4. XFS Online Fsck Design¶
This document captures the design of the online filesystem check feature for XFS. The purpose of this document is threefold:
To help kernel distributors understand exactly what the XFS online fsck feature is, and issues about which they should be aware.
To help people reading the code to familiarize themselves with the relevant concepts and design points before they start digging into the code.
To help developers maintaining the system by capturing the reasons supporting higher level decision making.
As the online fsck code is merged, the links in this document to topic branches will be replaced with links to code.
This document is licensed under the terms of the GNU Public License, v2. The primary author is Darrick J. Wong.
This design document is split into seven parts. Part 1 defines what fsck tools are and the motivations for writing a new one. Parts 2 and 3 present a high level overview of how online fsck process works and how it is tested to ensure correct functionality. Part 4 discusses the user interface and the intended usage modes of the new program. Parts 5 and 6 show off the high level components and how they fit together, and then present case studies of how each repair function actually works. Part 7 sums up what has been discussed so far and speculates about what else might be built atop online fsck.
4.1. 1. What is a Filesystem Check?¶
A Unix filesystem has four main responsibilities:
Provide a hierarchy of names through which application programs can associate arbitrary blobs of data for any length of time,
Virtualize physical storage media across those names, and
Retrieve the named data blobs at any time.
Examine resource usage.
Metadata directly supporting these functions (e.g. files, directories, space mappings) are sometimes called primary metadata. Secondary metadata (e.g. reverse mapping and directory parent pointers) support operations internal to the filesystem, such as internal consistency checking and reorganization. Summary metadata, as the name implies, condense information contained in primary metadata for performance reasons.
The filesystem check (fsck) tool examines all the metadata in a filesystem to look for errors. In addition to looking for obvious metadata corruptions, fsck also cross-references different types of metadata records with each other to look for inconsistencies. People do not like losing data, so most fsck tools also contains some ability to correct any problems found. As a word of caution -- the primary goal of most Linux fsck tools is to restore the filesystem metadata to a consistent state, not to maximize the data recovered. That precedent will not be challenged here.
Filesystems of the 20th century generally lacked any redundancy in the ondisk format, which means that fsck can only respond to errors by erasing files until errors are no longer detected. More recent filesystem designs contain enough redundancy in their metadata that it is now possible to regenerate data structures when non-catastrophic errors occur; this capability aids both strategies.
Note: |
System administrators avoid data loss by increasing the number of separate storage systems through the creation of backups; and they avoid downtime by increasing the redundancy of each storage system through the creation of RAID arrays. fsck tools address only the first problem. |
4.1.1. TLDR; Show Me the Code!¶
Code is posted to the kernel.org git trees as follows: kernel changes, userspace changes, and QA test changes. Each kernel patchset adding an online repair function will use the same branch name across the kernel, xfsprogs, and fstests git repos.
4.1.2. Existing Tools¶
The online fsck tool described here will be the third tool in the history of XFS (on Linux) to check and repair filesystems. Two programs precede it:
The first program, xfs_check, was created as part of the XFS debugger
(xfs_db) and can only be used with unmounted filesystems.
It walks all metadata in the filesystem looking for inconsistencies in the
metadata, though it lacks any ability to repair what it finds.
Due to its high memory requirements and inability to repair things, this
program is now deprecated and will not be discussed further.
The second program, xfs_repair, was created to be faster and more robust
than the first program.
Like its predecessor, it can only be used with unmounted filesystems.
It uses extent-based in-memory data structures to reduce memory consumption,
and tries to schedule readahead IO appropriately to reduce I/O waiting time
while it scans the metadata of the entire filesystem.
The most important feature of this tool is its ability to respond to
inconsistencies in file metadata and directory tree by erasing things as needed
to eliminate problems.
Space usage metadata are rebuilt from the observed file metadata.
4.1.3. Problem Statement¶
The current XFS tools leave several problems unsolved:
User programs suddenly lose access to the filesystem when unexpected shutdowns occur as a result of silent corruptions in the metadata. These occur unpredictably and often without warning.
Users experience a total loss of service during the recovery period after an unexpected shutdown occurs.
Users experience a total loss of service if the filesystem is taken offline to look for problems proactively.
Data owners cannot check the integrity of their stored data without reading all of it. This may expose them to substantial billing costs when a linear media scan performed by the storage system administrator might suffice.
System administrators cannot schedule a maintenance window to deal with corruptions if they lack the means to assess filesystem health while the filesystem is online.
Fleet monitoring tools cannot automate periodic checks of filesystem health when doing so requires manual intervention and downtime.
Users can be tricked into doing things they do not desire when malicious actors exploit quirks of Unicode to place misleading names in directories.
Given this definition of the problems to be solved and the actors who would benefit, the proposed solution is a third fsck tool that acts on a running filesystem.
This new third program has three components: an in-kernel facility to check
metadata, an in-kernel facility to repair metadata, and a userspace driver
program to drive fsck activity on a live filesystem.
xfs_scrub is the name of the driver program.
The rest of this document presents the goals and use cases of the new fsck
tool, describes its major design points in connection to those goals, and
discusses the similarities and differences with existing tools.
Note: |
Throughout this document, the existing offline fsck tool can also be
referred to by its current name “ |
The naming hierarchy is broken up into objects known as directories and files and the physical space is split into pieces known as allocation groups. Sharding enables better performance on highly parallel systems and helps to contain the damage when corruptions occur. The division of the filesystem into principal objects (allocation groups and inodes) means that there are ample opportunities to perform targeted checks and repairs on a subset of the filesystem.
While this is going on, other parts continue processing IO requests. Even if a piece of filesystem metadata can only be regenerated by scanning the entire system, the scan can still be done in the background while other file operations continue.
In summary, online fsck takes advantage of resource sharding and redundant metadata to enable targeted checking and repair operations while the system is running. This capability will be coupled to automatic system management so that autonomous self-healing of XFS maximizes service availability.
4.2. 2. Theory of Operation¶
Because it is necessary for online fsck to lock and scan live metadata objects,
online fsck consists of three separate code components.
The first is the userspace driver program xfs_scrub, which is responsible
for identifying individual metadata items, scheduling work items for them,
reacting to the outcomes appropriately, and reporting results to the system
administrator.
The second and third are in the kernel, which implements functions to check
and repair each type of online fsck work item.
Note: |
For brevity, this document shortens the phrase “online fsck work item” to “scrub item”. |
Scrub item types are delineated in a manner consistent with the Unix design philosophy, which is to say that each item should handle one aspect of a metadata structure, and handle it well.
4.2.1. Scope¶
In principle, online fsck should be able to check and to repair everything that the offline fsck program can handle. However, online fsck cannot be running 100% of the time, which means that latent errors may creep in after a scrub completes. If these errors cause the next mount to fail, offline fsck is the only solution. This limitation means that maintenance of the offline fsck tool will continue. A second limitation of online fsck is that it must follow the same resource sharing and lock acquisition rules as the regular filesystem. This means that scrub cannot take any shortcuts to save time, because doing so could lead to concurrency problems. In other words, online fsck is not a complete replacement for offline fsck, and a complete run of online fsck may take longer than online fsck. However, both of these limitations are acceptable tradeoffs to satisfy the different motivations of online fsck, which are to minimize system downtime and to increase predictability of operation.
4.2.2. Phases of Work¶
The userspace driver program xfs_scrub splits the work of checking and
repairing an entire filesystem into seven phases.
Each phase concentrates on checking specific types of scrub items and depends
on the success of all previous phases.
The seven phases are as follows:
Collect geometry information about the mounted filesystem and computer, discover the online fsck capabilities of the kernel, and open the underlying storage devices.
Check allocation group metadata, all realtime volume metadata, and all quota files. Each metadata structure is scheduled as a separate scrub item. If corruption is found in the inode header or inode btree and
xfs_scrubis permitted to perform repairs, then those scrub items are repaired to prepare for phase 3. Repairs are implemented by using the information in the scrub item to resubmit the kernel scrub call with the repair flag enabled; this is discussed in the next section. Optimizations and all other repairs are deferred to phase 4.Check all metadata of every file in the filesystem. Each metadata structure is also scheduled as a separate scrub item. If repairs are needed and
xfs_scrubis permitted to perform repairs, and there were no problems detected during phase 2, then those scrub items are repaired immediately. Optimizations, deferred repairs, and unsuccessful repairs are deferred to phase 4.All remaining repairs and scheduled optimizations are performed during this phase, if the caller permits them. Before starting repairs, the summary counters are checked and any necessary repairs are performed so that subsequent repairs will not fail the resource reservation step due to wildly incorrect summary counters. Unsuccessful repairs are requeued as long as forward progress on repairs is made somewhere in the filesystem. Free space in the filesystem is trimmed at the end of phase 4 if the filesystem is clean.
By the start of this phase, all primary and secondary filesystem metadata must be correct. Summary counters such as the free space counts and quota resource counts are checked and corrected. Directory entry names and extended attribute names are checked for suspicious entries such as control characters or confusing Unicode sequences appearing in names.
If the caller asks for a media scan, read all allocated and written data file extents in the filesystem. The ability to use hardware-assisted data file integrity checking is new to online fsck; neither of the previous tools have this capability. If media errors occur, they will be mapped to the owning files and reported.
Re-check the summary counters and presents the caller with a summary of space usage and file counts.
This allocation of responsibilities will be revisited later in this document.
4.2.3. Steps for Each Scrub Item¶
The kernel scrub code uses a three-step strategy for checking and repairing the one aspect of a metadata object represented by a scrub item:
The scrub item of interest is checked for corruptions; opportunities for optimization; and for values that are directly controlled by the system administrator but look suspicious. If the item is not corrupt or does not need optimization, resource are released and the positive scan results are returned to userspace. If the item is corrupt or could be optimized but the caller does not permit this, resources are released and the negative scan results are returned to userspace. Otherwise, the kernel moves on to the second step.
The repair function is called to rebuild the data structure. Repair functions generally choose rebuild a structure from other metadata rather than try to salvage the existing structure. If the repair fails, the scan results from the first step are returned to userspace. Otherwise, the kernel moves on to the third step.
In the third step, the kernel runs the same checks over the new metadata item to assess the efficacy of the repairs. The results of the reassessment are returned to userspace.
4.2.4. Classification of Metadata¶
Each type of metadata object (and therefore each type of scrub item) is classified as follows:
4.2.4.1. Primary Metadata¶
Metadata structures in this category should be most familiar to filesystem users either because they are directly created by the user or they index objects created by the user Most filesystem objects fall into this class:
Free space and reference count information
Inode records and indexes
Storage mapping information for file data
Directories
Extended attributes
Symbolic links
Quota limits
Scrub obeys the same rules as regular filesystem accesses for resource and lock acquisition.
Primary metadata objects are the simplest for scrub to process. The principal filesystem object (either an allocation group or an inode) that owns the item being scrubbed is locked to guard against concurrent updates. The check function examines every record associated with the type for obvious errors and cross-references healthy records against other metadata to look for inconsistencies. Repairs for this class of scrub item are simple, since the repair function starts by holding all the resources acquired in the previous step. The repair function scans available metadata as needed to record all the observations needed to complete the structure. Next, it stages the observations in a new ondisk structure and commits it atomically to complete the repair. Finally, the storage from the old data structure are carefully reaped.
