID) and Section 2.10.7.3 (Protection from Unauthorized State Changes)) has the side benefit of providing SSV-derived RPCSEC_GSS contexts (Section 2.10.8 (The SSV GSS Mechanism)).
| TOC |
As described to this point in the specification, the state model of NFSv4.1 is vulnerable to an attacker that sends a SEQUENCE operation with a forged session id and with a slot id that it expects the legitimate client to use next. When the legitimate client uses the slot id with the same sequence number, the server returns the attacker's result from the reply cache which disrupts the legitimate client and thus denies service to it. Similarly an attacker could send a CREATE_SESSION with a forged client ID to create a new session associated with the client ID. The attacker could send requests using the new session that change locking state, such as LOCKU operations to release locks the legitimate client has acquired. Setting a security policy on the file which requires RPCSEC_GSS credentials when manipulating the file's state is one potential work around, but has the disadvantage of preventing a legitimate client from releasing state when RPCSEC_GSS is required to do so, but a GSS context cannot be obtained (possibly because the user has logged off the client).
NFSv4.1 provides three options to a client for state protection which are specified when a client creates a client ID via EXCHANGE_ID (Section 18.35 (Operation 42: EXCHANGE_ID - Instantiate Client ID)).
The first (SP4_NONE) is to simply waive state protection.
The other two options (SP4_MACH_CRED and SP4_SSV) share several traits:
The SP4_MACH_CRED state protection option uses a machine credential where the principal that creates the client ID, must also be the principal that performs client ID and session maintenance operations. The security of the machine credential state protection approach depends entirely on safe guarding the per-machine credential. Assuming a proper safe guard, using the per-machine credential for operations like CREATE_SESSION, BIND_CONN_TO_SESSION, DESTROY_SESSION, and DESTROY_CLIENTID will prevent an attacker from associating a rogue connection with a session, or associating a rogue session with a client ID.
There are at least three scenarios for the SP4_MACH_CRED option:
The SP4_SSV protection option uses a Secret State Verifier (SSV) which is shared between a client and server. The SSV serves as the secret key for an internal (that is, internal to NFSv4.1) GSS mechanism that uses the secret key for Message Integrity Code (MIC) and Wrap tokens (Section 2.10.8 (The SSV GSS Mechanism)). The SP4_SSV protection option is intended for the client that has multiple users, and the system administrator does not wish to configure a permanent machine credential for each client. The SSV is established on the server via SET_SSV (see Section 18.47 (Operation 54: SET_SSV - Update SSV for a Client ID)). To prevent eavesdropping, a client SHOULD send SET_SSV via RPCSEC_GSS with the privacy service. Several aspects of the SSV make it intractable for an attacker to guess the SSV, and thus associate rogue connections with a session, and rogue sessions with a client ID:
Here are the types of attacks that can be attempted by an attacker named Eve on a victim named Bob, and how SP4_SSV protection foils each attack:
In summary, an attacker's disruption of state when SP4_SSV protection is in use is limited to the formative period of a client ID, its first session, and the establishment of the SSV. Once a non-malicious user uses the client ID, the client quickly detects any hijack and rectifies the situation. Once a non-malicious user successfully modifies the SSV, the attacker cannot use NFSv4.1 operations to disrupt the non-malicious user.
Note that neither the SP4_MACH_CRED nor SP4_SSV protection approaches prevent hijacking of a transport connection that has previously been associated with a session. If the goal of a counter threat strategy is to prevent connection hijacking, the use of IPsec is RECOMMENDED.
If a connection hijack occurs, the hijacker could in theory change locking state and negatively impact the service to legitimate clients. However if the server is configured to require the use of RPCSEC_GSS with integrity or privacy on the affected file objects, and if EXCHGID4_FLAG_BIND_PRINC_STATEID capability (Section 18.35 (Operation 42: EXCHANGE_ID - Instantiate Client ID)), is in force, this will thwart unauthorized attempts to change locking state.
| TOC |
The SSV provides the secret key for a mechanism that NFSv4.1 uses for state protection. Contexts for this mechanism are not established via the RPCSEC_GSS protocol. Instead, the contexts are automatically created when EXCHANGE_ID specifies SP4_SSV protection. The only tokens defined are the PerMsgToken (emitted by GSS_GetMIC) and the SealedMessage token (emitted by GSS_Wrap).
The mechanism OID for the SSV mechanism is: iso.org.dod.internet.private.enterprise.Michael Eisler.nfs.ssv_mech (1.3.6.1.4.1.28882.1.1). While the SSV mechanism does not define any initial context tokens, the OID can be used to let servers indicate that the SSV mechanism is acceptable whenever the client sends a SECINFO or SECINFO_NO_NAME operation (see Section 2.6 (Security Service Negotiation)).
The SSV mechanism defines four subkeys derived from the SSV value. Each time SET_SSV is invoked the subkeys are recalculated by the client and server. The calculation of each of the four subkeys depends on each of the four respective ssv_subkey4 enumerated values. The calculation uses the HMAC [11] (Krawczyk, H., Bellare, M., and R. Canetti, “HMAC: Keyed-Hashing for Message Authentication,” February 1997.), algorithm, using the current SSV as the key, the one way hash algorithm as negotiated by EXCHANGE_ID, and the input text as represented by the XDR encoded enumeration of type ssv_subkey4.
/* Input for computing subkeys */
enum ssv_subkey4 {
SSV4_SUBKEY_MIC_I2T = 1,
SSV4_SUBKEY_MIC_T2I = 2,
SSV4_SUBKEY_SEAL_I2T = 3,
SSV4_SUBKEY_SEAL_T2I = 4
};
The subkey derived from SSV4_SUBKEY_MIC_I2T is used for calculating message integrity codes (MICs) that originate from the NFSv4.1 client, whether as part of a request over the fore channel, or a response over the backchannel. The subkey derived from SSV4_SUBKEY-MIST2I is used for MICs originating from the NFSv4.1 server. The subkey derived from SSV4_SUBKEY_SEAL_I2T is used for encryption text originating from the NFSv4.1 client and the subkey derived from SSV4_SUBKEY_SEAL_T2I is used for encryption text originating from the NFSv4.1 server.
The PerMsgToken description is based on an XDR definition:
/* Input for computing smt_hmac */
struct ssv_mic_plain_tkn4 {
uint32_t smpt_ssv_seq;
opaque smpt_orig_plain<>;
};
/* SSV GSS PerMsgToken token */
struct ssv_mic_tkn4 {
uint32_t smt_ssv_seq;
opaque smt_hmac<>;
};
The field smt_hmac is an HMAC calculated by using the subkey derived from SSV4_SUBKEY_MIC_I2T or SSV4_SUBKEY_MIC_T2I as the key, the one way hash algorithm as negotiated by EXCHANGE_ID, and the input text as represented by data of type ssv_mic_plain_tkn4. The field smpt_ssv_seq is the same as smt_ssv_seq. The field smpt_orig_plain is the "message" input passed to GSS_GetMIC() (see Section 2.3.1 of [7] (Linn, J., “Generic Security Service Application Program Interface Version 2, Update 1,” January 2000.)). The caller of GSS_GetMIC() provides a pointer to a buffer containing the plain text. The SSV mechanism's entry point for GSS_GetMIC() encodes this into an opaque array, and the encoding will include an initial four byte length, plus any necessary padding. Prepended to this will be the XDR encoded value of smpt_ssv_seq thus making up an XDR encoding of a value of data type ssv_mic_plain_tkn4, which in turn is the input into the HMAC.
The token emitted by GSS_GetMIC() is XDR encoded and of XDR data type ssv_mic_tkn4. The field smt_ssv_seq comes from the SSV sequence number which is equal to 1 after SET_SSV (Section 18.47 (Operation 54: SET_SSV - Update SSV for a Client ID)) is called the first time on a client ID. Thereafter, it is incremented on each SET_SSV. Thus smt_ssv_seq represents the version of the SSV at the time GSS_GetMIC() was called. As noted in Section 18.35 (Operation 42: EXCHANGE_ID - Instantiate Client ID), the client and server can maintain multiple concurrent versions of the SSV. This allows the SSV to be changed without serializing all RPC calls that use the SSV mechanism with SET_SSV operations. Once the HMAC is calculated, it is XDR encoded into smt_hmac, which will include an initial four byte length, and any necessary padding. Prepended to this will be the XDR encoded value of smt_ssv_seq.
The SealedMessage description is based on an XDR definition:
/* Input for computing ssct_encr_data and ssct_hmac */
struct ssv_seal_plain_tkn4 {
opaque sspt_confounder<>;
uint32_t sspt_ssv_seq;
opaque sspt_orig_plain<>;
opaque sspt_pad<>;
};
/* SSV GSS SealedMessage token */
struct ssv_seal_cipher_tkn4 {
uint32_t ssct_ssv_seq;
opaque ssct_iv<>;
opaque ssct_encr_data<>;
opaque ssct_hmac<>;
};
The token emitted by GSS_Wrap() is XDR encoded and of XDR data type ssv_seal_cipher_tkn4.
The ssct_ssv_seq field has the same meaning as smt_ssv_seq.
The ssct_encr_data field is the result of encrypting a value of the XDR encoded data type ssv_seal_plain_tkn4. The encryption key is the subkey derived from SSV4_SUBKEY_SEAL_I2T or SSV4_SUBKEY_SEAL_T2I, and the encryption algorithm is that negotiated by EXCHANGE_ID.
The ssct_iv field is the initialization vector (IV) for the encryption algorithm (if applicable) and is sent in clear text. The content and size of the IV MUST comply with specification of the encryption algorithm. For example, the id-aes256-CBC algorithm MUST use a 16 byte initialization vector (IV) which MUST be unpredictable for each instance of a value of type ssv_seal_plain_tkn4 that is encrypted with a particular SSV key.
The ssct_hmac field is the result of computing an HMAC using value of the XDR encoded data type ssv_seal_plain_tkn4 as the input text. The key is the subkey derived from SSV4_SUBKEY_MIC_I2T or SSV4_SUBKEY_MIC_T2I, and the one way hash algorithm is that negotiated by EXCHANGE_ID.
The sspt_confounder field is a random value.
The sspt_ssv_seq field is the same as ssvt_ssv_seq.
The field sspt_orig_plain field is the original plaintext and is the "input_message" input passed to GSS_Wrap() (see Section 2.3.3 of [7] (Linn, J., “Generic Security Service Application Program Interface Version 2, Update 1,” January 2000.)). As with the handling of the plaintext by the SSV mechanism's GSS_GetMIC() entry point, the entry point for GSS_Wrap() expects a pointer to the plaintext, and will XDR encode an opaque array into sspt_orig_plain representing the plain text, along with the other fields of an instance of data type ssv_seal_plain_tkn4.
The sspt_pad field is present to support encryption algorithms that require inputs to be in fixed sized blocks. The content of sspt_pad is zero filled except for the length. Beware that the XDR encoding of ssv_seal_plain_tkn4 contains three variable length arrays, and so each array consumes four bytes for an array length, and each array that follows the length is always padded to a multiple of four bytes per the XDR standard.
For example suppose the encryption algorithm uses 16 byte blocks, and the sspt_confounder is three bytes long, and the sspt_orig_plain field is 15 bytes long. The XDR encoding of sspt_confounder uses eight bytes (4 + 3 + 1 byte pad), the XDR encoding of sspt_ssv_seq uses four bytes, the XDR encoding of sspt_orig_plain uses 20 bytes (4 + 15 + 1 byte pad), and the smallest XDR encoding of the sspt_pad field is four bytes. This totals 36 bytes. The next multiple of 16 is 48, thus the length field of sspt_pad needs to be set to 12 bytes, or a total encoding of 16 bytes. The total number of XDR encoded bytes is thus 8 + 4 + 20 + 16 = 48.
GSS_Wrap() emits a token that is an XDR encoding of a value of data type ssv_seal_cipher_tkn4. Note that regardless whether the caller of GSS_Wrap() requests confidentiality or not, the token always has confidentiality. This is because the SSV mechanism is for RPCSEC_GSS, and RPCSEC_GSS never produces GSS_wrap() tokens without confidentiality.
There is one SSV per client ID. Effectively there is a single GSS context for a client ID / SSV pair. All SSV mechanism RPCSEC_GSS handles of a client ID / SSV pair share the same GSS context. SSV GSS contexts do not expire except when the SSV is destroyed (causes would include the client ID being destroyed or a server restart). Since one purpose of context expiration is to replace keys that have been in use for "too long" hence vulnerable to compromise by brute force or accident, the client can replace the SSV key by sending periodic SET_SSV operations, by cycling through different users' RPCSEC_GSS credentials. This way the SSV is replaced without destroying the SSV's GSS contexts.
SSV RPCSEC_GSS handles can be expired or deleted by the server at any time and the EXCHANGE_ID operation can be used to create more SSV RPCSEC_GSS handles. Expiration of SSV RPCSEC_GSS handles does not imply that the SSV or its GSS context have expired.
The client MUST establish an SSV via SET_SSV before the SSV GSS context can be used to emit tokens from GSS_Wrap() and GSS_GetMIC(). If SET_SSV has not been successfully called, attempts to emit tokens MUST fail.
The SSV mechanism does not support replay detection and sequencing in its tokens because RPCSEC_GSS does not use those features (See Section 5.2.2 "Context Creation Requests" in [4] (Eisler, M., Chiu, A., and L. Ling, “RPCSEC_GSS Protocol Specification,” September 1997.)).
| TOC |
| TOC |
The server has the primary obligation to monitor the state of backchannel resources that the client has created for the server (RPCSEC_GSS contexts and backchannel connections). If these resources vanish, the server takes action as specified in Section 2.10.11.2 (Events Requiring Server Action).
| TOC |
The client SHOULD honor the following obligations in order to utilize the session:
| TOC |
If the client does not have a client ID, the client sends EXCHANGE_ID to establish a client ID. If it opts for SP4_MACH_CRED or SP4_SSV protection, in the spo_must_enforce list of operations, it SHOULD at minimum specify: CREATE_SESSION, DESTROY_SESSION, BIND_CONN_TO_SESSION, BACKCHANNEL_CTL, and DESTROY_CLIENTID. If opts for SP4_SSV protection, the client needs to ask for SSV-based RPCSEC_GSS handles.
The client uses the client ID to send a CREATE_SESSION on a connection to the server. The results of CREATE_SESSION indicate whether the server will persist the session reply cache through a server restarted or not, and the client notes this for future reference.
If the client specified SP4_SSV state protection when the client ID was created, then it SHOULD send SET_SSV in the first COMPOUND after the session is created. Each time a new principal goes to use the client ID, it SHOULD send a SET_SSV again.
If the client wants to use delegations, layouts, directory notifications, or any other state that requires a backchannel, then it must add a connection to the backchannel if CREATE_SESSION did not already do so. The client creates a connection, and calls BIND_CONN_TO_SESSION to associate the connection with the session and the session's backchannel. If CREATE_SESSION did not already do so, the client MUST tell the server what security is required in order for the client to accept callbacks. The client does this via BACKCHANNEL_CTL. If the client selected SP4_MACH_CRED or SP4_SSV protection when it called EXCHANGE_ID, then the client SHOULD specify that the backchannel use RPCSEC_GSS contexts for security.
If the client wants to use additional connections for the backchannel, then it must call BIND_CONN_TO_SESSION on each connection it wants to use with the session. If the client wants to use additional connections for the fore channel, then it must call BIND_CONN_TO_SESSION if it specified SP4_SSV or SP4_MACH_CRED state protection when the client ID was created.
At this point the session has reached steady state.
| TOC |
The server MAY maintain a session inactivity timer for each session. If the session inactivity timer expires, then the server MAY destroy the session. To avoid losing a session due to inactivity, the client MUST renew the session inactivity timer. The length of session inactivity timer MUST NOT be less than the lease_time attribute (Section 5.8.1.11 (Attribute 10: lease_time)). As with lease renewal (Section 8.3 (Lease Renewal)), when the server receives a SEQUENCE operation, it resets the session inactivity timer, and MUST NOT allow the timer to expire while the rest of the operations in the COMPOUND procedure's request are still executing. Once the last operation has finished, the server MUST set the session inactivity timer to expire no sooner that the sum of the current time and the value of the lease_time attribute.
| TOC |
| TOC |
The following events require client action to recover.
| TOC |
If all RPCSEC_GSS contexts granted by the client to the server for callback use have expired, the client MUST establish a new context via BACKCHANNEL_CTL. The sr_status_flags field of the SEQUENCE results indicates when callback contexts are nearly expired, or fully expired (see Section 18.46.3 (DESCRIPTION)).
| TOC |
If the client loses the last connection of the session, and if wants to retain the session, then it must create a new connection, and if, when the client ID was created, BIND_CONN_TO_SESSION was specified in the spo_must_enforce list, the client MUST use BIND_CONN_TO_SESSION to associate the connection with the session.
If there was a request outstanding at the time the of connection loss, then if client wants to continue to use the session it MUST retry the request, as described in Section 2.10.5.2 (Retry and Replay of Reply). Note that it is not necessary to retry requests over a connection with the same source network address or the same destination network address as the lost connection. As long as the session id, slot id, and sequence id in the retry match that of the original request, the server will recognize the request as a retry if it executed the request prior to disconnect.
If the connection that was lost was the last one associated with the backchannel, and the client wants to retain the backchannel and/or not put recallable state subject to revocation, the client must reconnect, and if it does, it MUST associate the connection to the session and backchannel via BIND_CONN_TO_SESSION. The server SHOULD indicate when it has no callback connection via the sr_status_flags result from SEQUENCE.
| TOC |
Via the sr_status_flags result of the SEQUENCE operation or other means, the client will learn if some or all of the RPCSEC_GSS contexts it assigned to the backchannel have been lost. If the client wants to the retain the backchannel and/or not put recallable state subjection to revocation, the client must use BACKCHANNEL_CTL to assign new contexts.
| TOC |
The replier might lose a record of the session. Causes include:
Loss of reply cache is equivalent to loss of session. The replier indicates loss of session to the requester by returning NFS4ERR_BADSESSION on the next operation that uses the session id that refers to the lost session.
After an event like a server restart, the client may have lost its connections. The client assumes for the moment that the session has not been lost. It reconnects, and if it specified connection association enforcement when the session was created, it invokes BIND_CONN_TO_SESSION using the session id. Otherwise, it invokes SEQUENCE. If BIND_CONN_TO_SESSION or SEQUENCE returns NFS4ERR_BADSESSION, the client knows the session was lost. If the connection survives session loss, then the next SEQUENCE operation the client sends over the connection will get back NFS4ERR_BADSESSION. The client again knows the session was lost.
When the client detects session loss, it must call CREATE_SESSION to recover. Any non-idempotent operations that were in progress may have been performed on the server at the time of session loss. The client has no general way to recover from this.
Note that loss of session does not imply loss of lock, open, delegation, or layout state because locks, opens, delegations, and layouts are tied to the client ID and depend on the client ID, not the session. Nor does loss of lock, open, delegation, or layout state imply loss of session state, because the session depends on the client ID; loss of client ID however does imply loss of session, lock, open, delegation, and layout state. See Section 8.4.2 (Server Failure and Recovery). A session can survive a server restart, but lock recovery may still be needed.
It is possible CREATE_SESSION will fail with NFS4ERR_STALE_CLIENTID (for example the server restarts and does not preserve client ID state). If so, the client needs to call EXCHANGE_ID, followed by CREATE_SESSION.
| TOC |
The following events require server action to recover.
| TOC |
As described in Section 18.35 (Operation 42: EXCHANGE_ID - Instantiate Client ID), a restarted client sends EXCHANGE_ID in such a way it causes the server to delete any sessions it had.
| TOC |
If a client crashes and never comes back, it will never send EXCHANGE_ID with its old client owner. Thus the server has session state that will never be used again. After an extended period of time and if the server has resource constraints, it MAY destroy the old session as well as locking state.
| TOC |
To the server, the extended network partition may be no different from a client crash with no restart (see Section 2.10.11.2.2 (Client Crash with No Restart)). Unless the server can discern that there is a network partition, it is free to treat the situation as if the client has crashed permanently.
| TOC |
If there were callback requests outstanding at the time of a connection loss, then the server MUST retry the request, as described in Section 2.10.5.2 (Retry and Replay of Reply). Note that it is not necessary to retry requests over a connection with the same source network address or the same destination network address as the lost connection. As long as the session id, slot id, and sequence id in the retry match that of the original request, the callback target will recognize the request as a retry even if it did see the request prior to disconnect.
If the connection lost is the last one associated with the backchannel, then the server MUST indicate that in the sr_status_flags field of every SEQUENCE reply until the backchannel is reestablished. There are two situations each of which use different status flags: no connectivity for the session's backchannel, and no connectivity for any session backchannel of the client. See Section 18.46 (Operation 53: SEQUENCE - Supply per-procedure sequencing and control) for a description of the appropriate flags in sr_status_flags.
| TOC |
The server SHOULD monitor when the number RPCSEC_GSS contexts assigned to the backchannel reaches one, and when that one context is near expiry (i.e. between one and two periods of lease time), indicate so in the sr_status_flags field of all SEQUENCE replies. The server MUST indicate when the all of the backchannel's assigned RPCSEC_GSS contexts have expired in the sr_status_flags field of all SEQUENCE replies.
| TOC |
A client and server can potentially be a non-pNFS implementation, a metadata server implementation, a data server implementation, or two or three types of implementations. The EXCHGID4_FLAG_USE_NON_PNFS, EXCHGID4_FLAG_USE_PNFS_MDS, and EXCHGID4_FLAG_USE_PNFS_DS flags (not mutually exclusive) are passed in the EXCHANGE_ID arguments and results to allow the client to indicate how it wants to use sessions created under the client ID, and to allow the server to indicate how it will allow the sessions to be used. See Section 13.1 (Client ID and Session Considerations) for pNFS sessions considerations.
| TOC |
The syntax and semantics to describe the data types of the NFSv4.1 protocol are defined in the XDR RFC4506 (Eisler, M., “XDR: External Data Representation Standard,” May 2006.) [2] and RPC RFC1831 (Srinivasan, R., “RPC: Remote Procedure Call Protocol Specification Version 2,” August 1995.) [3] documents. The next sections build upon the XDR data types to define constants, types and structures specific to this protocol. The full list of XDR data types is in [12] (Shepler, S., Eisler, M., and D. Noveck, “NFSv4 Minor Version 1 XDR Description A Work in Progress,” December 2007.).
| TOC |
const NFS4_FHSIZE = 128; const NFS4_VERIFIER_SIZE = 8; const NFS4_OPAQUE_LIMIT = 1024; const NFS4_SESSIONID_SIZE = 16; const NFS4_INT64_MAX = 0x7fffffffffffffff; const NFS4_UINT64_MAX = 0xffffffffffffffff; const NFS4_INT32_MAX = 0x7fffffff; const NFS4_UINT32_MAX = 0xffffffff; const NFS4_MAXFILELEN = 0xffffffffffffffff; const NFS4_MAXFILEOFF = 0xfffffffffffffffe;
Except where noted, all these constants are defined in bytes.
| TOC |
These are the base NFSv4.1 data types.
| Data Type | Definition |
|---|---|
| int32_t | typedef int int32_t; |
| uint32_t | typedef unsigned int uint32_t; |
| int64_t | typedef hyper int64_t; |
| uint64_t | typedef unsigned hyper uint64_t; |
| attrlist4 | typedef opaque attrlist4<>; |
| Used for file/directory attributes. | |
| bitmap4 | typedef uint32_t bitmap4<>; |
| Used in attribute array encoding. | |
| changeid4 | typedef uint64_t changeid4; |
| Used in the definition of change_info4. | |
| clientid4 | typedef uint64_t clientid4; |
| Shorthand reference to client identification. | |
| count4 | typedef uint32_t count4; |
| Various count parameters (READ, WRITE, COMMIT). | |
| length4 | typedef uint64_t length4; |
| Describes LOCK lengths. | |
| mode4 | typedef uint32_t mode4; |
| Mode attribute data type. | |
| nfs_cookie4 | typedef uint64_t nfs_cookie4; |
| Opaque cookie value for READDIR. | |
| nfs_fh4 | typedef opaque nfs_fh4<NFS4_FHSIZE>; |
| Filehandle definition. | |
| nfs_ftype4 | enum nfs_ftype4; |
| Various defined file types. | |
| nfsstat4 | enum nfsstat4; |
| Return value for operations. | |
| offset4 | typedef uint64_t offset4; |
| Various offset designations (READ, WRITE, LOCK, COMMIT). | |
| qop4 | typedef uint32_t qop4; |
| Quality of protection designation in SECINFO. | |
| sec_oid4 | typedef opaque sec_oid4<>; |
| Security Object Identifier. The sec_oid4 data type is not really opaque. Instead it contains an ASN.1 OBJECT IDENTIFIER as used by GSS-API in the mech_type argument to GSS_Init_sec_context. See [7] (Linn, J., “Generic Security Service Application Program Interface Version 2, Update 1,” January 2000.) for details. | |
| sequenceid4 | typedef uint32_t sequenceid4; |
| Sequence number used for various session operations (EXCHANGE_ID, CREATE_SESSION, SEQUENCE, CB_SEQUENCE). | |
| seqid4 | typedef uint32_t seqid4; |
| Sequence identifier used for file locking. | |
| sessionid4 | typedef opaque sessionid4[NFS4_SESSIONID_SIZE]; |
| Session identifier. | |
| slotid4 | typedef uint32_t slotid4; |
| Sequencing artifact for various session operations (SEQUENCE, CB_SEQUENCE). | |
| utf8string | typedef opaque utf8string<>; |
| UTF-8 encoding for strings. | |
| utf8str_cis | typedef utf8string utf8str_cis; |
| Case-insensitive UTF-8 string. | |
| utf8str_cs | typedef utf8string utf8str_cs; |
| Case-sensitive UTF-8 string. | |
| utf8str_mixed | typedef utf8string utf8str_mixed; |
| UTF-8 strings with a case sensitive prefix and a case insensitive suffix. | |
| component4 | typedef utf8str_cs component4; |
| Represents path name components. | |
| linktext4 | typedef utf8str_cs linktext4; |
| Symbolic link contents. | |
| pathname4 | typedef component4 pathname4<>; |
| Represents path name for fs_locations. | |
| verifier4 | typedef opaque verifier4[NFS4_VERIFIER_SIZE]; |
| Verifier used for various operations (COMMIT, CREATE, EXCHANGE_ID, OPEN, READDIR, WRITE) NFS4_VERIFIER_SIZE is defined as 8. |
End of Base Data Types
| Table 1 |
| TOC |
| TOC |
struct nfstime4 {
int64_t seconds;
uint32_t nseconds;
};
The nfstime4 data type gives the number of seconds and nanoseconds since midnight or 0 hour January 1, 1970 Coordinated Universal Time (UTC). Values greater than zero for the seconds field denote dates after the 0 hour January 1, 1970. Values less than zero for the seconds field denote dates before the 0 hour January 1, 1970. In both cases, the nseconds field is to be added to the seconds field for the final time representation. For example, if the time to be represented is one-half second before 0 hour January 1, 1970, the seconds field would have a value of negative one (-1) and the nseconds fields would have a value of one-half second (500000000). Values greater than 999,999,999 for nseconds are invalid.
This data type is used to pass time and date information. A server converts to and from its local representation of time when processing time values, preserving as much accuracy as possible. If the precision of timestamps stored for a file system object is less than defined, loss of precision can occur. An adjunct time maintenance protocol is RECOMMENDED to reduce client and server time skew.
| TOC |
enum time_how4 {
SET_TO_SERVER_TIME4 = 0,
SET_TO_CLIENT_TIME4 = 1
};
| TOC |
union settime4 switch (time_how4 set_it) {
case SET_TO_CLIENT_TIME4:
nfstime4 time;
default:
void;
};
The time_how4 and settime4 data types are used for setting timestamps in file object attributes. If set_it is SET_TO_SERVER_TIME4, then the server uses its local representation of time for the time value.
| TOC |
struct specdata4 {
uint32_t specdata1; /* major device number */
uint32_t specdata2; /* minor device number */
};
This data type represents the device numbers for the device file types NF4CHR and NF4BLK.
| TOC |
struct fsid4 {
uint64_t major;
uint64_t minor;
};
| TOC |
struct change_policy4 {
uint64_t cp_major;
uint64_t cp_minor;
};
The chg_policy4 data type is used for the change_policy RECOMMENDED attribute. It provides change sequencing indication analogous to the change attribute. To enable the server to present a value valid across server re-initialization without requiring persistent storage, two 64-bit quantities are used, allowing one to be a server instance id and the second to be incremented non-persistently, within a given server instance.
| TOC |
struct fattr4 {
bitmap4 attrmask;
attrlist4 attr_vals;
};
The fattr4 data type is used to represent file and directory attributes.
The bitmap is a counted array of 32 bit integers used to contain bit values. The position of the integer in the array that contains bit n can be computed from the expression (n / 32) and its bit within that integer is (n mod 32).
0 1 +-----------+-----------+-----------+-- | count | 31 .. 0 | 63 .. 32 | +-----------+-----------+-----------+--
| TOC |
struct change_info4 {
bool atomic;
changeid4 before;
changeid4 after;
};
This data type is used with the CREATE, LINK, OPEN, REMOVE, and RENAME operations to let the client know the value of the change attribute for the directory in which the target file system object resides.
| TOC |
struct netaddr4 {
/* see struct rpcb in RFC 1833 */
string na_r_netid<>; /* network id */
string na_r_addr<>; /* universal address */
};
The netaddr4 data type is used to identify TCP/IP based endpoints. The r_netid and r_addr fields are specified in RFC1833 (Srinivasan, R., “Binding Protocols for ONC RPC Version 2,” August 1995.) [26], but they are underspecified in RFC1833 (Srinivasan, R., “Binding Protocols for ONC RPC Version 2,” August 1995.) [26] as far as what they should look like for specific protocols. The next section clarifies this.
| TOC |
For TCP over IPv4 and for UDP over IPv4, the format of r_addr is the US-ASCII string:
h1.h2.h3.h4.p1.p2
The prefix, "h1.h2.h3.h4", is the standard textual form for representing an IPv4 address, which is always four bytes long. Assuming big-endian ordering, h1, h2, h3, and h4, are respectively, the first through fourth bytes each converted to ASCII-decimal. The suffix, "p1.p2", is a textual form for representing a TCP and UDP service port. Assuming big-endian ordering, p1 and p2 are, respectively, the first and second bytes each converted to ASCII-decimal. For example, if a host, in big-endian order, has an address of 0x0A010307 and there is a service listening on, in big endian order, port 0x020F (decimal 527), then the complete universal address is "10.1.3.7.2.15".
For TCP over IPv4 the value of r_netid is the string "tcp". For UDP over IPv4 the value of r_netid is the string "udp". That this document specifies the universal address and netid for UDP/IPv6 does not imply that UDP/IPv4 is a legal transport for NFSv4.1 (see Section 2.9 (Transport Layers)).
| TOC |
For TCP over IPv6 and for UDP over IPv6, the format of r_addr is the US-ASCII string:
x1:x2:x3:x4:x5:x6:x7:x8.p1.p2
The suffix "p1.p2" is the service port, and is computed the same way as with universal addresses for TCP and UDP over IPv4. The prefix, "x1:x2:x3:x4:x5:x6:x7:x8", is the preferred textual form for representing an IPv6 address as defined in Section 2.2 of RFC3513 (Hinden, R. and S. Deering, “IP Version 6 Addressing Architecture,” April 2003.) [13]. Additionally, the two alternative forms specified in Section 2.2 of RFC3513 are also acceptable.
For TCP over IPv6 the value of r_netid is the string "tcp6". For UDP over IPv6 the value of r_netid is the string "udp6". That this document specifies the universal address and netid for UDP/IPv6 does not imply that UDP/IPv6 is a legal transport for NFSv4.1 (see Section 2.9 (Transport Layers)).
| TOC |
struct state_owner4 {
clientid4 clientid;
opaque owner<NFS4_OPAQUE_LIMIT>;
};
typedef state_owner4 open_owner4;
typedef state_owner4 lock_owner4;
The state_owner4 data type is the base type for the open_owner4 Section 3.3.10.1 (open_owner4) and lock_owner4 Section 3.3.10.2 (lock_owner4).
| TOC |
This data type is used to identify the owner of open state.
| TOC |
This structure is used to identify the owner of byte-range locking state.
| TOC |
struct open_to_lock_owner4 {
seqid4 open_seqid;
stateid4 open_stateid;
seqid4 lock_seqid;
lock_owner4 lock_owner;
};
This data type is used for the first LOCK operation done for an open_owner4. It provides both the open_stateid and lock_owner such that the transition is made from a valid open_stateid sequence to that of the new lock_stateid sequence. Using this mechanism avoids the confirmation of the lock_owner/lock_seqid pair since it is tied to established state in the form of the open_stateid/open_seqid.
| TOC |
struct stateid4 {
uint32_t seqid;
opaque other[12];
};
This data type is used for the various state sharing mechanisms between the client and server. The client never modifies a value of data type stateid. The starting value of the seqid field is undefined. The server is required to increment the seqid field by one (1) at each transition of the stateid. This is important since the client will inspect the seqid in OPEN stateids to determine the order of OPEN processing done by the server.
| TOC |
enum layouttype4 {
LAYOUT4_NFSV4_1_FILES = 0x1,
LAYOUT4_OSD2_OBJECTS = 0x2,
LAYOUT4_BLOCK_VOLUME = 0x3
};
This data type indicates what type of layout is being used. The file server advertises the layout types it supports through the fs_layout_type file system attribute (Section 5.12.1 (Attribute 62: fs_layout_type)). A client asks for layouts of a particular type in LAYOUTGET, and processes those layouts in its layout-type-specific logic.
The layouttype4 data type is 32 bits in length. The range represented by the layout type is split into three parts. Type 0x0 is reserved. Types within the range 0x00000001-0x7FFFFFFF are globally unique and are assigned according to the description in Section 22.4 (Defining New Layout Types); they are maintained by IANA. Types within the range 0x80000000-0xFFFFFFFF are site specific and for private use only.
The LAYOUT4_NFSV4_1_FILES enumeration specifies that the NFSv4.1 file layout type, as defined in Section 13 (PNFS: NFSv4.1 File Layout Type), is to be used. The LAYOUT4_OSD2_OBJECTS enumeration specifies that the object layout, as defined in [30] (Halevy, B., Welch, B., and J. Zelenka, “Object-based pNFS Operations,” April 2008.), is to be used. Similarly, the LAYOUT4_BLOCK_VOLUME enumeration specifies that the block/volume layout, as defined in [31] (Black, D., Fridella, S., and J. Glasgow, “pNFS Block/Volume Layout,” April 2008.), is to be used.
| TOC |
const NFS4_DEVICEID4_SIZE = 16; typedef opaque deviceid4[NFS4_DEVICEID4_SIZE];
Layout information includes device IDs that specify a storage device through a compact handle. Addressing and type information is obtained with the GETDEVICEINFO operation. Device IDs are not guaranteed to be valid across metadata server restarts. A device ID is unique per client ID and layout type. See Section 12.2.10 (Device IDs) for more details.
| TOC |
struct device_addr4 {
layouttype4 da_layout_type;
opaque da_addr_body<>;
};
The device address is used to set up a communication channel with the storage device. Different layout types will require different data types to define how they communicate with storage devices. The opaque da_addr_body field must be interpreted based on the specified da_layout_type field.
This document defines the device address for the NFSv4.1 file layout (see Section 13.3 (File Layout Data Types)), which identifies a storage device by network IP address and port number. This is sufficient for the clients to communicate with the NFSv4.1 storage devices, and may be sufficient for other layout types as well. Device types for object storage devices and block storage devices (e.g., SCSI volume labels) will be defined by their respective layout specifications.
| TOC |
struct layout_content4 {
layouttype4 loc_type;
opaque loc_body<>;
};
The loc_body field must be interpreted based on the layout type (loc_type). This document defines the loc_body for the NFSv4.1 file layout type is defined; see Section 13.3 (File Layout Data Types) for its definition.
| TOC |
struct layout4 {
offset4 lo_offset;
length4 lo_length;
layoutiomode4 lo_iomode;
layout_content4 lo_content;
};
The layout4 data type defines a layout for a file. The layout type specific data is opaque within lo_content. Since layouts are sub-dividable, the offset and length together with the file's filehandle, the client ID, iomode, and layout type, identify the layout.
| TOC |
struct layoutupdate4 {
layouttype4 lou_type;
opaque lou_body<>;
};
The layoutupdate4 data type is used by the client to return updated layout information to the metadata server via the LAYOUTCOMMIT (Section 18.42 (Operation 49: LAYOUTCOMMIT - Commit writes made using a layout)) operation. This data type provides a channel to pass layout type specific information (in field lou_body) back to the metadata server. E.g., for the block/volume layout type this could include the list of reserved blocks that were written. The contents of the opaque lou_body argument are determined by the layout type. The NFSv4.1 file-based layout does not use this data type; if lou_type is LAYOUT4_NFSV4_1_FILES, the lou_body field MUST have a zero length.
| TOC |
struct layouthint4 {
layouttype4 loh_type;
opaque loh_body<>;
};
The layouthint4 data type is used by the client to pass in a hint about the type of layout it would like created for a particular file. It is the data type specified by the layout_hint attribute described in Section 5.12.4 (Attribute 63: layout_hint). The metadata server may ignore the hint, or may selectively ignore fields within the hint. This hint should be provided at create time as part of the initial attributes within OPEN. The loh_body field is specific to the type of layout (loh_type). The NFSv4.1 file-based layout uses the nfsv4_1_file_layouthint4 data type as defined in Section 13.3 (File Layout Data Types).
| TOC |
enum layoutiomode4 {
LAYOUTIOMODE4_READ = 1,
LAYOUTIOMODE4_RW = 2,
LAYOUTIOMODE4_ANY = 3
};
The iomode specifies whether the client intends to just read or both read and write the data represented by the layout. While the LAYOUTIOMODE4_ANY iomode MUST NOT be used in the arguments to the LAYOUTGET operation, it MAY be used in the arguments to the LAYOUTRETURN and CB_LAYOUTRECALL operations. The LAYOUTIOMODE4_ANY iomode specifies that layouts pertaining to both LAYOUTIOMODE4_READ and LAYOUTIOMODE4_RW iomodes are being returned or recalled, respectively. The metadata server's use of the iomode may depend on the layout type being used. The storage devices MAY validate I/O accesses against the iomode and reject invalid accesses.
| TOC |
struct nfs_impl_id4 {
utf8str_cis nii_domain;
utf8str_cs nii_name;
nfstime4 nii_date;
};
This data type is used to identify client and server implementation details. The nii_domain field is the DNS domain name that the implementer is associated with. The nii_name field is the product name of the implementation and is completely free form. It is RECOMMENDED that the nii_name be used to distinguish machine architecture, machine platforms, revisions, versions, and patch levels. The nii_date field is the timestamp of when the software instance was published or built.
| TOC |
struct threshold_item4 {
layouttype4 thi_layout_type;
bitmap4 thi_hintset;
opaque thi_hintlist<>;
};
This data type contains a list of hints specific to a layout type for helping the client determine when it should send I/O directly through the metadata server versus the storage devices. The data type consists of the layout type (thi_layout_type), a bitmap (thi_hintset) describing the set of hints supported by the server (they may differ based on the layout type), and a list of hints (thi_hintlist), whose content is determined by the hintset bitmap. See the mdsthreshold attribute for more details.
The thi_hintset field is a bitmap of the following values:
| name | # | Data Type | Description |
|---|---|---|---|
| threshold4_read_size | 0 | length4 | The file size below which it is RECOMMENDED to read data through the MDS. |
| threshold4_write_size | 1 | length4 | The file size below which it is RECOMMENDED to write data through the MDS. |
| threshold4_read_iosize | 2 | length4 | For read I/O sizes below this threshold it is RECOMMENDED to read data through the MDS |
| threshold4_write_iosize | 3 | length4 | For write I/O sizes below this threshold it is RECOMMENDED to write data through the MDS |
| TOC |
struct mdsthreshold4 {
threshold_item4 mth_hints<>;
};
This data type holds an array of elements of data type threshold_item4, each of which is valid for a particular layout type. An array is necessary because a server can support multiple layout types for a single file.
| TOC |
The filehandle in the NFS protocol is a per server unique identifier for a file system object. The contents of the filehandle are opaque to the client. Therefore, the server is responsible for translating the filehandle to an internal representation of the file system object.
| TOC |
The operations of the NFS protocol are defined in terms of one or more filehandles. Therefore, the client needs a filehandle to initiate communication with the server. With the NFSv3 protocol RFC1813 (Callaghan, B., Pawlowski, B., and P. Staubach, “NFS Version 3 Protocol Specification,” June 1995.) [22], there exists an ancillary protocol to obtain this first filehandle. The MOUNT protocol, RPC program number 100005, provides the mechanism of translating a string based file system path name to a filehandle which can then be used by the NFS protocols.
The MOUNT protocol has deficiencies in the area of security and use via firewalls. This is one reason that the use of the public filehandle was introduced in RFC2054 (Callaghan, B., “WebNFS Client Specification,” October 1996.) [32] and RFC2055 (Callaghan, B., “WebNFS Server Specification,” October 1996.) [33]. With the use of the public filehandle in combination with the LOOKUP operation in the NFSv3 protocol, it has been demonstrated that the MOUNT protocol is unnecessary for viable interaction between NFS client and server.
Therefore, the NFSv4.1 protocol will not use an ancillary protocol for translation from string based path names to a filehandle. Two special filehandles will be used as starting points for the NFS client.
| TOC |
The first of the special filehandles is the ROOT filehandle. The ROOT filehandle is the "conceptual" root of the file system name space at the NFS server. The client uses or starts with the ROOT filehandle by employing the PUTROOTFH operation. The PUTROOTFH operation instructs the server to set the "current" filehandle to the ROOT of the server's file tree. Once this PUTROOTFH operation is used, the client can then traverse the entirety of the server's file tree with the LOOKUP operation. A complete discussion of the server name space is in the Section 7 (Single-server Namespace).
| TOC |
The second special filehandle is the PUBLIC filehandle. Unlike the ROOT filehandle, the PUBLIC filehandle may be bound or represent an arbitrary file system object at the server. The server is responsible for this binding. It may be that the PUBLIC filehandle and the ROOT filehandle refer to the same file system object. However, it is up to the administrative software at the server and the policies of the server administrator to define the binding of the PUBLIC filehandle and server file system object. The client may not make any assumptions about this binding. The client uses the PUBLIC filehandle via the PUTPUBFH operation.
| TOC |
In the NFSv3 protocol, there was one type of filehandle with a single set of semantics. This type of filehandle is termed "persistent" in NFSv4.1. The semantics of a persistent filehandle remain the same as before. A new type of filehandle introduced in NFSv4.1 is the "volatile" filehandle, which attempts to accommodate certain server environments.