Because xfs_scrub locks a primary object for the duration of the repair,
this is effectively an offline repair operation performed on a subset of the
filesystem.
This minimizes the complexity of the repair code because it is not necessary to
handle concurrent updates from other threads, nor is it necessary to access
any other part of the filesystem.
As a result, indexed structures can be rebuilt very quickly, and programs
trying to access the damaged structure will be blocked until repairs complete.
The only infrastructure needed by the repair code are the staging area for
observations and a means to write new structures to disk.
Despite these limitations, the advantage that online repair holds is clear:
targeted work on individual shards of the filesystem avoids total loss of
service.
This mechanism is described in section 2.1 (“Off-Line Algorithm”) of V. Srinivasan and M. J. Carey, “Performance of On-Line Index Construction Algorithms”, Extending Database Technology, pp. 293-309, 1992.
Most primary metadata repair functions stage their intermediate results in an in-memory array prior to formatting the new ondisk structure, which is very similar to the list-based algorithm discussed in section 2.3 (“List-Based Algorithms”) of Srinivasan. However, any data structure builder that maintains a resource lock for the duration of the repair is always an offline algorithm.
4.2.4.2. Secondary Metadata¶
Metadata structures in this category reflect records found in primary metadata, but are only needed for online fsck or for reorganization of the filesystem.
Secondary metadata include:
Reverse mapping information
Directory parent pointers
This class of metadata is difficult for scrub to process because scrub attaches
to the secondary object but needs to check primary metadata, which runs counter
to the usual order of resource acquisition.
Frequently, this means that full filesystems scans are necessary to rebuild the
metadata.
Check functions can be limited in scope to reduce runtime.
Repairs, however, require a full scan of primary metadata, which can take a
long time to complete.
Under these conditions, xfs_scrub cannot lock resources for the entire
duration of the repair.
Instead, repair functions set up an in-memory staging structure to store observations. Depending on the requirements of the specific repair function, the staging index will either have the same format as the ondisk structure or a design specific to that repair function. The next step is to release all locks and start the filesystem scan. When the repair scanner needs to record an observation, the staging data are locked long enough to apply the update. While the filesystem scan is in progress, the repair function hooks the filesystem so that it can apply pending filesystem updates to the staging information. Once the scan is done, the owning object is re-locked, the live data is used to write a new ondisk structure, and the repairs are committed atomically. The hooks are disabled and the staging staging area is freed. Finally, the storage from the old data structure are carefully reaped.
Introducing concurrency helps online repair avoid various locking problems, but comes at a high cost to code complexity. Live filesystem code has to be hooked so that the repair function can observe updates in progress. The staging area has to become a fully functional parallel structure so that updates can be merged from the hooks. Finally, the hook, the filesystem scan, and the inode locking model must be sufficiently well integrated that a hook event can decide if a given update should be applied to the staging structure.
In theory, the scrub implementation could apply these same techniques for primary metadata, but doing so would make it massively more complex and less performant. Programs attempting to access the damaged structures are not blocked from operation, which may cause application failure or an unplanned filesystem shutdown.
Inspiration for the secondary metadata repair strategy was drawn from section 2.4 of Srinivasan above, and sections 2 (“NSF: Inded Build Without Side-File”) and 3.1.1 (“Duplicate Key Insert Problem”) in C. Mohan, “Algorithms for Creating Indexes for Very Large Tables Without Quiescing Updates”, 1992.
The sidecar index mentioned above bears some resemblance to the side file method mentioned in Srinivasan and Mohan. Their method consists of an index builder that extracts relevant record data to build the new structure as quickly as possible; and an auxiliary structure that captures all updates that would be committed to the index by other threads were the new index already online. After the index building scan finishes, the updates recorded in the side file are applied to the new index. To avoid conflicts between the index builder and other writer threads, the builder maintains a publicly visible cursor that tracks the progress of the scan through the record space. To avoid duplication of work between the side file and the index builder, side file updates are elided when the record ID for the update is greater than the cursor position within the record ID space.
To minimize changes to the rest of the codebase, XFS online repair keeps the replacement index hidden until it’s completely ready to go. In other words, there is no attempt to expose the keyspace of the new index while repair is running. The complexity of such an approach would be very high and perhaps more appropriate to building new indices.
Future Work Question: Can the full scan and live update code used to facilitate a repair also be used to implement a comprehensive check?
Answer: In theory, yes. Check would be much stronger if each scrub function employed these live scans to build a shadow copy of the metadata and then compared the shadow records to the ondisk records. However, doing that is a fair amount more work than what the checking functions do now. The live scans and hooks were developed much later. That in turn increases the runtime of those scrub functions.
4.2.4.3. Summary Information¶
Metadata structures in this last category summarize the contents of primary metadata records. These are often used to speed up resource usage queries, and are many times smaller than the primary metadata which they represent.
Examples of summary information include:
Summary counts of free space and inodes
File link counts from directories
Quota resource usage counts
Check and repair require full filesystem scans, but resource and lock acquisition follow the same paths as regular filesystem accesses.
The superblock summary counters have special requirements due to the underlying implementation of the incore counters, and will be treated separately. Check and repair of the other types of summary counters (quota resource counts and file link counts) employ the same filesystem scanning and hooking techniques as outlined above, but because the underlying data are sets of integer counters, the staging data need not be a fully functional mirror of the ondisk structure.
Inspiration for quota and file link count repair strategies were drawn from sections 2.12 (“Online Index Operations”) through 2.14 (“Incremental View Maintenance”) of G. Graefe, “Concurrent Queries and Updates in Summary Views and Their Indexes”, 2011.
Since quotas are non-negative integer counts of resource usage, online quotacheck can use the incremental view deltas described in section 2.14 to track pending changes to the block and inode usage counts in each transaction, and commit those changes to a dquot side file when the transaction commits. Delta tracking is necessary for dquots because the index builder scans inodes, whereas the data structure being rebuilt is an index of dquots. Link count checking combines the view deltas and commit step into one because it sets attributes of the objects being scanned instead of writing them to a separate data structure. Each online fsck function will be discussed as case studies later in this document.
4.2.5. Risk Management¶
During the development of online fsck, several risk factors were identified that may make the feature unsuitable for certain distributors and users. Steps can be taken to mitigate or eliminate those risks, though at a cost to functionality.
Decreased performance: Adding metadata indices to the filesystem increases the time cost of persisting changes to disk, and the reverse space mapping and directory parent pointers are no exception. System administrators who require the maximum performance can disable the reverse mapping features at format time, though this choice dramatically reduces the ability of online fsck to find inconsistencies and repair them.
Incorrect repairs: As with all software, there might be defects in the software that result in incorrect repairs being written to the filesystem. Systematic fuzz testing (detailed in the next section) is employed by the authors to find bugs early, but it might not catch everything. The kernel build system provides Kconfig options (
CONFIG_XFS_ONLINE_SCRUBandCONFIG_XFS_ONLINE_REPAIR) to enable distributors to choose not to accept this risk. The xfsprogs build system has a configure option (--enable-scrub=no) that disables building of thexfs_scrubbinary, though this is not a risk mitigation if the kernel functionality remains enabled.Inability to repair: Sometimes, a filesystem is too badly damaged to be repairable. If the keyspaces of several metadata indices overlap in some manner but a coherent narrative cannot be formed from records collected, then the repair fails. To reduce the chance that a repair will fail with a dirty transaction and render the filesystem unusable, the online repair functions have been designed to stage and validate all new records before committing the new structure.
Misbehavior: Online fsck requires many privileges -- raw IO to block devices, opening files by handle, ignoring Unix discretionary access control, and the ability to perform administrative changes. Running this automatically in the background scares people, so the systemd background service is configured to run with only the privileges required. Obviously, this cannot address certain problems like the kernel crashing or deadlocking, but it should be sufficient to prevent the scrub process from escaping and reconfiguring the system. The cron job does not have this protection.
Fuzz Kiddiez: There are many people now who seem to think that running automated fuzz testing of ondisk artifacts to find mischievous behavior and spraying exploit code onto the public mailing list for instant zero-day disclosure is somehow of some social benefit. In the view of this author, the benefit is realized only when the fuzz operators help to fix the flaws, but this opinion apparently is not widely shared among security “researchers”. The XFS maintainers’ continuing ability to manage these events presents an ongoing risk to the stability of the development process. Automated testing should front-load some of the risk while the feature is considered EXPERIMENTAL.
Many of these risks are inherent to software programming. Despite this, it is hoped that this new functionality will prove useful in reducing unexpected downtime.
4.3. 3. Testing Plan¶
As stated before, fsck tools have three main goals:
Detect inconsistencies in the metadata;
Eliminate those inconsistencies; and
Minimize further loss of data.
Demonstrations of correct operation are necessary to build users’ confidence that the software behaves within expectations. Unfortunately, it was not really feasible to perform regular exhaustive testing of every aspect of a fsck tool until the introduction of low-cost virtual machines with high-IOPS storage. With ample hardware availability in mind, the testing strategy for the online fsck project involves differential analysis against the existing fsck tools and systematic testing of every attribute of every type of metadata object. Testing can be split into four major categories, as discussed below.
4.3.1. Integrated Testing with fstests¶
The primary goal of any free software QA effort is to make testing as inexpensive and widespread as possible to maximize the scaling advantages of community. In other words, testing should maximize the breadth of filesystem configuration scenarios and hardware setups. This improves code quality by enabling the authors of online fsck to find and fix bugs early, and helps developers of new features to find integration issues earlier in their development effort.
The Linux filesystem community shares a common QA testing suite,
fstests, for
functional and regression testing.
Even before development work began on online fsck, fstests (when run on XFS)
would run both the xfs_check and xfs_repair -n commands on the test and
scratch filesystems between each test.
This provides a level of assurance that the kernel and the fsck tools stay in
alignment about what constitutes consistent metadata.
During development of the online checking code, fstests was modified to run
xfs_scrub -n between each test to ensure that the new checking code
produces the same results as the two existing fsck tools.
To start development of online repair, fstests was modified to run
xfs_repair to rebuild the filesystem’s metadata indices between tests.
This ensures that offline repair does not crash, leave a corrupt filesystem
after it exists, or trigger complaints from the online check.
This also established a baseline for what can and cannot be repaired offline.
To complete the first phase of development of online repair, fstests was
modified to be able to run xfs_scrub in a “force rebuild” mode.
This enables a comparison of the effectiveness of online repair as compared to
the existing offline repair tools.
4.3.2. General Fuzz Testing of Metadata Blocks¶
XFS benefits greatly from having a very robust debugging tool, xfs_db.
Before development of online fsck even began, a set of fstests were created
to test the rather common fault that entire metadata blocks get corrupted.
This required the creation of fstests library code that can create a filesystem
containing every possible type of metadata object.
Next, individual test cases were created to create a test filesystem, identify
a single block of a specific type of metadata object, trash it with the
existing blocktrash command in xfs_db, and test the reaction of a
particular metadata validation strategy.
This earlier test suite enabled XFS developers to test the ability of the in-kernel validation functions and the ability of the offline fsck tool to detect and eliminate the inconsistent metadata. This part of the test suite was extended to cover online fsck in exactly the same manner.
In other words, for a given fstests filesystem configuration:
For each metadata object existing on the filesystem:
Write garbage to it
Test the reactions of:
The kernel verifiers to stop obviously bad metadata
Offline repair (
xfs_repair) to detect and fixOnline repair (
xfs_scrub) to detect and fix
4.3.3. Targeted Fuzz Testing of Metadata Records¶
The testing plan for online fsck includes extending the existing fs testing
infrastructure to provide a much more powerful facility: targeted fuzz testing
of every metadata field of every metadata object in the filesystem.
xfs_db can modify every field of every metadata structure in every
block in the filesystem to simulate the effects of memory corruption and
software bugs.
Given that fstests already contains the ability to create a filesystem
containing every metadata format known to the filesystem, xfs_db can be
used to perform exhaustive fuzz testing!
For a given fstests filesystem configuration:
For each metadata object existing on the filesystem...
For each record inside that metadata object...
For each field inside that record...
For each conceivable type of transformation that can be applied to a bit field...
Clear all bits
Set all bits
Toggle the most significant bit
Toggle the middle bit
Toggle the least significant bit
Add a small quantity
Subtract a small quantity
Randomize the contents
...test the reactions of:
The kernel verifiers to stop obviously bad metadata
Offline checking (
xfs_repair -n)Offline repair (
xfs_repair)Online checking (
xfs_scrub -n)Online repair (
xfs_scrub)Both repair tools (
xfs_scruband thenxfs_repairif online repair doesn’t succeed)
This is quite the combinatoric explosion!
Fortunately, having this much test coverage makes it easy for XFS developers to
check the responses of XFS’ fsck tools.
Since the introduction of the fuzz testing framework, these tests have been
used to discover incorrect repair code and missing functionality for entire
classes of metadata objects in xfs_repair.
The enhanced testing was used to finalize the deprecation of xfs_check by
confirming that xfs_repair could detect at least as many corruptions as
the older tool.
These tests have been very valuable for xfs_scrub in the same ways -- they
allow the online fsck developers to compare online fsck against offline fsck,
and they enable XFS developers to find deficiencies in the code base.
Proposed patchsets include general fuzzer improvements, fuzzing baselines, and improvements in fuzz testing comprehensiveness.
4.3.4. Stress Testing¶
A unique requirement to online fsck is the ability to operate on a filesystem
concurrently with regular workloads.
Although it is of course impossible to run xfs_scrub with zero observable
impact on the running system, the online repair code should never introduce
inconsistencies into the filesystem metadata, and regular workloads should
never notice resource starvation.
To verify that these conditions are being met, fstests has been enhanced in
the following ways:
For each scrub item type, create a test to exercise checking that item type while running
fsstress.For each scrub item type, create a test to exercise repairing that item type while running
fsstress.Race
fsstressandxfs_scrub -nto ensure that checking the whole filesystem doesn’t cause problems.Race
fsstressandxfs_scrubin force-rebuild mode to ensure that force-repairing the whole filesystem doesn’t cause problems.Race
xfs_scrubin check and force-repair mode againstfsstresswhile freezing and thawing the filesystem.Race
xfs_scrubin check and force-repair mode againstfsstresswhile remounting the filesystem read-only and read-write.The same, but running
fsxinstead offsstress. (Not done yet?)
Success is defined by the ability to run all of these tests without observing any unexpected filesystem shutdowns due to corrupted metadata, kernel hang check warnings, or any other sort of mischief.
Proposed patchsets include general stress testing and the evolution of existing per-function stress testing.
4.4. 4. User Interface¶
The primary user of online fsck is the system administrator, just like offline repair. Online fsck presents two modes of operation to administrators: A foreground CLI process for online fsck on demand, and a background service that performs autonomous checking and repair.
4.4.1. Checking on Demand¶
For administrators who want the absolute freshest information about the
metadata in a filesystem, xfs_scrub can be run as a foreground process on
a command line.
The program checks every piece of metadata in the filesystem while the
administrator waits for the results to be reported, just like the existing
xfs_repair tool.
Both tools share a -n option to perform a read-only scan, and a -v
option to increase the verbosity of the information reported.
A new feature of xfs_scrub is the -x option, which employs the error
correction capabilities of the hardware to check data file contents.
The media scan is not enabled by default because it may dramatically increase
program runtime and consume a lot of bandwidth on older storage hardware.
The output of a foreground invocation is captured in the system log.
The xfs_scrub_all program walks the list of mounted filesystems and
initiates xfs_scrub for each of them in parallel.
It serializes scans for any filesystems that resolve to the same top level
kernel block device to prevent resource overconsumption.
4.4.2. Background Service¶
To reduce the workload of system administrators, the xfs_scrub package
provides a suite of systemd timers and services that
run online fsck automatically on weekends by default.
The background service configures scrub to run with as little privilege as
possible, the lowest CPU and IO priority, and in a CPU-constrained single
threaded mode.
This can be tuned by the systemd administrator at any time to suit the latency
and throughput requirements of customer workloads.
The output of the background service is also captured in the system log.
If desired, reports of failures (either due to inconsistencies or mere runtime
errors) can be emailed automatically by setting the EMAIL_ADDR environment
variable in the following service files:
xfs_scrub_fail@.servicexfs_scrub_media_fail@.servicexfs_scrub_all_fail.service
The decision to enable the background scan is left to the system administrator. This can be done by enabling either of the following services:
xfs_scrub_all.timeron systemd systemsxfs_scrub_all.cronon non-systemd systems
This automatic weekly scan is configured out of the box to perform an additional media scan of all file data once per month. This is less foolproof than, say, storing file data block checksums, but much more performant if application software provides its own integrity checking, redundancy can be provided elsewhere above the filesystem, or the storage device’s integrity guarantees are deemed sufficient.
The systemd unit file definitions have been subjected to a security audit
(as of systemd 249) to ensure that the xfs_scrub processes have as little
access to the rest of the system as possible.
This was performed via systemd-analyze security, after which privileges
were restricted to the minimum required, sandboxing was set up to the maximal
extent possible with sandboxing and system call filtering; and access to the
filesystem tree was restricted to the minimum needed to start the program and
access the filesystem being scanned.
The service definition files restrict CPU usage to 80% of one CPU core, and
apply as nice of a priority to IO and CPU scheduling as possible.
This measure was taken to minimize delays in the rest of the filesystem.
No such hardening has been performed for the cron job.
Proposed patchset: Enabling the xfs_scrub background service.
4.4.3. Health Reporting¶
XFS caches a summary of each filesystem’s health status in memory.
The information is updated whenever xfs_scrub is run, or whenever
inconsistencies are detected in the filesystem metadata during regular
operations.
System administrators should use the health command of xfs_spaceman to
download this information into a human-readable format.
If problems have been observed, the administrator can schedule a reduced
service window to run the online repair tool to correct the problem.
Failing that, the administrator can decide to schedule a maintenance window to
run the traditional offline repair tool to correct the problem.
Future Work Question: Should the health reporting integrate with the new inotify fs error notification system? Would it be helpful for sysadmins to have a daemon to listen for corruption notifications and initiate a repair?
Answer: These questions remain unanswered, but should be a part of the conversation with early adopters and potential downstream users of XFS.
Proposed patchsets include wiring up health reports to correction returns and preservation of sickness info during memory reclaim.
4.5. 5. Kernel Algorithms and Data Structures¶
This section discusses the key algorithms and data structures of the kernel code that provide the ability to check and repair metadata while the system is running. The first chapters in this section reveal the pieces that provide the foundation for checking metadata. The remainder of this section presents the mechanisms through which XFS regenerates itself.
4.5.1. Self Describing Metadata¶
Starting with XFS version 5 in 2012, XFS updated the format of nearly every ondisk block header to record a magic number, a checksum, a universally “unique” identifier (UUID), an owner code, the ondisk address of the block, and a log sequence number. When loading a block buffer from disk, the magic number, UUID, owner, and ondisk address confirm that the retrieved block matches the specific owner of the current filesystem, and that the information contained in the block is supposed to be found at the ondisk address. The first three components enable checking tools to disregard alleged metadata that doesn’t belong to the filesystem, and the fourth component enables the filesystem to detect lost writes.
Whenever a file system operation modifies a block, the change is submitted to the log as part of a transaction. The log then processes these transactions marking them done once they are safely persisted to storage. The logging code maintains the checksum and the log sequence number of the last transactional update. Checksums are useful for detecting torn writes and other discrepancies that can be introduced between the computer and its storage devices. Sequence number tracking enables log recovery to avoid applying out of date log updates to the filesystem.
These two features improve overall runtime resiliency by providing a means for the filesystem to detect obvious corruption when reading metadata blocks from disk, but these buffer verifiers cannot provide any consistency checking between metadata structures.
For more information, please see the documentation for XFS Self Describing Metadata
4.5.2. Reverse Mapping¶
The original design of XFS (circa 1993) is an improvement upon 1980s Unix filesystem design. In those days, storage density was expensive, CPU time was scarce, and excessive seek time could kill performance. For performance reasons, filesystem authors were reluctant to add redundancy to the filesystem, even at the cost of data integrity. Filesystems designers in the early 21st century choose different strategies to increase internal redundancy -- either storing nearly identical copies of metadata, or more space-efficient encoding techniques.
For XFS, a different redundancy strategy was chosen to modernize the design: a secondary space usage index that maps allocated disk extents back to their owners. By adding a new index, the filesystem retains most of its ability to scale well to heavily threaded workloads involving large datasets, since the primary file metadata (the directory tree, the file block map, and the allocation groups) remain unchanged. Like any system that improves redundancy, the reverse-mapping feature increases overhead costs for space mapping activities. However, it has two critical advantages: first, the reverse index is key to enabling online fsck and other requested functionality such as free space defragmentation, better media failure reporting, and filesystem shrinking. Second, the different ondisk storage format of the reverse mapping btree defeats device-level deduplication because the filesystem requires real redundancy.
Sidebar: |
A criticism of adding the secondary index is that it does nothing to improve the robustness of user data storage itself. This is a valid point, but adding a new index for file data block checksums increases write amplification by turning data overwrites into copy-writes, which age the filesystem prematurely. In keeping with thirty years of precedent, users who want file data integrity can supply as powerful a solution as they require. As for metadata, the complexity of adding a new secondary index of space usage is much less than adding volume management and storage device mirroring to XFS itself. Perfection of RAID and volume management are best left to existing layers in the kernel. |
The information captured in a reverse space mapping record is as follows:
struct xfs_rmap_irec {
xfs_agblock_t rm_startblock; /* extent start block */
xfs_extlen_t rm_blockcount; /* extent length */
uint64_t rm_owner; /* extent owner */
uint64_t rm_offset; /* offset within the owner */
unsigned int rm_flags; /* state flags */
};
The first two fields capture the location and size of the physical space, in units of filesystem blocks. The owner field tells scrub which metadata structure or file inode have been assigned this space. For space allocated to files, the offset field tells scrub where the space was mapped within the file fork. Finally, the flags field provides extra information about the space usage -- is this an attribute fork extent? A file mapping btree extent? Or an unwritten data extent?