The volatile filehandle type was introduced to address server functionality or implementation issues which make correct implementation of a persistent filehandle infeasible. Some server environments do not provide a file system level invariant that can be used to construct a persistent filehandle. The underlying server file system may not provide the invariant or the server's file system programming interfaces may not provide access to the needed invariant. Volatile filehandles may ease the implementation of server functionality such as hierarchical storage management or file system reorganization or migration. However, the volatile filehandle increases the implementation burden for the client.
Since the client will need to handle persistent and volatile filehandles differently, a file attribute is defined which may be used by the client to determine the filehandle types being returned by the server.
| TOC |
The filehandle contains all the information the server needs to distinguish an individual file. To the client, the filehandle is opaque. The client stores filehandles for use in a later request and can compare two filehandles from the same server for equality by doing a byte-by-byte comparison. However, the client MUST NOT otherwise interpret the contents of filehandles. If two filehandles from the same server are equal, they MUST refer to the same file. Servers SHOULD try to maintain a one-to-one correspondence between filehandles and files but this is not required. Clients MUST use filehandle comparisons only to improve performance, not for correct behavior. All clients need to be prepared for situations in which it cannot be determined whether two filehandles denote the same object and in such cases, avoid making invalid assumptions which might cause incorrect behavior. Further discussion of filehandle and attribute comparison in the context of data caching is presented in the Section 10.3.4 (Data Caching and File Identity).
As an example, in the case that two different path names when traversed at the server terminate at the same file system object, the server SHOULD return the same filehandle for each path. This can occur if a hard link is used to create two file names which refer to the same underlying file object and associated data. For example, if paths /a/b/c and /a/d/c refer to the same file, the server SHOULD return the same filehandle for both path names traversals.
| TOC |
A persistent filehandle is defined as having a fixed value for the lifetime of the file system object to which it refers. Once the server creates the filehandle for a file system object, the server MUST accept the same filehandle for the object for the lifetime of the object. If the server restarts, the NFS server must honor the same filehandle value as it did in the server's previous instantiation. Similarly, if the file system is migrated, the new NFS server must honor the same filehandle as the old NFS server.
The persistent filehandle will be become stale or invalid when the file system object is removed. When the server is presented with a persistent filehandle that refers to a deleted object, it MUST return an error of NFS4ERR_STALE. A filehandle may become stale when the file system containing the object is no longer available. The file system may become unavailable if it exists on removable media and the media is no longer available at the server or the file system in whole has been destroyed or the file system has simply been removed from the server's name space (i.e. unmounted in a UNIX environment).
| TOC |
A volatile filehandle does not share the same longevity characteristics of a persistent filehandle. The server may determine that a volatile filehandle is no longer valid at many different points in time. If the server can definitively determine that a volatile filehandle refers to an object that has been removed, the server should return NFS4ERR_STALE to the client (as is the case for persistent filehandles). In all other cases where the server determines that a volatile filehandle can no longer be used, it should return an error of NFS4ERR_FHEXPIRED.
The REQUIRED attribute "fh_expire_type" is used by the client to determine what type of filehandle the server is providing for a particular file system. This attribute is a bitmask with the following values:
- FH4_PERSISTENT
- The value of FH4_PERSISTENT is used to indicate a persistent filehandle, which is valid until the object is removed from the file system. The server will not return NFS4ERR_FHEXPIRED for this filehandle. FH4_PERSISTENT is defined as a value in which none of the bits specified below are set.
- FH4_VOLATILE_ANY
- The filehandle may expire at any time, except as specifically excluded (i.e. FH4_NO_EXPIRE_WITH_OPEN).
- FH4_NOEXPIRE_WITH_OPEN
- May only be set when FH4_VOLATILE_ANY is set. If this bit is set, then the meaning of FH4_VOLATILE_ANY is qualified to exclude any expiration of the filehandle when it is open.
- FH4_VOL_MIGRATION
- The filehandle will expire as a result of a file system transition (migration or replication), in those case in which the continuity of filehandle use is not specified by handle class information within the fs_locations_info attribute. When this bit is set, clients without access to fs_locations_info information should assume filehandles will expire on file system transitions.
- FH4_VOL_RENAME
- The filehandle will expire during rename. This includes a rename by the requesting client or a rename by any other client. If FH4_VOL_ANY is set, FH4_VOL_RENAME is redundant.
Servers which provide volatile filehandles that may expire while open (i.e. if FH4_VOL_MIGRATION or FH4_VOL_RENAME is set or if FH4_VOLATILE_ANY is set and FH4_NOEXPIRE_WITH_OPEN not set), should deny a RENAME or REMOVE that would affect an OPEN file of any of the components leading to the OPEN file. In addition, the server should deny all RENAME or REMOVE requests during the grace period upon server restart.
Servers which provide volatile filehandles that may expire while open require special care as regards handling of RENAMESs and REMOVEs. This situation can arise if FH4_VOL_MIGRATION or FH4_VOL_RENAME is set, if FH4_VOLATILE_ANY is set and FH4_NOEXPIRE_WITH_OPEN not set, or if a non-readonly file system has a transition target in a different handle class. In these cases, the server should deny a RENAME or REMOVE that would affect an OPEN file of any of the components leading to the OPEN file. In addition, the server should deny all RENAME or REMOVE requests during the grace period, in order to make sure that reclaims of files where filehandles may have expired do not do a reclaim for the wrong file.
Volatile filehandles are especially suitable for implementation of the pseudo file systems used to bridge exports. See Section 7.5 (Filehandle Volatility) for a discussion of this.
| TOC |
A volatile filehandle, while opaque to the client could contain:
[volatile bit = 1 | server boot time | slot | generation number]
When the client presents a volatile filehandle, the server makes the following checks, which assume that the check for the volatile bit has passed. If the server boot time is less than the current server boot time, return NFS4ERR_FHEXPIRED. If slot is out of range, return NFS4ERR_BADHANDLE. If the generation number does not match, return NFS4ERR_FHEXPIRED.
When the server restarts, the table is gone (it is volatile).
If volatile bit is 0, then it is a persistent filehandle with a different structure following it.
| TOC |
If possible, the client SHOULD recover from the receipt of an NFS4ERR_FHEXPIRED error. The client must take on additional responsibility so that it may prepare itself to recover from the expiration of a volatile filehandle. If the server returns persistent filehandles, the client does not need these additional steps.
For volatile filehandles, most commonly the client will need to store the component names leading up to and including the file system object in question. With these names, the client should be able to recover by finding a filehandle in the name space that is still available or by starting at the root of the server's file system name space.
If the expired filehandle refers to an object that has been removed from the file system, obviously the client will not be able to recover from the expired filehandle.
It is also possible that the expired filehandle refers to a file that has been renamed. If the file was renamed by another client, again it is possible that the original client will not be able to recover. However, in the case that the client itself is renaming the file and the file is open, it is possible that the client may be able to recover. The client can determine the new path name based on the processing of the rename request. The client can then regenerate the new filehandle based on the new path name. The client could also use the compound operation mechanism to construct a set of operations like:
RENAME A B
LOOKUP B
GETFH
Note that the COMPOUND procedure does not provide atomicity. This example only reduces the overhead of recovering from an expired filehandle.
| TOC |
To meet the requirements of extensibility and increased interoperability with non-UNIX platforms, attributes must be handled in a flexible manner. The NFSv3 fattr3 structure contains a fixed list of attributes that not all clients and servers are able to support or care about. The fattr3 structure can not be extended as new needs arise and it provides no way to indicate non-support. With the NFSv4.1 protocol, the client is able query what attributes the server supports and construct requests with only those supported attributes (or a subset thereof).
To this end, attributes are divided into three groups: REQUIRED, RECOMMENDED, and named. Both REQUIRED and RECOMMENDED attributes are supported in the NFSv4.1 protocol by a specific and well-defined encoding and are identified by number. They are requested by setting a bit in the bit vector sent in the GETATTR request; the server response includes a bit vector to list what attributes were returned in the response. New REQUIRED or RECOMMENDED attributes may be added to the NFSv4 protocol as part of a new minor version by publishing a standards-track RFC which allocates a new attribute number value and defines the encoding for the attribute. See Section 2.7 (Minor Versioning) for further discussion.
Named attributes are accessed by the new OPENATTR operation, which accesses a hidden directory of attributes associated with a file system object. OPENATTR takes a filehandle for the object and returns the filehandle for the attribute hierarchy. The filehandle for the named attributes is a directory object accessible by LOOKUP or READDIR and contains files whose names represent the named attributes and whose data bytes are the value of the attribute. For example:
| LOOKUP | "foo" | ; look up file |
| GETATTR | attrbits | |
| OPENATTR | ; access foo's named attributes | |
| LOOKUP | "x11icon" | ; look up specific attribute |
| READ | 0,4096 | ; read stream of bytes |
Named attributes are intended for data needed by applications rather than by an NFS client implementation. NFS implementors are strongly encouraged to define their new attributes as RECOMMENDED attributes by bringing them to the IETF standards-track process.
The set of attributes which are classified as REQUIRED is deliberately small since servers must do whatever it takes to support them. A server should support as many of the RECOMMENDED attributes as possible but by their definition, the server is not required to support all of them. Attributes are deemed REQUIRED if the data is both needed by a large number of clients and is not otherwise reasonably computable by the client when support is not provided on the server.
Note that the hidden directory returned by OPENATTR is a convenience for protocol processing. The client should not make any assumptions about the server's implementation of named attributes and whether the underlying file system at the server has a named attribute directory or not. Therefore, operations such as SETATTR and GETATTR on the named attribute directory are undefined.
| TOC |
These MUST be supported by every NFSv4.1 client and server in order to ensure a minimum level of interoperability. The server MUST store and return these attributes and the client MUST be able to function with an attribute set limited to these attributes. With just the REQUIRED attributes some client functionality may be impaired or limited in some ways. A client may ask for any of these attributes to be returned by setting a bit in the GETATTR request and the server must return their value.
| TOC |
These attributes are understood well enough to warrant support in the NFSv4.1 protocol. However, they may not be supported on all clients and servers. A client may ask for any of these attributes to be returned by setting a bit in the GETATTR request but must handle the case where the server does not return them. A client may ask for the set of attributes the server supports and SHOULD NOT request attributes the server does not support. A server should be tolerant of requests for unsupported attributes and simply not return them rather than considering the request an error. It is expected that servers will support all attributes they comfortably can and only fail to support attributes which are difficult to support in their operating environments. A server should provide attributes whenever they don't have to "tell lies" to the client. For example, a file modification time should be either an accurate time or should not be supported by the server. This will not always be comfortable to clients but the client is better positioned decide whether and how to fabricate or construct an attribute or whether to do without the attribute.
| TOC |
These attributes are not supported by direct encoding in the NFSv4 protocol but are accessed by string names rather than numbers and correspond to an uninterpreted stream of bytes which are stored with the file system object. The name space for these attributes may be accessed by using the OPENATTR operation. The OPENATTR operation returns a filehandle for a virtual "named attribute directory" and further perusal and modification of the name space may be done using operations that work on more typical directories. In particular, READDIR may be used to get a list of such named attributes and LOOKUP and OPEN may select a particular attribute. Creation of a new named attribute may be the result of an OPEN specifying file creation.
Once an OPEN is done, named attributes may be examined and changed by normal READ and WRITE operations using the filehandles and stateids returned by OPEN.
Named attributes and the named attribute directory may have their own (non-named) attributes. Each of objects must have all of the REQUIRED attributes and may have additional RECOMMENDED attributes. However, the set of attributes for named attributes and the named attribute directory need not be as large as, and typically will not be as large as that for other objects in that file system.
Named attributes and the named attribute directory may be the target of delegations (in the case of the named attribute directory these will be directory delegations). However, since granting of delegations or not is within the server's discretion, a server need not support delegations on named attributes or the named attribute directory.
It is RECOMMENDED that servers support arbitrary named attributes. A client should not depend on the ability to store any named attributes in the server's file system. If a server does support named attributes, a client which is also able to handle them should be able to copy a file's data and metadata with complete transparency from one location to another; this would imply that names allowed for regular directory entries are valid for named attribute names as well.
In NFSv4.1, the structure of named attribute directories is restricted in a number of ways, in order to prevent the development of non-interoperable implementations in which some servers support a fully general hierarchical directory structure for named attributes while others support a limited set, but fully adequate to the feature's goals. In such an environment, clients or applications might come to depend on non-portable extensions. The restrictions are:
Names of attributes will not be controlled by this document or other IETF standards track documents. See Section 22.1 (Named Attribute Definitions) for further discussion.
| TOC |
Each of the REQUIRED and RECOMMENDED attributes can be classified in one of three categories: per server, per file system, or per file system object. Note that it is possible that some per file system attributes may vary within the file system. See the "homogeneous" attribute for its definition. Note that the attributes time_access_set and time_modify_set are not listed in this section because they are write-only attributes corresponding to time_access and time_modify, and are used in a special instance of SETATTR.
lease_time
supported_attrs, suppattr_exclcreat, fh_expire_type, link_support, symlink_support, unique_handles, aclsupport, cansettime, case_insensitive, case_preserving, chown_restricted, files_avail, files_free, files_total, fs_locations, homogeneous, maxfilesize, maxname, maxread, maxwrite, no_trunc, space_avail, space_free, space_total, time_delta, change_policy, fs_status, fs_layout_type, fs_locations_info, fs_charset_cap
type, change, size, named_attr, fsid, rdattr_error, filehandle, acl, archive, fileid, hidden, maxlink, mimetype, mode, numlinks, owner, owner_group, rawdev, space_used, system, time_access, time_backup, time_create, time_metadata, time_modify, mounted_on_fileid, dir_notif_delay, dirent_notif_delay, dacl, sacl, layout_type, layout_hint, layout_blksize, layout_alignment, mdsthreshold, retention_get, retention_set, retentevt_get, retentevt_set, retention_hold, mode_set_masked
For quota_avail_hard, quota_avail_soft, and quota_used see their definitions below for the appropriate classification.
| TOC |
Some REQUIRED and RECOMMENDED attributes are set-only, i.e. they can be set via SETATTR but not retrieved via GETATTR. Similarly, some REQUIRED and RECOMMENDED attributes are get-only, i.e. they can be retrieved GETATTR but not set via SETATTR. If a client attempts to set a get-only attribute or get a set-only attributes, the server MUST return NFS4ERR_INVAL.
| TOC |
The list of REQUIRED attributes appears in Table 4. The meaning of hte columns of the table are:
| Name | Id | Data Type | Acc | Defined in: |
|---|---|---|---|---|
| supported_attrs | 0 | bitmap4 | R | Section 5.8.1.1 (Attribute 0: supported_attrs) |
| type | 1 | nfs_ftype4 | R | Section 5.8.1.2 (Attribute 1: type) |
| fh_expire_type | 2 | uint32_t | R | Section 5.8.1.3 (Attribute 2: fh_expire_type) |
| change | 3 | uint64_t | R | Section 5.8.1.4 (Attribute 3: change) |
| size | 4 | uint64_t | R W | Section 5.8.1.5 (Attribute 4: size) |
| link_support | 5 | bool | R | Section 5.8.1.6 (Attribute 5: link_support) |
| symlink_support | 6 | bool | R | Section 5.8.1.7 (Attribute 6: symlink_support) |
| named_attr | 7 | bool | R | Section 5.8.1.8 (Attribute 7: named_attr) |
| fsid | 8 | fsid4 | R | Section 5.8.1.9 (Attribute 8: fsid) |
| unique_handles | 9 | bool | R | Section 5.8.1.10 (Attribute 9: unique_handles) |
| lease_time | 10 | nfs_lease4 | R | Section 5.8.1.11 (Attribute 10: lease_time) |
| rdattr_error | 11 | enum | R | Section 5.8.1.12 (Attribute 11: rdattr_error) |
| filehandle | 19 | nfs_fh4 | R | Section 5.8.1.13 (Attribute 19: filehandle) |
| suppattr_exclcreat | 75 | bitmap4 | R | Section 5.8.1.14 (Attribute 75: suppattr_exclcreat) |
| Table 4 |
| TOC |
The RECOMMENDED attributes are defined in Table 5. The meanings of the column headers are the same as Table 4; see Section 5.6 (REQUIRED Attributes - List and Definition References) for the meanings.
| Table 5 |
* fs_locations_info4
| TOC |
| TOC |
| TOC |
The bit vector which would retrieve all REQUIRED and RECOMMENDED attributes that are supported for this object. The scope of this attribute applies to all objects with a matching fsid.
| TOC |
Designates the type of an object in terms of one of a number of special constants:
Within the explanatory text and operation descriptions, the following phrases will be used with the meanings given below:
| TOC |
Server uses this to specify filehandle expiration behavior to the client. See Section 4 (Filehandles) for additional description.
| TOC |
A value created by the server that the client can use to determine if file data, directory contents or attributes of the object have been modified. The server may return the object's time_metadata attribute for this attribute's value but only if the file system object can not be updated more frequently than the resolution of time_metadata.
| TOC |
The size of the object in bytes.
| TOC |
True, if the object's file system supports hard links.
| TOC |
True, if the object's file system supports symbolic links.
| TOC |
True, if this object has named attributes. In other words, object has a non-empty named attribute directory.
| TOC |
Unique file system identifier for the file system holding this object. fsid contains major and minor components each of which are of data type uint64_t.
| TOC |
True, if two distinct filehandles guaranteed to refer to two different file system objects.
| TOC |
Duration of leases at server in seconds.
| TOC |
Error returned from getattr during readdir.
| TOC |
The filehandle of this object (primarily for readdir requests).
| TOC |
The bit vector which would set all REQUIRED and RECOMMENDED attributes that are supported by the EXCLUSIVE4_1 method of file creation via the OPEN operation. The scope of this attribute applies to all objects with a matching fsid.
| TOC |
The definitions of most of the RECOMMENDED attributes follow. Collections that share a common category are defined in other sections.
| TOC |
True, if this file has been archived since the time of last modification (deprecated in favor of time_backup).
| TOC |
True, if the server able to change the times for a file system object as specified in a SETATTR operation.
| TOC |
True, if file name comparisons on this file system are case insensitive.
| TOC |
True, if file name case on this file system is preserved.
| TOC |
A value created by the server that the client can use to determine if some server policy related to the current file system has been subject to change. If the value remains the same then the client can be sure that the values of the attributes related to fs location and the fss_type field of the fs_status attribute have not changed. On the other hand, a change in this value does necessarily imply a change in policy. It is up to the client to interrogate the server to determine if some policy relevant to it has changed. See Section 3.3.6 (chg_policy4) for details.
This attribute MUST change when the value returned by the fs_locations or fs_locations_info attribute changes, when a file system goes from read-only to writable or vice versa, or when the allowable set of security flavors for the file system or any part thereof is changed.
| TOC |
If TRUE, the server will reject any request to change either the owner or the group associated with a file if the caller is not a privileged user (for example, "root" in UNIX operating environments or in Windows 2000 the "Take Ownership" privilege).
| TOC |
A number uniquely identifying the file within the file system.
| TOC |
File slots available to this user on the file system containing this object - this should be the smallest relevant limit.
| TOC |
Free file slots on the file system containing this object - this should be the smallest relevant limit.
| TOC |
Total file slots on the file system containing this object.
| TOC |
Character set capabilities for this file system. See Section 14.4 (UTF-8 Capabilities).
| TOC |
Locations where this file system may be found. If the server returns NFS4ERR_MOVED as an error, this attribute MUST be supported.
| TOC |
Full function file system location.
| TOC |
Generic file system type information.
| TOC |
True, if the file is considered hidden with respect to the Windows API.
| TOC |
True, if this object's file system is homogeneous, i.e. are per file system attributes the same for all file system's objects.
| TOC |
Maximum supported file size for the file system of this object.
| TOC |
Maximum number of links for this object.
| TOC |
Maximum file name size supported for this object.
| TOC |
Maximum read size supported for this object.
| TOC |
Maximum write size supported for this object. This attribute SHOULD be supported if the file is writable. Lack of this attribute can lead to the client either wasting bandwidth or not receiving the best performance.
| TOC |
MIME body type/subtype of this object.
| TOC |
Like fileid, but if the target filehandle is the root of a file system, this attribute represents the fileid of the underlying directory.
UNIX-based operating environments connect a file system into the namespace by connecting (mounting) the file system onto the existing file object (the mount point, usually a directory) of an existing file system. When the mount point's parent directory is read via an API like readdir(), the return results are directory entries, each with a component name and a fileid. The fileid of the mount point's directory entry will be different from the fileid that the stat() system call returns. The stat() system call is returning the fileid of the root of the mounted file system, whereas readdir() is returning the fileid stat() would have returned before any file systems were mounted on the mount point.
Unlike NFSv3, NFSv4.1 allows a client's LOOKUP request to cross other file systems. The client detects the file system crossing whenever the filehandle argument of LOOKUP has an fsid attribute different from that of the filehandle returned by LOOKUP. A UNIX-based client will consider this a "mount point crossing". UNIX has a legacy scheme for allowing a process to determine its current working directory. This relies on readdir() of a mount point's parent and stat() of the mount point returning fileids as previously described. The mounted_on_fileid attribute corresponds to the fileid that readdir() would have returned as described previously.
While the NFSv4.1 client could simply fabricate a fileid corresponding to what mounted_on_fileid provides (and if the server does not support mounted_on_fileid, the client has no choice), there is a risk that the client will generate a fileid that conflicts with one that is already assigned to another object in the file system. Instead, if the server can provide the mounted_on_fileid, the potential for client operational problems in this area is eliminated.
If the server detects that there is no mounted point at the target file object, then the value for mounted_on_fileid that it returns is the same as that of the fileid attribute.
The mounted_on_fileid attribute is RECOMMENDED, so the server SHOULD provide it if possible, and for a UNIX-based server, this is straightforward. Usually, mounted_on_fileid will be requested during a READDIR operation, in which case it is trivial (at least for UNIX-based servers) to return mounted_on_fileid since it is equal to the fileid of a directory entry returned by readdir(). If mounted_on_fileid is requested in a GETATTR operation, the server should obey an invariant that has it returning a value that is equal to the file object's entry in the object's parent directory, i.e. what readdir() would have returned. Some operating environments allow a series of two or more file systems to be mounted onto a single mount point. In this case, for the server to obey the aforementioned invariant, it will need to find the base mount point, and not the intermediate mount points.
| TOC |
If this attribute is TRUE, then if the client uses a file name longer than name_max, an error will be returned instead of the name being truncated.
| TOC |
Number of hard links to this object.
| TOC |
The string name of the owner of this object.
| TOC |
The string name of the group ownership of this object.
| TOC |
The value in bytes which represents the amount of additional disk space beyond the current allocation that can be allocated to this file or directory before further allocations will be refused. It is understood that this space may be consumed by allocations to other files or directories.
| TOC |
The value in bytes which represents the amount of additional disk space that can be allocated to this file or directory before the user may reasonably be warned. It is understood that this space may be consumed by allocations to other files or directories though there is a rule as to which other files or directories.
| TOC |
The value in bytes which represent the amount of disc
space used by this file or directory and possibly a
number of other similar files or directories, where the
set of "similar" meets at least the criterion that
allocating space to any file or directory in the set
will reduce the "quota_avail_hard" of every other file
or directory in the set.
Note that there may be a number of distinct but
overlapping sets of files or directories for which a
quota_used value is maintained. E.g. "all files with a
given owner", "all files with a given group owner". etc.
The server is at liberty to choose any of those sets but
should do so in a repeatable way. The rule may be
configured per file system or may be "choose the set with
the smallest quota".
| TOC |
Raw device identifier; the UNIX device major/minor node information. If the value of type is not NF4BLK or NF4CHR, the value returned SHOULD NOT be considered useful.
| TOC |
Disk space in bytes available to this user on the file system containing this object - this should be the smallest relevant limit.
| TOC |
Free disk space in bytes on the file system containing this object - this should be the smallest relevant limit.
| TOC |
Total disk space in bytes on the file system containing this object.
| TOC |
Number of file system bytes allocated to this object.
| TOC |
This attribute is TRUE if this file is a "system" file with respect to the Windows operating environment.
| TOC |
The time_access attribute represents the time of last access to the object by a read that was satisfied by the server. The notion of what is an "access" depends on server's operating environment and/or the server's file system semantics. For example, for servers obeying POSIX semantics, time_access would be updated only by the READLINK, READ, and READDIR operations and not any of the operations that modify the content of the object. Of course, setting the corresponding time_access_set attribute is another way to modify the time_access attribute.
Whenever the file object resides on a writable file system, the server should make best efforts to record time_access into stable storage. However, to mitigate the performance effects of doing so, and most especially whenever the server is satisfying the read of the object's content from its cache, the server MAY cache access time updates and lazily write them to stable storage. It is also acceptable to give administrators of the server the option to disable time_access updates.
| TOC |
Set the time of last access to the object. SETATTR use only.
| TOC |
The time of last backup of the object.
| TOC |
The time of creation of the object. This attribute does not have any relation to the traditional UNIX file attribute "ctime" or "change time".
| TOC |
Smallest useful server time granularity.
| TOC |
The time of last metadata modification of the object.
| TOC |
The time of last modification to the object.
| TOC |
Set the time of last modification to the object. SETATTR use only.
| TOC |
The RECOMMENDED attributes "owner" and "owner_group" (and also users and groups within the "acl" attribute) are represented in terms of a UTF-8 string. To avoid a representation that is tied to a particular underlying implementation at the client or server, the use of the UTF-8 string has been chosen. Note that section 6.1 of RFC2624 (Shepler, S., “NFS Version 4 Design Considerations,” June 1999.) [34] provides additional rationale. It is expected that the client and server will have their own local representation of owner and owner_group that is used for local storage or presentation to the end user. Therefore, it is expected that when these attributes are transferred between the client and server that the local representation is translated to a syntax of the form "user@dns_domain". This will allow for a client and server that do not use the same local representation the ability to translate to a common syntax that can be interpreted by both.
Similarly, security principals may be represented in different ways by different security mechanisms. Servers normally translate these representations into a common format, generally that used by local storage, to serve as a means of identifying the users corresponding to these security principals. When these local identifiers are translated to the form of the owner attribute, associated with files created by such principals they identify, in a common format, the users associated with each corresponding set of security principals.
The translation used to interpret owner and group strings is not specified as part of the protocol. This allows various solutions to be employed. For example, a local translation table may be consulted that maps between a numeric id to the user@dns_domain syntax. A name service may also be used to accomplish the translation. A server may provide a more general service, not limited by any particular translation (which would only translate a limited set of possible strings) by storing the owner and owner_group attributes in local storage without any translation or it may augment a translation method by storing the entire string for attributes for which no translation is available while using the local representation for those cases in which a translation is available.
Servers that do not provide support for all possible values of the owner and owner_group attributes, SHOULD return an error (NFS4ERR_BADOWNER) when a string is presented that has no translation, as the value to be set for a SETATTR of the owner, owner_group, or acl attributes. When a server does accept an owner or owner_group value as valid on a SETATTR (and similarly for the owner and group strings in an acl), it is promising to return that same string when a corresponding GETATTR is done. Configuration changes (including changes from the mapping of the string to the local representation) and ill-constructed name translations (those that contain aliasing) may make that promise impossible to honor. Servers should make appropriate efforts to avoid a situation in which these attributes have their values changed when no real change to ownership has occurred.
The "dns_domain" portion of the owner string is meant to be a DNS domain name. For example, user@ietf.org. Servers should accept as valid a set of users for at least one domain. A server may treat other domains as having no valid translations. A more general service is provided when a server is capable of accepting users for multiple domains, or for all domains, subject to security constraints.
In the case where there is no translation available to the client or server, the attribute value must be constructed without the "@". Therefore, the absence of the @ from the owner or owner_group attribute signifies that no translation was available at the sender and that the receiver of the attribute should not use that string as a basis for translation into its own internal format. Even though the attribute value can not be translated, it may still be useful. In the case of a client, the attribute string may be used for local display of ownership.
To provide a greater degree of compatibility with NFSv3, which identified users and groups by 32-bit unsigned user identifiers and group identifiers, owner and group strings that consist of decimal numeric values with no leading zeros can be given a special interpretation by clients and servers which choose to provide such support. The receiver may treat such a user or group string as representing the same user as would be represented by an NFSv3 uid or gid having the corresponding numeric value. A server is not obligated to accept such a string, but may return an NFS4ERR_BADOWNER instead. To avoid this mechanism being used to subvert user and group translation, so that a client might pass all of the owners and groups in numeric form, a server SHOULD return an NFS4ERR_BADOWNER error when there is a valid translation for the user or owner designated in this way. In that case, the client must use the appropriate name@domain string and not the special form for compatibility.
The owner string "nobody" may be used to designate an anonymous user, which will be associated with a file created by a security principal that cannot be mapped through normal means to the owner attribute.
| TOC |
With respect to the case_insensitive and case_preserving attributes, each UCS-4 character (which UTF-8 encodes) has a "long descriptive name" RFC1345 (Simonsen, K., “Character Mnemonics and Character Sets,” June 1992.) [35] which may or may not include the word "CAPITAL" or "SMALL". The presence of SMALL or CAPITAL allows an NFS server to implement unambiguous and efficient table driven mappings for case insensitive comparisons, and non-case-preserving storage. For general character handling and internationalization issues, see Section 14 (Internationalization).
| TOC |
As described in Section 18.39 (Operation 46: GET_DIR_DELEGATION - Get a directory delegation), the client can request a minimum delay for notifications of changes to attributes, but the server is free to ignore what the client requests. The client can determine in advance what notification delays the server will accept by issuing a GETATTR for either or both of two directory notification attributes. When the client calls the GET_DIR_DELEGATION operation and asks for attribute change notifications, it should request notification delays that are no less than the values in the server-provided attributes.
| TOC |
The dir_notif_delay attribute is the minimum number of seconds the server will delay before notifying the client of a change to the directory's attributes.
| TOC |
The dirent_notif_delay attribute is the minimum number of seconds the server will delay before notifying the client of a change to a file object that has an entry in the directory.
| TOC |
| TOC |
The fs_layout_type attribute (see Section 3.3.13 (layouttype4)) applies to a file system and indicates what layout types are supported by the file system. When the client encounters a new fsid, the client SHOULD obtain the value for the fs_layout_type attribute associated with the new file system. This attribute is used by the client to determine if the layout types supported by the server match any of the client's supported layout types.
| TOC |
When a client holds layouts on files of a file system, the layout_alignment attribute indicates the preferred alignment for I/O to files on that file system. Where possible, the client should send READ and WRITE operations with offsets that are whole multiples of the layout_alignment attribute.
| TOC |
When a client holds layouts on files of a file system, the layout_blksize attribute indicates the preferred block size for I/O to files on that file system. Where possible, the client should send READ operations with a count argument that is a whole multiple of layout_blksize, and WRITE operations with a data argument of size that is a whole multiple of layout_blksize.
| TOC |
The layout_hint attribute (see Section 3.3.19 (layouthint4)) may be set on newly created files to influence the metadata server's choice for the file's layout. If possible, this attribute is one of those set in the initial attributes within the OPEN operation. The metadata server may choose to ignore this attribute. The layout_hint attribute is a sub-set of the layout structure returned by LAYOUTGET. For example, instead of specifying particular devices, this would be used to suggest the stripe width of a file. The server implementation determines which fields within the layout will be used.
| TOC |
This attribute lists the layout type(s) available for a file. The value returned by the server is for informational purposes only. The client will use the LAYOUTGET operation to obtain the information needed in order to perform I/O. For example, the specific device information for the file and its layout.
| TOC |
This attribute is a server provided hint used to communicate to the client when it is more efficient to send READ and WRITE operations to the metadata server or the data server. The two types of thresholds described are file size thresholds and I/O size thresholds. If a file's size is smaller than the file size threshold, data accesses SHOULD be sent to the metadata server. If an I/O request has a length that is below the I/O size threshold, the I/O SHOULD be sent to the metadata server. Each threshold type is specified separately for READ and WRITE.
The server MAY provide both types of thresholds for a file. If both file size and I/O size are provided, the client SHOULD reach or exceed both thresholds before issuing its READ or WRITE requests to the data server. Alternatively, if only one of the specified thresholds are reached or exceeded, the I/O requests are sent to the metadata server.
For each threshold type, a value of 0 indicates no READ or WRITE should be sent to the metadata server, while a value of all 1s indicates all READS or WRITES should be sent to the metadata server.
The attribute is available on a per filehandle basis. If the current filehandle refers to a non-pNFS file or directory, the metadata server should return an attribute that is representative of the filehandle's file system. It is suggested that this attribute is queried as part of the OPEN operation. Due to dynamic system changes, the client should not assume that the attribute will remain constant for any specific time period, thus it should be periodically refreshed.
| TOC |
Retention is a concept whereby a file object can be placed in an immutable, undeletable, unrenamable state for a fixed or infinite duration of time. Once in this "retained" state, the file cannot be moved out of the state until the duration of retention has been reached.
When retention is enabled, retention MUST extend to the data of the file, and the name of file. The server MAY extend retention to any other property of the file, including any subset of REQUIRED, RECOMMENDED, and named attributes, with the exceptions noted in this section.
Servers MAY support or not support retention on any file object type.
The five retention attributes are explained in the next subsections.
| TOC |
If retention is enabled for the associated file, this attribute's value represents the retention begin time of the file object. This attribute's value is only readable with the GETATTR operation and MUST NOT be modified by the SETATTR operation (Section 5.5 (Set-Only and Get-Only Attributes)). The value of the attribute consists of:
const RET4_DURATION_INFINITE = 0xffffffffffffffff;
struct retention_get4 {
uint64_t rg_duration;
nfstime4 rg_begin_time<1>;
};
The field rg_duration is the duration in seconds indicating how long the file will be retained once retention is enabled. The field rg_begin_time is an array of up to one absolute time value. If the array is zero length, no beginning retention time has been established, and retention is not enabled. If rg_duration is equal to RET4_DURATION_INFINITE, the file, once retention is enabled, will be retained for an infinite duration.
If (as soon as) rg_duration is zero, then rg_begin_time will be of zero length, and again, retention is not (no longer) enabled.
| TOC |
This attribute is used to set the retention duration and optionally enable retention for the associated file object. This attribute is only modifiable via the SETATTR operation and MUST NOT be retrieved by the GETATTR operation (Section 5.5 (Set-Only and Get-Only Attributes)). This attribute corresponds to retention_get. The value of the attribute consists of:
struct retention_set4 {
bool rs_enable;
uint64_t rs_duration<1>;
};
If the client sets rs_enable to TRUE, then it is enabling retention on the file object with the begin time of retention starting from the server's current time and date. The duration of the retention can also be provided if the rs_duration array is of length one. The duration is the time in seconds from the begin time of retention, and if set to RET4_DURATION_INFINITE, the file is to be retained forever. If retention is enabled, with no duration specified in either this SETATTR or a previous SETATTR, the duration defaults to zero seconds. The server MAY restrict the enabling of retention or the duration of retention on the basis of the ACE4_WRITE_RETENTION ACL permission. The enabling of retention MUST NOT prevent the enabling of event-based retention nor the modification of the retention_hold attribute.
The following rules apply to both the retention_set and retentevt_set attributes.
| TOC |
Get the event-based retention duration, and if enabled, the event-based retention begin time of the file object. This attribute is like retention_get but refers to event-based retention. The event that triggers event-based retention is not defined by the NFSv4.1 specification.
| TOC |
Set the event-based retention duration, and optionally enable event-based retention on the file object. This attribute corresponds to retentevt_get, is like retention_set, but refers to event-based retention. When event based retention is set, the file MUST be retained even if non-event-based retention has been set, and the duration of non-event-based retention has been reached. Conversely, when non-event-based retention has been set, the file MUST be retained even if event-based retention has been set, and the duration of event-based retention has been reached. The server MAY restrict the enabling of event-based retention or the duration of event-based retention on the basis of the ACE4_WRITE_RETENTION ACL permission. The enabling of event-based retention MUST NOT prevent the enabling of non-event-based retention nor the modification of the retention_hold attribute.
| TOC |
Get or set administrative retention holds, one hold per bit position.
This attribute allows one to 64 administrative holds, one hold per bit on the attribute. If retention_hold is not zero, then the file MUST NOT be deleted, renamed, or modified, even if the duration on enabled event or non-event-based retention has been reached. The server MAY restrict the modification of retention_hold on the basis of the ACE4_WRITE_RETENTION_HOLD ACL permission. The enabling of administration retention holds does not prevent the enabling of event-based or non-event-based retention.
If the principal attempting to change retention_hold does not have ACE4_WRITE_RETENTION_HOLD permissions, the attempt MUST fail with NFS4ERR_ACCESS.
| TOC |
Access Control Lists (ACLs) are file attributes that specify fine grained access control. This chapter covers the "acl", "dacl", "sacl", "aclsupport", "mode", "mode_set_masked" file attributes, and their interactions. Note that file attributes may apply to any file system object.
| TOC |
ACLs and modes represent two well established models for specifying permissions. This chapter specifies requirements that attempt to meet the following goals:
| TOC |
| TOC |
The NFSv4.1 ACL attribute contains an array of access control entries (ACEs) that are associated with the file system object. Although the client can read and write the acl attribute, the server is responsible for using the ACL to perform access control. The client can use the OPEN or ACCESS operations to check access without modifying or reading data or metadata.
The NFS ACE structure is defined as follows:
typedef uint32_t acetype4;
typedef uint32_t aceflag4;
typedef uint32_t acemask4;
struct nfsace4 {
acetype4 type;
aceflag4 flag;
acemask4 access_mask;
utf8str_mixed who;
};
To determine if a request succeeds, the server processes each nfsace4 entry in order. Only ACEs which have a "who" that matches the requester are considered. Each ACE is processed until all of the bits of the requester's access have been ALLOWED. Once a bit (see below) has been ALLOWED by an ACCESS_ALLOWED_ACE, it is no longer considered in the processing of later ACEs. If an ACCESS_DENIED_ACE is encountered where the requester's access still has unALLOWED bits in common with the "access_mask" of the ACE, the request is denied. When the ACL is fully processed, if there are bits in the requester's mask that have not been ALLOWED or DENIED, access is denied.
Unlike the ALLOW and DENY ACE types, the ALARM and AUDIT ACE types do not affect a requester's access, and instead are for triggering events as a result of a requester's access attempt. Therefore, AUDIT and ALARM ACEs are processed only after processing ALLOW and DENY ACEs.
The NFSv4.1 ACL model is quite rich. Some server platforms may provide access control functionality that goes beyond the UNIX-style mode attribute, but which is not as rich as the NFS ACL model. So that users can take advantage of this more limited functionality, the server may support the acl attributes by mapping between its ACL model and the NFSv4.1 ACL model. Servers must ensure that the ACL they actually store or enforce is at least as strict as the NFSv4 ACL that was set. It is tempting to accomplish this by rejecting any ACL that falls outside the small set that can be represented accurately. However, such an approach can render ACLs unusable without special client-side knowledge of the server's mapping, which defeats the purpose of having a common NFSv4 ACL protocol. Therefore servers should accept every ACL that they can without compromising security. To help accomplish this, servers may make a special exception, in the case of unsupported permission bits, to the rule that bits not ALLOWED or DENIED by an ACL must be denied. For example, a UNIX-style server might choose to silently allow read attribute permissions even though an ACL does not explicitly allow those permissions. (An ACL that explicitly denies permission to read attributes should still be rejected.)
The situation is complicated by the fact that a server may have multiple modules that enforce ACLs. For example, the enforcement for NFSv4.1 access may be different from, but not weaker than, the enforcement for local access, and both may be different from the enforcement for access through other protocols such as SMB. So it may be useful for a server to accept an ACL even if not all of its modules are able to support it.
The guiding principle with regard to NFSv4 access is that the server must not accept ACLs that appear to make access to the file more restrictive than it really is.
| TOC |
The constants used for the type field (acetype4) are as follows:
const ACE4_ACCESS_ALLOWED_ACE_TYPE = 0x00000000; const ACE4_ACCESS_DENIED_ACE_TYPE = 0x00000001; const ACE4_SYSTEM_AUDIT_ACE_TYPE = 0x00000002; const ACE4_SYSTEM_ALARM_ACE_TYPE = 0x00000003;
Only the ALLOWED and DENIED bits types may be used in the dacl attribute, and only the AUDIT and ALARM bits may be used in the sacl attribute. All four are permitted in the acl attribute.
| Value | Abbreviation | Description |
|---|---|---|
| ACE4_ACCESS_ALLOWED_ACE_TYPE | ALLOW | Explicitly grants the access defined in acemask4 to the file or directory. |
| ACE4_ACCESS_DENIED_ACE_TYPE | DENY | Explicitly denies the access defined in acemask4 to the file or directory. |
| ACE4_SYSTEM_AUDIT_ACE_TYPE | AUDIT | LOG (in a system dependent way) any access attempt to a file or directory which uses any of the access methods specified in acemask4. |
| ACE4_SYSTEM_ALARM_ACE_TYPE | ALARM | Generate a system ALARM (system dependent) when any access attempt is made to a file or directory for the access methods specified in acemask4. |
The "Abbreviation" column denotes how the types will be referred to throughout the rest of this chapter.
| TOC |
A server need not support all of the above ACE types. This attribute indicates which ACE types are supported for the current file system. The bitmask constants used to represent the above definitions within the aclsupport attribute are as follows:
const ACL4_SUPPORT_ALLOW_ACL = 0x00000001; const ACL4_SUPPORT_DENY_ACL = 0x00000002; const ACL4_SUPPORT_AUDIT_ACL = 0x00000004; const ACL4_SUPPORT_ALARM_ACL = 0x00000008;
Servers which support either the ALLOW or DENY ACE type SHOULD support both ALLOW and DENY ACE types.
Clients should not attempt to set an ACE unless the server claims support for that ACE type. If the server receives a request to set an ACE that it cannot store, it MUST reject the request with NFS4ERR_ATTRNOTSUPP. If the server receives a request to set an ACE that it can store but cannot enforce, the server SHOULD reject the request with NFS4ERR_ATTRNOTSUPP.
Support for any of the ACL attributes is optional (albeit, RECOMMENDED). However, a server that supports either of the new ACL attributes (dacl or sacl) MUST allow use of the new ACL attributes to access all of the ACE types which it supports. In other words, if such a server supports ALLOW or DENY ACEs, then it MUST support the dacl attribute, and if it supports AUDIT or ALARM ACEs, then it MUST support the sacl attribute.
| TOC |
The bitmask constants used for the access mask field are as follows:
const ACE4_READ_DATA = 0x00000001; const ACE4_LIST_DIRECTORY = 0x00000001; const ACE4_WRITE_DATA = 0x00000002; const ACE4_ADD_FILE = 0x00000002; const ACE4_APPEND_DATA = 0x00000004; const ACE4_ADD_SUBDIRECTORY = 0x00000004; const ACE4_READ_NAMED_ATTRS = 0x00000008; const ACE4_WRITE_NAMED_ATTRS = 0x00000010; const ACE4_EXECUTE = 0x00000020; const ACE4_DELETE_CHILD = 0x00000040; const ACE4_READ_ATTRIBUTES = 0x00000080; const ACE4_WRITE_ATTRIBUTES = 0x00000100; const ACE4_WRITE_RETENTION = 0x00000200; const ACE4_WRITE_RETENTION_HOLD = 0x00000400; const ACE4_DELETE = 0x00010000; const ACE4_READ_ACL = 0x00020000; const ACE4_WRITE_ACL = 0x00040000; const ACE4_WRITE_OWNER = 0x00080000; const ACE4_SYNCHRONIZE = 0x00100000;
Note that some masks have coincident values, for example, ACE4_READ_DATA and ACE4_LIST_DIRECTORY. The mask entries ACE4_LIST_DIRECTORY, ACE4_ADD_FILE, and ACE4_ADD_SUBDIRECTORY are intended to be used with directory objects, while ACE4_READ_DATA, ACE4_WRITE_DATA, and ACE4_APPEND_DATA are intended to be used with non-directory objects.
| TOC |
- ACE4_READ_DATA
- Operation(s) affected:
- READ
- OPEN
- Discussion:
Permission to read the data of the file.
Servers SHOULD allow a user the ability to read the data of the file when only the ACE4_EXECUTE access mask bit is allowed.- ACE4_LIST_DIRECTORY
- Operation(s) affected:
- READDIR
- Discussion:
Permission to list the contents of a directory.- ACE4_WRITE_DATA
- Operation(s) affected:
- WRITE
- OPEN
- SETATTR of size
- Discussion:
Permission to modify a file's data.- ACE4_ADD_FILE
- Operation(s) affected:
- CREATE
- LINK
- OPEN
- RENAME
- Discussion:
Permission to add a new file in a directory. The CREATE operation is affected when nfs_ftype4 is NF4LNK, NF4BLK, NF4CHR, NF4SOCK, or NF4FIFO. (NF4DIR is not listed because it is covered by ACE4_ADD_SUBDIRECTORY.) OPEN is affected when used to create a regular file. LINK and RENAME are always affected.- ACE4_APPEND_DATA
- Operation(s) affected:
- WRITE
- OPEN
- SETATTR of size
- Discussion:
The ability to modify a file's data, but only starting at EOF. This allows for the notion of append-only files, by allowing ACE4_APPEND_DATA and denying ACE4_WRITE_DATA to the same user or group. If a file has an ACL such as the one described above and a WRITE request is made for somewhere other than EOF, the server SHOULD return NFS4ERR_ACCESS.- ACE4_ADD_SUBDIRECTORY
- Operation(s) affected:
- CREATE
- RENAME
- Discussion:
Permission to create a subdirectory in a directory. The CREATE operation is affected when nfs_ftype4 is NF4DIR. The RENAME operation is always affected.- ACE4_READ_NAMED_ATTRS
- Operation(s) affected:
- OPENATTR
- Discussion:
Permission to read the named attributes of a file or to lookup the named attributes directory. OPENATTR is affected when it is not used to create a named attribute directory. This is when 1.) createdir is TRUE, but a named attribute directory already exists, or 2.) createdir is FALSE.- ACE4_WRITE_NAMED_ATTRS
- Operation(s) affected:
- OPENATTR
- Discussion:
Permission to write the named attributes of a file or to create a named attribute directory. OPENATTR is affected when it is used to create a named attribute directory. This is when createdir is TRUE and no named attribute directory exists. The ability to check whether or not a named attribute directory exists depends on the ability to look it up, therefore, users also need the ACE4_READ_NAMED_ATTRS permission in order to create a named attribute directory.- ACE4_EXECUTE
- Operation(s) affected:
- READ
- OPEN
- REMOVE
- RENAME
- LINK
- CREATE
- Discussion:
Permission to execute a file.
Servers SHOULD allow a user the ability to read the data of the file when only the ACE4_EXECUTE access mask bit is allowed. This is because there is no way to execute a file without reading the contents. Though a server may treat ACE4_EXECUTE and ACE4_READ_DATA bits identically when deciding to permit a READ operation, it SHOULD still allow the two bits to be set independently in ACLs, and MUST distinguish between them when replying to ACCESS operations. In particular, servers SHOULD NOT silently turn on one of the two bits when the other is set, as that would make it impossible for the client to correctly enforce the distinction between read and execute permissions.
As an example, following a SETATTR of the following ACL:
nfsuser:ACE4_EXECUTE:ALLOW
A subsequent GETATTR of ACL for that file SHOULD return:
nfsuser:ACE4_EXECUTE:ALLOW
Rather than:
nfsuser:ACE4_EXECUTE/ACE4_READ_DATA:ALLOW- ACE4_EXECUTE
- Operation(s) affected:
- LOOKUP
- Discussion:
Permission to traverse/search a directory.- ACE4_DELETE_CHILD
- Operation(s) affected:
- REMOVE
- RENAME
- Discussion:
Permission to delete a file or directory within a directory. See Section 6.2.1.3.2 (ACE4_DELETE vs. ACE4_DELETE_CHILD) for information on ACE4_DELETE and ACE4_DELETE_CHILD interact.- ACE4_READ_ATTRIBUTES
- Operation(s) affected:
- GETATTR of file system object attributes
- VERIFY
- NVERIFY
- READDIR
- Discussion:
The ability to read basic attributes (non-ACLs) of a file. On a UNIX system, basic attributes can be thought of as the stat level attributes. Allowing this access mask bit would mean the entity can execute "ls -l" and stat. If a READDIR operation requests attributes, this mask must be allowed for the READDIR to succeed.- ACE4_WRITE_ATTRIBUTES
- Operation(s) affected:
- SETATTR of time_access_set, time_backup,
- time_create, time_modify_set, mimetype, hidden, system
- Discussion:
Permission to change the times associated with a file or directory to an arbitrary value. Also permission to change the mimetype, hidden and system attributes. A user having ACE4_WRITE_DATA or ACE4_WRITE_ATTRIBUTES will be allowed to set the times associated with a file to the current server time.- ACE4_WRITE_RETENTION
- Operation(s) affected:
- SETATTR of retention_set, retentevt_set.
- Discussion:
Permission to modify the durations of event and non-event-based retention. Also permission to enable event and non-event-based retention. A server MAY behave such that setting ACE4_WRITE_ATTRIBUTES allows ACE4_WRITE_RETENTION.- ACE4_WRITE_RETENTION_HOLD
- Operation(s) affected:
- SETATTR of retention_hold.
- Discussion:
Permission to modify the administration retention holds. A server MAY map ACE4_WRITE_ATTRIBUTES to ACE_WRITE_RETENTION_HOLD.- ACE4_DELETE
- Operation(s) affected:
- REMOVE
- Discussion:
Permission to delete the file or directory. See Section 6.2.1.3.2 (ACE4_DELETE vs. ACE4_DELETE_CHILD) for information on ACE4_DELETE and ACE4_DELETE_CHILD interact.- ACE4_READ_ACL
- Operation(s) affected:
- GETATTR of acl, dacl, or sacl
- NVERIFY
- VERIFY
- Discussion:
Permission to read the ACL.- ACE4_WRITE_ACL
- Operation(s) affected:
- SETATTR of acl and mode
- Discussion:
Permission to write the acl and mode attributes.- ACE4_WRITE_OWNER
- Operation(s) affected:
- SETATTR of owner and owner_group
- Discussion:
Permission to write the owner and owner_group attributes. On UNIX systems, this is the ability to execute chown() and chgrp().- ACE4_SYNCHRONIZE
- Operation(s) affected:
- NONE
- Discussion:
Permission to access file locally at the server with synchronized reads and writes.
Server implementations need not provide the granularity of control that is implied by this list of masks. For example, POSIX-based systems might not distinguish ACE4_APPEND_DATA (the ability to append to a file) from ACE4_WRITE_DATA (the ability to modify existing contents); both masks would be tied to a single "write" permission. When such a server returns attributes to the client, it would show both ACE4_APPEND_DATA and ACE4_WRITE_DATA if and only if the write permission is enabled.
If a server receives a SETATTR request that it cannot accurately implement, it should err in the direction of more restricted access, except in the previously discussed cases of execute and read. For example, suppose a server cannot distinguish overwriting data from appending new data, as described in the previous paragraph. If a client submits an ALLOW ACE where ACE4_APPEND_DATA is set but ACE4_WRITE_DATA is not (or vice versa), the server should either turn off ACE4_APPEND_DATA or reject the request with NFS4ERR_ATTRNOTSUPP.
| TOC |
Two access mask bits govern the ability to delete a directory entry: ACE4_DELETE on the object itself (the "target"), and ACE4_DELETE_CHILD on the containing directory (the "parent").
Many systems also take the "sticky bit" (MODE4_SVTX) on a directory to allow unlink only to a user that owns either the target or the parent; on some such systems the decision also depends on whether the target is writable.
Servers SHOULD allow unlink if either ACE4_DELETE is permitted on the target, or ACE4_DELETE_CHILD is permitted on the parent. (Note that this is true even if the parent or target explicitly denies one of these permissions.)
If the ACLs in question neither explicitly ALLOW nor DENY either of the above, and if MODE4_SVTX is not set on the parent, then the server SHOULD allow the removal if and only if ACE4_ADD_FILE is permitted. In the case where MODE4_SVTX is set, the server may also require the remover to own either the parent or the target, or may require the target to be writable.
This allows servers to support something close to traditional unix-like semantics, with ACE4_ADD_FILE taking the place of the write bit.
| TOC |
The bitmask constants used for the flag field are as follows:
const ACE4_FILE_INHERIT_ACE = 0x00000001; const ACE4_DIRECTORY_INHERIT_ACE = 0x00000002; const ACE4_NO_PROPAGATE_INHERIT_ACE = 0x00000004; const ACE4_INHERIT_ONLY_ACE = 0x00000008; const ACE4_SUCCESSFUL_ACCESS_ACE_FLAG = 0x00000010; const ACE4_FAILED_ACCESS_ACE_FLAG = 0x00000020; const ACE4_IDENTIFIER_GROUP = 0x00000040; const ACE4_INHERITED_ACE = 0x00000080;
A server need not support any of these flags. If the server supports flags that are similar to, but not exactly the same as, these flags, the implementation may define a mapping between the protocol-defined flags and the implementation-defined flags.
For example, suppose a client tries to set an ACE with ACE4_FILE_INHERIT_ACE set but not ACE4_DIRECTORY_INHERIT_ACE. If the server does not support any form of ACL inheritance, the server should reject the request with NFS4ERR_ATTRNOTSUPP. If the server supports a single "inherit ACE" flag that applies to both files and directories, the server may reject the request (i.e., requiring the client to set both the file and directory inheritance flags). The server may also accept the request and silently turn on the ACE4_DIRECTORY_INHERIT_ACE flag.
| TOC |
- ACE4_FILE_INHERIT_ACE
- Any non-directory file in any sub-directory will get this ACE inherited.
- ACE4_DIRECTORY_INHERIT_ACE
- Can be placed on a directory and indicates that this ACE should be added to each new directory created.
If this flag is set in an ACE in an ACL attribute to be set on a non-directory file system object, the operation attempting to set the ACL SHOULD fail with NFS4ERR_ATTRNOTSUPP.- ACE4_INHERIT_ONLY_ACE
- Can be placed on a directory but does not apply to the directory; ALLOW and DENY ACEs with this bit set do not affect access to the directory, and AUDIT and ALARM ACEs with this bit set do not trigger log or alarm events. Such ACEs only take effect once they are applied (with this bit cleared) to newly created files and directories as specified by the above two flags.
If this flag is present on an ACE, but neither ACE4_DIRECTORY_INHERIT_ACE nor ACE4_FILE_INHERIT_ACE is present, then an operation attempting to set such an attribute SHOULD fail with NFS4ERR_ATTRNOTSUPP.- ACE4_NO_PROPAGATE_INHERIT_ACE
- Can be placed on a directory. This flag tells the server that inheritance of this ACE should stop at newly created child directories.
- ACE4_INHERITED_ACE
- Indicates that this ACE is inherited from a parent directory. A server that supports automatic inheritance will place this flag on any ACEs inherited from the parent directory when creating a new object. Client applications will use this to perform automatic inheritance. Clients and servers MUST clear this bit in the acl attribute; it may only be used in the dacl and sacl attributes.
- ACE4_SUCCESSFUL_ACCESS_ACE_FLAG
- ACE4_FAILED_ACCESS_ACE_FLAG
- The ACE4_SUCCESSFUL_ACCESS_ACE_FLAG (SUCCESS) and ACE4_FAILED_ACCESS_ACE_FLAG (FAILED) flag bits may be set only on ACE4_SYSTEM_AUDIT_ACE_TYPE (AUDIT) and ACE4_SYSTEM_ALARM_ACE_TYPE (ALARM) ACE types. If during the processing of the file's ACL, the server encounters an AUDIT or ALARM ACE that matches the principal attempting the OPEN, the server notes that fact, and the presence, if any, of the SUCCESS and FAILED flags encountered in the AUDIT or ALARM ACE. Once the server completes the ACL processing, it then notes if the operation succeeded or failed. If the operation succeeded, and if the SUCCESS flag was set for a matching AUDIT or ALARM ACE, then the appropriate AUDIT or ALARM event occurs. If the operation failed, and if the FAILED flag was set for the matching AUDIT or ALARM ACE, then the appropriate AUDIT or ALARM event occurs. Either or both of the SUCCESS or FAILED can be set, but if neither is set, the AUDIT or ALARM ACE is not useful.
- The previously described processing applies to ACCESS operations even when they return NFS4_OK. For the purposes of AUDIT and ALARM, we consider an ACCESS operation to be a "failure" if it fails to return a bit that was requested and supported.
- ACE4_IDENTIFIER_GROUP
- Indicates that the "who" refers to a GROUP as defined under UNIX or a GROUP ACCOUNT as defined under Windows. Clients and servers MUST ignore the ACE4_IDENTIFIER_GROUP flag on ACEs with a who value equal to one of the special identifiers outlined in Section 6.2.1.5 (ACE Who).
| TOC |
The "who" field of an ACE is an identifier that specifies the principal or principals to whom the ACE applies. It may refer to a user or a group, with the flag bit ACE4_IDENTIFIER_GROUP specifying which.
There are several special identifiers which need to be understood universally, rather than in the context of a particular DNS domain. Some of these identifiers cannot be understood when an NFS client accesses the server, but have meaning when a local process accesses the file. The ability to display and modify these permissions is permitted over NFS, even if none of the access methods on the server understands the identifiers.
| Who | Description |
|---|---|
| OWNER | The owner of the file |
| GROUP | The group associated with the file. |
| EVERYONE | The world, including the owner and owning group. |
| INTERACTIVE | Accessed from an interactive terminal. |
| NETWORK | Accessed via the network. |
| DIALUP | Accessed as a dialup user to the server. |
| BATCH | Accessed from a batch job. |
| ANONYMOUS | Accessed without any authentication. |
| AUTHENTICATED | Any authenticated user (opposite of ANONYMOUS) |
| SERVICE | Access from a system service. |
| Table 7 |
To avoid conflict, these special identifiers are distinguished by an appended "@" and should appear in the form "xxxx@" (with no domain name after the "@"). For example: ANONYMOUS@.
The ACE4_IDENTIFIER_GROUP flag MUST be ignored on entries with these special identifiers. When encoding entries with these special identifiers, the ACE4_IDENTIFIER_GROUP flag SHOULD be set to zero.
| TOC |
It is important to note that "EVERYONE@" is not equivalent to the UNIX "other" entity. This is because, by definition, UNIX "other" does not include the owner or owning group of a file. "EVERYONE@" means literally everyone, including the owner or owning group.
| TOC |
The dacl attribute is like the acl attribute, but dacl allows just ALLOW and DENY ACEs. The dacl attribute supports automatic inheritance (see Section 6.4.3.2 (Automatic Inheritance)).
| TOC |
The sacl attribute is like the acl attribute, but sacl allows just AUDIT and ALARM ACEs. The sacl attribute supports automatic inheritance (see Section 6.4.3.2 (Automatic Inheritance)).
| TOC |
The NFSv4.1 mode attribute is based on the UNIX mode bits. The following bits are defined:
const MODE4_SUID = 0x800; /* set user id on execution */ const MODE4_SGID = 0x400; /* set group id on execution */ const MODE4_SVTX = 0x200; /* save text even after use */ const MODE4_RUSR = 0x100; /* read permission: owner */ const MODE4_WUSR = 0x080; /* write permission: owner */ const MODE4_XUSR = 0x040; /* execute permission: owner */ const MODE4_RGRP = 0x020; /* read permission: group */ const MODE4_WGRP = 0x010; /* write permission: group */ const MODE4_XGRP = 0x008; /* execute permission: group */ const MODE4_ROTH = 0x004; /* read permission: other */ const MODE4_WOTH = 0x002; /* write permission: other */ const MODE4_XOTH = 0x001; /* execute permission: other */
Bits MODE4_RUSR, MODE4_WUSR, and MODE4_XUSR apply to the principal identified in the owner attribute. Bits MODE4_RGRP, MODE4_WGRP, and MODE4_XGRP apply to principals identified in the owner_group attribute but who are not identified in the owner attribute. Bits MODE4_ROTH, MODE4_WOTH, MODE4_XOTH apply to any principal that does not match that in the owner attribute, and does not have a group matching that of the owner_group attribute.
Bits within the mode other than those specified above are not defined by this protocol. A server MUST NOT return bits other than those defined above in a GETATTR or READDIR operation, and it MUST return NFS4ERR_INVAL if bits other than those defined above are set in a SETATTR, CREATE, OPEN, VERIFY or NVERIFY operation.
| TOC |
The mode_set_masked attribute is a write-only attribute that allows individual bits in the mode attribute to be set or reset, without changing others. It allows, for example, the bits MODE4_SUID, MODE4_SGID, and MODE4_SVTX to be modified while leaving unmodified any of the nine low-order mode bits devoted to permissions.
In such instances that the nine low-order bits are left unmodified, then neither the acl nor the dacl attribute should be automatically modified as discussed in Section 6.4.1 (Setting the mode and/or ACL Attributes).
The mode_set_masked attribute consists of two words each in the form of a mode4. The first consists of the value to be applied to the current mode value and the second is a mask. Only bits set to one in the mask word are changed (set or reset) in the file's mode. All other bits in the mode remain unchanged. Bits in the first word that correspond to bits which are zero in the mask are ignored, except that undefined bits are checked for validity and can result in NFS4ERR_INVAL as described below.
The mode_set_masked attribute is only valid in a SETATTR operation. If it is used in a CREATE or OPEN operation, the server MUST return NFS4ERR_INVAL.
Bits not defined as valid in the mode attribute are not valid in either word of the mode_set_masked attribute. The server MUST return NFS4ERR_INVAL if any of those are on in a SETATTR. If the mode and mode_set_masked attributes are both specified in the same SETATTR, the server MUST also return NFS4ERR_INVAL.
| TOC |
The requirements in this section will be referred to in future sections, especially Section 6.4 (Requirements).
| TOC |
| TOC |
The server uses the algorithm described in Section 6.2.1 (Attribute 12: acl) to determine whether an ACL allows access to an object. However, the ACL may not be the sole determiner of access. For example:
| TOC |
Clients SHOULD NOT do their own access checks based on their interpretation the ACL, but rather use the OPEN and ACCESS operations to do access checks. This allows the client to act on the results of having the server determine whether or not access should be granted based on its interpretation of the ACL.
Clients must be aware of situations in which an object's ACL will define a certain access even though the server will not enforce it. In general, but especially in these situations, the client needs to do its part in the enforcement of access as defined by the ACL. To do this, the client MAY send the appropriate ACCESS operation prior to servicing the request of the user or application in order to determine whether the user or application should be granted the access requested. For examples in which the ACL may define accesses that the server doesn't enforce see Section 6.3.1.1 (Server Considerations).
| TOC |
The following method can be used to calculate the MODE4_R*, MODE4_W* and MODE4_X* bits of a mode attribute, based upon an ACL.
First, for each of the special identifiers OWNER@, GROUP@, and EVERYONE@, evaluate the ACL in order, considering only ALLOW and DENY ACEs for the identifier EVERYONE@ and for the identifier under consideration. The result of the evaluation will be an NFSv4 ACL mask showing exactly which bits are permitted to that identifier.
Then translate the calculated mask for OWNER@, GROUP@, and EVERYONE@ into mode bits for, respectively, the user, group, and other, as follows:
| TOC |
Some server implementations also add bits permitted to named users and groups to the group bits (MODE4_RGRP, MODE4_WGRP, and MODE4_XGRP).
Implementations are discouraged from doing this, because it has been found to cause confusion for users who see members of a file's group denied access that the mode bits appear to allow. (The presence of DENY ACEs may also lead to such behavior, but DENY ACEs are expected to be more rarely used.)
The same user confusion seen when fetching the mode also results if setting the mode does not effectively control permissions for the owner, group, and other users; this motivates some of the requirements that follow.
| TOC |
The server that supports both mode and ACL must take care to synchronize the MODE4_*USR, MODE4_*GRP, and MODE4_*OTH bits with the ACEs which have respective who fields of "OWNER@", "GROUP@", and "EVERYONE@" so that the client can see semantically equivalent access permissions exist whether the client asks for owner, owner_group and mode attributes, or for just the ACL.
In this section, much is made of the methods in Section 6.3.2 (Computing a Mode Attribute from an ACL). Many requirements refer to this section. But note that the methods have behaviors specified with "SHOULD". This is intentional, to avoid invalidating existing implementations that compute the mode according to the withdrawn POSIX ACL draft (1003.1e draft 17), rather than by actual permissions on owner, group, and other.
| TOC |
In the case where a server supports the sacl or dacl attribute, in addition to the acl attribute, the server MUST fail a request to set the acl attribute simultaneously with a dacl or sacl attribute. The error to be given is NFS4ERR_ATTRNOTSUP.
| TOC |
When any of the nine low-order mode bits are subject to change, either because the mode attribute was set or because the mode_set_masked attribute was set and the mask included one or more bits from the nine low-order mode bits, and no ACL attribute is explicitly set, the acl and dacl attributes must be modified in accordance with the updated value of those bits. This must happen even if the value of the low-order bits is the same after the mode is set as before.
Note that any AUDIT or ALARM ACEs (hence any ACEs in the sacl attribute) are unaffected by changes to the mode.
In cases in which the permissions bits are subject to change, the acl and dacl attributes MUST be modified such that the mode computed via the method in Section 6.3.2 (Computing a Mode Attribute from an ACL) yields the low-order nine bits (MODE4_R*, MODE4_W*, MODE4_X*) of the mode attribute as modified by the attribute change. The ACL attributes SHOULD also be modified such that:
Access mask bits other those listed above, appearing in ALLOW ACEs, MAY also be disabled.
Note that ACEs with the flag ACE4_INHERIT_ONLY_ACE set do not affect the permissions of the ACL itself, nor do ACEs of the type AUDIT and ALARM. As such, it is desirable to leave these ACEs unmodified when modifying the ACL attributes.
Also note that the requirement may be met by discarding the acl and dacl, in favor of an ACL that represents the mode and only the mode. This is permitted, but it is preferable for a server to preserve as much of the ACL as possible without violating the above requirements. Discarding the ACL makes it effectively impossible for a file created with a mode attribute to inherit an ACL (see Section 6.4.3 (Creating New Objects)).
| TOC |
When setting the acl or dacl and not setting the mode or mode_set_masked attributes, the permission bits of the mode need to be derived from the ACL. In this case, the ACL attribute SHOULD be set as given. The nine low-order bits of the mode attribute (MODE4_R*, MODE4_W*, MODE4_X*) MUST be modified to match the result of the method Section 6.3.2 (Computing a Mode Attribute from an ACL). The three high-order bits of the mode (MODE4_SUID, MODE4_SGID, MODE4_SVTX) SHOULD remain unchanged.
| TOC |
When setting both the mode (includes use of either the mode attribute or the mode_set_masked attribute) and the acl or dacl attributes in the same operation, the attributes MUST be applied in this order: mode (or mode_set_masked), then ACL. The mode-related attribute is set as given, then the ACL attribute is set as given, possibly changing the final mode, as described above in Section 6.4.1.2 (Setting ACL and not mode).
| TOC |
This section applies only to servers that support both the mode and ACL attributes.
Some server implementations may have a concept of "objects without ACLs", meaning that all permissions are granted and denied according to the mode attribute, and that no ACL attribute is stored for that object. If an ACL attribute is requested of such a server, the server SHOULD return an ACL that does not conflict with the mode; that is to say, the ACL returned SHOULD represent the nine low-order bits of the mode attribute (MODE4_R*, MODE4_W*, MODE4_X*) as described in Section 6.3.2 (Computing a Mode Attribute from an ACL).
For other server implementations, the ACL attribute is always present for every object. Such servers SHOULD store at least the three high-order bits of the mode attribute (MODE4_SUID, MODE4_SGID, MODE4_SVTX). The server SHOULD return a mode attribute if one is requested, and the low-order nine bits of the mode (MODE4_R*, MODE4_W*, MODE4_X*) MUST match the result of applying the method in Section 6.3.2 (Computing a Mode Attribute from an ACL) to the ACL attribute.
| TOC |
If a server supports any ACL attributes, it may use the ACL attributes on the parent directory to compute an initial ACL attribute for a newly created object. This will be referred to as the inherited ACL within this section. The act of adding one or more ACEs to the inherited ACL that are based upon ACEs in the parent directory's ACL will be referred to as inheriting an ACE within this section.
Implementors should standardize on what the behavior of CREATE and OPEN must be depending on the presence or absence of the mode and ACL attributes.
| TOC |
If the object being created is not a directory, the inherited ACL SHOULD NOT inherit ACEs from the parent directory ACL unless the ACE4_FILE_INHERIT_FLAG is set.
If the object being created is a directory, the inherited ACL should inherit all inheritable ACEs from the parent directory, those that have ACE4_FILE_INHERIT_ACE or ACE4_DIRECTORY_INHERIT_ACE flag set. If the inheritable ACE has ACE4_FILE_INHERIT_ACE set, but ACE4_DIRECTORY_INHERIT_ACE is clear, the inherited ACE on the newly created directory MUST have the ACE4_INHERIT_ONLY_ACE flag set to prevent the directory from being affected by ACEs meant for non-directories.
When a new directory is created, the server MAY split any inherited ACE which is both inheritable and effective (in other words, which has neither ACE4_INHERIT_ONLY_ACE nor ACE4_NO_PROPAGATE_INHERIT_ACE set), into two ACEs, one with no inheritance flags, and one with ACE4_INHERIT_ONLY_ACE set. (In the case of a dacl or sacl attribute, both of those ACEs SHOULD also have the ACE4_INHERITED_ACE flag set.) This makes it simpler to modify the effective permissions on the directory without modifying the ACE which is to be inherited to the new directory's children.
| TOC |
The acl attribute consists only of an array of ACEs, but the sacl (Attribute 59: sacl) and dacl (Attribute 58: dacl) attributes also include an additional flag field.
struct nfsacl41 {
aclflag4 na41_flag;
nfsace4 na41_aces<>;
};
The flag field applies to the entire sacl or dacl; three flag values are defined:
const ACL4_AUTO_INHERIT = 0x00000001; const ACL4_PROTECTED = 0x00000002; const ACL4_DEFAULTED = 0x00000004;
and all other bits must be cleared. The ACE4_INHERITED_ACE flag may be set in the ACEs of the sacl or dacl (whereas it must always be cleared in the acl).
Together these features allow a server to support automatic inheritance, which we now explain in more detail.
Inheritable ACEs are normally inherited by child objects only at the time that the child objects are created; later modifications to inheritable ACEs do not result in modifications to inherited ACEs on descendents.
However, the dacl and sacl provide an OPTIONAL mechanism which allows a client application to propagate changes to inheritable ACEs to an entire directory hierarchy.
A server that supports this performs inheritance at object creation time in the normal way, and SHOULD set the ACE4_INHERITED_ACE flag on any inherited ACEs as they are added to the new object.
A client application such as an ACL editor may then propagate changes to inheritable ACEs on a directory by recursively traversing that directory's descendants and modifying each ACL encountered to remove any ACEs with the ACE4_INHERITED_ACE flag and to replace them by the new inheritable ACEs (also with the ACE4_INHERITED_ACE flag set). It uses the existing ACE inheritance flags in the obvious way to decide which ACEs to propagate. (Note that it may encounter further inheritable ACEs when descending the directory hierarchy, and that those will also need to be taken into account when propagating inheritable ACEs to further descendants.)
The reach of this propagation may be limited in two ways: first, automatic inheritance is not performed from any directory ACL that has the ACL4_AUTO_INHERIT flag cleared; and second, automatic inheritance stops wherever an ACL with the ACL4_PROTECTED flag is set, preventing modification of that ACL and also (if the ACL is set on a directory) of the ACL on any of the object's descendants.
This propagation is performed independently for the sacl and the dacl attributes; thus the ACL4_AUTO_INHERIT and ACL4_PROTECTED flags may be independently set for the sacl and the dacl, and propagation of one type of acl may continue down a hierarchy even where propagation of the other acl has stopped.
New objects should be created with a dacl and a sacl that both have the ACL4_PROTECTED flag cleared and the ACL4_AUTO_INHERIT flag set to the same value as that on, respectively, the sacl or dacl of the parent object.
Both the dacl and sacl attributes are RECOMMENDED, and a server may support one without supporting the other.
A server that supports both the old acl attribute and one or both of the new dacl or sacl attributes must do so in such a way as to keep all three attributes consistent with each other. Thus the ACEs reported in the acl attribute should be the union of the ACEs reported in the dacl and sacl attributes, except that the ACE4_INHERITED_ACE flag must be cleared from the ACEs in the acl. And of course a client that queries only the acl will be unable to determine the values of the sacl or dacl flag fields.
When a client performs a SETATTR for the acl attribute, the server SHOULD set the ACL4_PROTECTED flag to true on both the sacl and the dacl. By using the acl attribute, as opposed to the dacl or sacl attributes, the client signals that it may not understand automatic inheritance, and thus cannot be trusted to set an ACL for which automatic inheritance would make sense.
When a client application queries an ACL, modifies it, and sets it again, it should leave any ACEs marked with ACE4_INHERITED_ACE unchanged, in their original order, at the end of the ACL. If the application is unable to do this, it should set the ACL4_PROTECTED flag. This behavior is not enforced by servers, but violations of this rule may lead to unexpected results when applications perform automatic inheritance.
If a server also supports the mode attribute, it SHOULD set the mode in such a way that leaves inherited ACEs unchanged, in their original order, at the end of the ACL. If it is unable to do so, it SHOULD set the ACL4_PROTECTED flag on the file's dacl.
Finally, in the case where the request that creates a new file or directory does not also set permissions for that file or directory, and there are also no ACEs to inherit from the parent's directory, then the server's choice of ACL for the new object is implementation-dependent. In this case, the server SHOULD set the ACL4_DEFAULTED flag on the ACL it chooses for the new object. An application performing automatic inheritance takes the ACL4_DEFAULTED flag as a sign that the ACL should be completely replaced by one generated using the automatic inheritance rules.
| TOC |
This chapter describes the NFSv4 single-server namespace. Single-server namespaces may be presented directly to clients, or they may be used as a basis to form larger multi-server namespaces (e.g. site-wide or organization-wide) to be presented to clients, as described in Section 11 (Multi-Server Namespace).
| TOC |
On a UNIX server, the namespace describes all the files reachable by pathnames under the root directory or "/". On a Windows server the namespace constitutes all the files on disks named by mapped disk letters. NFS server administrators rarely make the entire server's file system namespace available to NFS clients. More often portions of the namespace are made available via an "export" feature. In previous versions of the NFS protocol, the root filehandle for each export is obtained through the MOUNT protocol; the client sent a string that identified the export name within the namespace and the server returned the root filehandle for that export. The MOUNT protocol also provided an EXPORTS procedure that enumerated server's exports.
| TOC |
The NFSv4.1 protocol provides a root filehandle that clients can use to obtain filehandles for the exports of a particular server, via a series of LOOKUP operations within a COMPOUND, to traverse a path. A common user experience is to use a graphical user interface (perhaps a file "Open" dialog window) to find a file via progressive browsing through a directory tree. The client must be able to move from one export to another export via single-component, progressive LOOKUP operations.
This style of browsing is not well supported by the NFSv3 protocol. In NFSv3, the client expects all LOOKUP operations to remain within a single server file system. For example, the device attribute will not change. This prevents a client from taking namespace paths that span exports.
In the case of NFSv3, an automounter on the client can obtain a snapshot of the server's namespace using the EXPORTS procedure of the MOUNT protocol. If it understands the server's pathname syntax, it can create an image of the server's namespace on the client. The parts of the namespace that are not exported by the server are filled in with directories that might be constructed similarly to a NFSv4.1 "pseudo file system" (see Section 7.3 (Server Pseudo File System)) that allows the user to browse from one mounted file system to another. There is a drawback to this representation of the server's namespace on the client: it is static. If the server administrator adds a new export the client will be unaware of it.
| TOC |
NFSv4.1 servers avoid this namespace inconsistency by presenting all the exports for a given server within the framework of a single namespace, for that server. An NFSv4.1 client uses LOOKUP and READDIR operations to browse seamlessly from one export to another.
Where there are portions of the server namespace that are not exported, clients require some way of traversing those portions to reach actual exported file systems. A technique that servers may use to provide for this is to bridge unexported portion of the namespace via a "pseudo file system" that provides a view of exported directories only. A pseudo file system has a unique fsid and behaves like a normal, read-only file system.
Based on the construction of the server's namespace, it is possible that multiple pseudo file systems may exist. For example,
/a pseudo file system
/a/b real file system
/a/b/c pseudo file system
/a/b/c/d real file system
Each of the pseudo file systems is considered a separate entity and therefore MUST have its own fsid, unique among all the fsids for that server.
| TOC |
Certain operating environments are sometimes described as having "multiple roots". In such environments individual file systems are commonly represented by disk or volume names. NFSv4 servers for these platforms can construct a pseudo file system above these root names so that disk letters or volume names are simply directory names in the pseudo root.
| TOC |
The nature of the server's pseudo file system is that it is a logical representation of file system(s) available from the server. Therefore, the pseudo file system is most likely constructed dynamically when the server is first instantiated. It is expected that the pseudo file system may not have an on disk counterpart from which persistent filehandles could be constructed. Even though it is preferable that the server provide persistent filehandles for the pseudo file system, the NFS client should expect that pseudo file system filehandles are volatile. This can be confirmed by checking the associated "fh_expire_type" attribute for those filehandles in question. If the filehandles are volatile, the NFS client must be prepared to recover a filehandle value (e.g. with a series of LOOKUP operations) when receiving an error of NFS4ERR_FHEXPIRED.
Because it is quite likely that servers will implement pseudo file systems using volatile filehandles, clients need to be prepared for them, rather than assuming that all filehandles will be persistent.
| TOC |
If the server's root file system is exported, one might conclude that a pseudo file system is unneeded. This not necessarily so. Assume the following file systems on a server:
/ fs1 (exported)
/a fs2 (not exported)
/a/b fs3 (exported)
Because fs2 is not exported, fs3 cannot be reached with simple LOOKUPs. The server must bridge the gap with a pseudo file system.
| TOC |
The server file system environment may be constructed in such a way that one file system contains a directory which is 'covered' or mounted upon by a second file system. For example:
/a/b (file system 1)
/a/b/c/d (file system 2)
The pseudo file system for this server may be constructed to look like:
/ (place holder/not exported)
/a/b (file system 1)
/a/b/c/d (file system 2)
It is the server's responsibility to present the pseudo file system that is complete to the client. If the client sends a lookup request for the path "/a/b/c/d", the server's response is the filehandle of the root of the file system "/a/b/c/d". In previous versions of the NFS protocol, the server would respond with the filehandle of directory "/a/b/c/d" within the file system "/a/b".
The NFS client will be able to determine if it crosses a server mount point by a change in the value of the "fsid" attribute.
| TOC |
Because NFSv4 clients possess the ability to change the security mechanisms used, after determining what is allowed, by using SECINFO and SECINFO_NONAME, the server SHOULD NOT present a different view of the namespace based on the security mechanism being used by a client. Instead, it should present a consistent view and return NFS4ERR_WRONGSEC if an attempt is made to access data with an inappropriate security mechanism.
If security considerations make it necessary to hide the existence of a particular file system, as opposed to all of the data within it, the server can apply the security policy of a shared resource in the server's namespace to components of the resource's ancestors. For example:
/ (place holder/not exported)
/a/b (file system 1)
/a/b/MySecretProject (file system 2)
The /a/b/MySecretProject directory is a real file system and is the shared resource. Suppose the security policy for /a/b/MySecretProject is Kerberos with integrity and it is desired to limit knowledge of the existence of this file system. In this case, the server should apply the same security policy to /a/b. This allows for knowledge of the existence of a file system to be secured when desirable.
For the case of the use of multiple, disjoint security mechanisms in the server's resources, applying that sort of policy would result in the higher-level file system not being accessible using any security flavor, which would make the that higher-level file system inaccessible. Therefore, that sort of configuration is not compatible with hiding the existence (as opposed to the contents) from clients using multiple disjoint sets of security flavors.
In other circumstances, a desirable policy is for the security of a particular object in the server's namespace should include the union of all security mechanisms of all direct descendants. A common and convenient practice, unless strong security requirements dictate otherwise, is to make all of the pseudo file system accessible by all of the valid security mechanisms.
Where there is concern about the security of data on the network, clients should use strong security mechanisms to access the pseudo file system in order to prevent man-in-the-middle attacks.
| TOC |
Integrating locking into the NFS protocol necessarily causes it to be stateful. With the inclusion of such features as share reservations, file and directory delegations, recallable layouts, and support for mandatory byte-range locking, the protocol becomes substantially more dependent on proper management of state than the traditional combination of NFS and NLM [36] (The Open Group, “Protocols for Interworking: XNFS, Version 3W, ISBN 1-85912-184-5,” February 1998.). These features include expanded locking facilities, which provide some measure of interclient exclusion, but the state also offers features not readily providable using a stateless model. There are three components to making this state manageable:
In this model, the server owns the state information. The client requests changes in locks and the server responds with the changes made. Non-client-initiated changes in locking state are infrequent. The client receives prompt notification of such changes and can adjust its view of the locking state to reflect the server's changes.
Individual pieces of state created by the server and passed to the client at its request are represented by 128-bit stateids. These stateids may represent a particular open file, a set of byte-range locks held by a particular owner, or a recallable delegation of privileges to access a file in particular ways, or at a particular location.
In all cases, there is a transition from the most general information which represents a client as a whole to the eventual lightweight stateid used for most client and server locking interactions. The details of this transition will vary with the type of object but it always starts with a client ID.
| TOC |
A client must establish a client ID (see Section 2.4 (Client Identifiers and Client Owners)) and then one or more sessionids (see Section 2.10 (Session)) before performing any operations to open, lock, delegate, or obtain a layout for a file object. Each session id is associated with a specific client ID, and thus serves as a shorthand reference to an NFSv4.1 client.
For some types of locking interactions, the client will represent some number of internal locking entities called "owners", which normally correspond to processes internal to the client. For other types of locking-related objects, such as delegations and layouts, no such intermediate entities are provided for, and the locking-related objects are considered to be transferred directly between the server and a unitary client.
| TOC |
When the server grants a lock of any type (including opens, byte-range locks, delegations, and layouts) it responds with a unique stateid, that represents a set of locks (often a single lock) for the same file, of the same type, and sharing the same ownership characteristics. Thus opens of the same file by different open-owners each have an identifying stateid. Similarly, each set of byte-range locks on a file owned by a specific lock-owner has its own identifying stateid. Delegations and layouts also have associated stateids by which they may be referenced. The stateid is used as a shorthand reference to a lock or set of locks and given a stateid the server can determine the associated state-owner or state-owners (in the case of an open-owner/lock-owner pair) and the associated filehandle. When stateids are used, the current filehandle must be the one associated with that stateid.
All stateids associated with a given client ID are associated with a common lease which represents the claim of those stateids and the objects they represent to be maintained by the server. See Section 8.3 (Lease Renewal) for a discussion of leases.
The server may assign stateids independently for different clients. A stateid with the same bit pattern for one client may designate an entirely different set of locks for a different client. The stateid is always interpreted with respect to the client ID associated with the current session. Stateids apply to all sessions associated with the given client ID and the client may use a stateid obtained from one session on another session associated with the same client ID.
| TOC |
With the exception of special stateids (see Section 8.2.3 (Special Stateids)), each stateid represents locking objects of one of a set of types defined by the NFSv4.1 protocol. Note that in all these cases, where we speak of guarantee, it is understood there are situations such as a client restart, or lock revocation, that allow the guarantee to be voided.
| TOC |
Stateids are divided into two fields, a 96-bit "other" field identifying the specific set of locks and a 32-bit "seqid" sequence value. Except in the case of special stateids (see Section 8.2.3 (Special Stateids)), a particular value of the "other" field denotes a set of locks of the same type (for example byte-range locks, opens, delegations, or layouts), for a specific file or directory, and sharing the same ownership characteristics. The seqid designates a specific instance of such a set of locks, and is incremented to indicate changes in such a set of locks, either by the addition or deletion of locks from the set, a change in the byte-range they apply to, or an upgrade or downgrade in the type of one or more locks.
When such a set of locks is first created the server returns a stateid with seqid value of one. On subsequent operations which modify the set of locks the server is required to increment the seqid field by one (1) whenever it returns a stateid for the same state-owner/file/type combination and there is some change in the set of locks actually designated. In this case the server will return a stateid with an other field the same as previously used for that state-owner/file/type combination, with an incremented seqid field. This pattern continues until the seqid is incremented past NFS4_UINT32_MAX, and one (not zero) is the next seqid value.
The purpose of the incrementing of the seqid is to allow the server to communicate to the client the order in which operations that modified locking state associated with a stateid have been processed and to make it possible for the client to send requests that are conditional on the set of locks not having changed since the stateid in question was returned.
Except for layout stateids (Section 12.5.3 (Layout Stateid)) when a client sends a stateid to the server, it has two choices with regard to the seqid sent. It may set the seqid to zero to indicate to the server that it wishes the most up-to-date seqid for that stateid's "other" field to be used. This would be the common choice in the case of a stateid sent with a READ or WRITE operation. It also may set a non-zero value in which case the server checks if that seqid is the correct one. In that case the server is required to return NFS4ERR_OLD_STATEID if the seqid is lower than the most current value and NFS4ERR_BAD_STATEID if the seqid is greater than the most current value. This would be the common choice in the case of stateids sent with a CLOSE or OPEN_DOWNGRADE. Because OPENs may be sent in parallel for the same owner, a client might close a file without knowing that an OPEN upgrade had been done by the server, changing the lock in question. If CLOSE were sent with a zero seqid, the OPEN upgrade would be canceled before the client even received an indication that an upgrade had happened.