Online filesystem checking judges the consistency of each primary metadata record by comparing its information against all other space indices. The reverse mapping index plays a key role in the consistency checking process because it contains a centralized alternate copy of all space allocation information. Program runtime and ease of resource acquisition are the only real limits to what online checking can consult. For example, a file data extent mapping can be checked against:
The absence of an entry in the free space information.
The absence of an entry in the inode index.
The absence of an entry in the reference count data if the file is not marked as having shared extents.
The correspondence of an entry in the reverse mapping information.
There are several observations to make about reverse mapping indices:
Reverse mappings can provide a positive affirmation of correctness if any of the above primary metadata are in doubt. The checking code for most primary metadata follows a path similar to the one outlined above.
Proving the consistency of secondary metadata with the primary metadata is difficult because that requires a full scan of all primary space metadata, which is very time intensive. For example, checking a reverse mapping record for a file extent mapping btree block requires locking the file and searching the entire btree to confirm the block. Instead, scrub relies on rigorous cross-referencing during the primary space mapping structure checks.
Consistency scans must use non-blocking lock acquisition primitives if the required locking order is not the same order used by regular filesystem operations. For example, if the filesystem normally takes a file ILOCK before taking the AGF buffer lock but scrub wants to take a file ILOCK while holding an AGF buffer lock, scrub cannot block on that second acquisition. This means that forward progress during this part of a scan of the reverse mapping data cannot be guaranteed if system load is heavy.
In summary, reverse mappings play a key role in reconstruction of primary metadata. The details of how these records are staged, written to disk, and committed into the filesystem are covered in subsequent sections.
4.5.3. Checking and Cross-Referencing¶
The first step of checking a metadata structure is to examine every record contained within the structure and its relationship with the rest of the system. XFS contains multiple layers of checking to try to prevent inconsistent metadata from wreaking havoc on the system. Each of these layers contributes information that helps the kernel to make three decisions about the health of a metadata structure:
Is a part of this structure obviously corrupt (
XFS_SCRUB_OFLAG_CORRUPT) ?Is this structure inconsistent with the rest of the system (
XFS_SCRUB_OFLAG_XCORRUPT) ?Is there so much damage around the filesystem that cross-referencing is not possible (
XFS_SCRUB_OFLAG_XFAIL) ?Can the structure be optimized to improve performance or reduce the size of metadata (
XFS_SCRUB_OFLAG_PREEN) ?Does the structure contain data that is not inconsistent but deserves review by the system administrator (
XFS_SCRUB_OFLAG_WARNING) ?
The following sections describe how the metadata scrubbing process works.
4.5.3.1. Metadata Buffer Verification¶
The lowest layer of metadata protection in XFS are the metadata verifiers built into the buffer cache. These functions perform inexpensive internal consistency checking of the block itself, and answer these questions:
Does the block belong to this filesystem?
Does the block belong to the structure that asked for the read? This assumes that metadata blocks only have one owner, which is always true in XFS.
Is the type of data stored in the block within a reasonable range of what scrub is expecting?
Does the physical location of the block match the location it was read from?
Does the block checksum match the data?
The scope of the protections here are very limited -- verifiers can only establish that the filesystem code is reasonably free of gross corruption bugs and that the storage system is reasonably competent at retrieval. Corruption problems observed at runtime cause the generation of health reports, failed system calls, and in the extreme case, filesystem shutdowns if the corrupt metadata force the cancellation of a dirty transaction.
Every online fsck scrubbing function is expected to read every ondisk metadata block of a structure in the course of checking the structure. Corruption problems observed during a check are immediately reported to userspace as corruption; during a cross-reference, they are reported as a failure to cross-reference once the full examination is complete. Reads satisfied by a buffer already in cache (and hence already verified) bypass these checks.
4.5.3.2. Internal Consistency Checks¶
After the buffer cache, the next level of metadata protection is the internal record verification code built into the filesystem. These checks are split between the buffer verifiers, the in-filesystem users of the buffer cache, and the scrub code itself, depending on the amount of higher level context required. The scope of checking is still internal to the block. These higher level checking functions answer these questions:
Does the type of data stored in the block match what scrub is expecting?
Does the block belong to the owning structure that asked for the read?
If the block contains records, do the records fit within the block?
If the block tracks internal free space information, is it consistent with the record areas?
Are the records contained inside the block free of obvious corruptions?
Record checks in this category are more rigorous and more time-intensive. For example, block pointers and inumbers are checked to ensure that they point within the dynamically allocated parts of an allocation group and within the filesystem. Names are checked for invalid characters, and flags are checked for invalid combinations. Other record attributes are checked for sensible values. Btree records spanning an interval of the btree keyspace are checked for correct order and lack of mergeability (except for file fork mappings). For performance reasons, regular code may skip some of these checks unless debugging is enabled or a write is about to occur. Scrub functions, of course, must check all possible problems.
4.5.3.3. Validation of Userspace-Controlled Record Attributes¶
Various pieces of filesystem metadata are directly controlled by userspace. Because of this nature, validation work cannot be more precise than checking that a value is within the possible range. These fields include:
Superblock fields controlled by mount options
Filesystem labels
File timestamps
File permissions
File size
File flags
Names present in directory entries, extended attribute keys, and filesystem labels
Extended attribute key namespaces
Extended attribute values
File data block contents
Quota limits
Quota timer expiration (if resource usage exceeds the soft limit)
4.5.3.4. Cross-Referencing Space Metadata¶
After internal block checks, the next higher level of checking is cross-referencing records between metadata structures. For regular runtime code, the cost of these checks is considered to be prohibitively expensive, but as scrub is dedicated to rooting out inconsistencies, it must pursue all avenues of inquiry. The exact set of cross-referencing is highly dependent on the context of the data structure being checked.
The XFS btree code has keyspace scanning functions that online fsck uses to cross reference one structure with another. Specifically, scrub can scan the key space of an index to determine if that keyspace is fully, sparsely, or not at all mapped to records. For the reverse mapping btree, it is possible to mask parts of the key for the purposes of performing a keyspace scan so that scrub can decide if the rmap btree contains records mapping a certain extent of physical space without the sparsenses of the rest of the rmap keyspace getting in the way.
Btree blocks undergo the following checks before cross-referencing:
Does the type of data stored in the block match what scrub is expecting?
Does the block belong to the owning structure that asked for the read?
Do the records fit within the block?
Are the records contained inside the block free of obvious corruptions?
Are the name hashes in the correct order?
Do node pointers within the btree point to valid block addresses for the type of btree?
Do child pointers point towards the leaves?
Do sibling pointers point across the same level?
For each node block record, does the record key accurate reflect the contents of the child block?
Space allocation records are cross-referenced as follows:
Any space mentioned by any metadata structure are cross-referenced as follows:
Does the reverse mapping index list only the appropriate owner as the owner of each block?
Are none of the blocks claimed as free space?
If these aren’t file data blocks, are none of the blocks claimed as space shared by different owners?
Btree blocks are cross-referenced as follows:
Everything in class 1 above.
If there’s a parent node block, do the keys listed for this block match the keyspace of this block?
Do the sibling pointers point to valid blocks? Of the same level?
Do the child pointers point to valid blocks? Of the next level down?
Free space btree records are cross-referenced as follows:
Everything in class 1 and 2 above.
Does the reverse mapping index list no owners of this space?
Is this space not claimed by the inode index for inodes?
Is it not mentioned by the reference count index?
Is there a matching record in the other free space btree?
Inode btree records are cross-referenced as follows:
Everything in class 1 and 2 above.
Is there a matching record in free inode btree?
Do cleared bits in the holemask correspond with inode clusters?
Do set bits in the freemask correspond with inode records with zero link count?
Inode records are cross-referenced as follows:
Everything in class 1.
Do all the fields that summarize information about the file forks actually match those forks?
Does each inode with zero link count correspond to a record in the free inode btree?
File fork space mapping records are cross-referenced as follows:
Everything in class 1 and 2 above.
Is this space not mentioned by the inode btrees?
If this is a CoW fork mapping, does it correspond to a CoW entry in the reference count btree?
Reference count records are cross-referenced as follows:
Everything in class 1 and 2 above.
Within the space subkeyspace of the rmap btree (that is to say, all records mapped to a particular space extent and ignoring the owner info), are there the same number of reverse mapping records for each block as the reference count record claims?
Proposed patchsets are the series to find gaps in refcount btree, inode btree, and rmap btree records; to find mergeable records; and to improve cross referencing with rmap before starting a repair.
4.5.3.5. Checking Extended Attributes¶
Extended attributes implement a key-value store that enable fragments of data to be attached to any file. Both the kernel and userspace can access the keys and values, subject to namespace and privilege restrictions. Most typically these fragments are metadata about the file -- origins, security contexts, user-supplied labels, indexing information, etc.
Names can be as long as 255 bytes and can exist in several different namespaces. Values can be as large as 64KB. A file’s extended attributes are stored in blocks mapped by the attr fork. The mappings point to leaf blocks, remote value blocks, or dabtree blocks. Block 0 in the attribute fork is always the top of the structure, but otherwise each of the three types of blocks can be found at any offset in the attr fork. Leaf blocks contain attribute key records that point to the name and the value. Names are always stored elsewhere in the same leaf block. Values that are less than 3/4 the size of a filesystem block are also stored elsewhere in the same leaf block. Remote value blocks contain values that are too large to fit inside a leaf. If the leaf information exceeds a single filesystem block, a dabtree (also rooted at block 0) is created to map hashes of the attribute names to leaf blocks in the attr fork.
Checking an extended attribute structure is not so straightforward due to the lack of separation between attr blocks and index blocks. Scrub must read each block mapped by the attr fork and ignore the non-leaf blocks:
Walk the dabtree in the attr fork (if present) to ensure that there are no irregularities in the blocks or dabtree mappings that do not point to attr leaf blocks.
Walk the blocks of the attr fork looking for leaf blocks. For each entry inside a leaf:
Validate that the name does not contain invalid characters.
Read the attr value. This performs a named lookup of the attr name to ensure the correctness of the dabtree. If the value is stored in a remote block, this also validates the integrity of the remote value block.
4.5.3.6. Checking and Cross-Referencing Directories¶
The filesystem directory tree is a directed acylic graph structure, with files constituting the nodes, and directory entries (dirents) constituting the edges. Directories are a special type of file containing a set of mappings from a 255-byte sequence (name) to an inumber. These are called directory entries, or dirents for short. Each directory file must have exactly one directory pointing to the file. A root directory points to itself. Directory entries point to files of any type. Each non-directory file may have multiple directories point to it.