When a stateid is sent by the server to client as part of a callback operation, it is not subject to checking for a current seqid and returning NFS4ERR_OLD_STATEID. This is because the client is not in a position to know the most up-to-date seqid and thus cannot verify it. Unless specially noted, the seqid value for a stateid sent by the server to the client as part of a callback is required to be zero with NFS4ERR_BAD_STATEID returned if it is not.
In making comparisons between seqids, both by the client in determining the order of operations and by the server in determining whether the NFS4ERR_OLD_STATEID is to be returned, the possibility of the seqid being swapped around past the NFS4_UINT32_MAX value needs to be taken into account. When two seqid values are being compared, the total count of slots for all sessions associated with the current client is used to do this. When one seqid value is less that this total slot count and another seqid value is greater than NFS4_UINT32_MAX minus the total slot count, the former is to be treated as lower than the later, despite the fact that it is numerically greater.
| TOC |
Stateid values whose "other" field is either all zeros or all ones are reserved. They may not be assigned by the server but have special meanings defined by the protocol. The particular meaning depends on whether the "other" field is all zeros or all ones and the specific value of the "seqid" field.
The following combinations of "other" and "seqid" are defined in NFSv4.1:
If a stateid value is used which has all zero or all ones in the "other" field, but does not match one of the cases above, the server MUST return the error NFS4ERR_BAD_STATEID.
Special stateids, unlike other stateids, are not associated with individual client IDs or filehandles and can be used with all valid client IDs and filehandles. In the case of a special stateid designating the current stateid, the current stateid value substituted for the special stateid is associated with a particular client ID and filehandle, and so, if it is used where current filehandle does not match that associated with the current stateid, the operation to which the stateid is passed will return NFS4ERR_BAD_STATEID.
| TOC |
Stateids must remain valid until either a client restart or a server restart or until the client returns all of the locks associated with the stateid by means of an operation such as CLOSE or DELEGRETURN. If the locks are lost due to revocation the stateid remains a valid designation of that revoked state until the client frees it by using FREE_STATEID. Stateids associated with byte-range locks are an exception. They remain valid even if a LOCKU frees all remaining locks, so long as the open file with which they are associated remains open, unless the client does a FREE_STATEID to cause the stateid to be freed.
It should be noted that there are situations in which the client's locks become invalid, without the client requesting they be returned. These include lease expiration and a number of forms of lock revocation within the lease period. It is important to note that in these situations, the stateid remains valid and the client can use it to determine the disposition of the associated lost locks.
An "other" value must never be reused for a different purpose (i.e. different filehandle, owner, or type of locks) within the context of a single client ID. A server may retain the "other" value for the same purpose beyond the point where it may otherwise be freed but if it does so, it must maintain "seqid" continuity with previous values.
One mechanism that may be used to satisfy the requirement that the server recognize invalid and out-of-date stateids is for the server to divide the "other" field of the stateid into two fields.
And then store in each table entry,
With this information, an incoming stateid can be validated and the appropriate error returned when necessary. Special and non-special stateids are handled separately. (See Section 8.2.3 (Special Stateids) for a discussion of special stateids.)
Note that stateids are implicitly qualified by the current client ID, as derived from the client ID associated with the current session. Note however, that the semantics of the session will prevent stateids associated with a previous client or server instance from being analyzed by this procedure.
If server restart has resulted in an invalid client ID or a session id which is invalid, SEQUENCE will return an error and the operation that takes a stateid as an argument will never be processed.
If there has been a server restart where there is a persistent session, and all leased state has been lost, then the session in question will, although valid, be marked as dead, and any operation not satisfied by means of the reply cache will receive the error NFS4ERR_DEADSESSION, and thus not be processed as indicated below.
When a stateid is being tested, and the "other" field is all zeros or all ones, a check that the "other" and "seqid" fields match a defined combination for a special stateid is done and the results determined as follows:
When a stateid is being tested, and the "other" field is neither all zeros or all ones, the following procedure could be used to validate an incoming stateid and return an appropriate error, when necessary, assuming that the "other" field would be divided into a table index and an entry generation.
| TOC |
Clients performing I/O operations need to select an appropriate stateid based on the locks (including opens and delegations) held by the client and the various types of state-owners issuing the I/O requests. SETATTR operations which change the file size are treated like I/O operations in this regard.
The following rules, applied in order of decreasing priority, govern the selection of the appropriate stateid. In following these rules, the client will only consider locks of which it has actually received notification by an appropriate operation response or callback. Note that the rules are slightly different in the case of I/O to data servers when file layouts are being used (see Section 13.9.1 (Global Stateid Requirements)).
Ignoring these rules may result in situations in which the server does not have information necessary to properly process the request. For example, when mandatory byte-range locks are in effect, if the stateid does not indicate the proper lockowner, via a lock stateid, a request might be avoidably rejected.
The server however should not try to enforce these ordering rules and should use whatever information is available to proper process I/O requests. In particular, when a client has a delegation for a given file, it SHOULD take note of this fact in processing a request, even if it is sent with a special stateid.
| TOC |
Because each operation is associated with a session id and from that the clientid can be determined, operations do not need to include a stateid for the server to be able to determine whether the they should cause a delegation to be recalled or are to be treated as done within the scope of the delegation.
In the case of SETATTR operations, a stateid is present. In cases other than those which set the file size, the client may send either a special stateid or, when a delegation is held for the file in question, a delegation stateid. While the server SHOULD validate the stateid and may use the stateid to optimize the determination as to whether a delegation is held, it SHOULD note the presence of a delegation even when a special stateid is sent, and MUST accept a valid delegation stateid when sent.
| TOC |
The purpose of a lease is to allow the client to indicate to the server, in a low-overhead way, that it is active, and thus that the server is to retain the client's locks. This arrangement allows the server to remove stale locking-related objects that are held by a client that has crashed or is otherwise unreachable, once the relevant lease expires. This in turn allows other clients to obtain conflicting locks without being delayed indefinitely by inactive or unreachable clients. It is not a mechanism for cache consistency and lease renewals may not be denied if the lease interval has not expired.
Since each session is associated with a specific client (identified by the client's client ID), any operation sent on that session is an indication that the associated client is reachable. When a request is sent for a given session, successful execution of a SEQUENCE operation (or successful retrieval of the result of SEQUENCE from the reply cache) on an unexpired lease will result in the lease being implicitly renewed, for the standard renewal period (equal to the lease_time attribute).
If the client ID's lease has not expired when the server receives a SEQUENCE operation, then the server MUST renew the lease. If the client ID's lease has expired when the server receives a SEQUENCE operation, the server MAY renew the lease; this depends on whether any state was revoked as a result of the client's failure to renew the lease before expiration.
Absent other activity that would renew the lease, a COMPOUND consisting of a single SEQUENCE operation will suffice. The client should also take communication-related delays into account and take steps to ensure that the renewal messages actually reach the server in good time. For example:
If the server renews the lease upon receiving a SEQUENCE operation, the server MUST NOT allow the lease to expire while the rest of the operations in the COMPOUND procedure's request are still executing. Once the last operation has finished, and the response to COMPOUND has been sent, the server MUST set the lease to expire no sooner than the sum of current time and the value of the lease_time attribute.
A client ID's lease can expire when it has been at least the lease interval (lease_time) since the last lease-renewing SEQUENCE operation was sent on any of the client ID's sessions and there are no active COMPOUND operations on any such sessions.
Because the SEQUENCE operation is the basic mechanism to renew a lease, and because if must be done at least once for each lease period, it is the natural mechanism whereby the server will inform the client of changes in the lease status that the client needs to be informed of. The client should inspect the status flags (sr_status_flags) returned by sequence and take the appropriate action (see Section 18.46.3 (DESCRIPTION) for details).
| TOC |
A critical requirement in crash recovery is that both the client and the server know when the other has failed. Additionally, it is required that a client sees a consistent view of data across server restarts. All READ and WRITE operations that may have been queued within the client or network buffers must wait until the client has successfully recovered the locks protecting the READ and WRITE operations. Any that reach the server before the server can safely determine that the client has recovered enough locking state to be sure that such operations can be safely processed must be rejected. This will happen because either:
| TOC |
In the event that a client fails, the server may release the client's locks when the associated lease has expired. Conflicting locks from another client may only be granted after this lease expiration. As discussed in Section 8.3 (Lease Renewal), when a client has not failed and re-establishes its lease before expiration occurs, requests for conflicting locks will not be granted.
To minimize client delay upon restart, lock requests are associated with an instance of the client by a client-supplied verifier. This verifier is part of the client_owner4 sent in the initial EXCHANGE_ID call made by the client. The server returns a client ID as a result of the EXCHANGE_ID operation. The client then confirms the use of the client ID by establishing a session associated with that client ID (see Section 18.36.3 (DESCRIPTION) for a description how this is done). All locks, including opens, byte-range locks, delegations, and layouts obtained by sessions using that client ID are associated with that client ID.
Since the verifier will be changed by the client upon each initialization, the server can compare a new verifier to the verifier associated with currently held locks and determine that they do not match. This signifies the client's new instantiation and subsequent loss (upon confirmation of the new client ID) of locking state. As a result, the server is free to release all locks held which are associated with the old client ID which was derived from the old verifier. At this point conflicting locks from other clients, kept waiting while the lease had not yet expired, can be granted. In addition, all stateids associated with the old client ID can also be freed, as they are no longer reference-able.
Note that the verifier must have the same uniqueness properties as the verifier for the COMMIT operation.
| TOC |
If the server loses locking state (usually as a result of a restart), it must allow clients time to discover this fact and re-establish the lost locking state. The client must be able to re-establish the locking state without having the server deny valid requests because the server has granted conflicting access to another client. Likewise, if there is a possibility that clients have not yet re-established their locking state for a file, and that such locking state might make it invalid to perform READ or WRITE operations, for example through the establishment of mandatory locks, the server must disallow READ and WRITE operations for that file.
A client can determine that loss of locking state has occurred via several methods.
| TOC |
When state information and the associated locks are lost as a result of a server restart, the protocol must provide a way to cause that state to be re-established. The approach used is to define, for most types of locking state (layouts are an exception), a request whose function is to allow the client to re-establish on the server a lock first obtained from a previous instance. Generally these requests are variants of the requests normally used to create locks of that type and are referred to as "reclaim-type" requests and the process of re-establishing such locks is referred to as "reclaiming" them.
Because each client must have an opportunity to reclaim all of the locks that it has without the possibility that some other client will be granted a conflicting lock, a special period called the "grace period" is devoted to the reclaim process. During this period, requests creating client IDs and sessions are handled normally, but locking requests are subject to special restrictions. Only reclaim-type locking requests are allowed, unless the server can reliably determine (through state persistently maintained across restart instances), that granting any such lock cannot possibly conflict with a subsequent reclaim. When a request is made to obtain a new lock (i.e. not a reclaim-type request) during the grace period and such a determination cannot be made, the server must return the error NFS4ERR_GRACE.
Once a session is established using the new client ID, the client will use reclaim-type locking requests (e.g. LOCK requests with reclaim set to TRUE and OPEN operations with a claim type of CLAIM_PREVIOUS; see Section 9.11 (Reclaim of Open and Byte-Range Locks)) to re-establish its locking state. Once this is done, or if there is no such locking state to reclaim, the client sends a global RECLAIM_COMPLETE operation, i.e. one with the rca_one_fs argument set to FALSE, to indicate that it has reclaimed all of the locking state that it will reclaim. Once a client sends such a RECLAIM_COMPLETE operation, it may attempt non-reclaim locking operations, although it may get NFS4ERR_GRACE errors the operations until the period of special handling is over. See Section 11.7.7 (Lock State and File System Transitions) for a discussion of the analogous handling lock reclamation in the case of file systems transitioning from server to server.
During the grace period, the server must reject READ and WRITE operations and non-reclaim locking requests (i.e. other LOCK and OPEN operations) with an error of NFS4ERR_GRACE, unless it can guarantee that these may be done safely, as described below.
The grace period may last until all clients which are known to possibly have had locks have done a global RECLAIM_COMPLETE operation, indicating that they have finished reclaiming the locks they held before the server restart. This means that a client which has done a RECLAIM_COMPLETE must be prepared to receive an NFS4ERR_GRACE when attempting to acquire new locks. In order for the server to know that all clients with possible prior lock state have done a RECLAIM_COMPLETE, the server must maintain in stable storage a list of clients which may have such locks. The server may also terminate the grace period before all clients have done a global RECLAIM_COMPLETE. The server SHOULD NOT terminate the grace period before a time equal to the lease period in order to give clients an opportunity to find out about the server restart, as a result of issuing requests on associated sessions with a frequency governed by the lease time. Note that when a client does not issue such requests (or they are issued by the client but not received by the server), it is possible for the grace period to expire before the client finds out that the server restart has occurred.
Some additional time in order to allow a client to establish a new client ID and session and to effect lock reclaims may be added to the lease time. Note that analogous rules apply to file system-specific grace periods discussed in Section 11.7.7 (Lock State and File System Transitions).
If the server can reliably determine that granting a non-reclaim request will not conflict with reclamation of locks by other clients, the NFS4ERR_GRACE error does not have to be returned even within the grace period, although NFS4ERR_GRACE must always be returned to clients attempting a non-reclaim lock request before doing their own global RECLAIM_COMPLETE. For the server to be able to service READ and WRITE operations during the grace period, it must again be able to guarantee that no possible conflict could arise between a potential reclaim locking request and the READ or WRITE operation. If the server is unable to offer that guarantee, the NFS4ERR_GRACE error must be returned to the client.
For a server to provide simple, valid handling during the grace period, the easiest method is to simply reject all non-reclaim locking requests and READ and WRITE operations by returning the NFS4ERR_GRACE error. However, a server may keep information about granted locks in stable storage. With this information, the server could determine if a regular lock or READ or WRITE operation can be safely processed.
For example, if the server maintained on stable storage summary information on whether mandatory locks exist, either mandatory byte-range locks, or share reservations specifying deny modes, many requests could be allowed during the grace period. If it is known that no such share reservations exist, OPEN request that do not specify deny modes may be safely granted. If, in addition, it is known that no mandatory byte-range locks exist, either through information stored on stable storage or simply because the server does not support such locks, READ and WRITE requests may be safely processed during the grace period. Another important case is where it is known that no mandatory byte-range locks exist, either because the server does not provide support for them, or because their absence is known from persistently recorded data. In this case, READ and WRITE operations specifying stateids derived from reclaim-type operation may be validly processed during the grace period because the fact of the valid reclaim ensures that no lock subsequently granted can prevent the I/O.
To reiterate, for a server that allows non-reclaim lock and I/O requests to be processed during the grace period, it MUST determine that no lock subsequently reclaimed will be rejected and that no lock subsequently reclaimed would have prevented any I/O operation processed during the grace period.
Clients should be prepared for the return of NFS4ERR_GRACE errors for non-reclaim lock and I/O requests. In this case the client should employ a retry mechanism for the request. A delay (on the order of several seconds) between retries should be used to avoid overwhelming the server. Further discussion of the general issue is included in [37] (Floyd, S. and V. Jacobson, “The Synchronization of Periodic Routing Messages,” April 1994.). The client must account for the server that can perform I/O and non-reclaim locking requests within the grace period as well as those that cannot do so.
A reclaim-type locking request outside the server's grace period can only succeed if the server can guarantee that no conflicting lock or I/O request has been granted since restart.
A server may, upon restart, establish a new value for the lease period. Therefore, clients should, once a new client ID is established, refetch the lease_time attribute and use it as the basis for lease renewal for the lease associated with that server. However, the server must establish, for this restart event, a grace period at least as long as the lease period for the previous server instantiation. This allows the client state obtained during the previous server instance to be reliably re-established.
| TOC |
If the duration of a network partition is greater than the lease period provided by the server, the server will not have received a lease renewal from the client. If this occurs, the server may free all locks held for the client, or it may allow the lock state to remain for a considerable period, subject to the constraint that if a request for a conflicting lock is made, locks associated with an expired lease do not prevent such a conflicting lock from being granted but MUST be revoked as necessary so as not to interfere with such conflicting requests.
If the server chooses to delay freeing of lock state until there is a conflict, it may either free all of the clients locks once there is a conflict, or it may only revoke the minimum set of locks necessary to allow conflicting requests. When it adopts the finer-grained approach, it must revoke all locks associated with a given stateid, even if the conflict is with only a subset of locks.
When the server chooses to free all of a client's lock state, either immediately upon lease expiration, or a result of the first attempt to obtain a conflicting a lock, the server may report the loss of lock state in a number of ways.
The server may choose to invalidate the session and the associated client ID. In this case, once the client can communicate with the server, it will receive an NFS4ERR_BADSESSION error. Upon attempting to create a new session, it would get an NFS4ERR_STALE_CLIENTID. Upon creating the new client ID and new session it would attempt to reclaim locks not be allowed to do so by the server.
Another possibility is for the server to maintain the session and client ID but for all stateids held by the client to become invalid or stale. Once the client can reach the server after such a network partition, the status returned by the SEQUENCE operation will indicate a loss of locking state, i.e. the flag SEQ4_STATUS_EXPIRED_ALL_STATE_REVOKED will be set in sr_status_flags. In addition, all I/O submitted by the client with the now invalid stateids will fail with the server returning the error NFS4ERR_EXPIRED. Once the client learns of the loss of locking state, it will suitably notify the applications that held the invalidated locks. The client should then take action to free invalidated stateids, either by establishing a new client ID using a new verifier or by doing a FREE_STATEID operation to release each of the invalidated stateids.
When the server adopts a finer-grained approach to revocation of locks when lease have expired, only a subset of stateids will normally become invalid during a network partition. When the client can communicate with the server after such a network partition heals, the status returned by the SEQUENCE operation will indicate a partial loss of locking state (SEQ4_STATUS_EXPIRED_SOME_STATE_REVOKED). In addition, operations, including I/O submitted by the client, with the now invalid stateids will fail with the server returning the error NFS4ERR_EXPIRED. Once the client learns of the loss of locking state, it will use the TEST_STATEID operation on all of its stateids to determine which locks have been lost and then suitably notify the applications that held the invalidated locks. The client can then release the invalidated locking state and acknowledge the revocation of the associated locks by doing a FREE_STATEID operation on each of the invalidated stateids.
When a network partition is combined with a server restart, there are edge conditions that place requirements on the server in order to avoid silent data corruption following the server restart. Two of these edge conditions are known, and are discussed below.
The first edge condition arises as a result of the scenarios such as the following:
Thus, at the final step, the server has erroneously granted client A's lock reclaim. If client B modified the object the lock was protecting, client A will experience object corruption.
The second known edge condition arises in situations such as the following:
As with the first edge condition, the final step of the scenario of the second edge condition has the server erroneously granting client A's lock reclaim.
Solving the first and second edge conditions requires that the server either always assumes after it restarts that some edge condition occurs, and thus return NFS4ERR_NO_GRACE for all reclaim attempts, or that the server record some information in stable storage. The amount of information the server records in stable storage is in inverse proportion to how harsh the server intends to be whenever edge conditions arise. The server that is completely tolerant of all edge conditions will record in stable storage every lock that is acquired, removing the lock record from stable storage only when the lock is released. For the two edge conditions discussed above, the harshest a server can be, and still support a grace period for reclaims, requires that the server record in stable storage information some minimal information. For example, a server implementation could, for each client, save in stable storage a record containing:
Assuming the above record keeping, for the first edge condition, after the server restarts, the record that client A's lease expired means that another client could have acquired a conflicting byte-range lock, share reservation, or delegation. Hence the server must reject a reclaim from client A with the error NFS4ERR_NO_GRACE.
For the second edge condition, after the server restarts for a second time, the indication that the client had not completed its reclaims at the time at which the grace period ended means that the server must reject a reclaim from client A with the error NFS4ERR_NO_GRACE.
When either edge condition occurs, the client's attempt to reclaim locks will result in the error NFS4ERR_NO_GRACE. When this is received, or after the client restarts with no lock state, the client will send a global RECLAIM_COMPLETE. When the RECLAIM_COMPLETE is received, the server and client are again in agreement regarding reclaimable locks and both booleans in persistent storage can be reset, to be set again only when there is a subsequent event that causes lock reclaim operations to be questionable.
Regardless of the level and approach to record keeping, the server MUST implement one of the following strategies (which apply to reclaims of share reservations, byte-range locks, and delegations):
A mandate for the client's handling of the NFS4ERR_NO_GRACE error is outside the scope of this specification, since the strategies for such handling are very dependent on the client's operating environment. However, one potential approach is described below.
When the client receives NFS4ERR_NO_GRACE, it could examine the change attribute of the objects the client is trying to reclaim state for, and use that to determine whether to re-establish the state via normal OPEN or LOCK requests. This is acceptable provided the client's operating environment allows it. In other words, the client implementor is advised to document for his users the behavior. The client could also inform the application that its byte-range lock or share reservations (whether they were delegated or not) have been lost, such as via a UNIX signal, a GUI pop-up window, etc. See Section 10.5 (Data Caching and Revocation) for a discussion of what the client should do for dealing with unreclaimed delegations on client state.
For further discussion of revocation of locks see Section 8.5 (Server Revocation of Locks).
| TOC |
At any point, the server can revoke locks held by a client and the client must be prepared for this event. When the client detects that its locks have been or may have been revoked, the client is responsible for validating the state information between itself and the server. Validating locking state for the client means that it must verify or reclaim state for each lock currently held.
The first occasion of lock revocation is upon server restart. Note that this includes situations in which sessions are persistent and locking state is lost. In this class of instances, the client will receive an error (NFS4ERR_STALE_CLIENTID on an operation that takes client ID, usually as part of recovery in response to a problem with the current session) and the client will proceed with normal crash recovery as described in the Section 8.4.2.1 (State Reclaim).
The second occasion of lock revocation is the inability to renew the lease before expiration, as discussed in Section 8.4.3 (Network Partitions and Recovery). While this is considered a rare or unusual event, the client must be prepared to recover. The server is responsible for determining the precise consequences of the lease expiration, informing the client of the scope of the lock revocation decided upon. The client then uses the status information provided by the server in the SEQUENCE results (field sr_status_flags, see Section 18.46.3 (DESCRIPTION)) to synchronize its locking state with that of the server, in order to recover.
The third occasion of lock revocation can occur as a result of revocation of locks within the lease period, either because of administrative intervention, or because a recallable lock (a delegation or layout) was not returned within the lease period after having been recalled. While these are considered rare events, they are possible and the client must be prepared to deal with them. When either of these events occur, the client finds out about the situation through the status returned by the SEQUENCE operation. Any use of stateids associated with locks revoked during the lease period will receive the error NFS4ERR_ADMIN_REVOKED or NFS4ERR_DELEG_REVOKED, as appropriate.
In all situations in which a subset of locking state may have been revoked, which include all cases in which locking state is revoked within the lease period, it is up to the client to determine which locks have been revoked and which have not. It does this by using the TEST_STATEID operation on the appropriate set of stateids. Once the set of revoked locks has been determined, the applications can be notified, and the invalidated stateids can be freed and lock revocation acknowledged by using FREE_STATEID.
| TOC |
When determining the time period for the server lease, the usual lease tradeoffs apply. Short leases are good for fast server recovery at a cost of increased operations to effect lease renewal (when there are no other operations during the period to effect lease renewal as a side-effect). Long leases are certainly kinder and gentler to servers trying to handle very large numbers of clients. The number of extra requests to effect lock renewal drops in inverse proportion to the lease time. The disadvantages of long leases include the possibility of slower recovery after certain failures. After server failure, a longer grace period may be required when some clients do not promptly reclaim their locks and do a global RECLAIM_COMPLETE. In the event of client failure, there can be a longer period for leases to expire thus forcing conflicting requests to wait.
Long leases are practical if the server is can store lease state in non-volatile memory. Upon recovery, the server can reconstruct the lease state from its non-volatile memory and continue operation with its clients and therefore long leases would not be an issue.
| TOC |
To avoid the need for synchronized clocks, lease times are granted by the server as a time delta. However, there is a requirement that the client and server clocks do not drift excessively over the duration of the lease. There is also the issue of propagation delay across the network which could easily be several hundred milliseconds as well as the possibility that requests will be lost and need to be retransmitted.
To take propagation delay into account, the client should subtract it from lease times (e.g. if the client estimates the one-way propagation delay as 200 milliseconds, then it can assume that the lease is already 200 milliseconds old when it gets it). In addition, it will take another 200 milliseconds to get a response back to the server. So the client must send a lease renewal or write data back to the server at least 400 milliseconds before the lease would expire.
The server's lease period configuration should take into account the network distance of the clients that will be accessing the server's resources. It is expected that the lease period will take into account the network propagation delays and other network delay factors for the client population. Since the protocol does not allow for an automatic method to determine an appropriate lease period, the server's administrator may have to tune the lease period.
| TOC |
There are a number of operations and fields within existing operations that no longer have a function in NFSv4.1. In one way or another, these changes are all due to the implementation of sessions which provides client context and exactly once semantics as a base feature of the protocol, separate from locking itself.
The following NFSv4.0 operations MUST NOT be implemented in NFSv4.1. The server MUST return NFS4ERR_NOTSUPP if these operations are found in an NFSv4.1 COMPOUND.
Also, there are a number of fields, present in existing operations related to locking that have no use in minor version one. They were used in minor version zero to perform functions now provided in a different fashion.
Such vestigial fields in existing operations have no function in NFSv4.1 and are ignored by the server. Note that client IDs in operations new to NFSv4.1 (such as CREATE_SESSION and DESTROY_CLIENTID) are not ignored.
| TOC |
To support Win32 share reservations it is necessary to provide operations which atomically open or create files. Having a separate share/unshare operation would not allow correct implementation of the Win32 OpenFile API. In order to correctly implement share semantics, the previous NFS protocol mechanisms used when a file is opened or created (LOOKUP, CREATE, ACCESS) need to be replaced. The NFSv4.1 protocol defines an OPEN operation which is capable of atomically looking up, creating, and locking a file on the server.
| TOC |
It is assumed that manipulating a byte-range lock is rare when compared to READ and WRITE operations. It is also assumed that server restarts and network partitions are relatively rare. Therefore it is important that the READ and WRITE operations have a lightweight mechanism to indicate if they possess a held lock. A byte-range lock request contains the heavyweight information required to establish a lock and uniquely define the owner of the lock.
| TOC |
When opening a file or requesting a byte-range lock, the client must specify an identifier which represents the owner of the requested lock. This identifier is in the form of a state-owner, represented in the protocol by a state_owner4, a variable-length opaque array which, when concatenated with the current client ID uniquely defines the owner of lock managed by the client. This may be a thread id, process id, or other unique value.
Owners of opens and owners of byte-range locks are separate entities and remain separate even if the same opaque arrays are used to designate owners of each. The protocol distinguishes between open-owners (represented by open_owner4 structures) and lock-owners (represented by lock_owner4 structures).
Each open is associated with a specific open-owner while each byte-range lock is associated with a lock-owner and an open-owner, the latter being the open-owner associated with the open file under which the LOCK operation was done. Delegations and layouts, on the other hand, are not associated with a specific owner but are associated with the client as a whole (identified by a client ID).
| TOC |
All READ, WRITE and SETATTR operations contain a stateid. For the purposes of this section, SETATTR operations which change the size attribute of a file are treated as if they are writing the area between the old and new size (i.e. the range truncated or added to the file by means of the SETATTR), even where SETATTR is not explicitly mentioned in the text. The stateid passed to one of these operations must be one that represents an open, a set of byte-range locks, or a delegation, or it may be a special stateid representing anonymous access or the special bypass stateid.
If the state-owner performs a READ or WRITE in a situation in which it has established a byte-range lock or share reservation on the server (any OPEN constitutes a share reservation) the stateid (previously returned by the server) must be used to indicate what locks, including both byte-range locks and share reservations, are held by the state-owner. If no state is established by the client, either byte-range lock or share reservation, a special stateid for anonymous state (zero as "other" and "seqid") is used. (See Section 8.2.3 (Special Stateids) for a description of 'special' stateids in general.) Regardless whether a stateid for anonymous state or a stateid returned by the server is used, if there is a conflicting share reservation or mandatory byte-range lock held on the file, the server MUST refuse to service the READ or WRITE operation.
Share reservations are established by OPEN operations and by their nature are mandatory in that when the OPEN denies READ or WRITE operations, that denial results in such operations being rejected with error NFS4ERR_LOCKED. Byte-range locks may be implemented by the server as either mandatory or advisory, or the choice of mandatory or advisory behavior may be determined by the server on the basis of the file being accessed (for example, some UNIX-based servers support a "mandatory lock bit" on the mode attribute such that if set, byte-range locks are required on the file before I/O is possible). When byte-range locks are advisory, they only prevent the granting of conflicting lock requests and have no effect on READs or WRITEs. Mandatory byte-range locks, however, prevent conflicting I/O operations. When they are attempted, they are rejected with NFS4ERR_LOCKED. When the client gets NFS4ERR_LOCKED on a file it knows it has the proper share reservation for, it will need to send a LOCK request on the region of the file that includes the region the I/O was to be performed on, with an appropriate locktype (i.e. READ*_LT for a READ operation, WRITE*_LT for a WRITE operation).
Note that for UNIX environments that support mandatory file locking, the distinction between advisory and mandatory locking is subtle. In fact, advisory and mandatory byte-range locks are exactly the same in so far as the APIs and requirements on implementation. If the mandatory lock attribute is set on the file, the server checks to see if the lock-owner has an appropriate shared (read) or exclusive (write) byte-range lock on the region it wishes to read or write to. If there is no appropriate lock, the server checks if there is a conflicting lock (which can be done by attempting to acquire the conflicting lock on behalf of the lock-owner, and if successful, release the lock after the READ or WRITE is done), and if there is, the server returns NFS4ERR_LOCKED.
For Windows environments, byte-range locks are always mandatory, so the server always checks for byte-range locks during I/O requests.
Thus, the NFSv4.1 LOCK operation does not need to distinguish between advisory and mandatory byte-range locks. It is the NFSv4.1 server's processing of the READ and WRITE operations that introduces the distinction.
Every stateid which is validly passed to READ, WRITE or SETATTR, with the exception of special stateid values, defines an access mode for the file (i.e. READ, WRITE, or READ-WRITE)
When a READ, WRITE, or SETATTR (which specifies the size attribute) is done, the operation is subject to checking against the access mode to verify that the operation is appropriate given the stateid with which the operation is associated.
In the case of WRITE-type operations (i.e. WRITEs and SETATTRs which set size), the server MUST verify that the access mode allows writing and MUST return an NFS4ERR_OPENMODE error if it does not. In the case, of READ, the server may perform the corresponding check on the access mode, or it may choose to allow READ on opens for WRITE only, to accommodate clients whose write implementation may unavoidably do reads (e.g. due to buffer cache constraints). However, even if READs are allowed in these circumstances, the server MUST still check for locks that conflict with the READ (e.g. another open specify denial of READs). Note that a server which does enforce the access mode check on READs need not explicitly check for conflicting share reservations since the existence of OPEN for read access guarantees that no conflicting share reservation can exist.
The read bypass special stateid (all bits of "other" and "seqid" set to one) indicates a desire to bypass locking checks. The server MAY allow READ operations to bypass locking checks at the server, when this special stateid is used. However, WRITE operations with this special stateid value MUST NOT bypass locking checks and are treated exactly the same as if a special stateid for anonymous state were used.
A lock may not be granted while a READ or WRITE operation using one of the special stateids is being performed and the scope of the lock to be granted would conflict with the READ or WRITE operation. This can occur when:
When a client holds a delegation, it needs to ensure that the stateid sent conveys the association of operation with the delegation, to avoid the delegation from being avoidably recalled. When the delegation stateid, or a stateid open associated with that delegation, or a stateid representing byte-range locks derived form such an open is used, the server knows that the READ, WRITE, or SETATTR does not conflict with the delegation, but is sent under the aegis of the delegation. Even though it is possible for the server to determine from the client ID (via the session id) that the client does in fact have a delegation, the server is not obliged to check this, so using a special stateid can result in avoidable recall of the delegation.
| TOC |
The protocol allows a lock-owner to request a lock with a byte range and then either upgrade, downgrade, or unlock a sub-range of the initial lock, or a range that consists of a range which overlaps, fully or partially, that initial lock or a combination of a set of existing locks for the same lock-owner. It is expected that this will be an uncommon type of request. In any case, servers or server file systems may not be able to support sub-range lock semantics. In the event that a server receives a locking request that represents a sub-range of current locking state for the lock-owner, the server is allowed to return the error NFS4ERR_LOCK_RANGE to signify that it does not support sub-range lock operations. Therefore, the client should be prepared to receive this error and, if appropriate, report the error to the requesting application.
The client is discouraged from combining multiple independent locking ranges that happen to be adjacent into a single request since the server may not support sub-range requests and for reasons related to the recovery of file locking state in the event of server failure. As discussed in Section 8.4.2 (Server Failure and Recovery), the server may employ certain optimizations during recovery that work effectively only when the client's behavior during lock recovery is similar to the client's locking behavior prior to server failure.
| TOC |
If a client has a write lock on a byte-range, it can request an atomic downgrade of the lock to a read lock via the LOCK request, by setting the type to READ_LT. If the server supports atomic downgrade, the request will succeed. If not, it will return NFS4ERR_LOCK_NOTSUPP. The client should be prepared to receive this error, and if appropriate, report the error to the requesting application.
If a client has a read lock on a byte-range, it can request an atomic upgrade of the lock to a write lock via the LOCK request by setting the type to WRITE_LT or WRITEW_LT. If the server does not support atomic upgrade, it will return NFS4ERR_LOCK_NOTSUPP. If the upgrade can be achieved without an existing conflict, the request will succeed. Otherwise, the server will return either NFS4ERR_DENIED or NFS4ERR_DEADLOCK. The error NFS4ERR_DEADLOCK is returned if the client sent the LOCK request with the type set to WRITEW_LT and the server has detected a deadlock. The client should be prepared to receive such errors and if appropriate, report the error to the requesting application.
| TOC |
When a lock or unlock request is done, passing a stateid, the stateid returned has the same "other" value and a "seqid" value that is incremented to reflect the occurrence of the lock or unlock request. The server MUST increment the value of the "seqid" field whenever there is any change to the locking status of any byte offset as described by any of locks covered by the stateid. A change in locking status includes a change from locked to unlocked or the reverse or a change from being locked for read to being locked for write or the reverse.
When there is no such change, as, for example when a range already locked for write is locked again for write, the server MAY increment the "seqid" value.
| TOC |
When the same file is opened by multiple open-owners, a client will have multiple open stateids for that file, each associated with a different open-owner. In that case, there can be multiple LOCK and LOCKU requests for the same lock-owner issued using the different open stateids, and so a situation may arise in which there are multiple stateids, each representing byte-range locks on the same file and held by the same lock-owner but each associated with a different open-owner.
In such a situation, the locking status of each byte (i.e. whether it is locked, the read or write mode of the lock and the lock-owner holding the lock) MUST reflect the last LOCK or LOCKU operation done for the lock-owner in question, independent of the stateid through which the request was issued.
When a byte is locked by the lock-owner in question, the open-owner to which that lock is assigned SHOULD be that of the open-owner associated with the stateid through which the last LOCK of that byte was done. When there is a change in the open-owner associated with locks for the stateid through which a LOCK or LOCKU was done, the "seqid" field of the stateid MUST be incremented, even if the locking, in terms of lock-owners has not changed. When there is a change to the set of locked bytes associated with a different stateid for the same lock-owner, i.e. associated with a different open-owner, the "seqid" value for that stateid MUST NOT be incremented.
| TOC |
Some clients require the support of blocking locks. While NFSv4.1 provides a callback when a previously unavailable lock becomes available, this is an OPTIONAL feature and clients cannot depend on its presence. Clients need to be prepared to continually poll for the lock. This presents a fairness problem. Two of the lock types, READW and WRITEW, are used to indicate to the server that the client is requesting a blocking lock. When the callback is not used, the server should maintain an ordered list of pending blocking locks. When the conflicting lock is released, the server may wait for the period of time equal to lease_time for the first waiting client to re-request the lock. After the lease period expires, the next waiting client request is allowed the lock. Clients are required to poll at an interval sufficiently small that it is likely to acquire the lock in a timely manner. The server is not required to maintain a list of pending blocked locks as it is used to increase fairness and not correct operation. Because of the unordered nature of crash recovery, storing of lock state to stable storage would be required to guarantee ordered granting of blocking locks.
Servers may also note the lock types and delay returning denial of the request to allow extra time for a conflicting lock to be released, allowing a successful return. In this way, clients can avoid the burden of needlessly frequent polling for blocking locks. The server should take care in the length of delay in the event the client retransmits the request.
If a server receives a blocking lock request, denies it, and then later receives a nonblocking request for the same lock, which is also denied, then it should remove the lock in question from its list of pending blocking locks. Clients should use such a nonblocking request to indicate to the server that this is the last time they intend to poll for the lock, as may happen when the process requesting the lock is interrupted. This is a courtesy to the server, to prevent it from unnecessarily waiting a lease period before granting other lock requests. However, clients are not required to perform this courtesy, and servers must not depend on them doing so. Also, clients must be prepared for the possibility that this final locking request will be accepted.
When server indicates, via the flag OPEN4_RESULT_MAY_NOTIFY_LOCK, that CB_NOTIFY_LOCK callbacks will be done for the current open file, the client should take notice of this, but, since this is a hint, cannot rely on a CB_NOTIFY_LOCK always being done. A client may reasonably reduce the frequency with which it polls for a denied lock, since the greater latency that might occur is likely to be eliminated given a prompt callback, but it still needs to poll. When it receives a CB_NOTIFY_LOCK it should promptly try to obtain the lock, but it should be aware that other clients may polling and the server is under no obligation to reserve the lock for that particular client.
| TOC |
A share reservation is a mechanism to control access to a file. It is a separate and independent mechanism from byte-range locking. When a client opens a file, it sends an OPEN operation to the server specifying the type of access required (READ, WRITE, or BOTH) and the type of access to deny others (deny NONE, READ, WRITE, or BOTH). If the OPEN fails the client will fail the application's open request.
Pseudo-code definition of the semantics:
if (request.access == 0) {
return (NFS4ERR_INVAL)
} else {
if ((request.access & file_state.deny)) ||
(request.deny & file_state.access)) {
return (NFS4ERR_DENIED)
}
return (NFS4ERR_OK);
When doing this checking of share reservations on OPEN, the current file_state used in the algorithm includes bits that reflect all current opens, including those for the open-owner making the new OPEN request.
The constants used for the OPEN and OPEN_DOWNGRADE operations for the access and deny fields are as follows:
const OPEN4_SHARE_ACCESS_READ = 0x00000001; const OPEN4_SHARE_ACCESS_WRITE = 0x00000002; const OPEN4_SHARE_ACCESS_BOTH = 0x00000003; const OPEN4_SHARE_DENY_NONE = 0x00000000; const OPEN4_SHARE_DENY_READ = 0x00000001; const OPEN4_SHARE_DENY_WRITE = 0x00000002; const OPEN4_SHARE_DENY_BOTH = 0x00000003;
| TOC |
To provide correct share semantics, a client MUST use the OPEN operation to obtain the initial filehandle and indicate the desired access and what access, if any, to deny. Even if the client intends to use a special stateid for anonymous state or read bypass, it must still obtain the filehandle for the regular file with the OPEN operation so the appropriate share semantics can be applied. For clients that do not have a deny mode built into their open programming interfaces, deny equal to NONE should be used.
The OPEN operation with the CREATE flag, also subsumes the CREATE operation for regular files as used in previous versions of the NFS protocol. This allows a create with a share to be done atomically.
The CLOSE operation removes all share reservations held by the open-owner on that file. If byte-range locks are held, the client SHOULD release all locks before issuing a CLOSE. The server MAY free all outstanding locks on CLOSE but some servers may not support the CLOSE of a file that still has byte-range locks held. The server MUST return failure, NFS4ERR_LOCKS_HELD, if any locks would exist after the CLOSE.
The LOOKUP operation will return a filehandle without establishing any lock state on the server. Without a valid stateid, the server will assume the client has the least access. For example, a file opened with deny READ/WRITE using a filehandle obtained through LOOKUP could only be read using the special read bypass stateid and could not be written at all because it would not have a valid stateid and the special anonymous stateid would not be allowed access.
| TOC |
When an OPEN is done for a file and the open-owner for which the open is being done already has the file open, the result is to upgrade the open file status maintained on the server to include the access and deny bits specified by the new OPEN as well as those for the existing OPEN. The result is that there is one open file, as far as the protocol is concerned, and it includes the union of the access and deny bits for all of the OPEN requests completed. The open is represented by s single stateid whose "other" values matches that of the original open, and whose "seqid" value is incremented to reflect the occurrence of the upgrade. The increment is required in cases in which the "upgrade" results in no change to the open mode (e.g. an OPEN is done for read when the existing open file is opened for read-write). Only a single CLOSE will be done to reset the effects of both OPENs. The client may use the stateid returned by the OPEN effecting the upgrade or with a stateid sharing the same "other" field and a seqid of zero, although care needs to be taken as far as upgrades which happen while the CLOSE is pending. Note that the client, when issuing the OPEN, may not know that the same file is in fact being opened. The above only applies if both OPENs result in the OPENed object being designated by the same filehandle.
When the server chooses to export multiple filehandles corresponding to the same file object and returns different filehandles on two different OPENs of the same file object, the server MUST NOT "OR" together the access and deny bits and coalesce the two open files. Instead the server must maintain separate OPENs with separate stateids and will require separate CLOSEs to free them.
When multiple open files on the client are merged into a single open file object on the server, the close of one of the open files (on the client) may necessitate change of the access and deny status of the open file on the server. This is because the union of the access and deny bits for the remaining opens may be smaller (i.e. a proper subset) than previously. The OPEN_DOWNGRADE operation is used to make the necessary change and the client should use it to update the server so that share reservation requests by other clients are handled properly. The stateid returned has the same "other" field as that passed to the server. The "seqid" value in the returned stateid MUST be incremented, even is situation in which there is no change the access and deny bits for the file.
| TOC |
Unlike the case of NFSv4.0, in which OPEN operations for the same open-owner are inherently serialized because of the owner-based seqid, multiple OPENs for the same open-owner may be done in parallel. When clients do this, they may encounter situations in which, because of the existence of hard links, two OPEN operations may turn out to open the same file, with a later OPEN performed being an upgrade of the first, with this fact only visible to the client once the operations complete.
In this situation, clients may determine the order in which the OPENs were performed by examining the stateids returned by the OPENs. Stateids that share a common value of the "other" field can be recognized as having opened the same file, with the order of the operations determinable from the order of the "seqid" fields, mod any possible wraparound of the 32-bit field.