In XFS, directories are implemented as a file containing up to three 32GB partitions. The first partition contains directory entry data blocks. Each data block contains variable-sized records associating a user-provided name with an inumber and, optionally, a file type. If the directory entry data grows beyond one block, the second partition (which exists as post-EOF extents) is populated with a block containing free space information and an index that maps hashes of the dirent names to directory data blocks in the first partition. This makes directory name lookups very fast. If this second partition grows beyond one block, the third partition is populated with a linear array of free space information for faster expansions. If the free space has been separated and the second partition grows again beyond one block, then a dabtree is used to map hashes of dirent names to directory data blocks.
Checking a directory is pretty straightforward:
Walk the dabtree in the second partition (if present) to ensure that there are no irregularities in the blocks or dabtree mappings that do not point to dirent blocks.
Walk the blocks of the first partition looking for directory entries. Each dirent is checked as follows:
Does the name contain no invalid characters?
Does the inumber correspond to an actual, allocated inode?
Does the child inode have a nonzero link count?
If a file type is included in the dirent, does it match the type of the inode?
If the child is a subdirectory, does the child’s dotdot pointer point back to the parent?
If the directory has a second partition, perform a named lookup of the dirent name to ensure the correctness of the dabtree.
Walk the free space list in the third partition (if present) to ensure that the free spaces it describes are really unused.
Checking operations involving parents and file link counts are discussed in more detail in later sections.
4.5.3.6.1. Checking Directory/Attribute Btrees¶
As stated in previous sections, the directory/attribute btree (dabtree) index maps user-provided names to improve lookup times by avoiding linear scans. Internally, it maps a 32-bit hash of the name to a block offset within the appropriate file fork.
The internal structure of a dabtree closely resembles the btrees that record fixed-size metadata records -- each dabtree block contains a magic number, a checksum, sibling pointers, a UUID, a tree level, and a log sequence number. The format of leaf and node records are the same -- each entry points to the next level down in the hierarchy, with dabtree node records pointing to dabtree leaf blocks, and dabtree leaf records pointing to non-dabtree blocks elsewhere in the fork.
Checking and cross-referencing the dabtree is very similar to what is done for space btrees:
Does the type of data stored in the block match what scrub is expecting?
Does the block belong to the owning structure that asked for the read?
Do the records fit within the block?
Are the records contained inside the block free of obvious corruptions?
Are the name hashes in the correct order?
Do node pointers within the dabtree point to valid fork offsets for dabtree blocks?
Do leaf pointers within the dabtree point to valid fork offsets for directory or attr leaf blocks?
Do child pointers point towards the leaves?
Do sibling pointers point across the same level?
For each dabtree node record, does the record key accurate reflect the contents of the child dabtree block?
For each dabtree leaf record, does the record key accurate reflect the contents of the directory or attr block?
4.5.3.7. Cross-Referencing Summary Counters¶
XFS maintains three classes of summary counters: available resources, quota resource usage, and file link counts.
In theory, the amount of available resources (data blocks, inodes, realtime extents) can be found by walking the entire filesystem. This would make for very slow reporting, so a transactional filesystem can maintain summaries of this information in the superblock. Cross-referencing these values against the filesystem metadata should be a simple matter of walking the free space and inode metadata in each AG and the realtime bitmap, but there are complications that will be discussed in more detail later.
Quota usage and file link count checking are sufficiently complicated to warrant separate sections.
4.5.3.8. Post-Repair Reverification¶
After performing a repair, the checking code is run a second time to validate the new structure, and the results of the health assessment are recorded internally and returned to the calling process. This step is critical for enabling system administrator to monitor the status of the filesystem and the progress of any repairs. For developers, it is a useful means to judge the efficacy of error detection and correction in the online and offline checking tools.
4.5.4. Eventual Consistency vs. Online Fsck¶
Complex operations can make modifications to multiple per-AG data structures with a chain of transactions. These chains, once committed to the log, are restarted during log recovery if the system crashes while processing the chain. Because the AG header buffers are unlocked between transactions within a chain, online checking must coordinate with chained operations that are in progress to avoid incorrectly detecting inconsistencies due to pending chains. Furthermore, online repair must not run when operations are pending because the metadata are temporarily inconsistent with each other, and rebuilding is not possible.
Only online fsck has this requirement of total consistency of AG metadata, and should be relatively rare as compared to filesystem change operations. Online fsck coordinates with transaction chains as follows:
For each AG, maintain a count of intent items targeting that AG. The count should be bumped whenever a new item is added to the chain. The count should be dropped when the filesystem has locked the AG header buffers and finished the work.
When online fsck wants to examine an AG, it should lock the AG header buffers to quiesce all transaction chains that want to modify that AG. If the count is zero, proceed with the checking operation. If it is nonzero, cycle the buffer locks to allow the chain to make forward progress.
This may lead to online fsck taking a long time to complete, but regular filesystem updates take precedence over background checking activity. Details about the discovery of this situation are presented in the next section, and details about the solution are presented after that.
4.5.4.1. Discovery of the Problem¶
Midway through the development of online scrubbing, the fsstress tests uncovered a misinteraction between online fsck and compound transaction chains created by other writer threads that resulted in false reports of metadata inconsistency. The root cause of these reports is the eventual consistency model introduced by the expansion of deferred work items and compound transaction chains when reverse mapping and reflink were introduced.
Originally, transaction chains were added to XFS to avoid deadlocks when unmapping space from files. Deadlock avoidance rules require that AGs only be locked in increasing order, which makes it impossible (say) to use a single transaction to free a space extent in AG 7 and then try to free a now superfluous block mapping btree block in AG 3. To avoid these kinds of deadlocks, XFS creates Extent Freeing Intent (EFI) log items to commit to freeing some space in one transaction while deferring the actual metadata updates to a fresh transaction. The transaction sequence looks like this:
The first transaction contains a physical update to the file’s block mapping structures to remove the mapping from the btree blocks. It then attaches to the in-memory transaction an action item to schedule deferred freeing of space. Concretely, each transaction maintains a list of
struct xfs_defer_pendingobjects, each of which maintains a list ofstruct xfs_extent_free_itemobjects. Returning to the example above, the action item tracks the freeing of both the unmapped space from AG 7 and the block mapping btree (BMBT) block from AG 3. Deferred frees recorded in this manner are committed in the log by creating an EFI log item from thestruct xfs_extent_free_itemobject and attaching the log item to the transaction. When the log is persisted to disk, the EFI item is written into the ondisk transaction record. EFIs can list up to 16 extents to free, all sorted in AG order.The second transaction contains a physical update to the free space btrees of AG 3 to release the former BMBT block and a second physical update to the free space btrees of AG 7 to release the unmapped file space. Observe that the physical updates are resequenced in the correct order when possible. Attached to the transaction is a an extent free done (EFD) log item. The EFD contains a pointer to the EFI logged in transaction #1 so that log recovery can tell if the EFI needs to be replayed.
If the system goes down after transaction #1 is written back to the filesystem but before #2 is committed, a scan of the filesystem metadata would show inconsistent filesystem metadata because there would not appear to be any owner of the unmapped space. Happily, log recovery corrects this inconsistency for us -- when recovery finds an intent log item but does not find a corresponding intent done item, it will reconstruct the incore state of the intent item and finish it. In the example above, the log must replay both frees described in the recovered EFI to complete the recovery phase.
There are subtleties to XFS’ transaction chaining strategy to consider:
Log items must be added to a transaction in the correct order to prevent conflicts with principal objects that are not held by the transaction. In other words, all per-AG metadata updates for an unmapped block must be completed before the last update to free the extent, and extents should not be reallocated until that last update commits to the log.
AG header buffers are released between each transaction in a chain. This means that other threads can observe an AG in an intermediate state, but as long as the first subtlety is handled, this should not affect the correctness of filesystem operations.
Unmounting the filesystem flushes all pending work to disk, which means that offline fsck never sees the temporary inconsistencies caused by deferred work item processing.
In this manner, XFS employs a form of eventual consistency to avoid deadlocks and increase parallelism.
During the design phase of the reverse mapping and reflink features, it was decided that it was impractical to cram all the reverse mapping updates for a single filesystem change into a single transaction because a single file mapping operation can explode into many small updates:
The block mapping update itself
A reverse mapping update for the block mapping update
Fixing the freelist
A reverse mapping update for the freelist fix
A shape change to the block mapping btree
A reverse mapping update for the btree update
Fixing the freelist (again)
A reverse mapping update for the freelist fix
An update to the reference counting information
A reverse mapping update for the refcount update
Fixing the freelist (a third time)
A reverse mapping update for the freelist fix
Freeing any space that was unmapped and not owned by any other file
Fixing the freelist (a fourth time)
A reverse mapping update for the freelist fix
Freeing the space used by the block mapping btree
Fixing the freelist (a fifth time)
A reverse mapping update for the freelist fix
Free list fixups are not usually needed more than once per AG per transaction chain, but it is theoretically possible if space is very tight. For copy-on-write updates this is even worse, because this must be done once to remove the space from a staging area and again to map it into the file!
To deal with this explosion in a calm manner, XFS expands its use of deferred work items to cover most reverse mapping updates and all refcount updates. This reduces the worst case size of transaction reservations by breaking the work into a long chain of small updates, which increases the degree of eventual consistency in the system. Again, this generally isn’t a problem because XFS orders its deferred work items carefully to avoid resource reuse conflicts between unsuspecting threads.
However, online fsck changes the rules -- remember that although physical updates to per-AG structures are coordinated by locking the buffers for AG headers, buffer locks are dropped between transactions. Once scrub acquires resources and takes locks for a data structure, it must do all the validation work without releasing the lock. If the main lock for a space btree is an AG header buffer lock, scrub may have interrupted another thread that is midway through finishing a chain. For example, if a thread performing a copy-on-write has completed a reverse mapping update but not the corresponding refcount update, the two AG btrees will appear inconsistent to scrub and an observation of corruption will be recorded. This observation will not be correct. If a repair is attempted in this state, the results will be catastrophic!
Several other solutions to this problem were evaluated upon discovery of this flaw and rejected:
Add a higher level lock to allocation groups and require writer threads to acquire the higher level lock in AG order before making any changes. This would be very difficult to implement in practice because it is difficult to determine which locks need to be obtained, and in what order, without simulating the entire operation. Performing a dry run of a file operation to discover necessary locks would make the filesystem very slow.
Make the deferred work coordinator code aware of consecutive intent items targeting the same AG and have it hold the AG header buffers locked across the transaction roll between updates. This would introduce a lot of complexity into the coordinator since it is only loosely coupled with the actual deferred work items. It would also fail to solve the problem because deferred work items can generate new deferred subtasks, but all subtasks must be complete before work can start on a new sibling task.
Teach online fsck to walk all transactions waiting for whichever lock(s) protect the data structure being scrubbed to look for pending operations. The checking and repair operations must factor these pending operations into the evaluations being performed. This solution is a nonstarter because it is extremely invasive to the main filesystem.
4.5.4.2. Intent Drains¶
Online fsck uses an atomic intent item counter and lock cycling to coordinate with transaction chains. There are two key properties to the drain mechanism. First, the counter is incremented when a deferred work item is queued to a transaction, and it is decremented after the associated intent done log item is committed to another transaction. The second property is that deferred work can be added to a transaction without holding an AG header lock, but per-AG work items cannot be marked done without locking that AG header buffer to log the physical updates and the intent done log item. The first property enables scrub to yield to running transaction chains, which is an explicit deprioritization of online fsck to benefit file operations. The second property of the drain is key to the correct coordination of scrub, since scrub will always be able to decide if a conflict is possible.
For regular filesystem code, the drain works as follows:
Call the appropriate subsystem function to add a deferred work item to a transaction.
The function calls
xfs_defer_drain_bumpto increase the counter.When the deferred item manager wants to finish the deferred work item, it calls
->finish_itemto complete it.The
->finish_itemimplementation logs some changes and callsxfs_defer_drain_dropto decrease the sloppy counter and wake up any threads waiting on the drain.The subtransaction commits, which unlocks the resource associated with the intent item.
For scrub, the drain works as follows:
Lock the resource(s) associated with the metadata being scrubbed. For example, a scan of the refcount btree would lock the AGI and AGF header buffers.
If the counter is zero (
xfs_defer_drain_busyreturns false), there are no chains in progress and the operation may proceed.Otherwise, release the resources grabbed in step 1.
Wait for the intent counter to reach zero (
xfs_defer_drain_intents), then go back to step 1 unless a signal has been caught.
To avoid polling in step 4, the drain provides a waitqueue for scrub threads to be woken up whenever the intent count drops to zero.
The proposed patchset is the scrub intent drain series.
4.5.4.3. Static Keys (aka Jump Label Patching)¶
Online fsck for XFS separates the regular filesystem from the checking and
repair code as much as possible.
However, there are a few parts of online fsck (such as the intent drains, and
later, live update hooks) where it is useful for the online fsck code to know
what’s going on in the rest of the filesystem.
Since it is not expected that online fsck will be constantly running in the
background, it is very important to minimize the runtime overhead imposed by
these hooks when online fsck is compiled into the kernel but not actively
running on behalf of userspace.
Taking locks in the hot path of a writer thread to access a data structure only
to find that no further action is necessary is expensive -- on the author’s
computer, this have an overhead of 40-50ns per access.
Fortunately, the kernel supports dynamic code patching, which enables XFS to
replace a static branch to hook code with nop sleds when online fsck isn’t
running.
This sled has an overhead of however long it takes the instruction decoder to
skip past the sled, which seems to be on the order of less than 1ns and
does not access memory outside of instruction fetching.
When online fsck enables the static key, the sled is replaced with an unconditional branch to call the hook code. The switchover is quite expensive (~22000ns) but is paid entirely by the program that invoked online fsck, and can be amortized if multiple threads enter online fsck at the same time, or if multiple filesystems are being checked at the same time. Changing the branch direction requires taking the CPU hotplug lock, and since CPU initialization requires memory allocation, online fsck must be careful not to change a static key while holding any locks or resources that could be accessed in the memory reclaim paths. To minimize contention on the CPU hotplug lock, care should be taken not to enable or disable static keys unnecessarily.
Because static keys are intended to minimize hook overhead for regular filesystem operations when xfs_scrub is not running, the intended usage patterns are as follows:
The hooked part of XFS should declare a static-scoped static key that defaults to false. The
DEFINE_STATIC_KEY_FALSEmacro takes care of this. The static key itself should be declared as astaticvariable.When deciding to invoke code that’s only used by scrub, the regular filesystem should call the
static_branch_unlikelypredicate to avoid the scrub-only hook code if the static key is not enabled.The regular filesystem should export helper functions that call
static_branch_incto enable andstatic_branch_decto disable the static key. Wrapper functions make it easy to compile out the relevant code if the kernel distributor turns off online fsck at build time.Scrub functions wanting to turn on scrub-only XFS functionality should call the
xchk_fsgates_enablefrom the setup function to enable a specific hook. This must be done before obtaining any resources that are used by memory reclaim. Callers had better be sure they really need the functionality gated by the static key; theTRY_HARDERflag is useful here.
Online scrub has resource acquisition helpers (e.g. xchk_perag_lock) to
handle locking AGI and AGF buffers for all scrubber functions.
If it detects a conflict between scrub and the running transactions, it will
try to wait for intents to complete.
If the caller of the helper has not enabled the static key, the helper will
return -EDEADLOCK, which should result in the scrub being restarted with the
TRY_HARDER flag set.
The scrub setup function should detect that flag, enable the static key, and
try the scrub again.
Scrub teardown disables all static keys obtained by xchk_fsgates_enable.
For more information, please see the kernel documentation of Static Keys.
4.5.5. Pageable Kernel Memory¶
Some online checking functions work by scanning the filesystem to build a shadow copy of an ondisk metadata structure in memory and comparing the two copies. For online repair to rebuild a metadata structure, it must compute the record set that will be stored in the new structure before it can persist that new structure to disk. Ideally, repairs complete with a single atomic commit that introduces a new data structure. To meet these goals, the kernel needs to collect a large amount of information in a place that doesn’t require the correct operation of the filesystem.
Kernel memory isn’t suitable because:
Allocating a contiguous region of memory to create a C array is very difficult, especially on 32-bit systems.
Linked lists of records introduce double pointer overhead which is very high and eliminate the possibility of indexed lookups.
Kernel memory is pinned, which can drive the system into OOM conditions.
The system might not have sufficient memory to stage all the information.
At any given time, online fsck does not need to keep the entire record set in
memory, which means that individual records can be paged out if necessary.
Continued development of online fsck demonstrated that the ability to perform
indexed data storage would also be very useful.
Fortunately, the Linux kernel already has a facility for byte-addressable and
pageable storage: tmpfs.
In-kernel graphics drivers (most notably i915) take advantage of tmpfs files
to store intermediate data that doesn’t need to be in memory at all times, so
that usage precedent is already established.
Hence, the xfile was born!
Historical Sidebar: |
The first edition of online repair inserted records into a new btree as it found them, which failed because filesystem could shut down with a built data structure, which would be live after recovery finished. The second edition solved the half-rebuilt structure problem by storing everything in memory, but frequently ran the system out of memory. The third edition solved the OOM problem by using linked lists, but the memory overhead of the list pointers was extreme. |
4.5.5.1. xfile Access Models¶
A survey of the intended uses of xfiles suggested these use cases:
Arrays of fixed-sized records (space management btrees, directory and extended attribute entries)
Sparse arrays of fixed-sized records (quotas and link counts)
Large binary objects (BLOBs) of variable sizes (directory and extended attribute names and values)
Staging btrees in memory (reverse mapping btrees)
Arbitrary contents (realtime space management)
To support the first four use cases, high level data structures wrap the xfile to share functionality between online fsck functions. The rest of this section discusses the interfaces that the xfile presents to four of those five higher level data structures. The fifth use case is discussed in the realtime summary case study.
XFS is very record-based, which suggests that the ability to load and store
complete records is important.
To support these cases, a pair of xfile_load and xfile_store
functions are provided to read and persist objects into an xfile that treat any
error as an out of memory error. For online repair, squashing error conditions
in this manner is an acceptable behavior because the only reaction is to abort
the operation back to userspace.
However, no discussion of file access idioms is complete without answering the question, “But what about mmap?” It is convenient to access storage directly with pointers, just like userspace code does with regular memory. Online fsck must not drive the system into OOM conditions, which means that xfiles must be responsive to memory reclamation. tmpfs can only push a pagecache folio to the swap cache if the folio is neither pinned nor locked, which means the xfile must not pin too many folios.
Short term direct access to xfile contents is done by locking the pagecache folio and mapping it into kernel address space. Object load and store uses this mechanism. Folio locks are not supposed to be held for long periods of time, so long term direct access to xfile contents is done by bumping the folio refcount, mapping it into kernel address space, and dropping the folio lock. These long term users must be responsive to memory reclaim by hooking into the shrinker infrastructure to know when to release folios.
The xfile_get_folio and xfile_put_folio functions are provided to
retrieve the (locked) folio that backs part of an xfile and to release it.
The only code to use these folio lease functions are the xfarray
sorting algorithms and the in-memory
btrees.
4.5.5.2. xfile Access Coordination¶
For security reasons, xfiles must be owned privately by the kernel.
They are marked S_PRIVATE to prevent interference from the security system,
must never be mapped into process file descriptor tables, and their pages must
never be mapped into userspace processes.
To avoid locking recursion issues with the VFS, all accesses to the shmfs file
are performed by manipulating the page cache directly.
xfile writers call the ->write_begin and ->write_end functions of the
xfile’s address space to grab writable pages, copy the caller’s buffer into the
page, and release the pages.
xfile readers call shmem_read_mapping_page_gfp to grab pages directly
before copying the contents into the caller’s buffer.
In other words, xfiles ignore the VFS read and write code paths to avoid
having to create a dummy struct kiocb and to avoid taking inode and
freeze locks.
tmpfs cannot be frozen, and xfiles must not be exposed to userspace.
If an xfile is shared between threads to stage repairs, the caller must provide its own locks to coordinate access. For example, if a scrub function stores scan results in an xfile and needs other threads to provide updates to the scanned data, the scrub function must provide a lock for all threads to share.
4.5.5.3. Arrays of Fixed-Sized Records¶
In XFS, each type of indexed space metadata (free space, inodes, reference counts, file fork space, and reverse mappings) consists of a set of fixed-size records indexed with a classic B+ tree. Directories have a set of fixed-size dirent records that point to the names, and extended attributes have a set of fixed-size attribute keys that point to names and values. Quota counters and file link counters index records with numbers. During a repair, scrub needs to stage new records during the gathering step and retrieve them during the btree building step.
Although this requirement can be satisfied by calling the read and write
methods of the xfile directly, it is simpler for callers for there to be a
higher level abstraction to take care of computing array offsets, to provide
iterator functions, and to deal with sparse records and sorting.
The xfarray abstraction presents a linear array for fixed-size records atop
the byte-accessible xfile.
4.5.5.3.1. Array Access Patterns¶
Array access patterns in online fsck tend to fall into three categories. Iteration of records is assumed to be necessary for all cases and will be covered in the next section.
The first type of caller handles records that are indexed by position.
Gaps may exist between records, and a record may be updated multiple times
during the collection step.
In other words, these callers want a sparse linearly addressed table file.
The typical use case are quota records or file link count records.
Access to array elements is performed programmatically via xfarray_load and
xfarray_store functions, which wrap the similarly-named xfile functions to
provide loading and storing of array elements at arbitrary array indices.
Gaps are defined to be null records, and null records are defined to be a
sequence of all zero bytes.
Null records are detected by calling xfarray_element_is_null.
They are created either by calling xfarray_unset to null out an existing
record or by never storing anything to an array index.
The second type of caller handles records that are not indexed by position
and do not require multiple updates to a record.