When the possibility exists that the client will send multiple OPENs for the same open-owner in parallel, it may be the case that an open upgrade may happen without the client knowing beforehand that this could happen. Because of this possibility, CLOSEs and OPEN_DOWNGRADEs, should generally be sent with a non-zero seqid in the stateid, to avoid the possibility that the status change associated with an open upgrade is not inadvertently lost.
| TOC |
Special forms of the LOCK and OPEN operations are provided when it is necessary to re-establish byte-range locks or opens after a server failure.
Reclaims of opens associated with delegations are discussed in Section 10.2.1 (Delegation Recovery).
| TOC |
Client-side caching of data, of file attributes, and of file names is essential to providing good performance with the NFS protocol. Providing distributed cache coherence is a difficult problem and previous versions of the NFS protocol have not attempted it. Instead, several NFS client implementation techniques have been used to reduce the problems that a lack of coherence poses for users. These techniques have not been clearly defined by earlier protocol specifications and it is often unclear what is valid or invalid client behavior.
The NFSv4.1 protocol uses many techniques similar to those that have been used in previous protocol versions. The NFSv4.1 protocol does not provide distributed cache coherence. However, it defines a more limited set of caching guarantees to allow locks and share reservations to be used without destructive interference from client side caching.
In addition, the NFSv4.1 protocol introduces a delegation mechanism which allows many decisions normally made by the server to be made locally by clients. This mechanism provides efficient support of the common cases where sharing is infrequent or where sharing is read-only.
| TOC |
Caching techniques used in previous versions of the NFS protocol have been successful in providing good performance. However, several scalability challenges can arise when those techniques are used with very large numbers of clients. This is particularly true when clients are geographically distributed which classically increases the latency for cache revalidation requests.
The previous versions of the NFS protocol repeat their file data cache validation requests at the time the file is opened. This behavior can have serious performance drawbacks. A common case is one in which a file is only accessed by a single client. Therefore, sharing is infrequent.
In this case, repeated reference to the server to find that no conflicts exist is expensive. A better option with regards to performance is to allow a client that repeatedly opens a file to do so without reference to the server. This is done until potentially conflicting operations from another client actually occur.
A similar situation arises in connection with file locking. Sending file lock and unlock requests to the server as well as the read and write requests necessary to make data caching consistent with the locking semantics (see Section 10.3.2 (Data Caching and File Locking)) can severely limit performance. When locking is used to provide protection against infrequent conflicts, a large penalty is incurred. This penalty may discourage the use of file locking by applications.
The NFSv4.1 protocol provides more aggressive caching strategies with the following design goals:
The appropriate requirements for the server are discussed in later sections in which specific forms of caching are covered (see Section 10.4 (Open Delegation)).
| TOC |
Recallable delegation of server responsibilities for a file to a client improves performance by avoiding repeated requests to the server in the absence of inter-client conflict. With the use of a "callback" RPC from server to client, a server recalls delegated responsibilities when another client engages in sharing of a delegated file.
A delegation is passed from the server to the client, specifying the object of the delegation and the type of delegation. There are different types of delegations but each type contains a stateid to be used to represent the delegation when performing operations that depend on the delegation. This stateid is similar to those associated with locks and share reservations but differs in that the stateid for a delegation is associated with a client ID and may be used on behalf of all the open-owners for the given client. A delegation is made to the client as a whole and not to any specific process or thread of control within it.
The backchannel is established by CREATE_SESSION and BIND_CONN_TO_SESSION, and the client is required to maintain it. Because the backchannel may be down, even temporarily, correct protocol operation does not depend on them. Preliminary testing of backchannel functionality by means of a CB_COMPOUND procedure with a single operation, CB_SEQUENCE, can be used to check the continuity of the backchannel. A server avoids delegating responsibilities until it has determined that the backchannel exists. Because the granting of a delegation is always conditional upon the absence of conflicting access, clients must not assume that a delegation will be granted and they must always be prepared for OPENs, WANT_DELEGATIONs, and GET_DIR_DELEGATIONs to be processed without any delegations being granted.
Once granted, a delegation behaves in many ways like a lock. There is an associated lease that is subject to renewal together with all of the other leases held by that client.
Unlike locks, an operation by a second client to a delegated file will cause the server to recall a delegation through a callback. For individual operations, we will describe, under IMPLEMENTATION, when such operations are required to effect a recall. A number of points should be noted, however.
Despite those caveats, the implementation sections for a number of operations describe situations in which delegation recall would be required under some common circumstances:
On recall, the client holding the delegation must flush modified state (such as modified data) to the server and return the delegation. The conflicting request will not be acted on until the recall is complete. The recall is considered complete when the client returns the delegation or the server times its wait for the delegation to be returned and revokes the delegation as a result of the timeout. In the interim, the server will either delay responding to conflicting requests or respond to them with NFSERR_DELAY. Following the resolution of the recall, the server has the information necessary to grant or deny the second client's request.
At the time the client receives a delegation recall, it may have substantial state that needs to be flushed to the server. Therefore, the server should allow sufficient time for the delegation to be returned since it may involve numerous RPCs to the server. If the server is able to determine that the client is diligently flushing state to the server as a result of the recall, the server may extend the usual time allowed for a recall. However, the time allowed for recall completion should not be unbounded.
An example of this is when responsibility to mediate opens on a given file is delegated to a client (see Section 10.4 (Open Delegation)). The server will not know what opens are in effect on the client. Without this knowledge the server will be unable to determine if the access and deny state for the file allows any particular open until the delegation for the file has been returned.
A client failure or a network partition can result in failure to respond to a recall callback. In this case, the server will revoke the delegation which in turn will render useless any modified state still on the client.
| TOC |
There are three situations that delegation recovery must deal with:
In the event the client restarts, the failure to renew the lease will result in the revocation of byte-range locks and share reservations. Delegations, however, may be treated a bit differently.
There will be situations in which delegations will need to be reestablished after a client restarts. The reason for this is the client may have file data stored locally and this data was associated with the previously held delegations. The client will need to reestablish the appropriate file state on the server.
To allow for this type of client recovery, the server MAY extend the period for delegation recovery beyond the typical lease expiration period. This implies that requests from other clients that conflict with these delegations will need to wait. Because the normal recall process may require significant time for the client to flush changed state to the server, other clients need be prepared for delays that occur because of a conflicting delegation. This longer interval would increase the window for clients to restart and consult stable storage so that the delegations can be reclaimed. For open delegations, such delegations are reclaimed using OPEN with a claim type of CLAIM_DELEGATE_PREV or CLAIM_DELEG_PREV_FH (See Section 10.5 (Data Caching and Revocation) and Section 18.16 (Operation 18: OPEN - Open a Regular File) for discussion of open delegation and the details of OPEN respectively).
A server MAY support claim types of CLAIM_DELEGATE_PREV and CLAIM_DELEG_PREV_FH, and if it does, it MUST NOT remove delegations upon a CREATE_SESSION that confirms a client ID created by EXCHANGE_ID, and instead MUST, for a period of time no less than that of the value of the lease_time attribute, maintain the client's delegations to allow time for the client to send CLAIM_DELEGATE_PREV requests. The server that supports CLAIM_DELEGATE_PREV and/or CLAIM_DELEG_PREV_FH MUST support the DELEGPURGE operation.
When the server restarts, delegations are reclaimed (using the OPEN operation with CLAIM_PREVIOUS) in a similar fashion to byte-range locks and share reservations. However, there is a slight semantic difference. In the normal case if the server decides that a delegation should not be granted, it performs the requested action (e.g. OPEN) without granting any delegation. For reclaim, the server grants the delegation but a special designation is applied so that the client treats the delegation as having been granted but recalled by the server. Because of this, the client has the duty to write all modified state to the server and then return the delegation. This process of handling delegation reclaim reconciles three principles of the NFSv4.1 protocol:
When a client needs to reclaim a delegation and there is no associated open, the client may use the CLAIM_PREVIOUS variant of the WANT_DELEGATION operation. However, since the server is not required to support this operation, an alternative is to reclaim via a dummy open together with the delegation using an OPEN of type CLAIM_PREVIOUS. The dummy open file can be released using a CLOSE to re-establish the original state to be reclaimed, a delegation without an associated open.
When a client has more than a single open associated with a delegation, state for those additional opens can be established using OPEN operations of type CLAIM_DELEGATE_CUR. When these are used to establish opens associated with reclaimed delegations, the server MUST allow them when made within the grace period.
When a network partition occurs, delegations are subject to freeing by the server when the lease renewal period expires. This is similar to the behavior for locks and share reservations. For delegations, however, the server may extend the period in which conflicting requests are held off. Eventually the occurrence of a conflicting request from another client will cause revocation of the delegation. A loss of the backchannel (e.g. by later network configuration change) will have the same effect. A recall request will fail and revocation of the delegation will result.
A client normally finds out about revocation of a delegation when it uses a stateid associated with a delegation and receives one of the errors NFS4EER_EXPIRED, NFS4ERR_ADMIN_REVOKED, or NFS4ERR_DELEG_REVOKED. It also may find out about delegation revocation after a client restart when it attempts to reclaim a delegation and receives that same error. Note that in the case of a revoked write open delegation, there are issues because data may have been modified by the client whose delegation is revoked and separately by other clients. See Section 10.5.1 (Revocation Recovery for Write Open Delegation) for a discussion of such issues. Note also that when delegations are revoked, information about the revoked delegation will be written by the server to stable storage (as described in Section 8.4.3 (Network Partitions and Recovery)). This is done to deal with the case in which a server restarts after revoking a delegation but before the client holding the revoked delegation is notified about the revocation.
| TOC |
When applications share access to a set of files, they need to be implemented so as to take account of the possibility of conflicting access by another application. This is true whether the applications in question execute on different clients or reside on the same client.
Share reservations and byte-range locks are the facilities the NFSv4.1 protocol provides to allow applications to coordinate access by using mutual exclusion facilities. The NFSv4.1 protocol's data caching must be implemented such that it does not invalidate the assumptions that those using these facilities depend upon.
| TOC |
In order to avoid invalidating the sharing assumptions that applications rely on, NFSv4.1 clients should not provide cached data to applications or modify it on behalf of an application when it would not be valid to obtain or modify that same data via a READ or WRITE operation.
Furthermore, in the absence of open delegation (see Section 10.4 (Open Delegation)), two additional rules apply. Note that these rules are obeyed in practice by many NFSv3 clients.
| TOC |
For those applications that choose to use file locking instead of share reservations to exclude inconsistent file access, there is an analogous set of constraints that apply to client side data caching. These rules are effective only if the file locking is used in a way that matches in an equivalent way the actual READ and WRITE operations executed. This is as opposed to file locking that is based on pure convention. For example, it is possible to manipulate a two-megabyte file by dividing the file into two one-megabyte regions and protecting access to the two regions by file locks on bytes zero and one. A lock for write on byte zero of the file would represent the right to do READ and WRITE operations on the first region. A lock for write on byte one of the file would represent the right to do READ and WRITE operations on the second region. As long as all applications manipulating the file obey this convention, they will work on a local file system. However, they may not work with the NFSv4.1 protocol unless clients refrain from data caching.
The rules for data caching in the file locking environment are:
Note that flushing data to the server and the invalidation of cached data must reflect the actual byte ranges locked or unlocked. Rounding these up or down to reflect client cache block boundaries will cause problems if not carefully done. For example, writing a modified block when only half of that block is within an area being unlocked may cause invalid modification to the region outside the unlocked area. This, in turn, may be part of a region locked by another client. Clients can avoid this situation by synchronously performing portions of write operations that overlap that portion (initial or final) that is not a full block. Similarly, invalidating a locked area which is not an integral number of full buffer blocks would require the client to read one or two partial blocks from the server if the revalidation procedure shows that the data which the client possesses may not be valid.
The data that is written to the server as a prerequisite to the unlocking of a region must be written, at the server, to stable storage. The client may accomplish this either with synchronous writes or by following asynchronous writes with a COMMIT operation. This is required because retransmission of the modified data after a server restart might conflict with a lock held by another client.
A client implementation may choose to accommodate applications which use byte-range locking in non-standard ways (e.g. using a byte-range lock as a global semaphore) by flushing to the server more data upon an LOCKU than is covered by the locked range. This may include modified data within files other than the one for which the unlocks are being done. In such cases, the client must not interfere with applications whose READs and WRITEs are being done only within the bounds of byte-range locks which the application holds. For example, an application locks a single byte of a file and proceeds to write that single byte. A client that chose to handle a LOCKU by flushing all modified data to the server could validly write that single byte in response to an unrelated unlock. However, it would not be valid to write the entire block in which that single written byte was located since it includes an area that is not locked and might be locked by another client. Client implementations can avoid this problem by dividing files with modified data into those for which all modifications are done to areas covered by an appropriate byte-range lock and those for which there are modifications not covered by a byte-range lock. Any writes done for the former class of files must not include areas not locked and thus not modified on the client.
| TOC |
Client side data caching needs to respect mandatory file locking when it is in effect. The presence of mandatory file locking for a given file is indicated when the client gets back NFS4ERR_LOCKED from a READ or WRITE on a file it has an appropriate share reservation for. When mandatory locking is in effect for a file, the client must check for an appropriate file lock for data being read or written. If a lock exists for the range being read or written, the client may satisfy the request using the client's validated cache. If an appropriate file lock is not held for the range of the read or write, the read or write request must not be satisfied by the client's cache and the request must be sent to the server for processing. When a read or write request partially overlaps a locked region, the request should be subdivided into multiple pieces with each region (locked or not) treated appropriately.
| TOC |
When clients cache data, the file data needs to be organized according to the file system object to which the data belongs. For NFSv3 clients, the typical practice has been to assume for the purpose of caching that distinct filehandles represent distinct file system objects. The client then has the choice to organize and maintain the data cache on this basis.
In the NFSv4.1 protocol, there is now the possibility to have significant deviations from a "one filehandle per object" model because a filehandle may be constructed on the basis of the object's pathname. Therefore, clients need a reliable method to determine if two filehandles designate the same file system object. If clients were simply to assume that all distinct filehandles denote distinct objects and proceed to do data caching on this basis, caching inconsistencies would arise between the distinct client side objects which mapped to the same server side object.
By providing a method to differentiate filehandles, the NFSv4.1 protocol alleviates a potential functional regression in comparison with the NFSv3 protocol. Without this method, caching inconsistencies within the same client could occur and this has not been present in previous versions of the NFS protocol. Note that it is possible to have such inconsistencies with applications executing on multiple clients but that is not the issue being addressed here.
For the purposes of data caching, the following steps allow an NFSv4.1 client to determine whether two distinct filehandles denote the same server side object:
| TOC |
When a file is being OPENed, the server may delegate further handling of opens and closes for that file to the opening client. Any such delegation is recallable, since the circumstances that allowed for the delegation are subject to change. In particular, the server may receive a conflicting OPEN from another client, the server must recall the delegation before deciding whether the OPEN from the other client may be granted. Making a delegation is up to the server and clients should not assume that any particular OPEN either will or will not result in an open delegation. The following is a typical set of conditions that servers might use in deciding whether OPEN should be delegated:
There are two types of open delegations, read and write. A read open delegation allows a client to handle, on its own, requests to open a file for reading that do not deny read access to others. Multiple read open delegations may be outstanding simultaneously and do not conflict. A write open delegation allows the client to handle, on its own, all opens. Only one write open delegation may exist for a given file at a given time and it is inconsistent with any read open delegations.
When a client has a read open delegation, it is assured that neither the contents, the attributes (with the exception of time_access), nor the names of any links to the file will change without its knowledge, so long as the delegation is held. When a client has a write open delegation, it may modify the file data locally since no other client will be accessing the file's data. The client holding a write delegation may only locally affect file attributes which are intimately connected with the file data: size, change, time_access, time_metadata, and time_modify. All other attributes must be reflected on the server.
When a client has an open delegation, it does not need to send OPENs or CLOSEs to the server. Instead the client may update the appropriate status internally. For a read open delegation, opens that cannot be handled locally (opens for write or that deny read access) must be sent to the server.
When an open delegation is made, the reply to the OPEN contains an open delegation structure which specifies the following:
The delegation stateid is separate and distinct from the stateid for the OPEN proper. The standard stateid, unlike the delegation stateid, is associated with a particular lock-owner and will continue to be valid after the delegation is recalled and the file remains open.
When a request internal to the client is made to open a file and an open delegation is in effect, it will be accepted or rejected solely on the basis of the following conditions. Any requirement for other checks to be made by the delegate should result in open delegation being denied so that the checks can be made by the server itself.
The nfsace4 passed with delegation can be used to avoid frequent ACCESS calls. The permission check should be as follows:
The server may return an nfsace4 that is more restrictive than the actual ACL of the file. This includes an nfsace4 that specifies denial of all access. Note that some common practices such as mapping the traditional user "root" to the user "nobody" may make it incorrect to return the actual ACL of the file in the delegation response.
The use of a delegation together with various other forms of caching creates the possibility that no server authentication and authorization will ever be performed for a given user since all of the user's requests might be satisfied locally. Where the client is depending on the server for authentication and authorization, the client should be sure authentication and authorization occurs for each user by use of the ACCESS operation. This should be the case even if an ACCESS operation would not be required otherwise. As mentioned before, the server may enforce frequent authentication by returning an nfsace4 denying all access with every open delegation.
| TOC |
An OPEN delegation allows much of the message overhead associated with the opening and closing files to be eliminated. An open when an open delegation is in effect does not require that a validation message be sent to the server. The continued endurance of the "read open delegation" provides a guarantee that no OPEN for write and thus no write has occurred. Similarly, when closing a file opened for write and if write open delegation is in effect, the data written does not have to be written to the server until the open delegation is recalled. The continued endurance of the open delegation provides a guarantee that no open and thus no read or write has been done by another client.
For the purposes of open delegation, READs and WRITEs done without an OPEN are treated as the functional equivalents of a corresponding type of OPEN. Although client SHOULD NOT use special stateids when an open exists, delegation handling on the server can use the client ID associated with the current session to determine if the operation has been done by the holder of the delegation, in which case, no recall is necessary, or by another client, in which case the delegation must be recalled and I/O not proceed until the delegation is recalled or revoked.
With delegations, a client is able to avoid writing data to the server when the CLOSE of a file is serviced. The file close system call is the usual point at which the client is notified of a lack of stable storage for the modified file data generated by the application. At the close, file data is written to the server and through normal accounting the server is able to determine if the available file system space for the data has been exceeded (i.e. server returns NFS4ERR_NOSPC or NFS4ERR_DQUOT). This accounting includes quotas. The introduction of delegations requires that a alternative method be in place for the same type of communication to occur between client and server.
In the delegation response, the server provides either the limit of the size of the file or the number of modified blocks and associated block size. The server must ensure that the client will be able to write modified data to the server of a size equal to that provided in the original delegation. The server must make this assurance for all outstanding delegations. Therefore, the server must be careful in its management of available space for new or modified data taking into account available file system space and any applicable quotas. The server can recall delegations as a result of managing the available file system space. The client should abide by the server's state space limits for delegations. If the client exceeds the stated limits for the delegation, the server's behavior is undefined.
Based on server conditions, quotas or available file system space, the server may grant write open delegations with very restrictive space limitations. The limitations may be defined in a way that will always force modified data to be flushed to the server on close.
With respect to authentication, flushing modified data to the server after a CLOSE has occurred may be problematic. For example, the user of the application may have logged off the client and unexpired authentication credentials may not be present. In this case, the client may need to take special care to ensure that local unexpired credentials will in fact be available. This may be accomplished by tracking the expiration time of credentials and flushing data well in advance of their expiration or by making private copies of credentials to assure their availability when needed.
| TOC |
When a client holds a write open delegation, lock operations are performed locally. This includes those required for mandatory file locking. This can be done since the delegation implies that there can be no conflicting locks. Similarly, all of the revalidations that would normally be associated with obtaining locks and the flushing of data associated with the releasing of locks need not be done.
When a client holds a read open delegation, lock operations are not performed locally. All lock operations, including those requesting non-exclusive locks, are sent to the server for resolution.
| TOC |
The server needs to employ special handling for a GETATTR where the target is a file that has a write open delegation in effect. The reason for this is that the client holding the write delegation may have modified the data and the server needs to reflect this change to the second client that submitted the GETATTR. Therefore, the client holding the write delegation needs to be interrogated. The server will use the CB_GETATTR operation. The only attributes that the server can reliably query via CB_GETATTR are size and change.
Since CB_GETATTR is being used to satisfy another client's GETATTR request, the server only needs to know if the client holding the delegation has a modified version of the file. If the client's copy of the delegated file is not modified (data or size), the server can satisfy the second client's GETATTR request from the attributes stored locally at the server. If the file is modified, the server only needs to know about this modified state. If the server determines that the file is currently modified, it will respond to the second client's GETATTR as if the file had been modified locally at the server.
Since the form of the change attribute is determined by the server and is opaque to the client, the client and server need to agree on a method of communicating the modified state of the file. For the size attribute, the client will report its current view of the file size. For the change attribute, the handling is more involved.
For the client, the following steps will be taken when receiving a write delegation:
For simplicity of implementation, the client MAY for each CB_GETATTR return the same value d. This is true even if, between successive CB_GETATTR operations, the client again modifies in the file's data or metadata in its cache. The client can return the same value because the only requirement is that the client be able to indicate to the server that the client holds modified data. Therefore, the value of d may always be c + 1.
While the change attribute is opaque to the client in the sense that it has no idea what units of time, if any, the server is counting change with, it is not opaque in that the client has to treat it as an unsigned integer, and the server has to be able to see the results of the client's changes to that integer. Therefore, the server MUST encode the change attribute in network order when sending it to the client. The client MUST decode it from network order to its native order when receiving it and the client MUST encode it network order when sending it to the server. For this reason, change is defined as an unsigned integer rather than an opaque array of bytes.
For the server, the following steps will be taken when providing a write delegation:
As discussed earlier in this section, the client MAY return the same cc value on subsequent CB_GETATTR calls, even if the file was modified in the client's cache yet again between successive CB_GETATTR calls. Therefore, the server must assume that the file has been modified yet again, and MUST take care to ensure that the new nsc it constructs and returns is greater than the previous nsc it returned. An example implementation's delegation record would satisfy this mandate by including a boolean field (let us call it "modified") that is set to FALSE when the delegation is granted, and an sc value set at the time of grant to the change attribute value. The modified field would be set to TRUE the first time cc != sc, and would stay TRUE until the delegation is returned or revoked. The processing for constructing nsc, time_modify, and time_metadata would use this pseudo code:
if (!modified) {
do CB_GETATTR for change and size;
if (cc != sc)
modified = TRUE;
} else {
do CB_GETATTR for size;
}
if (modified) {
sc = sc + 1;
time_modify = time_metadata = current_time;
update sc, time_modify, time_metadata into file's metadata;
}
This would return to the client (that sent GETATTR) the attributes it requested, but make sure size comes from what CB_GETATTR returned. The server would not update the file's metadata with the client's modified size.
In the case that the file attribute size is different than the server's current value, the server treats this as a modification regardless of the value of the change attribute retrieved via CB_GETATTR and responds to the second client as in the last step.
This methodology resolves issues of clock differences between client and server and other scenarios where the use of CB_GETATTR break down.
It should be noted that the server is under no obligation to use CB_GETATTR and therefore the server MAY simply recall the delegation to avoid its use.
| TOC |
The following events necessitate recall of an open delegation:
Whether a RENAME of a directory in the path leading to the file results in recall of an open delegation depends on the semantics of the server's file system. If that file system denies such RENAMEs when a file is open, the recall must be performed to determine whether the file in question is, in fact, open.
In addition to the situations above, the server may choose to recall open delegations at any time if resource constraints make it advisable to do so. Clients should always be prepared for the possibility of recall.
When a client receives a recall for an open delegation, it needs to update state on the server before returning the delegation. These same updates must be done whenever a client chooses to return a delegation voluntarily. The following items of state need to be dealt with:
In the case of write open delegation, file locking imposes some additional requirements. To precisely maintain the associated invariant, it is required to flush any modified data in any region for which a write lock was released while the write delegation was in effect. However, because the write open delegation implies no other locking by other clients, a simpler implementation is to flush all modified data for the file (as described just above) if any write lock has been released while the write open delegation was in effect.
An implementation need not wait until delegation recall (or deciding to voluntarily return a delegation) to perform any of the above actions, if implementation considerations (e.g. resource availability constraints) make that desirable. Generally, however, the fact that the actual open state of the file may continue to change makes it not worthwhile to send information about opens and closes to the server, except as part of delegation return. Only in the case of closing the open that resulted in obtaining the delegation would clients be likely to do this early, since, in that case, the close once done will not be undone. Regardless of the client's choices on scheduling these actions, all must be performed before the delegation is returned, including (when applicable) the close that corresponds to the open that resulted in the delegation. These actions can be performed either in previous requests or in previous operations in the same COMPOUND request.
| TOC |
A client may fail to respond to a recall for various reasons, such as a failure of the backchannel from server to the client. The client may be unaware of a failure in the backchannel. This lack of awareness could result in the client finding out long after the failure that its delegation has been revoked, and another client has modified the data for which the client had a delegation. This is especially a problem for the client that held a write delegation.
Status bits returned by SEQUENCE operations help to provide an alternate way of informing the client of issues regarding the status of the backchannel and of recalled delegations. When the backchannel is not available, the server returns the status bit SEQ4_STATUS_CB_PATH_DOWN on SEQUENCE operations. The client can react by attempting to re-establish the backchannel and by returning recallable objects if a backchannel cannot be successfully re-established.
Whether the backchannel is functioning or not, it may be that the recalled delegation is not returned. Note that the client's lease might still be renewed, even though the recalled delegation is not returned. In this situation, servers SHOULD revoke delegations that are not returned in a period of time equal to the lease period. This period of time should allow the client time to note the backchannel-down status and re-establish the backchannel.
When delegations are revoked, the server will return with the SEQ4_STATUS_RECALLABLE_STATE_REVOKED status bit set on subsequent SEQUENCE operations. The client should note this and then use TEST_STATEID to find which delegations have been recalled.
| TOC |
At the point a delegation is revoked, if there are associated opens on the client, these opens may or may not be revoked. If no lock or open is granted that is inconsistent with the existing open, the stateid for the open may remain valid, and be disconnected from the revoked delegation, just as would be the case if the delegation were returned.
For example, if an OPEN for read-write with DENY=NONE is associated with the delegation, granting of another such OPEN to a different client will revoke the delegation but need not revoke the OPEN, since no lock inconsistent with that OPEN has been granted. On the other hand, if an OPEN denying write is granted, then the existing open must be revoked.
When opens and/or locks are revoked, the applications holding these opens or locks need to be notified. This notification usually occurs by returning errors for READ/WRITE operations or when a close is attempted for the open file.
If no opens exist for the file at the point the delegation is revoked, then notification of the revocation is unnecessary. However, if there is modified data present at the client for the file, the user of the application should be notified. Unfortunately, it may not be possible to notify the user since active applications may not be present at the client. See Section 10.5.1 (Revocation Recovery for Write Open Delegation) for additional details.
| TOC |
In addition to providing delegations as part of the reply to OPEN operations, servers MAY provide delegations separate from open, via the OPTIONAL WANT_DELEGATION operation. This allows delegations to be obtained in advance of an OPEN that might benefit from them, for objects which are not a valid target of OPEN, or to deal with cases in which a delegation has been recalled and the client wants to make an attempt to re-establish it if the absence of use by other clients allows that.
The WANT_DELEGATION operation may be performed on any type of file object other than a directory.
When a delegation is obtained using WANT_DELEGATION, any open files for the same filehandle held by that client are to be treated as subordinate to the delegation, just as if they had been created using an OPEN of type CLAIM_DELEGATE_CUR. They are otherwise unchanged as to seqid, access and deny modes, and the relationship with byte-range locks. Similarly, existing byte-range locks subordinate to an open which becomes subordinate to a delegation, become indirectly subordinate to that new delegation.
The WANT_DELEGATION operation provides for delivery of delegations via callbacks, when the delegations are not immediately available. When a requested delegation is available, it is delivered to the client via a CB_PUSH_DELEG operation. When this happens, open files for the same filehandle become subordinate to the new delegation at the point at which the delegation is delivered , just as if they had been created using an OPEN of type CLAIM_DELEGATE_CUR. Similarly, for existing byte-range locks subordinate to an open.
| TOC |
When locks and delegations are revoked, the assumptions upon which successful caching depend are no longer guaranteed. For any locks or share reservations that have been revoked, the corresponding state-owner needs to be notified. This notification includes applications with a file open that has a corresponding delegation which has been revoked. Cached data associated with the revocation must be removed from the client. In the case of modified data existing in the client's cache, that data must be removed from the client without it being written to the server. As mentioned, the assumptions made by the client are no longer valid at the point when a lock or delegation has been revoked. For example, another client may have been granted a conflicting lock after the revocation of the lock at the first client. Therefore, the data within the lock range may have been modified by the other client. Obviously, the first client is unable to guarantee to the application what has occurred to the file in the case of revocation.
Notification to a state-owner will in many cases consist of simply returning an error on the next and all subsequent READs/WRITEs to the open file or on the close. Where the methods available to a client make such notification impossible because errors for certain operations may not be returned, more drastic action such as signals or process termination may be appropriate. The justification for this is that an invariant for which an application depends on may be violated. Depending on how errors are typically treated for the client operating environment, further levels of notification including logging, console messages, and GUI pop-ups may be appropriate.
| TOC |
Revocation recovery for a write open delegation poses the special issue of modified data in the client cache while the file is not open. In this situation, any client which does not flush modified data to the server on each close must ensure that the user receives appropriate notification of the failure as a result of the revocation. Since such situations may require human action to correct problems, notification schemes in which the appropriate user or administrator is notified may be necessary. Logging and console messages are typical examples.
If there is modified data on the client, it must not be flushed normally to the server. A client may attempt to provide a copy of the file data as modified during the delegation under a different name in the file system name space to ease recovery. Note that when the client can determine that the file has not been modified by any other client, or when the client has a complete cached copy of file in question, such a saved copy of the client's view of the file may be of particular value for recovery. In other case, recovery using a copy of the file based partially on the client's cached data and partially on the server copy as modified by other clients, will be anything but straightforward, so clients may avoid saving file contents in these situations or mark the results specially to warn users of possible problems.
Saving of such modified data in delegation revocation situations may be limited to files of a certain size or might be used only when sufficient disk space is available within the target file system. Such saving may also be restricted to situations when the client has sufficient buffering resources to keep the cached copy available until it is properly stored to the target file system.
| TOC |
This section pertains to the caching of a file's attributes on a client when that client does not hold a delegation on the file.
The attributes discussed in this section do not include named attributes. Individual named attributes are analogous to files and caching of the data for these needs to be handled just as data caching is for ordinary files. Similarly, LOOKUP results from an OPENATTR directory are to be cached on the same basis as any other pathnames and similarly for directory contents.
Clients may cache file attributes obtained from the server and use them to avoid subsequent GETATTR requests. Such caching is write through in that modification to file attributes is always done by means of requests to the server and should not be done locally and cached. The exception to this are modifications to attributes that are intimately connected with data caching. Therefore, extending a file by writing data to the local data cache is reflected immediately in the size as seen on the client without this change being immediately reflected on the server. Normally such changes are not propagated directly to the server but when the modified data is flushed to the server, analogous attribute changes are made on the server. When open delegation is in effect, the modified attributes may be returned to the server in reaction to a CB_RECALL call.
The result of local caching of attributes is that the attribute caches maintained on individual clients will not be coherent. Changes made in one order on the server may be seen in a different order on one client and in a third order on a different client.
The typical file system application programming interfaces do not provide means to atomically modify or interrogate attributes for multiple files at the same time. The following rules provide an environment where the potential incoherences mentioned above can be reasonably managed. These rules are derived from the practice of previous NFS protocols.
Note that if the full set of attributes to be cached is requested by READDIR, the results can be cached by the client on the same basis as attributes obtained via GETATTR.
A client may validate its cached version of attributes for a file by fetching just both the change and time_access attributes and assuming that if the change attribute has the same value as it did when the attributes were cached, then no attributes other than time_access have changed. The reason why time_access is also fetched is because many servers operate in environments where the operation that updates change does not update time_access. For example, POSIX file semantics do not update access time when a file is modified by the write system call. Therefore, the client that wants a current time_access value should fetch it with change during the attribute cache validation processing and update its cached time_access.
The client may maintain a cache of modified attributes for those attributes intimately connected with data of modified regular files (size, time_modify, and change). Other than those three attributes, the client MUST NOT maintain a cache of modified attributes. Instead, attribute changes are immediately sent to the server.
In some operating environments, the equivalent to time_access is expected to be implicitly updated by each read of the content of the file object. If an NFS client is caching the content of a file object, whether it is a regular file, directory, or symbolic link, the client SHOULD NOT update the time_access attribute (via SETATTR or a small READ or READDIR request) on the server with each read that is satisfied from cache. The reason is that this can defeat the performance benefits of caching content, especially since an explicit SETATTR of time_access may alter the change attribute on the server. If the change attribute changes, clients that are caching the content will think the content has changed, and will re-read unmodified data from the server. Nor is the client encouraged to maintain a modified version of time_access in its cache, since this would mean that the client will either eventually have to write the access time to the server with bad performance effects, or it would never update the server's time_access, thereby resulting in a situation where an application that caches access time between a close and open of the same file observes the access time oscillating between the past and present. The time_access attribute always means the time of last access to a file by a read that was satisfied by the server. This way clients will tend to see only time_access changes that go forward in time.
| TOC |
Some operating environments include the capability for an application to map a file's content into the application's address space. Each time the application accesses a memory location that corresponds to a block that has not been loaded into the address space, a page fault occurs and the file is read (or if the block does not exist in the file, the block is allocated and then instantiated in the application's address space).
As long as each memory mapped access to the file requires a page fault, the relevant attributes of the file that are used to detect access and modification (time_access, time_metadata, time_modify, and change) will be updated. However, in many operating environments, when page faults are not required these attributes will not be updated on reads or updates to the file via memory access (regardless whether the file is local file or is being access remotely). A client or server MAY fail to update attributes of a file that is being accessed via memory mapped I/O. This has several implications:
Here the challenge is for each client to resynchronize to get a correct view of the first page. In many operating environments, the virtual memory management systems on each client only know a page is modified, not that a subset of the page corresponding to the respective lock regions has been modified. So it is not possible for each client to do the right thing, which is to only write to the server that portion of the page that is locked. For example, if client A simply writes out the page, and then client B writes out the page, client A's data is lost.
Moreover, if mandatory locking is enabled on the file, then we have a different problem. When clients A and B execute the STORE instructions, the resulting page faults require a byte-range lock on the entire page. Each client then tries to extend their locked range to the entire page, which results in a deadlock. Communicating the NFS4ERR_DEADLOCK error to a STORE instruction is difficult at best.
If a client is locking the entire memory mapped file, there is no problem with advisory or mandatory byte-range locking, at least until the client unlocks a region in the middle of the file.
Given the above issues the following are permitted:
| TOC |
The NFSv4.1 directory delegation facility (described in Section 10.9 (Directory Delegations) below) is OPTIONAL for servers to implement. Even where it is implemented, it may not be always be functional because of resource availability issues or other constraints. Thus, it is important to understand how name and directory caching are done in the absence of directory delegations. Those topics are discussed in the next in Section 10.8.1 (Name Caching).
| TOC |
The results of LOOKUP and READDIR operations may be cached to avoid the cost of subsequent LOOKUP operations. Just as in the case of attribute caching, inconsistencies may arise among the various client caches. To mitigate the effects of these inconsistencies and given the context of typical file system APIs, an upper time boundary is maintained on how long a client name cache entry can be kept without verifying that the entry has not been made invalid by a directory change operation performed by another client.
When a client is not making changes to a directory for which there exist name cache entries, the client needs to periodically fetch attributes for that directory to ensure that it is not being modified. After determining that no modification has occurred, the expiration time for the associated name cache entries may be updated to be the current time plus the name cache staleness bound.
When a client is making changes to a given directory, it needs to determine whether there have been changes made to the directory by other clients. It does this by using the change attribute as reported before and after the directory operation in the associated change_info4 value returned for the operation. The server is able to communicate to the client whether the change_info4 data is provided atomically with respect to the directory operation. If the change values are provided atomically, the client has a basis for determining, given proper care, whether other clients are modifying the directory is question.
The simplest way to enable the client to make this determination is for the client to serialize all changes made to a specific directory. When this is done, and the server provides before and after values of the change attribute atomically, the client can simply compare the after value of the change attribute from one operation on a directory with the before value on the next subsequent operation modifying that directory. When these are equal, the client is assured that no other client is modifying the directory in question.
When such serialization is not used, and there may be multiple simultaneous outstanding operations modifying a single directory sent from a single client, making this sort of determination can be more complicated, since two such operations which are recognized as complete in a different order than they were actually performed, might give an appearance consistent with modification being made by another client. Where this appears to happen, the client needs to await the completion of all such modifications that were started previously, to see if the outstanding before and after change numbers can be sorted into a chain such that the before value of one change number matches the after value of a previous one, in a chain consistent with this client being the only one modifying the directory.
In either of these cases, the client is able to determine whether the directory is being modified by another client. If the comparison indicates that the directory was updated by another client, the name cache associated with the modified directory is purged from the client. If the comparison indicates no modification, the name cache can be updated on the client to reflect the directory operation and the associated timeout extended. The post-operation change value needs to be saved as the basis for future change_info4 comparisons.
As demonstrated by the scenario above, name caching requires that the client revalidate name cache data by inspecting the change attribute of a directory at the point when the name cache item was cached. This requires that the server update the change attribute for directories when the contents of the corresponding directory is modified. For a client to use the change_info4 information appropriately and correctly, the server must report the pre and post operation change attribute values atomically. When the server is unable to report the before and after values atomically with respect to the directory operation, the server must indicate that fact in the change_info4 return value. When the information is not atomically reported, the client should not assume that other clients have not changed the directory.
| TOC |
The results of READDIR operations may be used to avoid subsequent READDIR operations. Just as in the cases of attribute and name caching, inconsistencies may arise among the various client caches. To mitigate the effects of these inconsistencies, and given the context of typical file system APIs, the following rules should be followed:
The revalidation technique parallels that discussed in the case of name caching. When the client is not changing the directory in question, checking the change attribute of the directory with GETATTR is adequate. The lifetime of the cache entry can be extended at these checkpoints. When a client is modifying the directory, the client needs to use the change_info4 data to determine whether there are other clients modifying the directory. If it is determined that no other client modifications are occurring, the client may update its directory cache to reflect its own changes.
As demonstrated previously, directory caching requires that the client revalidate directory cache data by inspecting the change attribute of a directory at the point when the directory was cached. This requires that the server update the change attribute for directories when the contents of the corresponding directory is modified. For a client to use the change_info4 information appropriately and correctly, the server must report the pre and post operation change attribute values atomically. When the server is unable to report the before and after values atomically with respect to the directory operation, the server must indicate that fact in the change_info4 return value. When the information is not atomically reported, the client should not assume that other clients have not changed the directory.
| TOC |
| TOC |
Directory caching for the NFSv4.1 protocol, as previously described, is similar to file caching in previous versions. Clients typically cache directory information for a duration determined by the client. At the end of a predefined timeout, the client will query the server to see if the directory has been updated. By caching attributes, clients reduce the number of GETATTR calls made to the server to validate attributes. Furthermore, frequently accessed files and directories, such as the current working directory, have their attributes cached on the client so that some NFS operations can be performed without having to make an RPC call. By caching name and inode information about most recently looked up entries in a Directory Name Lookup Cache (DNLC), clients do not need to send LOOKUP calls to the server every time these files are accessed.
This caching approach works reasonably well at reducing network traffic in many environments. However, it does not address environments where there are numerous queries for files that do not exist. In these cases of "misses", the client sends requests to the server in order to provide reasonable application semantics and promptly detect the creation of new directory entries. Examples of high miss activity are compilation in software development environments. The current behavior of NFS limits its potential scalability and wide-area sharing effectiveness in these types of environments. Other distributed stateful file system architectures such as AFS and DFS have proven that adding state around directory contents can greatly reduce network traffic in high-miss environments.
Delegation of directory contents is an OPTIONAL feature of NFSv4.1. Directory delegations provide similar traffic reduction benefits as with file delegations. By allowing clients to cache directory contents (in a read-only fashion) while being notified of changes, the client can avoid making frequent requests to interrogate the contents of slowly-changing directories, reducing network traffic and improving client performance. It can also simplify the task of determining whether other clients are making changes to the directory when the client itself is making many changes to the directory and changes are not serialized.
Directory delegations allow improved namespace cache consistency to be achieved through delegations and synchronous recalls, in the absence of notifications. In addition, if time-based consistency is sufficient, asynchronous notifications can provide performance benefits for the client, and possibly the server, under some common operating conditions such as slowly-changing and/or very large directories.
| TOC |
NFSv4.1 introduces the GET_DIR_DELEGATION (Section 18.39 (Operation 46: GET_DIR_DELEGATION - Get a directory delegation)) operation to allow the client to ask for a directory delegation. The delegation covers directory attributes and all entries in the directory. If either of these change, the delegation will be recalled synchronously. The operation causing the recall will have to wait before the recall is complete. Any changes to directory entry attributes will not cause the delegation to be recalled.
In addition to asking for delegations, a client can also ask for notifications for certain events. These events include changes to the directory's attributes and/or its contents. If a client asks for notification for a certain event, the server will notify the client when that event occurs. This will not result in the delegation being recalled for that client. The notifications are asynchronous and provide a way of avoiding recalls in situations where a directory is changing enough that the pure recall model may not be effective while trying to allow the client to get substantial benefit. In the absence of notifications, once the delegation is recalled the client has to refresh its directory cache which might not be very efficient for very large directories.
The delegation is read-only and the client may not make changes to the directory other than by performing NFSv4.1 operations that modify the directory or the associated file attributes so that the server has knowledge of these changes. In order to keep the client namespace synchronized with the server, the server will, if the client has requested notifications, notify the client holding the delegation of the changes made as a result. This is to avoid any need for subsequent GETATTR or READDIR calls to the server. If a single client is holding the delegation and that client makes any changes to the directory (i.e. the changes are made via operations sent though a session associated with the client ID holding the delegation), the delegation will not be recalled. Multiple clients may hold a delegation on the same directory, but if any such client modifies the directory, the server MUST recall the delegation from the other clients, unless those clients have made provisions to be notified of that sort of modification.
Delegations can be recalled by the server at any time. Normally, the server will recall the delegation when the directory changes in a way that is not covered by the notification, or when the directory changes and notifications have not been requested. If another client removes the directory for which a delegation has been granted, the server will recall the delegation.
| TOC |
See Section 5.11 (Directory Notification Attributes) for a description of the attributes associated with directory notifications.
| TOC |
The server will recall the directory delegation by sending a callback to the client. It will use the same callback procedure as used for recalling file delegations. The server will recall the delegation when the directory changes in a way that is not covered by the notification. However the server need not recall the delegation if attributes of an entry within the directory change.
If the server notices that handing out a delegation for a directory is causing too many notifications to be sent out, it may decide not to hand out delegations for that directory, or recall those already granted. If a client tries to remove the directory for which a delegation has been granted, the server will recall all associated delegations.
The implementation sections for a number of operations describe situations in which notification or delegation recall would be required under some common circumstances. In this regard, a similar set of caveats to those listed in Section 10.2 (Delegation and Callbacks) apply.
| TOC |
Recovery from client or server restart for state on regular files has two main goals, avoiding the necessity of breaking application guarantees with respect to locked files and delivery of updates cached at the client. Neither of these goals applies to directories protected by read delegations and notifications. Thus, no provision is made for reclaiming directory delegations in the event of client or server restart. The client can simply establish a directory delegation in the same fashion as was done initially.
| TOC |
NFSv4.1 supports attributes that allow a namespace to extend beyond the boundaries of a single server. It is RECOMMENDED that clients and servers support construction of such multi-server namespaces. Use of such multi-server namespaces is OPTIONAL however, and for many purposes, single-server namespace are perfectly acceptable. Use of multi-server namespaces can provide many advantages, however, by separating a file system's logical position in a namespace from the (possibly changing) logistical and administrative considerations that result in particular file systems being located on particular servers.
| TOC |
NFSv4.1 contains RECOMMENDED attributes that allow file systems on one server to be associated with one or more instances of that file system on other servers. These attributes specify such file system instances by specifying a server address target (either as a DNS name representing one or more IP addresses or as a literal IP address) together with the path of that file system within the associated single-server namespace.
The fs_locations_info RECOMMENDED attribute allows specification of one or more file system instance locations where the data corresponding to a given file system may be found. This attribute provides to the client, in addition to information about file system instance locations, significant information about the various file system instance choices (e.g. priority for use, writability, currency, etc.). It also includes information to help the client efficiently effect as seamless a transition as possible among multiple file system instances, when and if that should be necessary.
The fs_locations RECOMMENDED attribute is inherited from NFSv4.0 and only allows specification of the file system locations where the data corresponding to a given file system may be found. Servers SHOULD make this attribute available whenever fs_locations_info is supported, but client use of fs_locations_info is to be preferred.
| TOC |
A given location in an NFSv4.1 namespace (typically but not necessarily a multi-server namespace) can have a number of file system instance locations associated with it (via the fs_locations or fs_locations_info attribute). There may also be an actual current file system at that location, accessible via normal namespace operations (e.g. LOOKUP). In this case, the file system is said to be "present" at that position in the namespace and clients will typically use it, reserving use of additional locations specified via the location-related attributes to situations in which the principal location is no longer available.
When there is no actual file system at the namespace location in question, the file system is said to be "absent". An absent file system contains no files or directories other than the root. Any reference to it, except to access a small set of attributes useful in determining alternate locations, will result in an error, NFS4ERR_MOVED. Note that if the server ever returns the error NFS4ERR_MOVED, it MUST support the fs_locations attribute and SHOULD support the fs_locations_info and fs_status attributes.
While the error name suggests that we have a case of a file system which once was present, and has only become absent later, this is only one possibility. A position in the namespace may be permanently absent with the set of file system(s) designated by the location attributes being the only realization. The name NFS4ERR_MOVED reflects an earlier, more limited conception of its function, but this error will be returned whenever the referenced file system is absent, whether it has moved or not.
Except in the case of GETATTR-type operations (to be discussed later), when the current filehandle at the start of an operation is within an absent file system, that operation is not performed and the error NFS4ERR_MOVED returned, to indicate that the file system is absent on the current server.
Because a GETFH cannot succeed if the current filehandle is within an absent file system, filehandles within an absent file system cannot be transferred to the client. When a client does have filehandles within an absent file system, it is the result of obtaining them when the file system was present, and having the file system become absent subsequently.
It should be noted that because the check for the current filehandle being within an absent file system happens at the start of every operation, operations that change the current filehandle so that it is within an absent file system will not result in an error. This allows such combinations as PUTFH-GETATTR and LOOKUP-GETATTR to be used to get attribute information, particularly location attribute information, as discussed below.
The RECOMMENDED file system attribute fs_status can be used to interrogate the present/absent status of a given file system.
| TOC |
When a file system is absent, most attributes are not available, but it is necessary to allow the client access to the small set of attributes that are available, and most particularly those that give information about the correct current locations for this file system, fs_locations and fs_locations_info.
| TOC |
As mentioned above, an exception is made for GETATTR in that attributes may be obtained for a filehandle within an absent file system. This exception only applies if the attribute mask contains at least one attribute bit that indicates the client is interested in a result regarding an absent file system: fs_locations, fs_locations_info, or fs_status. If none of these attributes is requested, GETATTR will result in an NFS4ERR_MOVED error.
When a GETATTR is done on an absent file system, the set of supported attributes is very limited. Many attributes, including those that are normally REQUIRED, will not be available on an absent file system. In addition to the attributes mentioned above (fs_locations, fs_locations_info, fs_status), the following attributes SHOULD be available on absent file systems, in the case of RECOMMENDED attributes at least to the same degree that they are available on present file systems.
- change_policy:
- This attribute is useful for absent file systems and can be helpful in summarizing to the client when any of the location-related attributes changes.
- fsid:
- This attribute should be provided so that the client can determine file system boundaries, including, in particular, the boundary between present and absent file systems. This value must be different from any other fsid on the current server and need have no particular relationship to fsids on any particular destination to which the client might be directed.
- mounted_on_fileid:
- For objects at the top of an absent file system this attribute needs to be available. Since the fileid is one which is within the present parent file system, there should be no need to reference the absent file system to provide this information.
Other attributes SHOULD NOT be made available for absent file systems, even when it is possible to provide them. The server should not assume that more information is always better and should avoid gratuitously providing additional information.
When a GETATTR operation includes a bit mask for one of the attributes fs_locations, fs_locations_info, or fs_status, but where the bit mask includes attributes which are not supported, GETATTR will not return an error, but will return the mask of the actual attributes supported with the results.
Handling of VERIFY/NVERIFY is similar to GETATTR in that if the attribute mask does not include fs_locations, fs_locations_info, or fs_status, the error NFS4ERR_MOVED will result. It differs in that any appearance in the attribute mask of an attribute not supported for an absent file system (and note that this will include some normally REQUIRED attributes), will also cause an NFS4ERR_MOVED result.
| TOC |
A READDIR performed when the current filehandle is within an absent file system will result in an NFS4ERR_MOVED error, since, unlike the case of GETATTR, no such exception is made for READDIR.
Attributes for an absent file system may be fetched via a READDIR for a directory in a present file system, when that directory contains the root directories of one or more absent file systems. In this case, the handling is as follows:
| TOC |
The location-bearing attributes (fs_locations and fs_locations_info), provide, together with the possibility of absent file systems, a number of important facilities in providing reliable, manageable, and scalable data access.
When a file system is present, these attributes can provide alternative locations, to be used to access the same data, in the event of server failures, communications problems, or other difficulties that make continued access to the current file system impossible or otherwise impractical. Under some circumstances multiple alternative locations may be used simultaneously to provide higher performance access to the file system in question. Provision of such alternate locations is referred to as "replication" although there are cases in which replicated sets of data are not in fact present, and the replicas are instead different paths to the same data.
When a file system is present and becomes absent, clients can be given the opportunity to have continued access to their data, at an alternate location. In this case, a continued attempt to use the data in the now-absent file system will result in an NFS4ERR_MOVED error and at that point the successor locations (typically only one but multiple choices are possible) can be fetched and used to continue access. Transfer of the file system contents to the new location is referred to as "migration", but it should be kept in mind that there are cases in which this term can be used, like "replication", when there is no actual data migration per se.
Where a file system was not previously present, specification of file system location provides a means by which file systems located on one server can be associated with a namespace defined by another server, thus allowing a general multi-server namespace facility. A designation of such a location, in place of an absent file system, is called a "referral".
Because client support for location-related attributes is OPTIONAL, a server may (but is not required to) take action to hide migration and referral events from such clients, by acting as a proxy, for example. The server can determine the presence of client support from the arguments of the EXCHANGE_ID operation (see Section 18.35.3 (DESCRIPTION)).
| TOC |
The fs_locations and fs_locations_info attributes provide alternative locations, to be used to access data in place of or in addition to the current file system instance. On first access to a file system, the client should obtain the value of the set of alternate locations by interrogating the fs_locations or fs_locations_info attribute, with the latter being preferred.
In the event that server failures, communications problems, or other difficulties make continued access to the current file system impossible or otherwise impractical, the client can use the alternate locations as a way to get continued access to its data. Depending on specific attributes of these alternate locations, as indicated within the fs_locations_info attribute, multiple locations may be used simultaneously, to provide higher performance through the exploitation of multiple paths between client and target file system.
The alternate locations may be physical replicas of the (typically read-only) file system data, or they may reflect alternate paths to the same server or provide for the use of various forms of server clustering in which multiple servers provide alternate ways of accessing the same physical file system. How these different modes of file system transition are represented within the fs_locations and fs_locations_info attributes and how the client deals with file system transition issues will be discussed in detail below.
Multiple server addresses, whether they are derived from a single entry with a DNS name representing a set of IP addresses, or from multiple entries each with its own server address may correspond to the same actual server. The fact that two addresses correspond to the same server is shown by a common so_major_id field within the eir_server_owner field returned by EXCHANGE_ID (see Section 18.35.3 (DESCRIPTION)). For a detailed discussion of how server address targets interact with the determination of server identity specified by the server owner field, see Section 11.5 (Location Entries and Server Identity).
| TOC |
When a file system is present and becomes absent, clients can be given the opportunity to have continued access to their data, at an alternate location, as specified by the fs_locations or fs_locations_info attribute. Typically, a client will be accessing the file system in question, get an NFS4ERR_MOVED error, and then use the fs_locations or fs_locations_info attribute to determine the new location of the data. When fs_locations_info is used, additional information will be available which will define the nature of the client's handling of the transition to a new server.
Such migration can be helpful in providing load balancing or general resource reallocation. The protocol does not specify how the file system will be moved between servers. It is anticipated that a number of different server-to-server transfer mechanisms might be used with the choice left to the server implementer. The NFSv4.1 protocol specifies the method used to communicate the migration event between client and server.
The new location may be an alternate communication path to the same server, or, in the case of various forms of server clustering, another server providing access to the same physical file system. The client's responsibilities in dealing with this transition depend on the specific nature of the new access path and how and whether data was in fact migrated. These issues will be discussed in detail below.
When multiple server addresses correspond to the same actual server, as shown by a common value for the so_major_id field of the eir_server_owner field returned by EXCHANGE_ID, the location or locations may designate alternate server addresses in the form of specific server network addresses. These could be used to access the file system in question at those addresses and when it is no longer accessible at the original address.
Although a single successor location is typical, multiple locations may be provided, together with information that allows priority among the choices to be indicated, via information in the fs_locations_info attribute. Where suitable clustering mechanisms make it possible to provide multiple identical file systems or paths to them, this allows the client the opportunity to deal with any resource or communications issues that might limit data availability.
When an alternate location is designated as the target for migration, it must designate the same data (with metadata being the same to the degree indicated by the fs_locations_info attribute). Where file systems are writable, a change made on the original file system must be visible on all migration targets. Where a file system is not writable but represents a read-only copy (possibly periodically updated) of a writable file system, similar requirements apply to the propagation of updates. Any change visible in the original file system must already be effected on all migration targets, to avoid any possibility, that a client in effecting a transition to the migration target will see any reversion in file system state.
| TOC |
Referrals provide a way of placing a file system in a location within the namespace essentially without respect to its physical location on a given server. This allows a single server or a set of servers to present a multi-server namespace that encompasses file systems located on multiple servers. Some likely uses of this include establishment of site-wide or organization-wide namespaces, or even knitting such together into a truly global namespace.
Referrals occur when a client determines, upon first referencing a position in the current namespace, that it is part of a new file system and that the file system is absent. When this occurs, typically by receiving the error NFS4ERR_MOVED, the actual location or locations of the file system can be determined by fetching the fs_locations or fs_locations_info attribute.
The locations-related attribute may designate a single file system location or multiple file system locations, to be selected based on the needs of the client. The server, in the fs_locations_info attribute may specify priorities to be associated with various file system location choices. The server may assign different priorities to different locations as reported to individual clients, in order to adapt to client physical location or to effect load balancing. When both read-only and read-write file systems are present, some of the read-only locations may not be absolutely up-to-date (as they would have to be in the case of replication and migration). Servers may also specify file system locations that include client-substituted variables so that different clients are referred to different file systems (with different data contents) based on client attributes such as CPU architecture.
When the fs_locations_info attribute indicates that there are multiple possible targets listed, the relationships among them may be important to the client in selecting the one to use. The same rules specified in Section 11.4.1 (File System Replication) defining the appropriate standards for the data propagation, apply to these multiple replicas as well. For example, the client might prefer a writable target on a server that has additional writable replicas to which it subsequently might switch. Note that, as distinguished from the case of replication, there is no need to deal with the case of propagation of updates made by the current client, since the current client has not accessed the file system in question.
Use of multi-server namespaces is enabled by NFSv4.1 but is not required. The use of multi-server namespaces and their scope will depend on the applications used, and system administration preferences.
Multi-server namespaces can be established by a single server providing a large set of referrals to all of the included file systems. Alternatively, a single multi-server namespace may be administratively segmented with separate referral file systems (on separate servers) for each separately-administered portion of the namespace. Any segment or the top-level referral file system may use replicated referral file systems for higher availability.
Generally, multi-server namespaces are for the most part uniform, in that the same data made available to one client at a given location in the namespace is made available to all clients at that location. There are however facilities provided which allow different clients to be directed to different sets of data, so as to adapt to such client characteristics as CPU architecture.
| TOC |
As mentioned above, a single location entry may have a server address target in the form of a DNS name which may represent multiple IP addresses, while multiple location entries may have their own server address targets, that reference the same server. Whether two IP addresses designate the same server is indicated by the existence of a common so_major_id field within the eir_server_owner field returned by EXCHANGE_ID (see Section 18.35.3 (DESCRIPTION)), subject to further verification, for details of which see Section 2.10.4 (Trunking).
When multiple addresses for the same server exist, the client may assume that for each file system in the namespace of a given server network address, there exist file systems at corresponding namespace locations for each of the other server network addresses. It may do this even in the absence of explicit listing in fs_locations and fs_locations_info. Such corresponding file system locations can be used as alternate locations, just as those explicitly specified via the fs_locations and fs_locations_info attributes. Where these specific addresses are explicitly designated in the fs_locations_info attribute, the conditions of use specified in this attribute (e.g. priorities, specification of simultaneous use) may limit the client's use of these alternate locations.
If a single location entry designates multiple server IP addresses, the client cannot assume that these addresses are multiple paths to the same server. In most case they will be, but the client MUST verify that before acting on that assumption. When two server addresses are designated by a single location entry and they correspond to different servers, this normally indicates some sort of misconfiguration, and so the client should avoid use such location entries when alternatives are available. When they are not, clients should pick one of IP addresses and use it, without using others that are not directed to the same server.
| TOC |
When clients make use of servers that implement referrals, replication, and migration, care should be taken so that a user who mounts a given file system that includes a referral or a relocated file system continues to see a coherent picture of that user-side file system despite the fact that it contains a number of server-side file systems which may be on different servers.
One important issue is upward navigation from the root of a server-side file system to its parent (specified as ".." in UNIX), in the case in which it transitions to that file system as a result of referral, migration, or a transition as a result of replication. When the client is at such a point, and it needs to ascend to the parent, it must go back to the parent as seen within the multi-server namespace rather issuing a LOOKUPP call to the server, which would result in the parent within that server's single-server namespace. In order to do this, the client needs to remember the filehandles that represent such file system roots, and use these instead of issuing a LOOKUPP to the current server. This will allow the client to present to applications a consistent namespace, where upward navigation and downward navigation are consistent.
Another issue concerns refresh of referral locations. When referrals are used extensively, they may change as server configurations change. It is expected that clients will cache information related to traversing referrals so that future client side requests are resolved locally without server communication. This is usually rooted in client-side name lookup caching. Clients should periodically purge this data for referral points in order to detect changes in location information. When the change_policy attribute changes for directories that hold referral entries or for the referral entries themselves, clients should consider any associated cached referral information to be out of date.
| TOC |
Transitions between file system instances, whether due to switching between replicas upon server unavailability, or in response to server-initiated migration events are best dealt with together. This is so even though for the server, pragmatic considerations will normally force different implementation strategies for planned and unplanned transitions. Even though the prototypical use cases of replication and migration contain distinctive sets of features, when all possibilities for these operations are considered, there is an underlying unity of these operations, from the client's point of view, that makes treating them together desirable.
A number of methods are possible for servers to replicate data and to track client state in order to allow clients to transition between file system instances with a minimum of disruption. Such methods vary between those that use inter-server clustering techniques to limit the changes seen by the client, to those that are less aggressive, use more standard methods of replicating data, and impose a greater burden on the client to adapt to the transition.
The NFSv4.1 protocol does not impose choices on clients and servers with regard to that spectrum of transition methods. In fact, there are many valid choices, depending on client and application requirements and their interaction with server implementation choices. The NFSv4.1 protocol does define the specific choices that can be made, how these choices are communicated to the client and how the client is to deal with any discontinuities.
In the sections below, references will be made to various possible server implementation choices as a way of illustrating the transition scenarios that clients may deal with. The intent here is not to define or limit server implementations but rather to illustrate the range of issues that clients may face.
In the discussion below, references will be made to a file system having a particular property or of two file systems (typically the source and destination) belonging to a common class of any of several types. Two file systems that belong to such a class share some important aspect of file system behavior that clients may depend upon when present, to easily effect a seamless transition between file system instances. Conversely, where the file systems do not belong to such a common class, the client has to deal with various sorts of implementation discontinuities which may cause performance or other issues in effecting a transition.
Where the fs_locations_info attribute is available, such file system classification data will be made directly available to the client (see Section 11.10 (The Attribute fs_locations_info) for details). When only fs_locations is available, default assumptions with regard to such classifications have to be inferred (see Section 11.9 (The Attribute fs_locations) for details).
In cases in which one server is expected to accept opaque values from the client that originated from another server, the servers SHOULD encode the "opaque" values in big endian byte order. If this is done, servers acting as replicas or immigrating file systems will be able to parse values like stateids, directory cookies, filehandles, etc. even if their native byte order is different from that of other servers cooperating in the replication and migration of the file system.
| TOC |
When a single file system may be accessed at multiple locations, whether this is because of an indication of file system identity as reported by the fs_locations or fs_locations_info attributes or because two file system instances have corresponding locations on server addresses which connect to the same server (as indicated by a common so_major_id field in the eir_server_owner field returned by EXCHANGE_ID), the client will, depending on specific circumstances as discussed below, either:
Which of these choices is possible, and how a transition is effected, is governed by equivalence classes of file system instances as reported by the fs_locations_info attribute, and, for file system instances in the same location within a multiple single-server namespace as indicated by the so_major_id field in the eir_server_owner field returned by EXCHANGE_ID.
| TOC |
When two file system instances have the same location within their respective single-server namespaces and those two server network addresses designate the same server (as indicated by the same so_major_id value in the eir_server_owner value returned in response to EXCHANGE_ID), those file systems instances can be treated as the same, and either used together simultaneously or serially with no transition activity required on the part of the client. In this case we refer to the transition as "transparent" and the client in transferring access from to the other is acting as it would in the event that communication is interrupted, with a new connection and possibly a new session being established to continue access to the same file system.
Whether simultaneous use of the two file system instances is valid is controlled by whether the fs_locations_info attribute shows the two instances as having the same simultaneous-use class. See Section 11.10.1 (The fs_locations_server4 Structure) for information about the definition of the various use classes, including the simultaneous-use class.
Note that for two such file systems, any information within the fs_locations_info attribute that indicates the need for special transition activity, i.e. the appearance of the two file system instances with different handle, fileid, write-verifier, change, readdir classes, indicates a serious problem and the client, if it allows transition to the file system instance at all, must not treat this as a transparent transition. The server SHOULD NOT indicate that these instances belong to different handle, fileid, write-verifier, change, readdir classes, whether the two instances are shown belonging to the same simultaneous-use class or not.
Where these conditions do not apply, a non-transparent file system instance transition is required with the details depending on the respective handle, fileid, write-verifier, change, readdir classes of the two file system instances and whether the two servers address in question have the same eir_server_scope value as reported by EXCHANGE_ID.
| TOC |
When the conditions in Section 11.7.2 (Simultaneous Use and Transparent Transitions) hold, in either of the following two cases, the client may use the two file system instances simultaneously.
In this case, the client may use both file system instances simultaneously, as representations of the same file system, whether that happens because the two network addresses connect to the same physical server or because different servers connect to clustered file systems and export their data in common. When simultaneous use is in effect, any change made to one file system instance must be immediately reflected in the other file system instance(s). Locks are treated as part of a common lease, associated with a common client ID. Depending on the details of the eir_server_owner returned by EXCHANGE_ID, the two server instances may be accessed by different sessions or a single session in common.
| TOC |
When the conditions in Section 11.7.2.1 (Simultaneous Use of File System Instances) hold and the fs_locations_info attribute explicitly shows the file system instances for these distinct network addresses as belonging to different simultaneous-use classes, the file system instances should not be used by the client simultaneously, but rather serially with one being used unless and until communication difficulties, lack of responsiveness, or an explicit migration event causes another file system instance (or set of file system instances sharing a common simultaneous-use class) to be used.
When a change of file system instance is to be done, the client will use the same client ID already in effect. If it already has connections to the new server address, these will be used. Otherwise new connections to existing sessions or new sessions associated with the existing client ID are established as indicated by the eir_server_owner returned by EXCHANGE_ID.
In all such transparent transition cases, the following apply:
| TOC |
There are a number of ways in which filehandles can be handled across a file system transition. These can be divided into two broad classes depending upon whether the two file systems across which the transition happens share sufficient state to effect some sort of continuity of file system handling.
When there is no such co-operation in filehandle assignment, the two file systems are reported as being in different handle classes. In this case, all filehandles are assumed to expire as part of the file system transition. Note that this behavior does not depend on fh_expire_type attribute and supersedes the specification of FH4_VOL_MIGRATION bit, which only affects behavior when fs_locations_info is not available.
When there is co-operation in filehandle assignment, the two file systems are reported as being in the same handle classes. In this case, persistent filehandles remain valid after the file system transition, while volatile filehandles (excluding those that are only volatile due to the FH4_VOL_MIGRATION bit) are subject to expiration on the target server.
| TOC |
In NFSv4.0, the issue of continuity of fileids in the event of a file system transition was not addressed. The general expectation had been that in situations in which the two file system instances are created by a single vendor using some sort of file system image copy, fileids will be consistent across the transition while in the analogous multi-vendor transitions they will not. This poses difficulties, especially for the client without special knowledge of the transition mechanisms adopted by the server. Note that although fileid is not a REQUIRED attribute, many servers support fileids and many clients provide API's that depend on fileids.
It is important to note that while clients themselves may have no trouble with a fileid changing as a result of a file system transition event, applications do typically have access to the fileid (e.g. via stat), and the result of this is that an application may work perfectly well if there is no file system instance transition or if any such transition is among instances created by a single vendor, yet be unable to deal with the situation in which a multi-vendor transition occurs, at the wrong time.
Providing the same fileids in a multi-vendor (multiple server vendors) environment has generally been held to be quite difficult. While there is work to be done, it needs to be pointed out that this difficulty is partly self-imposed. Servers have typically identified fileid with inode number, i.e. with a quantity used to find the file in question. This identification poses special difficulties for migration of a file system between vendors where assigning the same index to a given file may not be possible. Note here that a fileid is not required to be useful to find the file in question, only that it is unique within the given file system. Servers prepared to accept a fileid as a single piece of metadata and store it apart from the value used to index the file information can relatively easily maintain a fileid value across a migration event, allowing a truly transparent migration event.
In any case, where servers can provide continuity of fileids, they should, and the client should be able to find out that such continuity is available and take appropriate action. Information about the continuity (or lack thereof) of fileids across a file system transition is represented by specifying whether the file systems in question are of the same fileid class.
Note that when consistent fileids do not exist across a transition (either because there is no continuity of fileids or because fileid is not a supported attribute on one of instances involved), and there are no reliable filehandles across a transition event (either because there is no filehandle continuity or because the filehandles are volatile), the client is in a position where it cannot verify that files it was accessing before the transition are the same objects. It is forced to assume that no object has been renamed, and, unless there are guarantees that provide this (e.g. the file system is read-only), problems for applications may occur. Therefore, use of such configurations should be limited to situations where the problems that this may cause can be tolerated.
| TOC |
Since fsids are generally only unique within a per-server basis, it is likely that they will change during a file system transition. One exception is the case of transparent transitions, but in that case we have multiple network addresses that are defined as the same server (as specified by a common value of the so_major_id field of eir_server_owner). Clients should not make the fsids received from the server visible to applications since they may not be globally unique, and because they may change during a file system transition event. Applications are best served if they are isolated from such transitions to the extent possible.
Although normally, a single source file system will transition to a single target file system, there is a provision for splitting a single source file system into multiple target file systems, by specifying the FSLI4F_MULTI_FS flag.
| TOC |
When a file system transition is made and the fs_locations_info indicates that the file system in question may be split into multiple file systems (via the FSLI4F_MULTI_FS flag), the client SHOULD do GETATTRs to determine the fsid attribute on all known objects within the file system undergoing transition to determine the new file system boundaries.
Clients may maintain the fsids passed to existing applications by mapping all of the fsids for the descendent file systems to the common fsid used for the original file system.
Splitting a file system may be done on a transition between file systems of the same fileid class, since the fact that fileids are unique within the source file system ensure they will be unique in each of the target file systems.
| TOC |
Since the change attribute is defined as a server-specific one, change attributes fetched from one server are normally presumed to be invalid on another server. Such a presumption is troublesome since it would invalidate all cached change attributes, requiring refetching. Even more disruptive, the absence of any assured continuity for the change attribute means that even if the same value is retrieved on refetch no conclusions can drawn as to whether the object in question has changed. The identical change attribute could be merely an artifact of a modified file with a different change attribute construction algorithm, with that new algorithm just happening to result in an identical change value.
When the two file systems have consistent change attribute formats, and this fact is communicated to the client by reporting as in the same change class, the client may assume a continuity of change attribute construction and handle this situation just as it would be handled without any file system transition.
| TOC |
In a file system transition, the client needs to handle cases in which the two servers have cooperated in state management and in which they have not. Cooperation by two servers in state management requires coordination of client IDs. Before the client attempts to use a client ID associated with one server in a request to the server of the other file system, it must eliminate the possibility that two non-cooperating servers have assigned the same client ID by accident. The client needs to compare the eir_server_scope values returned by each server. If the scope values do not match, then the servers have not cooperated in state management. If the scope values match, then this indicates the servers have cooperated in assigning client IDs to the point that they will reject client IDs that refer to state they do not know about.
In the case of migration, the servers involved in the migration of a file system SHOULD transfer all server state from the original to the new server. When this is done, it must be done in a way that is transparent to the client. With replication, such a degree of common state is typically not the case. Clients, however should use the information provided by the eir_server_scope returned by EXCHANGE_ID to determine whether such sharing may be in effect, rather than making assumptions based on the reason for the transition.
This state transfer will reduce disruption to the client when a file system transition occurs. If the servers are successful in transferring all state, the client can attempt to establish sessions associated with the client ID used for the source file system instance. If the server accepts that as a valid client ID, then the client may use the existing stateids associated with that client ID for the old file system instance in connection with that same client ID in connection with the transitioned file system instance.
When the two servers belong to the same server scope, it does not mean that when dealing with the transition, the client will not have to reclaim state. However it does mean that the client may proceed using its current client ID when establishing communication with the new server and the new server will either recognize the client ID as valid, or reject it, in which case locks must be reclaimed by the client.
File systems co-operating in state management may actually share state or simply divide the id space so as to recognize (and reject as stale) each other's stateids and client IDs. Servers which do share state may not do so under all conditions or at all times. The requirement for the server is that if it cannot be sure in accepting a client ID that it reflects the locks the client was given, it must treat all associated state as stale and report it as such to the client.
When the two file system instances are on servers that do not share a server scope value, the client must establish a new client ID on the destination, if it does not have one already, and reclaim locks if possible. In this case, old stateids and client IDs should not be presented to the new server since there is no assurance that they will not conflict with IDs valid on that server.
In either case, when actual locks are not known to be maintained, the destination server may establish a grace period specific to the given file system, with non-reclaim locks being rejected for that file system, even though normal locks are being granted for other file systems. Clients should not infer the absence of a grace period for file systems being transitioned to a server from responses to requests for other file systems.
In the case of lock reclamation for a given file system after a file system transition, edge conditions can arise similar to those for reclaim after server restart (although in the case of the planned state transfer associated with migration, these can be avoided by securely recording lock state as part of state migration). Unless the destination server can guarantee that locks will not be incorrectly granted, the destination server should not allow lock reclaims and avoid establishing a grace period.
Once all locks have been reclaimed, or there were no locks to reclaim, the client indicates that there are no more reclaims to be done for the file system in question by issuing a RECLAIM_COMPLETE operation with the rca_one_fs parameter set to true. Once this has been done, non-reclaim locking operations may be done, and any subsequent request to do reclaims will be rejected with the error NFS4ERR_NO_GRACE.
Information about client identity may be propagated between servers in the form of client_owner4 and associated verifiers, under the assumption that the client presents the same values to all the servers with which it deals.
Servers are encouraged to provide facilities to allow locks to be reclaimed on the new server after a file system transition. Often, however, in cases in which the two servers do not share a server scope value, such facilities may not be available and client should be prepared to re-obtain locks, even though it is possible that the client may have its LOCK or OPEN request denied due to a conflicting lock.
The consequences of having no facilities available to reclaim locks on the sew server will depend on the type of environment. In some environments, such as the transition between read-only file systems, such denial of locks should not pose large difficulties in practice. When an attempt to re-establish a lock on a new server is denied, the client should treat the situation as if its original lock had been revoked. Note that when the lock is granted, the client cannot assume that no conflicting lock could have been granted in the interim. Where change attribute continuity is present, the client may check the change attribute to check for unwanted file modifications. Where even this is not available, and the file system is not read-only, a client may reasonably treat all pending locks as having been revoked.
| TOC |
In the case of lease renewal, the client may not be submitting requests for a file system that has been transferred to another server. This can occur because of the lease renewal mechanism. The client renews the lease associated with all file systems when submitting a request on an associated session, regardless of the specific file system being referenced.
In order for the client to schedule renewal of leases where there is locking state that may have been relocated to the new server, the client must find out about lease relocation before those leases expire. To accomplish this, the SEQUENCE operation will return the status bit SEQ4_STATUS_LEASE_MOVED, if responsibility for any of the locking state renewed has been transferred to a new server. This will continue until the client receives an NFS4ERR_MOVED error for each of the file systems for which there has been locking state relocation.
When a client receives an SEQ4_STATUS_LEASE_MOVED indication, it should perform an operation on each file system associated with the server where there is locking state for the current client associated with the file system in question. The client may choose to reference all file systems in the interests of simplicity but what is important is that it must reference all file systems for which there was locking state where that state moved. Once the client receives an NFS4ERR_MOVED error for each file system, the SEQ4_STATUS_LEASE_MOVED indication is cleared. The client can terminate the process of checking file systems once this indication is cleared (but only if the client has received a reply for all outstanding SEQUENCE requests on all sessions it has with the server), since there are no others for which locking state has moved.
A client may use GETATTR of the fs_status (or fs_locations_info) attribute on all of the file systems to get absence indications in a single (or a few) request(s), since absent file systems will not cause an error in this context. However, it still must do an operation which receives NFS4ERR_MOVED on each file system, in order to clear the SEQ4_STATUS_LEASE_MOVED indication is cleared.
Once the set of file systems with transferred locking state has been determined, the client can follow the normal process to obtain the new server information (through the fs_locations and fs_locations_info attributes) and perform renewal of those leases on the new server, unless information in fs_locations_info attribute shows that no state could have been transferred. If the server has not had state transferred to it transparently, the client will receive NFS4ERR_STALE_CLIENTID from the new server, as described above, and the client can then reclaim locks as is done in the event of server failure.
| TOC |
In order that the client may appropriately manage its leases in the case of a file system transition, the destination server must establish proper values for the lease_time attribute.
When state is transferred transparently, that state should include the correct value of the lease_time attribute. The lease_time attribute on the destination server must never be less than that on the source since this would result in premature expiration of leases granted by the source server. Upon transitions in which state is transferred transparently, the client is under no obligation to re-fetch the lease_time attribute and may continue to use the value previously fetched (on the source server).
If state has not been transferred transparently, either because the associated servers are shown as having different eir_server_scope strings or because the client ID is rejected when presented to the new server, the client should fetch the value of lease_time on the new (i.e. destination) server, and use it for subsequent locking requests. However the server must respect a grace period at least as long as the lease_time on the source server, in order to ensure that clients have ample time to reclaim their lock before potentially conflicting non-reclaimed locks are granted.
| TOC |
In a file system transition, the two file systems may be clustered in the handling of unstably written data. When this is the case, and the two file systems belong to the same write-verifier class, write verifiers returned from one system may be compared to those returned by the other and superfluous writes avoided.
When two file systems belong to different write-verifier classes, any verifier generated by one must not be compared to one provided by the other. Instead, it should be treated as not equal even when the values are identical.
| TOC |
In a file system transition, the two file systems may be consistent in their handling of READDIR cookies and verifiers. When this is the case, and the two file systems belong to the same readdir class, READDIR cookies and verifiers from one system may be recognized by the other and READDIR operations started on one server may be validly continued on the other, simply by presenting the cookie and verifier returned by a READDIR operation done on the first file system to the second.
When two file systems belong to different readdir classes, any READDIR cookie and verifier generated by one is not valid on the second, and must not be presented to that server by the client. The client should act as if the verifier was rejected.
| TOC |
When multiple replicas exist and are used simultaneously or in succession by a client, applications using them will normally expect that they contain data the same data or data which is consistent with the normal sorts of changes that are made by other clients updating the data of the file system. (with metadata being the same to the degree indicated by the fs_locations_info attribute). However, when multiple file systems are presented as replicas of one another, the precise relationship between the data of one and the data of another is not, as a general matter, specified by the NFSv4.1 protocol. It is quite possible to present as replicas file systems where the data of those file systems is sufficiently different that some applications have problems dealing with the transition between replicas. The namespace will typically be constructed so that applications can choose an appropriate level of support, so that in one position in the namespace a varied set of replicas will be listed while in another only those that are up-to-date may be considered replicas. The protocol does define three special cases of the relationship among replicas to be specified by the server and relied upon by clients:
| TOC |
Referrals are effected when an absent file system is encountered, and one or more alternate locations are made available by the fs_locations or fs_locations_info attributes. The client will typically get an NFS4ERR_MOVED error, fetch the appropriate location information and proceed to access the file system on a different server, even though it retains its logical position within the original namespace. Referrals differ from migration events in that they happen only when the client has not previously referenced the file system in question (so there is nothing to transition). Referrals can only come into effect when an absent file system is encountered at its root.
The examples given in the sections below are somewhat artificial in that an actual client will not typically do a multi-component lookup, but will have cached information regarding the upper levels of the name hierarchy. However, these example are chosen to make the required behavior clear and easy to put within the scope of a small number of requests, without getting unduly into details of how specific clients might choose to cache things.
| TOC |
Let us suppose that the following COMPOUND is sent in an environment in which /this/is/the/path is absent from the target server. This may be for a number of reasons. It may be the case that the file system has moved, or, it may be the case that the target server is functioning mainly, or solely, to refer clients to the servers on which various file systems are located.
Under the given circumstances, the following will be the result.
Given the failure of the GETFH, the client has the job of determining the root of the absent file system and where to find that file system, i.e. the server and path relative to that server's root fh. Note here that in this example, the client did not obtain filehandles and attribute information (e.g. fsid) for the intermediate directories, so that it would not be sure where the absent file system starts. It could be the case, for example, that /this/is/the is the root of the moved file system and that the reason that the lookup of "path" succeeded is that the file system was not absent on that operation but was moved between the last LOOKUP and the GETFH (since COMPOUND is not atomic). Even if we had the fsids for all of the intermediate directories, we could have no way of knowing that /this/is/the/path was the root of a new file system, since we don't yet have its fsid.
In order to get the necessary information, let us re-send the chain of LOOKUPs with GETFHs and GETATTRs to at least get the fsids so we can be sure where the appropriate file system boundaries are. The client could choose to get fs_locations_info at the same time but in most cases the client will have a good guess as to where file system boundaries are (because of where and where not NFS4ERR_MOVED was received) making fetching of fs_locations_info unnecessary.
- OP01:
- PUTROOTFH --> NFS_OK
- -
- Current fh is root of pseudo-fs.
- OP02:
- GETATTR(fsid) --> NFS_OK
- -
- Just for completeness. Normally, clients will know the fsid of the pseudo-fs as soon as they establish communication with a server.
- OP03:
- LOOKUP "this" --> NFS_OK
- OP04:
- GETATTR(fsid) --> NFS_OK
- -
- Get current fsid to see where file system boundaries are. The fsid will be that for the pseudo-fs in this example, so no boundary.
- OP05:
- GETFH --> NFS_OK
- -
- Current fh is for /this and is within pseudo-fs.
- OP06:
- LOOKUP "is" --> NFS_OK
- -
- Current fh is for /this/is and is within pseudo-fs.
- OP07:
- GETATTR(fsid) --> NFS_OK
- -
- Get current fsid to see where file system boundaries are. The fsid will be that for the pseudo-fs in this example, so no boundary.
- OP08:
- GETFH --> NFS_OK
- -
- Current fh is for /this/is and is within pseudo-fs.
- OP09:
- LOOKUP "the" --> NFS_OK
- -
- Current fh is for /this/is/the and is within pseudo-fs.
- OP10:
- GETATTR(fsid) --> NFS_OK
- -
- Get current fsid to see where file system boundaries are. The fsid will be that for the pseudo-fs in this example, so no boundary.
- OP11:
- GETFH --> NFS_OK
- -
- Current fh is for /this/is/the and is within pseudo-fs.
- OP12:
- LOOKUP "path" --> NFS_OK
- -
- Current fh is for /this/is/the/path and is within a new, absent file system, but ...
- -
- The client will never see the value of that fh
- OP13:
- GETATTR(fsid, fs_locations_info) --> NFS_OK
- -
- We are getting the fsid to know where the file system boundaries are. In this operation the fsid will be different than that of the parent directory (which in turn was retrieved in OP10). Note that the fsid we are given will not necessarily be preserved at the new location. That fsid might be different and in fact the fsid we have for this file system might be a valid fsid of a different file system on that new server.
- -
- In this particular case, we are pretty sure anyway that what has moved is /this/is/the/path rather than /this/is/the since we have the fsid of the latter and it is that of the pseudo-fs, which presumably cannot move. However, in other examples, we might not have this kind of information to rely on (e.g. /this/is/the might be a non-pseudo file system separate from /this/is/the/path), so we need to have another reliable source information on the boundary of the file system which is moved. If, for example, the file system "/this/is" had moved we would have a case of migration rather than referral and once the boundaries of the migrated file system was clear we could fetch fs_locations_info.
- -
- We are fetching fs_locations_info because the fact that we got an NFS4ERR_MOVED at this point means that it most likely that this is a referral and we need the destination. Even if it is the case that "/this/is/the" is a file system which has migrated, we will still need the location information for that file system.
- OP14:
- GETFH --> NFS4ERR_MOVED
- -
- Fails because current fh is in an absent file system at the start of the operation and the spec makes no exception for GETFH. Note that this means the server will never send the client a filehandle from within an absent file system.
Given the above, the client knows where the root of the absent file system is (/this/is/the/path), by noting where the change of fsid occurred (between "the" and "path"). The fs_locations_info attribute also gives the client the actual location of the absent file system, so that the referral can proceed. The server gives the client the bare minimum of information about the absent file system so that there will be very little scope for problems of conflict between information sent by the referring server and information of the file system's home. No filehandles and very few attributes are present on the referring server and the client can treat those it receives as basically transient information with the function of enabling the referral.
| TOC |
Another context in which a client may encounter referrals is when it does a READDIR on directory in which some of the sub-directories are the roots of absent file systems.
Suppose such a directory is read as follows:
In this case, because rdattr_error is not requested, fs_locations_info is not requested, and some of attributes cannot be provided, the result will be an NFS4ERR_MOVED error on the READDIR, with the detailed results as follows:
So now suppose that we re-send with rdattr_error:
The results will be:
So suppose we do another READDIR to get fs_locations_info (although we could have used a GETATTR directly, as in Section 11.8.1 (Referral Example (LOOKUP))).
The results would be:
The attributes for the directory entry with the component named "path" will only contain
The attributes for entry "path" will not contain size or time_modify because these attributes are not available within an absent file system.
| TOC |
The fs_locations attribute is structured in the following way:
struct fs_location4 {
utf8str_cis server<>;
pathname4 rootpath;
};
struct fs_locations4 {
pathname4 fs_root;
fs_location4 locations<>;
};
The fs_location4 data type is used to represent the location of a file system by providing a server name and the path to the root of the file system within that server's namespace. When a set of servers have corresponding file systems at the same path within their namespaces, an array of server names may be provided. An entry in the server array is a UTF-8 string and represents one of a traditional DNS host name, IPv4 address, or IPv6 address, or an zero-length string. A zero-length string SHOULD be used to indicate the current address being used for the RPC call. It is not a requirement that all servers that share the same rootpath be listed in one fs_location4 instance. The array of server names is provided for convenience. Servers that share the same rootpath may also be listed in separate fs_location4 entries in the fs_locations attribute.
The fs_locations4 data type and fs_locations attribute contain an array of such locations. Since the namespace of each server may be constructed differently, the "fs_root" field is provided. The path represented by fs_root represents the location of the file system in the current server's namespace, i.e. that of the server from which the fs_locations attribute was obtained. The fs_root path is meant to aid the client by clearly referencing the root of the file system whose locations are being reported, no matter what object within the current file system the current filehandle designates. The fs_root is simply the pathname the client used to reach the object on the current server, the object being that the fs_locations attribute applies to.
When the fs_locations attribute is interrogated and there are no alternate file system locations, the server SHOULD return a zero-length array of fs_location4 structures, together with a valid fs_root.
As an example, suppose there is a replicated file system located at two servers (servA and servB). At servA, the file system is located at path "/a/b/c". At, servB the file system is located at path "/x/y/z". If the client were to obtain the fs_locations value for the directory at "/a/b/c/d", it might not necessarily know that the file system's root is located in servA's namespace at "/a/b/c". When the client switches to servB, it will need to determine that the directory it first referenced at servA is now represented by the path "/x/y/z/d" on servB. To facilitate this, the fs_locations attribute provided by servA would have a fs_root value of "/a/b/c" and two entries in fs_locations. One entry in fs_locations will be for itself (servA) and the other will be for servB with a path of "/x/y/z". With this information, the client is able to substitute "/x/y/z" for the "/a/b/c" at the beginning of its access path and construct "/x/y/z/d" to use for the new server.