The typical use case here is rebuilding space btrees and key/value btrees.
These callers can add records to the array without caring about array indices
via the xfarray_append function, which stores a record at the end of the
array.
For callers that require records to be presentable in a specific order (e.g.
rebuilding btree data), the xfarray_sort function can arrange the sorted
records; this function will be covered later.
The third type of caller is a bag, which is useful for counting records.
The typical use case here is constructing space extent reference counts from
reverse mapping information.
Records can be put in the bag in any order, they can be removed from the bag
at any time, and uniqueness of records is left to callers.
The xfarray_store_anywhere function is used to insert a record in any
null record slot in the bag; and the xfarray_unset function removes a
record from the bag.
The proposed patchset is the big in-memory array.
4.5.5.3.2. Iterating Array Elements¶
Most users of the xfarray require the ability to iterate the records stored in the array. Callers can probe every possible array index with the following:
xfarray_idx_t i;
foreach_xfarray_idx(array, i) {
xfarray_load(array, i, &rec);
/* do something with rec */
}
All users of this idiom must be prepared to handle null records or must already know that there aren’t any.
For xfarray users that want to iterate a sparse array, the xfarray_iter
function ignores indices in the xfarray that have never been written to by
calling xfile_seek_data (which internally uses SEEK_DATA) to skip areas
of the array that are not populated with memory pages.
Once it finds a page, it will skip the zeroed areas of the page.
xfarray_idx_t i = XFARRAY_CURSOR_INIT;
while ((ret = xfarray_iter(array, &i, &rec)) == 1) {
/* do something with rec */
}
4.5.5.3.3. Sorting Array Elements¶
During the fourth demonstration of online repair, a community reviewer remarked that for performance reasons, online repair ought to load batches of records into btree record blocks instead of inserting records into a new btree one at a time. The btree insertion code in XFS is responsible for maintaining correct ordering of the records, so naturally the xfarray must also support sorting the record set prior to bulk loading.
4.5.5.3.3.1. Case Study: Sorting xfarrays¶
The sorting algorithm used in the xfarray is actually a combination of adaptive
quicksort and a heapsort subalgorithm in the spirit of
Sedgewick and
pdqsort, with customizations for the Linux
kernel.
To sort records in a reasonably short amount of time, xfarray takes
advantage of the binary subpartitioning offered by quicksort, but it also uses
heapsort to hedge against performance collapse if the chosen quicksort pivots
are poor.
Both algorithms are (in general) O(n * lg(n)), but there is a wide performance
gulf between the two implementations.
The Linux kernel already contains a reasonably fast implementation of heapsort. It only operates on regular C arrays, which limits the scope of its usefulness. There are two key places where the xfarray uses it:
Sorting any record subset backed by a single xfile page.
Loading a small number of xfarray records from potentially disparate parts of the xfarray into a memory buffer, and sorting the buffer.
In other words, xfarray uses heapsort to constrain the nested recursion of
quicksort, thereby mitigating quicksort’s worst runtime behavior.
Choosing a quicksort pivot is a tricky business. A good pivot splits the set to sort in half, leading to the divide and conquer behavior that is crucial to O(n * lg(n)) performance. A poor pivot barely splits the subset at all, leading to O(n2) runtime. The xfarray sort routine tries to avoid picking a bad pivot by sampling nine records into a memory buffer and using the kernel heapsort to identify the median of the nine.
Most modern quicksort implementations employ Tukey’s “ninther” to select a pivot from a classic C array. Typical ninther implementations pick three unique triads of records, sort each of the triads, and then sort the middle value of each triad to determine the ninther value. As stated previously, however, xfile accesses are not entirely cheap. It turned out to be much more performant to read the nine elements into a memory buffer, run the kernel’s in-memory heapsort on the buffer, and choose the 4th element of that buffer as the pivot. Tukey’s ninthers are described in J. W. Tukey, The ninther, a technique for low-effort robust (resistant) location in large samples, in Contributions to Survey Sampling and Applied Statistics, edited by H. David, (Academic Press, 1978), pp. 251–257.
The partitioning of quicksort is fairly textbook -- rearrange the record subset around the pivot, then set up the current and next stack frames to sort with the larger and the smaller halves of the pivot, respectively. This keeps the stack space requirements to log2(record count).
As a final performance optimization, the hi and lo scanning phase of quicksort keeps examined xfile pages mapped in the kernel for as long as possible to reduce map/unmap cycles. Surprisingly, this reduces overall sort runtime by nearly half again after accounting for the application of heapsort directly onto xfile pages.
4.5.5.4. Blob Storage¶
Extended attributes and directories add an additional requirement for staging
records: arbitrary byte sequences of finite length.
Each directory entry record needs to store entry name,
and each extended attribute needs to store both the attribute name and value.
The names, keys, and values can consume a large amount of memory, so the
xfblob abstraction was created to simplify management of these blobs
atop an xfile.
Blob arrays provide xfblob_load and xfblob_store functions to retrieve
and persist objects.
The store function returns a magic cookie for every object that it persists.
Later, callers provide this cookie to the xblob_load to recall the object.
The xfblob_free function frees a specific blob, and the xfblob_truncate
function frees them all because compaction is not needed.
The details of repairing directories and extended attributes will be discussed in a subsequent section about atomic file content exchanges. However, it should be noted that these repair functions only use blob storage to cache a small number of entries before adding them to a temporary ondisk file, which is why compaction is not required.
The proposed patchset is at the start of the extended attribute repair series.
4.5.5.5. In-Memory B+Trees¶
The chapter about secondary metadata mentioned that checking and repairing of secondary metadata commonly requires coordination between a live metadata scan of the filesystem and writer threads that are updating that metadata. Keeping the scan data up to date requires requires the ability to propagate metadata updates from the filesystem into the data being collected by the scan. This can be done by appending concurrent updates into a separate log file and applying them before writing the new metadata to disk, but this leads to unbounded memory consumption if the rest of the system is very busy. Another option is to skip the side-log and commit live updates from the filesystem directly into the scan data, which trades more overhead for a lower maximum memory requirement. In both cases, the data structure holding the scan results must support indexed access to perform well.
Given that indexed lookups of scan data is required for both strategies, online fsck employs the second strategy of committing live updates directly into scan data. Because xfarrays are not indexed and do not enforce record ordering, they are not suitable for this task. Conveniently, however, XFS has a library to create and maintain ordered reverse mapping records: the existing rmap btree code! If only there was a means to create one in memory.
Recall that the xfile abstraction represents memory pages as a
regular file, which means that the kernel can create byte or block addressable
virtual address spaces at will.
The XFS buffer cache specializes in abstracting IO to block-oriented address
spaces, which means that adaptation of the buffer cache to interface with
xfiles enables reuse of the entire btree library.
Btrees built atop an xfile are collectively known as xfbtrees.
The next few sections describe how they actually work.
The proposed patchset is the in-memory btree series.
4.5.5.5.1. Using xfiles as a Buffer Cache Target¶
Two modifications are necessary to support xfiles as a buffer cache target.
The first is to make it possible for the struct xfs_buftarg structure to
host the struct xfs_buf rhashtable, because normally those are held by a
per-AG structure.
The second change is to modify the buffer ioapply function to “read” cached
pages from the xfile and “write” cached pages back to the xfile.
Multiple access to individual buffers is controlled by the xfs_buf lock,
since the xfile does not provide any locking on its own.
With this adaptation in place, users of the xfile-backed buffer cache use
exactly the same APIs as users of the disk-backed buffer cache.
The separation between xfile and buffer cache implies higher memory usage since
they do not share pages, but this property could some day enable transactional
updates to an in-memory btree.
Today, however, it simply eliminates the need for new code.
4.5.5.5.2. Space Management with an xfbtree¶
Space management for an xfile is very simple -- each btree block is one memory page in size. These blocks use the same header format as an on-disk btree, but the in-memory block verifiers ignore the checksums, assuming that xfile memory is no more corruption-prone than regular DRAM. Reusing existing code here is more important than absolute memory efficiency.
The very first block of an xfile backing an xfbtree contains a header block. The header describes the owner, height, and the block number of the root xfbtree block.
To allocate a btree block, use xfile_seek_data to find a gap in the file.
If there are no gaps, create one by extending the length of the xfile.
Preallocate space for the block with xfile_prealloc, and hand back the
location.
To free an xfbtree block, use xfile_discard (which internally uses
FALLOC_FL_PUNCH_HOLE) to remove the memory page from the xfile.
4.5.5.5.3. Populating an xfbtree¶
An online fsck function that wants to create an xfbtree should proceed as follows:
Call
xfile_createto create an xfile.Call
xfs_alloc_memory_buftargto create a buffer cache target structure pointing to the xfile.Pass the buffer cache target, buffer ops, and other information to
xfbtree_initto initialize the passed instruct xfbtreeand write an initial root block to the xfile. Each btree type should define a wrapper that passes necessary arguments to the creation function. For example, rmap btrees definexfs_rmapbt_mem_createto take care of all the necessary details for callers.Pass the xfbtree object to the btree cursor creation function for the btree type. Following the example above,
xfs_rmapbt_mem_cursortakes care of this for callers.Pass the btree cursor to the regular btree functions to make queries against and to update the in-memory btree. For example, a btree cursor for an rmap xfbtree can be passed to the
xfs_rmap_*functions just like any other btree cursor. See the next section for information on dealing with xfbtree updates that are logged to a transaction.When finished, delete the btree cursor, destroy the xfbtree object, free the buffer target, and the destroy the xfile to release all resources.
4.5.5.5.4. Committing Logged xfbtree Buffers¶
Although it is a clever hack to reuse the rmap btree code to handle the staging structure, the ephemeral nature of the in-memory btree block storage presents some challenges of its own. The XFS transaction manager must not commit buffer log items for buffers backed by an xfile because the log format does not understand updates for devices other than the data device. An ephemeral xfbtree probably will not exist by the time the AIL checkpoints log transactions back into the filesystem, and certainly won’t exist during log recovery. For these reasons, any code updating an xfbtree in transaction context must remove the buffer log items from the transaction and write the updates into the backing xfile before committing or cancelling the transaction.
The xfbtree_trans_commit and xfbtree_trans_cancel functions implement
this functionality as follows:
Find each buffer log item whose buffer targets the xfile.
Record the dirty/ordered status of the log item.
Detach the log item from the buffer.
Queue the buffer to a special delwri list.
Clear the transaction dirty flag if the only dirty log items were the ones that were detached in step 3.
Submit the delwri list to commit the changes to the xfile, if the updates are being committed.
After removing xfile logged buffers from the transaction in this manner, the transaction can be committed or cancelled.
4.5.6. Bulk Loading of Ondisk B+Trees¶
As mentioned previously, early iterations of online repair built new btree structures by creating a new btree and adding observations individually. Loading a btree one record at a time had a slight advantage of not requiring the incore records to be sorted prior to commit, but was very slow and leaked blocks if the system went down during a repair. Loading records one at a time also meant that repair could not control the loading factor of the blocks in the new btree.