Note that: there is no requirement that the number of components in each rootpath be the same; there is no relation between the number of components in rootpath or fs_root; and the none of the components in each rootpath and fs_root have to be the same. In the above example, we could have had a third element in the locations array, with server equal to "servC", and rootpath equal to "/I/II", and a fourth element in locations with server equal to "servD", and rootpath equal to "/aleph/beth/gimel/daleth/he".
The relationship between fs_root to a rootpath is that the client replaces the pathname indicated in fs_root for the current server for the substitute indicated in rootpath for the new server.
For an example for a referred or migrated file system, suppose there is a file system located at serv1. At serv1, the file system is located at "/az/buky/vedi/glagoli". The client finds that object at "glagoli" has migrated (or is a referral). The client gets the fs_locations attribute, which contains an fs_root of "/az/buky/vedi/glagoli", and one element in the locations array, with server equal to "serv2", and rootpath equal to "/izhitsa/fita". The client replaces "/az/buky/vedi/glagoli" with "/izhitsa/fita", and uses the latter pathname on "serv2".
Thus, the server MUST return an fs_root that is equal to the path the client used to reach the object the fs_locations attribute applies to. Otherwise the client cannot determine the new path to use on the new server.
Since the fs_locations attribute lacks information defining various attributes of the various file system choices presented, it SHOULD only be interrogated and used when fs_locations_info is not available. When fs_locations is used, information about the specific locations should be assumed based on the following rules.
The following rules are general and apply irrespective of the context.
For other class assignments, handling of file system transitions depends on the reasons for the transition:
The specific choices reflect typical implementation patterns for failover and controlled migration respectively. Since other choices are possible and useful, this information is better obtained by using fs_locations_info. When a server implementation needs to communicate other choices, it MUST support the fs_locations_info attribute.
See Section 21 (Security Considerations) for a discussion on the recommendations for the security flavor to be used by any GETATTR operation that requests the "fs_locations" attribute.
| TOC |
The fs_locations_info attribute is intended as a more functional replacement for fs_locations which will continue to exist and be supported. Clients can use it to get a more complete set of information about alternative file system locations. When the server does not support fs_locations_info, fs_locations can be used to get a subset of the information. A server which supports fs_locations_info MUST support fs_locations as well.
There is additional information present in fs_locations_info, that is not available in fs_locations:
The fs_locations_info attribute is structured similarly to the fs_locations attribute. A top-level structure (fs_locations_info4) contains the entire attribute including the root pathname of the file system and an array of lower-level structures that define replicas that share a common root path on their respective servers. The lower-level structure in turn (fs_locations_item4) contains a specific pathname and information on one or more individual server replicas. For that last lowest-level fs_locations_info has a fs_locations_server4 structure that contains per-server-replica information in addition to the server name. This per-server-replica information includes a nominally opaque array, fls_info, in which specific pieces of information are located at the specific indices listed below.
The attribute will always contains at least a single fs_locations_server entry. Typically, this will be an entry with the FS4LIGF_CUR_REQ flag set, although in the case of a referral there will be no entry with that flag set.
It should be noted that fs_locations_info attributes returned by servers for various replicas may different for various reasons. One server may know about a set of replicas that are not know to other servers. Further, compatibility attributes may differ. Filehandles may by of the same class going from replica A to replica B but not going in the reverse direction. This may happen because the filehandles are the same but the server implementation for the server on which replica B may not have provision to note and report that equivalence.
The fs_locations_info attribute consists of a root pathname (fli_fs_root, just like fs_root in the fs_locations attribute), together with an array of fs_location_item4 structures. The fs_location_item4 structures in turn consist of a root pathname (fli_rootpath) together with an array (fli_entries) of elements of data type fs_locations_server4, all defined as follows.
/*
* Defines an individual server replica
*/
struct fs_locations_server4 {
int32_t fls_currency;
opaque fls_info<>;
utf8str_cis fls_server;
};
/*
* Byte indices of items within
* fls_info: flag fields, class numbers,
* bytes indicating ranks and orders.
*/
const FSLI4BX_GFLAGS = 0;
const FSLI4BX_TFLAGS = 1;
const FSLI4BX_CLSIMUL = 2;
const FSLI4BX_CLHANDLE = 3;
const FSLI4BX_CLFILEID = 4;
const FSLI4BX_CLWRITEVER = 5;
const FSLI4BX_CLCHANGE = 6;
const FSLI4BX_CLREADDIR = 7;
const FSLI4BX_READRANK = 8;
const FSLI4BX_WRITERANK = 9;
const FSLI4BX_READORDER = 10;
const FSLI4BX_WRITEORDER = 11;
/*
* Bits defined within the general flag byte.
*/
const FSLI4GF_WRITABLE = 0x01;
const FSLI4GF_CUR_REQ = 0x02;
const FSLI4GF_ABSENT = 0x04;
const FSLI4GF_GOING = 0x08;
const FSLI4GF_SPLIT = 0x10;
/*
* Bits defined within the transport flag byte.
*/
const FSLI4TF_RDMA = 0x01;
/*
* Defines a set of replicas sharing
* a common value of the root path
* with in the corresponding
* single-server namespaces.
*/
struct fs_locations_item4 {
fs_locations_server4 fli_entries<>;
pathname4 fli_rootpath;
};
/*
* Defines the overall structure of
* the fs_locations_info attribute.
*/
struct fs_locations_info4 {
uint32_t fli_flags;
int32_t fli_valid_for;
pathname4 fli_fs_root;
fs_locations_item4 fli_items<>;
};
/*
* Flag bits in fli_flags.
*/
const FSLI4IF_VAR_SUB = 0x00000001;
typedef fs_locations_info4 fattr4_fs_locations_info;
As noted above, the fs_locations_info attribute, when supported, may be requested of absent file systems without causing NFS4ERR_MOVED to be returned and it is generally expected that it will be available for both present and absent file systems even if only a single fs_locations_server4 entry is present, designating the current (present) file system, or two fs_locations_server4 entries designating the previous location of an absent file system (the one just referenced) and its successor location. Servers are strongly urged to support this attribute on all file systems if they support it on any file system.
The data presented in the fs_locations_info attribute may be obtained by the server in any number of ways, including specification by the administrator or by current protocols for transferring data among replicas and protocols not yet developed. NFSv4.1 only defines how this information is presented by the server to the client.
| TOC |
The fs_locations_server4 structure consists of the following items:
Data within the fls_info array is in the form of 8-bit data items with constants giving the offsets within the array of various values describing this particular file system instance. This style of definition was chosen, in preference to explicit XDR structure definitions for these values, for a number of reasons.
This encoding scheme can be adapted to the specification of multi-byte numeric values, even though none are currently defined. If extensions are made via standards-track RFC's, multi-byte quantities will be encoded as a range of bytes with a range of indices with the byte interpreted in big endian byte order. Further any such index assignments are constrained so that the relevant quantities will not cross XDR word boundaries.
The set of fls_info data is subject to expansion in a future minor version, or in a standard-track RFC, within the context of a single minor version. The server SHOULD NOT send and the client MUST NOT use indices within the fls_info array that are not defined in standards-track RFC's.
The fls_info array contains within it:
The general file system characteristics flag (at byte index FSLI4BX_GFLAGS) has the following bits defined within it:
The transport-flag field (at byte index FSLI4BX_TFLAGS) contains the following bits related to the transport capabilities of the specific file system.
Attribute continuity and file system identity information are expressed by defining equivalence relations on the sets of file systems presented to the client. Each such relation is expressed as a set of file system equivalence classes. For each relation, a file system has an 8-bit class number. Two file systems belong to the same class if both have identical non-zero class numbers. Zero is treated as non-matching. Most often, the relevant question for the client will be whether a given replica is identical-to/continuous-with the current one in a given respect but the information should be available also as to whether two other replicas match in that respect as well.
The following fields specify the file system's class numbers for the equivalence relations used in determining the nature of file system transitions. See Section 11.7 (Effecting File System Transitions) for details about how this information is to be used. Servers may assign these values as they wish, so long as file system instances that share the same value have the specified relationship to one another, conversely file systems which have the specified relationship to one another share a common class value. As each instance entry is added, the relationships of this instance to previously entered instances can be consulted and if one is found that bears the specified relationship, that entry's class value can be copied to the new entry. When no such previous entry exists, a new value for that byte index, not previously used can be selected, most likely by increment the value of the last class value assigned for that index.
Server-specified preference information is also provided via 8-bit values within the fls_info array. The values provide a rank and an order (see below) to be used with separate values specifiable for the cases of read-only and writable file systems. These values are compared for different file systems to establish the server-specified preference, with lower values indicating "more preferred".
Rank is used to express a strict server-imposed ordering on clients, with lower values indicating "more preferred." Clients should attempt to use all replicas with a given rank before they use one with a higher rank. Only if all of those file systems are unavailable should the client proceed to those of a higher rank. Because specifying a rank will override client preferences, servers should be conservative about using this mechanism, particularly when the environment is one in client communication characteristics are not tightly controlled and visible to the server.
Within a rank, the order value is used to specify the server's preference to guide the client's selection when the client's own preferences are not controlling, with lower values of order indicating "more preferred." If replicas are approximately equal in all respects, clients should defer to the order specified by the server. When clients look at server latency as part of their selection, they are free to use this criterion but it is suggested that when latency differences are not significant, the server- specified order should guide selection.
Depending on the potential need for write access by a given client, one of the pairs of rank and order values is used. The read rank and order should only be used if the client knows that only reading will ever be done or if it is prepared to switch to a different replica in the event that any write access capability is required in the future.
| TOC |
The fs_locations_info4 structure, encoding the fs_locations_info attribute, contains the following:
The FSLI4IF_VAR_SUB flag within fli_flags controls whether variable substitution is to be enabled. See Section 11.10.3 (The fs_locations_item4 Structure) for an explanation of variable substitution.
| TOC |
The fs_locations_item4 structure contains a pathname (in the field fli_rootpath) which encodes the path of the target file system replicas on the set of servers designated by the included fs_locations_server4 entries. The precise manner in which this target location is specified depends on the value of the FSLI4IF_VAR_SUB flag within the associated fs_locations_info4 structure.
If this flag is not set, then fli_rootpath simply designates the location of the target file system within each server's single-server namespace just as it does for the rootpath within the fs_location4 structure. When this bit is set, however, component entries of a certain form are subject to client-specific variable substitution so as to allow a degree of namespace non-uniformity in order to accommodate the selection of client-specific file system targets to adapt to different client architectures or other characteristics.
When such substitution is in effect a variable beginning with the string "${" and ending with the string "}" and containing a colon is to be replaced by the client-specific value associated with that variable. The string "unknown" should be used by the client when it has no value for such a variable. The pathname resulting from such substitutions is used to designate the target file system, so that different clients may have different file systems, corresponding to that location in the multi-server namespace.
As mentioned above, such substituted pathname variables contain a colon. The part before the colon is to be a DNS domain name with the part after being a case-insensitive alphanumeric string.
Where the domain is "ietf.org", only variable names defined in this document or subsequent standards-track RFC's are subject to such substitution. Organizations are free to use their domain names to create their own sets of client-specific variables, to be subject to such substitution. In case where such variables are intended to be used more broadly than a single organization, publication of an informational RFC defining such variables is RECOMMENDED.
The variable ${ietf.org:CPU_ARCH} is used to denote the CPU architecture object files are compiled. This specification does not limit the acceptable values (except that they must be valid UTF-8 strings) but such values as "x86", "x86_64" and "sparc" would be expected to be used in line with industry practice.
The variable ${ietf.org:OS_TYPE} is used to denote the operating system and thus the kernel and library API's for which code might be compiled. This specification does not limit the acceptable values (except that they must be valid UTF-8 strings) but such values as "linux" and "freebsd" would be expected to be used in line with industry practice.
The variable ${ietf.org:OS_VERSION} is used to denote the operating system version and the thus the specific details of versioned interfaces for which code might be compiled. This specification does not limit the acceptable values (except that they must be valid UTF-8 strings) but combinations of numbers and letters with interspersed dots would be expected to be used in line with industry practice, with the details of the version format depending on the specific value of the value of the variable ${ietf.org:OS_TYPE} with which it is used.
Use of these variable could result in direction of different clients to different file systems on the same server, as appropriate to particular clients. In cases in which the target file systems are located on different servers, a single server could serve as a referral point so that each valid combination of variable values would designate a referral hosted on a single server, with the targets of those referrals on a number of different servers.
Because namespace administration is affected by the values selected to substitute for various variables, clients should provide convenient means of determining what variable substitutions a client will implement, as well as, where appropriate, providing means to control the substitutions to be used. The exact means by which this will be done is outside the scope of this specification.
Although variable substitution is most suitable for use in the context of referrals, if may be used in the context of replication and migration. If it is used in these contexts, the server must ensure that no matter what values the client presents for the substituted variables, the result is always a valid successor file system instance to that from which a transition is occurring, i.e. that the data is identical or represents a later image of a writable file system.
Note that when fli_rootpath is a null pathname (that is, one with zero components), the file system designated is at the root of the specified server, whether the FSLI4IF_VAR_SUB flag within the associated fs_locations_info4 structure is set or not.
| TOC |
In an environment in which multiple copies of the same basic set of data are available, information regarding the particular source of such data and the relationships among different copies can be very helpful in providing consistent data to applications.
enum fs4_status_type {
STATUS4_FIXED = 1,
STATUS4_UPDATED = 2,
STATUS4_VERSIONED = 3,
STATUS4_WRITABLE = 4,
STATUS4_REFERRAL = 5
};
struct fs4_status {
bool fss_absent;
fs4_status_type fss_type;
utf8str_cs fss_source;
utf8str_cs fss_current;
int32_t fss_age;
nfstime4 fss_version;
};
The boolean fss_absent indicates whether the file system is currently absent. This value will be set if the file system was previously present and becomes absent, or if the file system has never been present and the type is STATUS4_REFERRAL. When this boolean is set and the type is not STATUS4_REFERRAL, the remaining information in the fs4_status reflects that last valid when the file system was present.
The fss_type field indicates the kind of file system image represented. This is of particular importance when using the version values to determine appropriate succession of file system images. When fss_absent is set, and the file system was previously present, the value of fss_type reflected is that when the file was last present. Five values are distinguished:
Note that in the STATUS4_UPDATED and STATUS4_VERSIONED cases, the server is responsible for the appropriate handling of locks that are inconsistent with external changes to delegations. If a server gives out delegations, they SHOULD be recalled before an inconsistent change made to data, and MUST be revoked if this is not possible. Similarly, if an open is inconsistent with data that is changed (the open denies WRITE and the data is changed), that lock SHOULD be considered administratively revoked.
The opaque strings fss_source and fss_current provide a way of presenting information about the source of the file system image being present. It is not intended that client do anything with this information other than make it available to administrative tools. It is intended that this information be helpful when researching possible problems with a file system image that might arise when it is unclear if the correct image is being accessed and if not, how that image came to be made. This kind of dianostic information will be helpful, if, as seems likely, copies of file systems are made in many different ways (e.g. simple user-level copies, file system-level point-in-time copies, clones of the underlying storage), under a variety of administrative arrangements. In such environments, determining how a given set of data was constructed can be very helpful in resolving problems.
The opaque string fss_source is used to indicate the source of a given file system with the expectation that tools capable of creating a file system image propagate this information, when that is possible. It is understood that this may not always be possible since a user-level copy may be thought of as creating a new data set and the tools used may have no mechanism to propagate this data. When a file system is initially created, it is desirable to associate with it data regarding how the file system was created, where it was created, by whom, etc. Making this information available in this attribute in a human-readable string form will be helpful for applications and system administrators and also serves to make it available when the original file system is used to make subsequent copies.
The opaque string fss_current should provide whatever information is available about the source of the current copy. Such information as the tool creating it, any relevant parameters to that tool, the time at which the copy was done, the user making the change, the server on which the change was made, etc. All information should be in a human-readable string form.
The field fss_age provides an indication of how out-of-date the file system currently is with respect to its ultimate data source (in case of cascading data updates). This complements the fls_currency field of fs_locations_server4 (see Section 11.10 (The Attribute fs_locations_info)) in the following way: the information in fls_currency gives a bound for how out of date the data in a file system might typically get, while the value in fss_age gives a bound on how out of date that data actually is. Negative values imply that no information is available. A zero means that this data is known to be current. A positive value means that this data is known to be no older than that number of seconds with respect to the ultimate data source. Using this value, the client may be able to decide that a data copy is too old, so that it may search for a newer version to use.
The fss_version field provides a version identification, in the form of a time value, such that successive versions always have later time values. When the fs_type is anything other than STATUS4_VERSIONED, the server may provide such a value but there is no guarantee as to its validity and clients will not use it except to provide additional information to add to fss_source and fss_current.
When fss_type is STATUS4_VERSIONED, servers SHOULD provide a value of version which progresses monotonically whenever any new version of the data is established. This allows the client, if reliable image progression is important to it, to fetch this attribute as part of each COMPOUND where data or metadata from the file system is used.
When it is important to the client to make sure that only valid successor images are accepted, it must make sure that it does not read data or metadata from the file system without updating its sense of the current state of the image, to avoid the possibility that the fs_status which the client holds will be one for an earlier image, and so accept a new file system instance which is later than that but still earlier than updated data read by the client.
In order to do this reliably, it must do a GETATTR of the fs_status attribute that follows any interrogation of data or metadata within the file system in question. Often this is most conveniently done by appending such a GETATTR after all other operations that reference a given file system. When errors occur between reading file system data and performing such a GETATTR, care must be exercised to make sure that the data in question is not used before obtaining the proper fs_status value. In this connection, when an OPEN is done within such a versioned file system and the associated GETATTR of fs_status is not successfully completed, the open file in question must not be accessed until that fs_status is fetched.
The procedure above will ensure that before using any data from the file system the client has in hand a newly-fetched current version of the file system image. Multiple values for multiple requests in flight can be resolved by assembling them into the required partial order (and the elements should form a total order within it) and using the last. The client may then, when switching among file system instances, decline to use an instance which does not have an fss_type of STATUS4_VERSIONED or whose fss_version field is earlier than the last one obtained from the predecessor file system instance.
| TOC |
| TOC |
pNFS is an OPTIONAL feature within NFSv4.1; the pNFS feature set allows direct client access to the storage devices containing file data. When file data for a single NFSv4 server is stored on multiple and/or higher throughput storage devices (by comparison to the server's throughput capability), the result can be significantly better file access performance. The relationship among multiple clients, a single server, and multiple storage devices for pNFS (server and clients have access to all storage devices) is shown in this diagram:
+-----------+
|+-----------+ +-----------+
||+-----------+ | |
||| | NFSv4.1 + pNFS | |
+|| Clients |<------------------------------>| Server |
+| | | |
+-----------+ | |
||| +-----------+
||| |
||| |
||| Storage +-----------+ |
||| Protocol |+-----------+ |
||+----------------||+-----------+ Control |
|+-----------------||| | Protocol|
+------------------+|| Storage |------------+
+| Devices |
+-----------+
| Figure 68 |
In this model, the clients, server, and storage devices are responsible for managing file access. This is in contrast to NFSv4 without pNFS where it is primarily the server's responsibility; some of this responsibility may be delegated to the client under strictly specified conditions.
pNFS takes the form of OPTIONAL operations that manage protocol objects called 'layouts' which contain data location information. The layout is managed in a similar fashion as NFSv4.1 data delegations are managed. For example, the layout is leased, recallable and revocable. However, layouts are distinct abstractions and are manipulated with new operations. When a client holds a layout, it is granted the ability to access the data location directly using the location information specified in the layout.
There are interactions between layouts and other NFSv4.1 abstractions such as data delegations and byte-range locking. Delegation issues are discussed in Section 12.5.5 (Recalling a Layout). Byte range locking issues are discussed in Section 12.2.9 (Layout Iomode) and Section 12.5.1 (Guarantees Provided by Layouts).
The NFSv4.1 pNFS feature has been structured to allow for a variety of storage protocols to be defined and used. As noted in the diagram above, the storage protocol is the method used by the client to store and retrieve data directly from the storage devices. The NFSv4.1 protocol directly defines one storage protocol, the NFSv4.1 storage type, and its use.
Examples of other storage protocols that could be used with NFSv4.1's pNFS are:
It is possible that various storage protocols are available to both client and server and it may be possible that a client and server do not have a matching storage protocol available to them. Because of this, the pNFS server MUST support normal NFSv4.1 access to any file accessible by the pNFS feature; this will allow for continued interoperability between a NFSv4.1 client and server.
| TOC |
NFSv4.1's pNFS feature partitions the file system protocol into two parts: metadata and data. Where data is the contents of a file and metadata is "everything else". The metadata functionality is implemented by a metadata server that supports pNFS and the operations described in (Section 18 (NFSv4.1 Operations)). The data functionality is implemented by a storage device that supports the storage protocol. A subset (defined in Section 13.6 (Operations Sent to NFSv4.1 Data Servers)) of NFSv4.1 itself is one such storage protocol. New terms are introduced to the NFSv4.1 nomenclature and existing terms are clarified to allow for the description of the pNFS feature.
| TOC |
Information about a file system object, such as its name, location within the namespace, owner, ACL and other attributes. Metadata may also include storage location information and this will vary based on the underlying storage mechanism that is used.
| TOC |
An NFSv4.1 server which supports the pNFS feature. A variety of architectural choices exists for the metadata server and its use of what file system information is held at the server. Some servers may contain metadata only for the file objects that reside at the metadata server while file data resides on the associated storage devices. Other metadata servers may hold both metadata and a varying degree of file data.
| TOC |
An NFSv4.1 client that supports pNFS operations and supports at least one storage protocol or layout type for performing I/O to storage devices.
| TOC |
A storage device stores a regular file's data, but leaves metadata management to the metadata server. A storage device could be another NFSv4.1 server, an object storage device (OSD), a block device accessed over a SAN (e.g., either FiberChannel or iSCSI SAN), or some other entity.
| TOC |
A storage protocol is the protocol used between the pNFS client and the storage device to access the file data.
| TOC |
The control protocol is used by the exported file system between the metadata server and storage devices. Specification of such protocols is outside the scope of the NFSv4.1 protocol. Such control protocols would be used to control activities such as the allocation and deallocation of storage and the management of state required by the storage devices to perform client access control.
A particular control protocol is not REQUIRED by NFSv4.1 but requirements are placed on the control protocol for maintaining attributes like modify time, the change attribute, and the end-of-file (EOF) position.
| TOC |
A layout describes the mapping of a file's data to the storage devices that hold the data. A layout is said to belong to a specific layout type (data type layouttype4, see Section 3.3.13 (layouttype4)). The layout type allows for variants to handle different storage protocols, such as those associated with block/volume [31] (Black, D., Fridella, S., and J. Glasgow, “pNFS Block/Volume Layout,” April 2008.), object [30] (Halevy, B., Welch, B., and J. Zelenka, “Object-based pNFS Operations,” April 2008.), and file (Section 13 (PNFS: NFSv4.1 File Layout Type)) layout types. A metadata server, along with its control protocol, MUST support at least one layout type. A private sub-range of the layout type name space is also defined. Values from the private layout type range MAY be used for internal testing or experimentation.
As an example, layout of the file layout type could be an array of tuples (e.g., deviceID, file_handle), along with a definition of how the data is stored across the devices (e.g., striping). A block/volume layout might be an array of tuples that store <deviceID, block_number, block count> along with information about block size and the associated file offset of the block number. An object layout might be an array of tuples <deviceID, objectID> and an additional structure (i.e., the aggregation map) that defines how the logical byte sequence of the file data is serialized into the different objects. Note that the actual layouts are typically more complex than these simple expository examples.
Requests for pNFS-related operations will often specify a layout type. Examples of such operations are GETDEVICEINFO and LAYOUTGET. The response for these operations will include structures such a device_addr4 or a layout4, each of which includes a layout type within it. The layout type sent by the server MUST always be the same one requested by the client. When a client sends a response that includes a different layout type, the client SHOULD ignore the response and behave as if the server had returned an error response.
| TOC |
A layout defines how a file's data is organized on one or more storage devices. There are many potential layout types; each of the layout types are differentiated by the storage protocol used to access data and in the aggregation scheme that lays out the file data on the underlying storage devices. A layout is precisely identified by the following tuple: <client ID, filehandle, layout type, iomode, range>; where filehandle refers to the filehandle of the file on the metadata server.
It is important to define when layouts overlap and/or conflict with each other. For two layouts with overlapping byte ranges to actually overlap each other, both layouts must be of the same layout type, correspond to the same filehandle, and have the same iomode. Layouts conflict when they overlap and differ in the content of the layout (i.e., the storage device/file mapping parameters differ). Note that differing iomodes do not lead to conflicting layouts. It is permissible for layouts with different iomodes, pertaining to the same byte range, to be held by the same client. An example of this would be copy-on-write functionality for a block/volume layout type.
| TOC |
The layout iomode (data type layoutiomode4, see Section 3.3.20 (layoutiomode4)) indicates to the metadata server the client's intent to perform either just read operations or a mixture of I/O possibly containing read and write operations. For certain layout types, it is useful for a client to specify this intent at LAYOUTGET (Section 18.43 (Operation 50: LAYOUTGET - Get Layout Information)) time. For example, block/volume based protocols, block allocation could occur when a READ/WRITE iomode is specified. A special LAYOUTIOMODE4_ANY iomode is defined and can only be used for LAYOUTRETURN and CB_LAYOUTRECALL, not for LAYOUTGET. It specifies that layouts pertaining to both READ and READ/WRITE iomodes are being returned or recalled, respectively.
A storage device may validate I/O with regards to the iomode; this is dependent upon storage device implementation and layout type. Thus, if the client's layout iomode is inconsistent with the I/O being performed, the storage device may reject the client's I/O with an error indicating a new layout with the correct I/O mode should be fetched. For example, if a client gets a layout with a READ iomode and performs a WRITE to a storage device, the storage device is allowed to reject that WRITE.
The iomode does not conflict with OPEN share modes or lock requests; open mode and lock conflicts are enforced as they are without the use of pNFS, and are logically separate from the pNFS layout level. As well, open modes and locks are the preferred method for restricting user access to data files. For example, an OPEN of read, deny-write does not conflict with a LAYOUTGET containing an iomode of READ/WRITE performed by another client. Applications that depend on writing into the same file concurrently may use byte-range locking to serialize their accesses.
| TOC |
The device ID (data type deviceid4, see Section 3.3.14 (deviceid4)) names a group of storage devices. The scope of a device ID is per pair of client ID and layout type. In practice, a significant amount of information may be required to fully address a storage device. Rather than embedding all such information in a layout, layouts embed device IDs. The NFSv4.1 operation GETDEVICEINFO (Section 18.40 (Operation 47: GETDEVICEINFO - Get Device Information)) is used to retrieve the complete address information (including all device addresses for the device ID) regarding the storage device according to its layout type and device ID. For example, the address of an NFSv4.1 data server or of an object storage device could be an IP address and port. The address of a block storage device could be a volume label.
Clients cannot expect the mapping between a device ID and its storage device address(es) to persist across metadata server restart. See Section 12.7.4 (Recovery from Metadata Server Restart) for a description of how recovery works in that situation.
A device ID lives as long as there is a layout referring to the device ID. If there are no layouts referring to the device ID, the server is free to delete the device ID any time. Once a device ID is deleted by the server, the server MUST NOT reuse the device ID for the same layout type and client ID again. This requirement is feasible because the device ID is 16 bytes long, leaving sufficient room to store a generation number if server's implementation requires most of the rest of the device ID's content to be reused. This requirement is necessary because otherwise the race conditions between asynchronous notification of device ID addition and deletion would be too difficult to sort out.
Device ID to device address mappings are not leased, and can be changed at any time. (Note that while device ID to device address mappings are likely to change after the metadata server restarts, the server is not required to change the mappings.) A server has two choices for changing mappings. It can recall all layouts referring to the device ID or it can use a notification mechanism.
The NFSv4.1 protocol has no optimal way to recall all layouts that referred to a particular device ID (unless the server associates a single device ID with a single fsid or a single client ID; in which case, CB_LAYOUTRECALL has options for recalling all layouts associated with the fsid, client ID pair or just the client ID).
Via a notification mechanism (see Section 20.12 (Operation 14: CB_NOTIFY_DEVICEID - Notify device ID changes)), device ID to device address mappings can change over the duration of server operation without recalling or revoking the layouts that refer to device ID. The notification mechanism can also delete a device ID, but only if the client has no layouts referring to the device ID. A notification of a change to a device ID to device address mapping will immediately or eventually invalidate some or all of the device ID's mappings. The server MUST support notifications and the client must request them before they can be used. For further information about the notification types Section 20.12 (Operation 14: CB_NOTIFY_DEVICEID - Notify device ID changes).
| TOC |
NFSv4.1 has several operations that are needed for pNFS servers, regardless of layout type or storage protocol. These operations are all sent to a metadata server and summarized here. While pNFS is an OPTIONAL feature, if pNFS is implemented, some operations are REQUIRED in order to comply with pNFS. See Section 17 (Operations: REQUIRED, RECOMMENDED, or OPTIONAL).
These are the fore channel pNFS operations:
- GETDEVICEINFO.
- As noted previously (Section 12.2.10 (Device IDs)), GETDEVICEINFO (Section 18.40 (Operation 47: GETDEVICEINFO - Get Device Information)) returns the mapping of device ID to storage device address.
- GETDEVICELIST
- (Section 18.41 (Operation 48: GETDEVICELIST - Get All Device Mappings for a File System)), allows clients to fetch all device IDs for a specific file system.
- LAYOUTGET
- (Section 18.43 (Operation 50: LAYOUTGET - Get Layout Information)) is used by a client to get a layout for a file.
- LAYOUTCOMMIT
- (Section 18.42 (Operation 49: LAYOUTCOMMIT - Commit writes made using a layout)) is used to inform the metadata server of the client's intent to commit data which has been written to the storage device; the storage device as originally indicated in the return value of LAYOUTGET.
- LAYOUTRETURN
- (Section 18.44 (Operation 51: LAYOUTRETURN - Release Layout Information)) is used to return layouts for a file, an FSID and for client ID.
These are the backchannel pNFS operations:
- CB_LAYOUTRECALL
- (Section 20.3 (Operation 5: CB_LAYOUTRECALL - Recall Layout from Client)) recalls a layout or all layouts belonging to a file system, or all layouts belonging to a client ID.
- CB_RECALL_ANY
- (Section 20.6 (Operation 8: CB_RECALL_ANY - Keep any N recallable objects)), tells a client that it needs to return some number of recallable objects, including layouts, to the metadata server.
- CB_RECALLABLE_OBJ_AVAIL
- (Section 20.7 (Operation 9: CB_RECALLABLE_OBJ_AVAIL - Signal Resources for Recallable Objects)) tells a client that a recallable object that it was denied (in case of pNFS, a layout, denied by LAYOUTGET) due to resource exhaustion, is now available.
- CB_NOTIFY_DEVICEID
- (Section 20.12 (Operation 14: CB_NOTIFY_DEVICEID - Notify device ID changes)) Notifies the client of changes to device IDs.
| TOC |
A number of attributes specific to pNFS are listed and described in Section 5.12 (pNFS Attribute Definitions)
| TOC |
| TOC |
Layouts grant to the client the ability to access data located at a storage device with the appropriate storage protocol. The client is guaranteed the layout will be recalled when one of two things occur; either a conflicting layout is requested or the state encapsulated by the layout becomes invalid and this can happen when an event directly or indirectly modifies the layout. When a layout is recalled and returned by the client, the client continues with the ability to access file data with normal NFSv4.1 operations through the metadata server. Only the ability to access the storage devices is affected.
The requirement of NFSv4.1, that all user access rights MUST be obtained through the appropriate open, lock, and access operations, is not modified with the existence of layouts. Layouts are provided to NFSv4.1 clients and user access still follows the rules of the protocol as if they did not exist. It is a requirement that for a client to access a storage device, a layout must be held by the client. If a storage device receives an I/O for a byte range for which the client does not hold a layout, the storage device SHOULD reject that I/O request. Note that the act of modifying a file for which a layout is held, does not necessarily conflict with the holding of the layout that describes the file being modified. Therefore, it is the requirement of the storage protocol or layout type that determines the necessary behavior. For example, block/volume layout types require that the layout's iomode agree with the type of I/O being performed.
Depending upon the layout type and storage protocol in use, storage device access permissions may be granted by LAYOUTGET and may be encoded within the type-specific layout. For an example of storage device access permissions see an object based protocol such as [40] (Weber, R., “Object-Based Storage Device Commands (OSD),” July 2004.). If access permissions are encoded within the layout, the metadata server SHOULD recall the layout when those permissions become invalid for any reason; for example when a file becomes unwritable or inaccessible to a client. Note, clients are still required to perform the appropriate access operations with open, lock and access as described above. The degree to which it is possible for the client to circumvent these access operations and the consequences of doing so must be clearly specified by the individual layout type specifications. In addition, these specifications must be clear about the requirements and non-requirements for the checking performed by the server.
In the presence of pNFS functionality, mandatory file locks MUST behave as they would without pNFS. Therefore, if mandatory file locks and layouts are provided simultaneously, the storage device MUST be able to enforce the mandatory file locks. For example, if one client obtains a mandatory lock and a second client accesses the storage device, the storage device MUST appropriately restrict I/O for the byte range of the mandatory file lock. If the storage device is incapable of providing this check in the presence of mandatory file locks, the metadata server then MUST NOT grant layouts and mandatory file locks simultaneously.
| TOC |
A client obtains a layout with the LAYOUTGET operation. The metadata server will grant layouts of a particular type (e.g., block/volume, object, or file). The client selects an appropriate layout type that the server supports and the client is prepared to use. The layout returned to the client might not exactly match the requested byte range as described in Section 18.43.3 (DESCRIPTION). As needed a client may make multiple LAYOUTGET requests; these might result in multiple overlapping, non-conflicting layouts (see Section 12.2.8 (Layout)).
In order to get a layout, the client must first have opened the file via the OPEN operation. When a client has no layout on a file, it MUST present a stateid as returned by OPEN, a delegation stateid, or a byte-range lock stateid in the loga_stateid argument. A successful LAYOUTGET result includes a layout stateid. The first successful LAYOUTGET processed by the server using a non-layout stateid as an argument MUST have the "seqid" field of the layout stateid in the response set to one. Thereafter, the client uses a layout stateid (see Section 12.5.3 (Layout Stateid)) on future invocations of LAYOUTGET on the file, and the "seqid" MUST NOT be set to zero. Once the layout has been retrieved, it can be held across multiple OPEN and CLOSE sequences. Therefore, a client may hold a layout for a file that is not currently open by any user on the client. This allows for the caching of layouts beyond CLOSE.
The storage protocol used by the client to access the data on the storage device is determined by the layout's type. The client is responsible for matching the layout type with an available method to interpret and use the layout. The method for this layout type selection is outside the scope of the pNFS functionality.
Although the metadata server is in control of the layout for a file, the pNFS client can provide hints to the server when a file is opened or created about the preferred layout type and aggregation schemes. pNFS introduces a layout_hint (Section 5.12.4 (Attribute 63: layout_hint)) attribute that the client can set at file creation time to provide a hint to the server for new files. Setting this attribute separately, after the file has been created might make it difficult, or impossible, for the server implementation to comply.
Because the EXCLUSIVE4 createmode4 does not allow the setting of attributes at file creation time, NFSv4.1 introduces the EXCLUSIVE4_1 createmode4, which does allow attributes to be set at file creation time. In addition, if the session is created with persistent reply caches, EXCLUSIVE4_1 is neither necessary nor allowed. Instead, GUARDED4 both works better and is prescribed. Table 18 in Section 18.16.3 (DESCRIPTION), summarizes how a client is allowed to send an exclusive create.
| TOC |
As with all other stateids, the layout stateid consists of a "seqid" and "other" field. Once a layout stateid is changed, the "other" field will stay constant unless the stateid is revoked, or the client returns all layouts on the file and the server disposes of the stateid. The "seqid" field is initially set to one, and is never zero on any NFSv4.1 operation that uses layout stateids, whether it is a fore channel or backchannel operation. After the layout stateid is established, the server increments by one the value of the "seqid" in each subsequent LAYOUTGET and LAYOUTRETURN response, and in each CB_LAYOUTRECALL request.
Given the design goal of pNFS to provide parallelism, the layout stateid differs from other stateid types in that the client is expected to send LAYOUTGET and LAYOUTRETURN operations in parallel. The "seqid" value is used by the client to properly sort responses to LAYOUTGET and LAYOUTRETURN. The "seqid" is also used to prevent race conditions between LAYOUTGET and CB_LAYOUTRECALL. Given the processing rules differ from layout stateids and other stateid types, only the pNFS sections of this document should be considered to determine proper layout stateid handling.
Once the client receives a layout stateid, it MUST use the correct "seqid" for subsequent LAYOUTGET or LAYOUTRETURN operations. The correct "seqid" is defined as the highest "seqid" value from responses of fully processed LAYOUTGET or LAYOUTRETURN operations or arguments of a fully processed CB_LAYOUTRECALL operation. Since the server is incrementing the "seqid" value on each layout operation, the client may determine the order of operation processing by inspecting the "seqid" value. In the case of overlapping layout ranges, the ordering information will provide the client the knowledge of which layout ranges are held. Note that overlapping layout ranges may occur because of the client's specific requests or because the server is allowed to expand the range of a requested layout and notify the client in the LAYOUTRETURN results. Additional layout stateid sequencing requirements are provided in Section 12.5.5.2 (Sequencing of Layout Operations).
The client's receipt of a "seqid" is not sufficient for subsequent use. The client must fully process the operations before the "seqid" can be used. For LAYOUTGET results, if the client is not using the forgetful model (Section 12.5.5.1 (Layout Recall Callback Robustness)), it MUST first update its record of what ranges of the file's layout it has before using the seqid. For LAYOUTRETURN results, the client MUST delete the range from its record of what ranges of the file's layout it had before using the seqid. For CB_LAYOUTRECALL arguments, the client MUST send a response to the recall before using the seqid. The fundamental requirement in client processing is that the "seqid" is used to provide the order of processing. LAYOUTGET results may be processed in parallel. LAYOUTRETURN results may be processed in parallel. LAYOUTGET and LAYOUTRETURN responses may be processed in parallel as long as the ranges do not overlap. CB_LAYOUTRECALL request processing MUST be processed in "seqid" order at all times.
Once a client has no more layouts on a file, the layout stateid is no longer valid, and MUST NOT be used. Any attempt to use such a layout stateid will result in NFS4ERR_BAD_STATEID.
| TOC |
Allowing for varying storage protocols capabilities, the pNFS protocol does not require the metadata server and storage devices to have a consistent view of file attributes and data location mappings. Data location mapping refers to aspects such as which offsets store data as opposed to storing holes (see Section 13.4.4 (Sparse and Dense Stripe Unit Packing) for a discussion). Related issues arise for storage protocols where a layout may hold provisionally allocated blocks where the allocation of those blocks does not survive a complete restart of both the client and server. Because of this inconsistency, it is necessary to re-synchronize the client with the metadata server and its storage devices and make any potential changes available to other clients. This is accomplished by use of the LAYOUTCOMMIT operation.
The LAYOUTCOMMIT operation is responsible for committing a modified layout to the metadata server. The data should be written and committed to the appropriate storage devices before the LAYOUTCOMMIT occurs. The scope of the LAYOUTCOMMIT operation depends on the storage protocol in use. It is important to note that the level of synchronization is from the point of view of the client which sent the LAYOUTCOMMIT. The updated state on the metadata server need only reflect the state as of the client's last operation previous to the LAYOUTCOMMIT. It is not REQUIRED to maintain a global view that accounts for other clients' I/O that may have occurred within the same time frame.
For block/volume-based layouts, LAYOUTCOMMIT may require updating the block list that comprises the file and committing this layout to stable storage. For file-layouts synchronization of attributes between the metadata and storage devices primarily the size attribute is required.
The control protocol is free to synchronize the attributes before it receives a LAYOUTCOMMIT, however upon successful completion of a LAYOUTCOMMIT, state that exists on the metadata server that describes the file MUST be in sync with the state existing on the storage devices that comprise that file as of the issuing client's last operation. Thus, a client that queries the size of a file between a WRITE to a storage device and the LAYOUTCOMMIT may observe a size that does not reflect the actual data written.
The client MUST have a layout in order to issue LAYOUTCOMMIT.
| TOC |
The change and time_modify attributes may be updated by the server when the LAYOUTCOMMIT operation is processed. The reason for this is that some layout types do not support the update of these attributes when the storage devices process I/O operations. If client has a layout with the LAYOUTIOMODE4_RW iomode on the file, the client MAY provide a suggested value to the server for time_modify within the arguments to LAYOUTCOMMIT. Based on the layout type, the provided value may or may not be used. The server should sanity check the client provided values before they are used. For example, the server should ensure that time does not flow backwards. The client always has the option to set time_modify through an explicit SETATTR operation.
For some layout protocols, the storage device is able to notify the metadata server of the occurrence of an I/O and as a result the change and time_modify attributes may be updated at the metadata server. For a metadata server that is capable of monitoring updates to the change and time_modify attributes, LAYOUTCOMMIT processing is not required to update the change attribute; in this case the metadata server must ensure that no further update to the data has occurred since the last update of the attributes; file-based protocols may have enough information to make this determination or may update the change attribute upon each file modification. This also applies for the time_modify attribute. If the server implementation is able to determine that the file has not been modified since the last time_modify update, the server need not update time_modify at LAYOUTCOMMIT. At LAYOUTCOMMIT completion, the updated attributes should be visible if that file was modified since the latest previous LAYOUTCOMMIT or LAYOUTGET.
| TOC |
The size of a file may be updated when the LAYOUTCOMMIT operation is used by the client. One of the fields in the argument to LAYOUTCOMMIT is loca_last_write_offset; this field indicates the highest byte offset written but not yet committed with the LAYOUTCOMMIT operation. The data type of loca_last_write_offset is newoffset4 and is switched on a boolean value, no_newoffset, that indicates if a previous write occurred or not. If no_newoffset is FALSE, an offset is not given. If the client has a layout with LAYOUTIOMODE4_RW iomode on the file, with an lo_offset and lo_length that overlaps loca_last_write_offset, then the client MAY set no_newoffset to TRUE and provide an offset that will update the file size. Keep in mind that offset is not the same as length, though they are related. For example, a loca_last_write_offset value of zero means that one byte was written at offset zero, and so the length of the file is at least one byte.
The metadata server may do one of the following:
The method chosen to update the file's size will depend on the storage device's and/or the control protocol's capabilities. For example, if the storage devices are block devices with no knowledge of file size, the metadata server must rely on the client to set the last write offset appropriately.
The results of LAYOUTCOMMIT contain a new size value in the form of a newsize4 union data type. If the file's size is set as a result of LAYOUTCOMMIT, the metadata server must reply with the new size; otherwise the new size is not provided. If the file size is updated, the metadata server SHOULD update the storage devices such that the new file size is reflected when LAYOUTCOMMIT processing is complete. For example, the client should be able to READ up to the new file size.
The client can extend the length of a file or truncate a file by sending a SETATTR operation to the metadata server with the size attribute specified. If the size specified is larger than the current size of the file, the file is "zero extended", i.e., zeroes are implicitly added between the file's previous EOF and the new EOF. (In many implementations the zero extended region of the file consists of unallocated holes in the file.) When the client writes past EOF via WRITE, the SETATTR operation does not need to be used.
| TOC |
The LAYOUTCOMMIT argument contains a loca_layoutupdate field (Section 18.42.1 (ARGUMENT)) of data type layoutupdate4 (Section 3.3.18 (layoutupdate4)). This argument is a layout type-specific structure. The structure can be used to pass arbitrary layout type-specific information from the client to the metadata server at LAYOUTCOMMIT time. For example, if using a block/volume layout, the client can indicate to the metadata server which reserved or allocated blocks the client used or did not use. The content of loca_layoutupdate (field lou_body) need not be the same layout type-specific content returned by LAYOUTGET (Section 18.43.2 (RESULT)) in the loc_body field of the lo_content field, of the logr_layout field. The content of loca_layoutupdate is defined by the layout type specification and is opaque to LAYOUTCOMMIT.
| TOC |
Since a layout protects a client's access to a file via a direct client-storage-device path, a layout need only be recalled when it is semantically unable to serve this function. Typically, this occurs when the layout no longer encapsulates the true location of the file over the byte range it represents. Any operation or action, such as server driven restriping or load balancing, that changes the layout will result in a recall of the layout. A layout is recalled by the CB_LAYOUTRECALL callback operation (see Section 20.3 (Operation 5: CB_LAYOUTRECALL - Recall Layout from Client)) and returned with LAYOUTRETURN Section 18.44 (Operation 51: LAYOUTRETURN - Release Layout Information). The CB_LAYOUTRECALL operation may recall a layout identified by a byte range, all the layouts associated with a file system (FSID), or all layouts associated with a client ID. Section 12.5.5.2 (Sequencing of Layout Operations) discusses sequencing issues surrounding the getting, returning, and recalling of layouts.
An iomode is also specified when recalling a layout. Generally, the iomode in the recall request must match the layout being returned; for example, a recall with an iomode of LAYOUTIOMODE4_RW should cause the client to only return LAYOUTIOMODE4_RW layouts and not LAYOUTIOMODE4_READ layouts. However, a special LAYOUTIOMODE4_ANY enumeration is defined to enable recalling a layout of any iomode; in other words, the client must return both read-only and read/write layouts.
A REMOVE operation SHOULD cause the metadata server to recall the layout to prevent the client from accessing a non-existent file and to reclaim state stored on the client. Since a REMOVE may be delayed until the last close of the file has occurred, the recall may also be delayed until this time. After the last reference on the file has been released and the file has been removed, the client should no longer be able to perform I/O using the layout. In the case of a files based layout, the data server SHOULD return NFS4ERR_STALE in response to any operation on the removed file.
Once a layout has been returned, the client MUST NOT send I/Os to the storage devices for the file, byte range, and iomode represented by the returned layout. If a client does send an I/O to a storage device for which it does not hold a layout, the storage device SHOULD reject the I/O.
Although pNFS does not alter the file data caching capabilities of clients, or their semantics, it recognizes that some clients may perform more aggressive write-behind caching to optimize the benefits provided by pNFS. However, write-behind caching may negatively affect the latency in returning a layout in response to a CB_LAYOUTRECALL; this is similar to file delegations and the impact that file data caching has on DELEGRETURN. Client implementations SHOULD limit the amount of unwritten data they have outstanding at any one time in order to prevent excessively long responses to CB_LAYOUTRECALL. Once a layout is recalled, a server MUST wait one lease period before taking further action. As soon as a lease period has past, the server may choose to fence the client's access to the storage devices if the server perceives the client has taken too long to return a layout. However, just as in the case of data delegation and DELEGRETURN, the server may choose to wait given that the client is showing forward progress on its way to returning the layout. This forward progress can take the form of successful interaction with the storage devices or sub-portions of the layout being returned by the client. The server can also limit exposure to these problems by limiting the byte ranges initially provided in the layouts and thus the amount of outstanding modified data.
| TOC |
It has been assumed thus far that pNFS client state for a file exactly matches the pNFS server state for that file and client regarding layout ranges and iomode. This assumption leads to the implication that any callback results in a LAYOUTRETURN or set of LAYOUTRETURNs that exactly match the range in the callback, since both client and server agree about the state being maintained. However, it can be useful if this assumption does not always hold. For example:
Thus, in light of the above, it is useful for a server to be able to send callbacks for layout ranges it has not granted to a client, and for a client to return ranges it does not hold. A pNFS client MUST always return layouts that comprise the full range specified by the recall. Note, the full recalled layout range need not be returned as part of a single operation, but may be returned in portions. This allows the client to stage the flushing of dirty data, layout commits, and returns. Also, it indicates to the metadata server that the client is making progress.
When a layout is returned, the client MUST NOT have any outstanding I/O requests to the storage devices involved in the layout. Rephrasing, the client MUST NOT return the layout while it has outstanding I/O requests to the storage device.
Even with this requirement for the client, it is possible that I/O requests may be presented to a storage device no longer allowed to perform them. Since the server has no strict control as to when the client will return the layout, the server may later decide to unilaterally revoke the client's access to the storage devices as provided by the layout. In choosing to revoke access, the server must deal with the possibility of lingering I/O request; those outstanding I/O requests are still in flight to storage devices identified by the revoked layout. All layout specifications MUST define whether unilateral layout revocation by the metadata server is supported; if it is, the specification must also describe how lingering writes are processed. For example, storage devices identified by the revoked layout could be fenced off from the client that held the layout.
In order to ensure client/server convergence with regard to layout state, the final LAYOUTRETURN operation in a sequence of LAYOUTRETURN operations for a particular recall, MUST specify the entire range being recalled, echoing the recalled layout type, iomode, recall/return type (FILE, FSID, or ALL), and byte range; even if layouts pertaining to partial ranges were previously returned. In addition, if the client holds no layouts that overlaps the range being recalled, the client should return the NFS4ERR_NOMATCHING_LAYOUT error code to CB_LAYOUTRECALL. This allows the server to update its view of the client's layout state.
| TOC |
As with other stateful operations, pNFS requires the correct sequencing of layout operations. PNFS uses the "seqid" in the layout stateid to provide the correct sequencing between regular operations and callbacks. It is the server's responsibility to avoid inconsistencies regarding the layouts provided and the client's responsibility to properly serialize its layout requests and layout returns.
| TOC |
One critical issue with regard to layout operations sequencing concerns callbacks. The protocol must defend against races between the reply to a LAYOUTGET or LAYOUTRETURN operation and a subsequent CB_LAYOUTRECALL. A client MUST NOT process a CB_LAYOUTRECALL that implies one or more outstanding LAYOUTGET or LAYOUTRETURN operations to which the client has not yet received a reply. The client detects such a CB_LAYOUTRECALL by examining the "seqid" field of the recall's layout stateid. If the "seqid" is not one higher than what the client currently has recorded, and the client has at least one LAYOUTGET and/or LAYOUTRETURN operation outstanding, the client knows the server sent the CB_LAYOUTRECALL after sending a response to an outstanding LAYOUTGET or LAYOUTRETURN. The client MUST wait before processing such a CB_LAYOUTRECALL until it processes all replies for outstanding LAYOUTGET and LAYOUTRETURN operations for the corresponding file with seqid less than the seqid given by CB_LAYOUTRECALL (lor_stateid, see Section 20.3 (Operation 5: CB_LAYOUTRECALL - Recall Layout from Client).)
In addition to the seqid-based mechanism, Section 2.10.5.3 (Resolving Server Callback Races) describes the sessions mechanism for allowing the client to detect callback race conditions and delay processing such a CB_LAYOUTRECALL. The server MAY reference conflicting operations in the CB_SEQUENCE that precedes the CB_LAYOUTRECALL. Because the server has already sent replies for these operations before issuing the callback, the replies may race with the CB_LAYOUTRECALL. The client MUST wait for all the referenced calls to complete and update its view of the layout state before processing the CB_LAYOUTRECALL.
| TOC |
The protocol allows the client to send concurrent LAYOUTGET and LAYOUTRETURN operations to the server. The protocol does not provide any means for the server to process the requests in the same order in which they were created. However, through the use of the "seqid" field in the layout stateid, the client can determine the order in which parallel outstanding operations were processed by the server. Thus, when a layout retrieved by an outstanding LAYOUTGET operation intersects with a layout returned by an outstanding LAYOUTRETURN on the same file, the order in which the two conflicting operations are processed determines the final state of the overlapping layout. The order is determined by the "seqid" returned in each operation: the operation with the higher seqid was executed later.
It is permissible for the client to send in parallel multiple LAYOUTGET operations for the same file or multiple LAYOUTRETURN operations for the same file, and a mix of both.
It is permissible for the client to use the current stateid (see Section 16.2.3.1.2 (Current Stateid)) for LAYOUTGET operations for example when compounding LAYOUTGETs or compounding OPEN and LAYOUTGETs. It is also permissible to use the current stateid when compounding LAYOUTRETURNs.
It is permissible for the client to use the current stateid when combining LAYOUTRETURN and LAYOUTGET operations for the same file in the same COMPOUND request since the server MUST process these in order. However, if a client does send such COMPOUND requests, it MUST NOT have more than one outstanding for the same file at the same time and MUST NOT have other LAYOUTGET or LAYOUTRETURN operations outstanding at the same time for that same file.
| TOC |
Consider a pNFS client that has sent a LAYOUTGET and before it receives the reply to LAYOUTGET, it receives a CB_LAYOUTRECALL for the same file with an overlapping range. There are two possibilities, which the client can distinguish via the layout stateid in the recall.
If these possibilities cannot be distinguished, a deadlock could result, as the client must wait for the LAYOUTGET response before processing the recall in the first case, but that response will not arrive until after the recall is processed in the second case. Note that in the first case, the "seqid" in the layout stateid of the recall is two greater than what the client has recorded and in the second case, the "seqid" is one greater than what the client has recorded. This allows the client to disambiguate between the two cases. The client thus knows precisely which possibility applies.
In case 1 the client knows it needs to wait for the LAYOUTGET response before processing the recall (or the client can return NFS4ERR_DELAY).
In case 2 the client will not wait for the LAYOUTGET response before processing the recall, because waiting would cause deadlock. Therefore, the action at the client will only require waiting in the case that the client has not yet seen the server's earlier responses to the LAYOUTGET operation(s).
The recall process can be considered completed when the final LAYOUTRETURN operation for the recalled range is completed. The LAYOUTRETURN uses the layout stateid (with seqid) specified in CB_LAYOUTRECALL. If the client uses multiple LAYOUTRETURNs in processing the recall, the first LAYOUTRETURN will use the layout stateid as specified in CB_LAYOUTRECALL. Subsequent LAYOUTRETURNs will use the highest seqid as is the usual case.
| TOC |
Consider a race from the metadata server's point of view. The metadata server has sent a CB_LAYOUTRECALL and receives an overlapping LAYOUTGET for the same file before the LAYOUTRETURN(s) that respond to the CB_LAYOUTRECALL. There are three cases:
| TOC |
The rules for layout stateid processing differ from other stateids in the protocol because the "seqid" value cannot be zero and the stateid's "seqid" value changes in a CB_LAYOUTRECALL operation. The non-zero requirement combined with the inherent parallelism of layout operations means that a set of LAYOUTGET and LAYOUTRETURN operations may contain the same value for "seqid". The server uses a slightly modified version of the modulo arithmetic as described in Section 2.10.5.1 (Slot Identifiers and Reply Cache) when incrementing the layout stateid's "seqid". The modification to that modulo arithmetic description is to not use zero. The modulo arithmetic is also used for the comparisons of "seqid" values in the processing of CB_LAYOUTRECALL events as described above in Section 12.5.5.2.1.3 (Server Considerations).
Just as the server validates the "seqid" in the event of CB_LAYOUTRECALL usage, as described in Section 12.5.5.2.1.3 (Server Considerations), the server also validates the "seqid" value to ensure that it is within an appropriate range. This range represents the degree of parallelism the server supports for layout stateids. If the client is sending multiple layout operations to the server in parallel, by definition, the "seqid" value in the supplied stateid will not be the current "seqid" as held by the server. The range of parallelism spans from the highest or current "seqid" to a "seqid" value in the past. To assist in the discussion, the server's current "seqid" value for a layout stateid is defined as: SERVER_CURRENT_SEQID. The lowest "seqid" value that is acceptable to the server is represented by PAST_SEQID. And the value for the range of valid "seqid"s or range of parallelism is VALID_SEQID_RANGE. Therefore, the following holds: VALID_SEQID_RANGE = SERVER_CURRENT_SEQID - PAST_SEQID. In the following, all arithmetic is the modulo arithmetic as described above.
The server MUST support a minimum VALID_SEQID_RANGE. The minimum is defined as: VALID_SEQID_RANGE = summation of 1..N of (ca_maxoperations(i) - 1) where N is the number of session fore channels and ca_maxoperations(i) is the value of the ca_maxoperations returned from CREATE_SESSION of the i'th session. The reason for minus 1 is to allow for the required SEQUENCE operation. The server MAY support a VALID_SEQID_RANGE value larger than the minimum. The maximum VALID_SEQID_RANGE is (2 ^ 32 - 2) (accounts for 0 not being a valid "seqid" value).
If the server finds the "seqid" is zero, the NFS4ERR_BAD_STATEID error is returned to the client. The server further validates the "seqid" to ensure it is within the range of parallelism, VALID_SEQID_RANGE. If the "seqid" value is outside of that range, the error NFS4ERR_OLD_STATEID is returned to the client. Upon receipt of NFS4ERR_OLD_STATEID, the client updates the stateid in the layout request based on processing of other layout requests and re-sends the operation to the server.
| TOC |
PNFS supports recalling and returning all layouts that are for files belonging to a particular fsid (LAYOUTRECALL4_FSID, LAYOUTRETURN4_FSID) or client ID (LAYOUTRECALL4_ALL, LAYOUTRETURN4_ALL). There are no "bulk" stateids, so detection of races via the seqid is not possible. The server MUST NOT initiate bulk recall while another recall is in progress, or the corresponding LAYOUTRETURN is in progress or pending. In the event the server sends a bulk recall while the client has pending or in progress LAYOUTRETURN, CB_LAYOUTRECALL, or LAYOUTGET, the client returns NFS4ERR_DELAY. In the event the client sends a LAYOUTGET or LAYOUTRETURN while a bulk recall is in progress, the server returns NFS4ERR_RECALLCONFLICT. If the client sends a LAYOUTGET or LAYOUTRETURN after the server receives NFS4ERR_DELAY from a bulk recall, then to ensure forward progress, the server MAY return NFS4ERR_RECALLCONFLICT.
Once a CB_LAYOUTRECALL of LAYOUTRECALL4_ALL is sent, the server MUST NOT allow the client to use any layout stateid except for LAYOUTCOMMIT operations. Once the client receives a CB_LAYOUTRECALL of LAYOUTRECALL4_ALL, it MUST NOT use any layout stateid except for LAYOUTCOMMIT operations. Once a LAYOUTRETURN of LAYOUTRETURN4_ALL is sent, all layout stateids granted to the client ID are freed. The client MUST NOT use the layout stateids again. It MUST use LAYOUTGET to obtain new layout stateids.
Once a CB_LAYOUTRECALL of LAYOUTRECALL4_FSID is sent, the server MUST NOT allow the client to use any layout stateid that refers to a file with the specified fsid except for LAYOUTCOMMIT operations. Once the client receives a CB_LAYOUTRECALL of LAYOUTRECALL4_ALL, it MUST NOT use any layout stateid that refers to a file with the specified fsid except for LAYOUTCOMMIT operations. Once a LAYOUTRETURN of LAYOUTRETURN4_FSID is sent, all layout stateids granted to the referenced fsid are freed. The client MUST NOT use the layout stateids for files with the referenced fsid again. It MUST use LAYOUTGET to obtain new layout stateids files with the referenced fsid.
If the server has sent a bulk CB_LAYOUTRECALL, and receives a LAYOUTGET, or a LAYOUTRETURN with a stateid, the server MUST return NFS4ERR_RECALLCONFLICT. If the server has sent a bulk CB_LAYOUTRECALL, and receives a LAYOUTRETURN with an lr_returntype that is not equal to the lor_recalltype of the CB_LAYOUTRECALL, the server MUST return NFS4ERR_RECALLCONFLICT.
| TOC |
Parallel NFS permits servers to revoke layouts from clients that fail to response to recalls and/or fail to renew their lease in time. Whether the server revokes the layout or not depends on the layout type, and what actions are taken with respect to the client's I/O to data servers is also layout type specific.
| TOC |
Asynchronous writes written through the metadata server may be propagated lazily to the storage devices. For data written asynchronously through the metadata server, a client performing a read at the appropriate storage device is not guaranteed to see the newly written data until a COMMIT occurs at the metadata server. While the write is pending, reads to the storage device may give out either the old data, the new data, or a mixture of new and old. Upon completion of a synchronous WRITE or COMMIT (for asynchronously written data), the metadata server MUST ensure that storage devices give out the new data and that the data has been written to stable storage. If the server implements its storage in any way such that it cannot obey these constraints, then it MUST recall the layouts to prevent reads being done that cannot be handled correctly. Note that the layouts MUST be recalled prior to the server responding to the associated WRITE operations.
| TOC |
This section describes the operations flow taken by a pNFS client to a metadata server and storage device.
When a pNFS client encounters a new FSID, it sends a GETATTR to the NFSv4.1 server for the fs_layout_type (Section 5.12.1 (Attribute 62: fs_layout_type)) attribute. If the attribute returns at least one layout type, and the layout types returned are among the set supported by the client, the client knows that pNFS is a possibility for the file system. If, from the server that returned the new FSID, the client does not have a client ID that came from an EXCHANGE_ID result that returned EXCHGID4_FLAG_USE_PNFS_MDS, it MUST send an EXCHANGE_ID to the server with the EXCHGID4_FLAG_USE_PNFS_MDS bit set. If the server's response does not have EXCHGID4_FLAG_USE_PNFS_MDS, then contrary to what the fs_layout_type attribute said, the server does not support pNFS, and the client will not be able use pNFS to that server; in this case, the server MUST return NFS4ERR_NOTSUPP in response to any pNFS operation.
The client then creates a session, requesting a persistent session, so that exclusive creates can be done with single round trip via the createmode4 of GUARDED4. If the session ends up not being persistent, the client will use EXCLUSIVE4_1 for exclusive creates.
If a file is to be created on a pNFS enabled file system, the client uses the OPEN operation. With the normal set of attributes that may be provided upon OPEN used for creation, there is an OPTIONAL layout_hint attribute. The client's use of layout_hint allows the client to express its preference for a layout type and its associated layout details. The use of a createmode4 of UNCHECKED4, GUARDED4, or EXCLUSIVE4_1 will allow the client to provide the layout_hint attribute at create time. The client MUST NOT use EXCLUSIVE4 (see Table 18). The client is RECOMMENDED to combine a GETATTR operation after the OPEN within the same COMPOUND. The GETATTR may then retrieve the layout_type attribute for the newly created file. The client will then know what layout type the server has chosen for the file and therefore what storage protocol the client must use.
If the client wants to open an existing file, then it also includes a GETATTR to determine what layout type the file supports.
The GETATTR in either the file creation or plain file open case can also include the layout_blksize and layout_alignment attributes so that the client can determine optimal offsets and lengths for I/O on the file.
Assuming the client supports the layout type returned by GETATTR and it chooses to use pNFS for data access, it then sends LAYOUTGET using the filehandle and stateid returned by OPEN, specifying the range it wants to do I/O on. The response is a layout, which may be a subset of the range for which the client asked. It also includes device IDs and a description of how data is organized (or in the case of writing, how data is to be organized) across the devices. The device IDs and data description are encoded in a format that is specific to the layout type, but the client is expected to understand.
When the client wants to send an I/O, it determines which device ID it needs to send the I/O command to by examining the data description in the layout. It then sends a GETDEVICEINFO to find the device address(es) of the device ID. The client then sends the I/O request one of device ID's device addresses, using the storage protocol defined for the layout type. Note that if a client has multiple I/Os to send, these I/O requests may be done in parallel.
If the I/O was a WRITE, then at some point the client may want to use LAYOUTCOMMIT to commit the modification time and the new size of the file (if it believes it extended the file size) to the metadata server and the modified data to the file system.
| TOC |
Recovery is complicated by the distributed nature of the pNFS protocol. In general, crash recovery for layouts is similar to crash recovery for delegations in the base NFSv4.1 protocol. However, the client's ability to perform I/O without contacting the metadata server introduces subtleties that must be handled correctly if the possibility of file system corruption is to be avoided.
| TOC |
Client recovery for layouts is similar to client recovery for other lock and delegation state. When an pNFS client restarts, it will lose all information about the layouts that it previously owned. There are two methods by which the server can reclaim these resources and allow otherwise conflicting layouts to be provided to other clients.
The first is through the expiry of the client's lease. If the client recovery time is longer than the lease period, the client's lease will expire and the server will know that state may be released. For layouts the server may release the state immediately upon lease expiry or it may allow the layout to persist awaiting possible lease revival, as long as no other layout conflicts.
The second is through the client restarting in less time than it takes for the lease period to expire. In such a case, the client will contact the server through the standard EXCHANGE_ID protocol. The server will find that the client's co_ownerid matches the co_ownerid of the previous client invocation, but that the verifier is different. The server uses this as a signal to release all layout state associated with the client's previous invocation. In this scenario, the data written by the client but not covered by a successful LAYOUTCOMMIT is in an undefined state; it may have been written or it may now be lost. This is acceptable behavior and it is the client's responsibility to use LAYOUTCOMMIT to achieve the desired level of stability.
| TOC |
If a client believes its lease has expired, it MUST NOT send I/O to the storage device until it has validated its lease. The client can send a SEQUENCE operation to the metadata server. If the SEQUENCE operation is successful, but sr_status_flag has SEQ4_STATUS_EXPIRED_ALL_STATE_REVOKED, SEQ4_STATUS_EXPIRED_SOME_STATE_REVOKED, or SEQ4_STATUS_ADMIN_STATE_REVOKED set, the client MUST NOT use currently held layouts. The client has two choices to recover from the lease expiration. First, for all modified but uncommitted data, write it to the metadata server using the FILE_SYNC4 flag for the WRITEs or WRITE and COMMIT. Second, the client reestablishes a client ID and session with the server and obtain new layouts and device ID to device address mappings for the modified data ranges and then write the data to the storage devices with the newly obtained layouts.
If sr_status_flags from the metadata server has SEQ4_STATUS_RESTART_RECLAIM_NEEDED set (or SEQUENCE returns NFS4ERR_STALE_CLIENTID, or SEQUENCE returns NFS4ERR_BAD_SESSION and CREATE_SESSION returns NFS4ERR_STALE_CLIENTID) then the metadata server has restarted, and the client SHOULD recover using the methods described in Section 12.7.4 (Recovery from Metadata Server Restart).
If sr_status_flags from the metadata server has SEQ4_STATUS_LEASE_MOVED set, then the client recovers by following the procedure described in Section 11.7.7.1 (Leases and File System Transitions). After that, the client may get an indication that the layout state was not moved with the file system. The client recovers as in the other applicable situations discussed in Paragraph 1 or Paragraph 2 of this section.
If sr_status_flags reports no loss of state, then the lease for the layouts the client has are valid and renewed, and the client can once again send I/O requests to the storage devices.
While clients SHOULD NOT send I/Os to storage devices that may extend past the lease expiration time period, this is not always possible; for example, an extended network partition that starts after the I/O is sent and does not heal until the I/O request is received by the storage device. Thus the metadata server and/or storage devices are responsible for protecting themselves from I/Os that are sent before the lease expires, but arrive after the lease expires. See Section 12.7.3 (Dealing with Loss of Layout State on the Metadata Server).
| TOC |
This is a description of the case where all of the following are true:
The metadata server and its storage devices MUST solve this by fencing the client. In other words, prevent the execution of I/O operations from the client to the storage devices after layout state loss. The details of how fencing is done are specific to the layout type. The solution for NFSv4.1 file-based layouts is described in (Section 13.11 (Layout Revocation and Fencing)), and for other layout types in their respective external specification documents.
| TOC |
The pNFS client will discover that the metadata server has restarted via the methods described in Section 8.4.2 (Server Failure and Recovery) and discussed in a pNFS-specific context in Paragraph 2, of Section 12.7.2 (Dealing with Lease Expiration on the Client). The client MUST stop using layouts and delete the device ID to device address mappings it previously received from the metadata server. Having done that, if the client wrote data to the storage device without committing the layouts via LAYOUTCOMMIT, then the client has additional work to do in order to have the client, metadata server and storage device(s) all synchronized on the state of the data.
| TOC |
Some of the recovery scenarios thus far noted that some operations, namely WRITE and LAYOUTGET might be permitted during the metadata server's grace period. The metadata server may allow these operations during its grace period. For LAYOUTGET, the metadata server must reliably determine that servicing such a request will not conflict with an impending LAYOUTCOMMIT reclaim request. For WRITE, it must reliably determine that it will not conflict with an impending OPEN; or a LOCK where the file has mandatory file locking enabled.
As mentioned previously, some operations, namely WRITE and LAYOUTGET may be rejected during the metadata server's grace period, because to provide simple, valid handling during the grace period, the easiest method is to simply reject all non-reclaim pNFS requests and WRITE operations by returning the NFS4ERR_GRACE error. However, depending on the storage protocol (which is specific to the layout type) and metadata server implementation, the metadata server may be able to determine that a particular request is safe. For example, a metadata server may save provisional allocation mappings for each file to stable storage, as well as information about potentially conflicting OPEN share modes and mandatory byte-range locks that might have been in effect at the time of restart, and use this information during the recovery grace period to determine that a WRITE request is safe.
| TOC |
Recovery from storage device restart is mostly dependent upon the layout type in use. However, there are a few general techniques a client can use if it discovers a storage device has crashed while holding modified, uncommitted data that was asynchronously written. First and foremost, it is important to realize that the client is the only one which has the information necessary to recover non-committed data; since, it holds the modified data and probably nothing else does. Second, the best solution is for the client to err on the side of caution and attempt to re-write the modified data through another path.
The client SHOULD immediately write the data to the metadata server, with the stable field in the WRITE4args set to FILE_SYNC4. Once it does this, there is no need to wait for the original storage device.
| TOC |
If the same physical hardware is used to implement both a metadata server and storage device, then the same hardware entity is to be understood to be implementing two distinct roles and it is important that it be clearly understood on behalf of which role the hardware is executing at any given time.
Two sub-cases can be distinguished.
| TOC |
pNFS separates file system metadata and data and provides access to both. There are pNFS-specific operations (listed in Section 12.3 (pNFS Operations)) that provide access to the metadata; all existing NFSv4.1 conventional (non-pNFS) security mechanisms and features apply to accessing the metadata. The combination of components in a pNFS system (see Figure 68) is required to preserve the security properties of NFSv4.1 with respect to an entity accessing storage device from a client, including security countermeasures to defend against threats that NFSv4.1 provides defenses for in environments where these threats are considered significant.
In some cases, the security countermeasures for connections to storage devices may take the form of physical isolation or a recommendation not to use pNFS in an environment. For example, it may be impractical to provide confidentiality protection for some storage protocols to protect against eavesdropping; in environments where eavesdropping on such protocols is of sufficient concern to require countermeasures, physical isolation of the communication channel (e.g., via direct connection from client(s) to storage device(s)) and/or a decision to forego use of pNFS (e.g., and fall back to conventional NFSv4.1) may be appropriate courses of action.
Where communication with storage devices is subject to the same threats as client to metadata server communication, the protocols used for that communication need to provide security mechanisms as strong as or no weaker than those available via RPSEC_GSS for NFSv4.1.
pNFS implementations MUST NOT remove NFSv4.1's access controls. The combination of clients, storage devices, and the metadata server are responsible for ensuring that all client to storage device file data access respects NFSv4.1's ACLs and file open modes. This entails performing both of these checks on every access in the client, the storage device, or both (as applicable; when the storage device is an NFSv4.1 server, the storage device is ultimately responsible for controlling access). If a pNFS configuration performs these checks only in the client, the risk of a misbehaving client obtaining unauthorized access is an important consideration in determining when it is appropriate to use such a pNFS configuration. Such layout types SHOULD NOT be used when client-only access checks do not provide sufficient assurance that NFSv4.1 access control is being applied correctly.
| TOC |
This section describes the semantics and format of NFSv4.1 file-based layouts for pNFS. NFSv4.1 file-based layouts uses the LAYOUT4_NFSV4_1_FILES layout type. The LAYOUT4_NFSV4_1_FILES type defines striping data across multiple NFSv4.1 data servers.
| TOC |
Sessions are a REQUIRED feature of NFSv4.1, and this extends to both the metadata server and file-based (NFSv4.1-based) data servers.
The role a server plays in pNFS is determined by the result it returns from EXCHANGE_ID. The roles are:
The client MAY request zero or more of EXCHGID4_FLAG_USE_NON_PNFS, EXCHGID4_FLAG_USE_PNFS_DS, or EXCHGID4_FLAG_USE_PNFS_MDS, even though some combinations (e.g. EXCHGID4_FLAG_USE_NON_PNFS | EXCHGID4_FLAG_USE_PNFS_MDS) are contradictory. The server however MUST only return the following acceptable combinations:
| Acceptable Results from EXCHANGE_ID |
|---|
| EXCHGID4_FLAG_USE_PNFS_MDS |
| EXCHGID4_FLAG_USE_PNFS_MDS | EXCHGID4_FLAG_USE_PNFS_DS |
| EXCHGID4_FLAG_USE_PNFS_DS |
| EXCHGID4_FLAG_USE_NON_PNFS |
| EXCHGID4_FLAG_USE_PNFS_DS | EXCHGID4_FLAG_USE_NON_PNFS |
As the above table implies, a server can have one or two roles. A server can be both a metadata server and a data server or it can be both a data server and non-metadata server. In addition to returning two roles in EXCHANGE_ID's results, and thus serving both roles via a common client ID, a server can serve two roles by returning a unique client ID and server owner for each role in each of two EXCHANGE_ID results, with each result indicating each role.
In the case of a server with concurrent PNFS roles that are served by a common client ID, if the EXCHANGE_ID request from the client has zero or a combination of the bits set in eia_flags, the server result should set bits which represent the higher of the acceptable combination of the server roles, with a preference to match the roles requested by the client. Thus if a client request has (EXCHGID4_FLAG_USE_NON_PNFS | EXCHGID4_FLAG_USE_PNFS_MDS | EXCHGID4_FLAG_USE_PNFS_DS) flags set, and the server is both a metadata server and a data server, serving both the roles by a common client ID, the server SHOULD return with (EXCHGID4_FLAG_USE_PNFS_MDS | EXCHGID4_FLAG_USE_PNFS_DS) set.
In the case of a server that has multiple concurrent PNFS roles, each role served by a unique client ID, if the client specifies zero or a combination of roles in the request, the server results SHOULD return only one of the roles from the combination specified by the client request. If the role specified by the server result does not match the intended use by the client, the client should send the EXCHANGE_ID specifying just the interested PNFS role.
If a pNFS metadata client gets a layout that refers it to an NFSv4.1 data server, it needs a client ID on that data server. If it does not yet have a client ID from the server that had the EXCHGID4_FLAG_USE_PNFS_DS flag set in the EXCHANGE_ID results, then the client must send an EXCHANGE_ID to the data server, using the same co_ownerid as it sent to the metadata server, with the EXCHGID4_FLAG_USE_PNFS_DS flag set in the arguments. If the server's EXCHANGE_ID results have EXCHGID4_FLAG_USE_PNFS_DS set, then the client may use the client ID to create sessions that will exchange pNFS data operations. The client ID returned by the data server has no relationship with the client ID returned by a metadata server unless the client IDs are equal and the server owners and server scopes of the data server and metadata server are equal.
In NFSv4.1, the session id in the SEQUENCE operation implies the client ID, which in turn might be used by the server to map the stateid to the right client/server pair. However, when a data server is presented with a READ or WRITE operation with a stateid, because the stateid is associated with client ID on a metadata server, and because the session id in the preceding SEQUENCE operation is tied to the client ID of the data server, the data server has no obvious way to determine the metadata server from the COMPOUND procedure, and thus has no way to validate the stateid. One RECOMMENDED approach is for pNFS servers to encode metadata server routing and/or identity information in the data server filehandles as returned in the layout.
If metadata server routing and/or identity information is encoded in data server filehandles, when the metadata server identity or location changes, the data server filehandles it gave out must become invalid (stale), and so the metadata server must first recall the layouts. Invalidating a data server filehandle does not render the NFS client's data cache invalid. The client's cache should map a data server filehandle to a metadata server filehandle, and a metadata server filehandle to cached data.
If a server is both a metadata server and a data server, the server might need to distinguish operations on files that are directed to the metadata server from those that are directed to the data server. It is RECOMMENDED that the values of the filehandles returned by the LAYOUTGET operation to be different than the value of the filehandle returned by the OPEN of the same file.
Another scenario is for the metadata server and the storage device to be distinct from one client's point of view, and the roles reversed from another client's point of view. For example, in the cluster file system model, a metadata server to one client may be a data server to another client. If NFSv4.1 is being used as the storage protocol, then pNFS servers need to encode the values of filehandles according to their specific roles.
| TOC |
Section 2.10.9.2 (Obligations of the Client) states that a client has to keep its lease renewed in order to prevent a session from being deleted by the server. If the reply to EXCHANGE_ID has just the EXCHGID4_FLAG_USE_PNFS_DS role set, then as noted in Section 13.6 (Operations Sent to NFSv4.1 Data Servers) the client will not be able to determine the data server's lease_time attribute, because GETATTR will not be permitted. Instead, the rule is that any time a client receives a layout referring it to a data server that returns just the EXCHGID4_FLAG_USE_PNFS_DS role, the client MAY assume that the lease_time attribute from the metadata server that returned the layout applies to the data server. Thus the data server MUST be aware of the values of all lease_time attributes of all metadata servers it is providing I/O for, and MUST use the maximum of all such lease_time values as the lease interval for all client IDs and sessions established on it.
For example, if one metadata server has a lease_time attribute of 20 seconds, and a second metadata server has a lease_time attribute of 10 seconds, then if both servers return layouts that refer to an EXCHGID4_FLAG_USE_PNFS_DS-only data server, the data server MUST renew a client's lease if the interval between two SEQUENCE operations on different COMPOUND requests is less than 20 seconds.
| TOC |
The following definitions apply to the LAYOUT4_NFSV4_1_FILES layout type, and may be applicable to other layout types.
- Unit.
- A unit is a fixed size quantity of data written to a data server.
- Pattern.
- A pattern is a method of distributing one or more equal sized units across a set of data servers. A pattern is iterated one or more times.
- Stripe.
- An stripe is a set of data distributed across a set of data servers in a pattern before that pattern repeats.
- Stripe Count.
- A stripe count is the number of units in a pattern.
- Stripe Width.
- A stripe width is the size of stripe in bytes. The stripe width = the stripe count * the size of the stripe unit.
Hereafter, this document will refer to a unit that is a written in a pattern as a "stripe unit".
A pattern may have more stripe units than data servers. If so, some data servers will have more than one stripe unit per stripe. A data server that has multiple stripe units per stripe MAY store each unit in a different data file (and depending on the implementation, will possibly assign a unique data filehandle to each data file).
| TOC |
The high level NFSv4.1 layout types are nfsv4_1_file_layouthint4, nfsv4_1_file_layout_ds_addr4, and nfsv4_1_file_layout4.
The SETATTR operation supports a layout hint attribute (Section 5.12.4 (Attribute 63: layout_hint)). When the client sets a layout hint (data type layouthint4) with a layout type of LAYOUT4_NFSV4_1_FILES (the loh_type field), the loh_body field contains a value of data type nfsv4_1_file_layouthint4.
const NFL4_UFLG_MASK = 0x0000003F;
const NFL4_UFLG_DENSE = 0x00000001;
const NFL4_UFLG_COMMIT_THRU_MDS = 0x00000002;
const NFL4_UFLG_STRIPE_UNIT_SIZE_MASK
= 0xFFFFFFC0;
typedef uint32_t nfl_util4;
enum filelayout_hint_care4 {
NFLH4_CARE_DENSE = NFL4_UFLG_DENSE,
NFLH4_CARE_COMMIT_THRU_MDS
= NFL4_UFLG_COMMIT_THRU_MDS,
NFLH4_CARE_STRIPE_UNIT_SIZE
= 0x00000040,
NFLH4_CARE_STRIPE_COUNT = 0x00000080
};
/* Encoded in the loh_body field of type layouthint4: */
struct nfsv4_1_file_layouthint4 {
uint32_t nflh_care;
nfl_util4 nflh_util;
count4 nflh_stripe_count;
};
The generic layout hint structure is described in Section 3.3.19 (layouthint4). The client uses the layout hint in the layout_hint (Section 5.12.4 (Attribute 63: layout_hint)) attribute to indicate the preferred type of layout to be used for a newly created file. The LAYOUT4_NFSV4_1_FILES layout type-specific content for the layout hint is composed of three fields. The first field, nflh_care, is a set of flags indicating which values of the hint the client cares about. If the NFLH4_CARE_DENSE flag is set, then the client indicates in the second field, nflh_util, a preference for how the data file is packed (Section 13.4.4 (Sparse and Dense Stripe Unit Packing)), which is controlled by the value of nflh_util & NFL4_UFLG_DENSE. If the NFLH4_CARE_COMMIT_THRU_MDS flag is set, then the client indicates a preference for whether the client should send COMMIT operations to the metadata server or data server (Section 13.7 (COMMIT Through Metadata Server)), which is controlled by the value of nflh_util & NFL4_UFLG_COMMIT_THRU_MDS. If the NFLH4_CARE_STRIPE_UNIT_SIZE flag is set, the client indicates its preferred stripe unit size, which is indicated in nflh_util & NFL4_UFLG_STRIPE_UNIT_SIZE_MASK (thus the stripe unit size MUST be a multiple of 64 bytes). The minimum stripe unit size is 64 bytes. If the NFLH4_CARE_STRIPE_COUNT flag is set, the client indicates in the third field, nflh_stripe_count, the stripe count. The stripe count multiplied by the stripe unit size is the stripe width.
When LAYOUTGET returns a LAYOUT4_NFSV4_1_FILES layout (indicated in the loc_type field of the lo_content field), the loc_body field of the lo_content field contains a value of data type nfsv4_1_file_layout4. Among other content, nfsv4_1_file_layout4 has a storage device ID (field nfl_deviceid) of data type deviceid4. The GETDEVICEINFO operation maps a device ID to a storage device address (type device_addr4). When GETDEVICEINFO returns a device address with a layout type of LAYOUT4_NFSV4_1_FILES (the da_layout_type field), the da_addr_body field contains a value of data type nfsv4_1_file_layout_ds_addr4.
typedef netaddr4 multipath_list4<>;
/* Encoded in the da_addr_body field of type device_addr4: */
struct nfsv4_1_file_layout_ds_addr4 {
uint32_t nflda_stripe_indices<>;
multipath_list4 nflda_multipath_ds_list<>;
};
The nfsv4_1_file_layout_ds_addr4 data type represents the device address. It is composed of two fields:
/* Encoded in the loc_body field of type layout_content4: */
struct nfsv4_1_file_layout4 {
deviceid4 nfl_deviceid;
nfl_util4 nfl_util;
uint32_t nfl_first_stripe_index;
offset4 nfl_pattern_offset;
nfs_fh4 nfl_fh_list<>;
};
The nfsv4_1_file_layout4 data type represents the layout. It is composed of the following fields:
The details on the interpretation of the layout are in Section 13.4 (Interpreting the File Layout).
| TOC |
| TOC |
To find the stripe unit number that corresponds to the client's logical file offset, the pattern offset must also be used. The i'th stripe unit (SUi) is:
relative_offset = file_offset - nfl_pattern_offset;
SUi = floor(relative_offset / stripe_unit_size);
| TOC |
When sparse packing is used, the algorithm for determining the filehandle and set of data server network addresses to write stripe unit i (SUi) to is:
stripe_count = number of elements in nflda_stripe_indices;
j = (SUi + nfl_first_stripe_index) % stripe_count;
idx = nflda_stripe_indices[j];
fh_count = number of elements in nfl_fh_list;
ds_count = number of elements in nflda_multipath_ds_list;
switch (fh_count) {
case ds_count:
fh = nfl_fh_list[idx];
break;
case 1:
fh = nfl_fh_list[0];
break;
case 0:
fh = filehandle returned by OPEN;
break;
default:
throw a fatal exception;
break;
}
address_list = nflda_multipath_ds_list[idx];
The client would then select a data server from address_list, and send a READ or WRITE operation using the filehandle specified in fh.
Consider the following example:
Suppose we have a device address consisting of seven data servers, arranged in three equivalence (Section 13.5 (Data Server Multipathing)) classes:
{ A, B, C, D }, { E }, { F, G }
Where A through G are network addresses.
Then
nflda_multipath_ds_list<> = { A, B, C, D }, { E }, { F, G }
i.e.
nflda_multipath_ds_list[0] = { A, B, C, D }
nflda_multipath_ds_list[1] = { E }
nflda_multipath_ds_list[2] = { F, G }
Suppose the striping index array is:
nflda_stripe_indices<> = { 2, 0, 1, 0 }
Now suppose the client gets a layout which has a device ID that maps to the above device address. The initial index,
nfl_first_stripe_index = 2,
and
nfl_fh_list = { 0x36, 0x87, 0x67 }.
If the client wants to write to SU0, the set of valid { network address, filehandle } combinations for SUi are determined by:
nfl_first_stripe_index = 2
So
idx = nflda_stripe_indices[(0 + 2) % 4]
= nflda_stripe_indices[2]
= 1
So
nflda_multipath_ds_list[1] = { E }
and
nfl_fh_list[1] = { 0x87 }
The client can thus write SU0 to { 0x87, { E }, }.
The destinations of the first thirteen storage units are:
| SUi | filehandle | data servers |
|---|---|---|
| 0 | 87 | E |
| 1 | 36 | A,B,C,D |
| 2 | 67 | F,G |
| 3 | 36 | A,B,C,D |
| & |