Fortunately, the venerable xfs_repair tool had a more efficient means for
rebuilding a btree index from a collection of records -- bulk btree loading.
This was implemented rather inefficiently code-wise, since xfs_repair
had separate copy-pasted implementations for each btree type.
To prepare for online fsck, each of the four bulk loaders were studied, notes were taken, and the four were refactored into a single generic btree bulk loading mechanism. Those notes in turn have been refreshed and are presented below.
4.5.6.1. Geometry Computation¶
The zeroth step of bulk loading is to assemble the entire record set that will
be stored in the new btree, and sort the records.
Next, call xfs_btree_bload_compute_geometry to compute the shape of the
btree from the record set, the type of btree, and any load factor preferences.
This information is required for resource reservation.
First, the geometry computation computes the minimum and maximum records that will fit in a leaf block from the size of a btree block and the size of the block header. Roughly speaking, the maximum number of records is:
maxrecs = (block_size - header_size) / record_size
The XFS design specifies that btree blocks should be merged when possible, which means the minimum number of records is half of maxrecs:
minrecs = maxrecs / 2
The next variable to determine is the desired loading factor. This must be at least minrecs and no more than maxrecs. Choosing minrecs is undesirable because it wastes half the block. Choosing maxrecs is also undesirable because adding a single record to each newly rebuilt leaf block will cause a tree split, which causes a noticeable drop in performance immediately afterwards. The default loading factor was chosen to be 75% of maxrecs, which provides a reasonably compact structure without any immediate split penalties:
default_load_factor = (maxrecs + minrecs) / 2
If space is tight, the loading factor will be set to maxrecs to try to avoid running out of space:
leaf_load_factor = enough space ? default_load_factor : maxrecs
Load factor is computed for btree node blocks using the combined size of the btree key and pointer as the record size:
maxrecs = (block_size - header_size) / (key_size + ptr_size)
minrecs = maxrecs / 2
node_load_factor = enough space ? default_load_factor : maxrecs
Once that’s done, the number of leaf blocks required to store the record set can be computed as:
leaf_blocks = ceil(record_count / leaf_load_factor)
The number of node blocks needed to point to the next level down in the tree is computed as:
n_blocks = (n == 0 ? leaf_blocks : node_blocks[n])
node_blocks[n + 1] = ceil(n_blocks / node_load_factor)
The entire computation is performed recursively until the current level only needs one block. The resulting geometry is as follows:
For AG-rooted btrees, this level is the root level, so the height of the new tree is
level + 1and the space needed is the summation of the number of blocks on each level.For inode-rooted btrees where the records in the top level do not fit in the inode fork area, the height is
level + 2, the space needed is the summation of the number of blocks on each level, and the inode fork points to the root block.For inode-rooted btrees where the records in the top level can be stored in the inode fork area, then the root block can be stored in the inode, the height is
level + 1, and the space needed is one less than the summation of the number of blocks on each level. This only becomes relevant when non-bmap btrees gain the ability to root in an inode, which is a future patchset and only included here for completeness.
4.5.6.2. Reserving New B+Tree Blocks¶
Once repair knows the number of blocks needed for the new btree, it allocates
those blocks using the free space information.
Each reserved extent is tracked separately by the btree builder state data.
To improve crash resilience, the reservation code also logs an Extent Freeing
Intent (EFI) item in the same transaction as each space allocation and attaches
its in-memory struct xfs_extent_free_item object to the space reservation.
If the system goes down, log recovery will use the unfinished EFIs to free the
unused space, the free space, leaving the filesystem unchanged.
Each time the btree builder claims a block for the btree from a reserved extent, it updates the in-memory reservation to reflect the claimed space. Block reservation tries to allocate as much contiguous space as possible to reduce the number of EFIs in play.
While repair is writing these new btree blocks, the EFIs created for the space reservations pin the tail of the ondisk log. It’s possible that other parts of the system will remain busy and push the head of the log towards the pinned tail. To avoid livelocking the filesystem, the EFIs must not pin the tail of the log for too long. To alleviate this problem, the dynamic relogging capability of the deferred ops mechanism is reused here to commit a transaction at the log head containing an EFD for the old EFI and new EFI at the head. This enables the log to release the old EFI to keep the log moving forwards.
EFIs have a role to play during the commit and reaping phases; please see the next section and the section about reaping for more details.
Proposed patchsets are the bitmap rework and the preparation for bulk loading btrees.
4.5.6.3. Writing the New Tree¶
This part is pretty simple -- the btree builder (xfs_btree_bulkload) claims
a block from the reserved list, writes the new btree block header, fills the
rest of the block with records, and adds the new leaf block to a list of
written blocks:
┌────┐
│leaf│
│RRR │
└────┘
Sibling pointers are set every time a new block is added to the level:
┌────┐ ┌────┐ ┌────┐ ┌────┐
│leaf│→│leaf│→│leaf│→│leaf│
│RRR │←│RRR │←│RRR │←│RRR │
└────┘ └────┘ └────┘ └────┘
When it finishes writing the record leaf blocks, it moves on to the node blocks To fill a node block, it walks each block in the next level down in the tree to compute the relevant keys and write them into the parent node:
┌────┐ ┌────┐
│node│──────→│node│
│PP │←──────│PP │
└────┘ └────┘
↙ ↘ ↙ ↘
┌────┐ ┌────┐ ┌────┐ ┌────┐
│leaf│→│leaf│→│leaf│→│leaf│
│RRR │←│RRR │←│RRR │←│RRR │
└────┘ └────┘ └────┘ └────┘
When it reaches the root level, it is ready to commit the new btree!:
┌─────────┐
│ root │
│ PP │
└─────────┘
↙ ↘
┌────┐ ┌────┐
│node│──────→│node│
│PP │←──────│PP │
└────┘ └────┘
↙ ↘ ↙ ↘
┌────┐ ┌────┐ ┌────┐ ┌────┐
│leaf│→│leaf│→│leaf│→│leaf│
│RRR │←│RRR │←│RRR │←│RRR │
└────┘ └────┘ └────┘ └────┘
The first step to commit the new btree is to persist the btree blocks to disk
synchronously.
This is a little complicated because a new btree block could have been freed
in the recent past, so the builder must use xfs_buf_delwri_queue_here to
remove the (stale) buffer from the AIL list before it can write the new blocks
to disk.
Blocks are queued for IO using a delwri list and written in one large batch
with xfs_buf_delwri_submit.
Once the new blocks have been persisted to disk, control returns to the individual repair function that called the bulk loader. The repair function must log the location of the new root in a transaction, clean up the space reservations that were made for the new btree, and reap the old metadata blocks:
Commit the location of the new btree root.
For each incore reservation:
Log Extent Freeing Done (EFD) items for all the space that was consumed by the btree builder. The new EFDs must point to the EFIs attached to the reservation to prevent log recovery from freeing the new blocks.
For unclaimed portions of incore reservations, create a regular deferred extent free work item to be free the unused space later in the transaction chain.
The EFDs and EFIs logged in steps 2a and 2b must not overrun the reservation of the committing transaction. If the btree loading code suspects this might be about to happen, it must call
xrep_defer_finishto clear out the deferred work and obtain a fresh transaction.
Clear out the deferred work a second time to finish the commit and clean the repair transaction.
The transaction rolling in steps 2c and 3 represent a weakness in the repair algorithm, because a log flush and a crash before the end of the reap step can result in space leaking. Online repair functions minimize the chances of this occurring by using very large transactions, which each can accommodate many thousands of block freeing instructions. Repair moves on to reaping the old blocks, which will be presented in a subsequent section after a few case studies of bulk loading.
4.5.6.3.1. Case Study: Rebuilding the Inode Index¶
The high level process to rebuild the inode index btree is:
Walk the reverse mapping records to generate
struct xfs_inobt_recrecords from the inode chunk information and a bitmap of the old inode btree blocks.Append the records to an xfarray in inode order.
Use the
xfs_btree_bload_compute_geometryfunction to compute the number of blocks needed for the inode btree. If the free space inode btree is enabled, call it again to estimate the geometry of the finobt.Allocate the number of blocks computed in the previous step.
Use
xfs_btree_bloadto write the xfarray records to btree blocks and generate the internal node blocks. If the free space inode btree is enabled, call it again to load the finobt.Commit the location of the new btree root block(s) to the AGI.
Reap the old btree blocks using the bitmap created in step 1.
Details are as follows.
The inode btree maps inumbers to the ondisk location of the associated
inode records, which means that the inode btrees can be rebuilt from the
reverse mapping information.
Reverse mapping records with an owner of XFS_RMAP_OWN_INOBT marks the
location of the old inode btree blocks.
Each reverse mapping record with an owner of XFS_RMAP_OWN_INODES marks the
location of at least one inode cluster buffer.
A cluster is the smallest number of ondisk inodes that can be allocated or
freed in a single transaction; it is never smaller than 1 fs block or 4 inodes.
For the space represented by each inode cluster, ensure that there are no
records in the free space btrees nor any records in the reference count btree.
If there are, the space metadata inconsistencies are reason enough to abort the
operation.
Otherwise, read each cluster buffer to check that its contents appear to be
ondisk inodes and to decide if the file is allocated
(xfs_dinode.i_mode != 0) or free (xfs_dinode.i_mode == 0).
Accumulate the results of successive inode cluster buffer reads until there is
enough information to fill a single inode chunk record, which is 64 consecutive
numbers in the inumber keyspace.
If the chunk is sparse, the chunk record may include holes.
Once the repair function accumulates one chunk’s worth of data, it calls
xfarray_append to add the inode btree record to the xfarray.
This xfarray is walked twice during the btree creation step -- once to populate
the inode btree with all inode chunk records, and a second time to populate the
free inode btree with records for chunks that have free non-sparse inodes.
The number of records for the inode btree is the number of xfarray records,
but the record count for the free inode btree has to be computed as inode chunk
records are stored in the xfarray.
The proposed patchset is the AG btree repair series.
4.5.6.3.2. Case Study: Rebuilding the Space Reference Counts¶
Reverse mapping records are used to rebuild the reference count information. Reference counts are required for correct operation of copy on write for shared file data. Imagine the reverse mapping entries as rectangles representing extents of physical blocks, and that the rectangles can be laid down to allow them to overlap each other. From the diagram below, it is apparent that a reference count record must start or end wherever the height of the stack changes. In other words, the record emission stimulus is level-triggered:
█ ███
██ █████ ████ ███ ██████
██ ████ ███████████ ████ █████████
████████████████████████████████ ███████████
^ ^ ^^ ^^ ^ ^^ ^^^ ^^^^ ^ ^^ ^ ^ ^
2 1 23 21 3 43 234 2123 1 01 2 3 0
The ondisk reference count btree does not store the refcount == 0 cases because the free space btree already records which blocks are free. Extents being used to stage copy-on-write operations should be the only records with refcount == 1. Single-owner file blocks aren’t recorded in either the free space or the reference count btrees.
The high level process to rebuild the reference count btree is:
Walk the reverse mapping records to generate
struct xfs_refcount_irecrecords for any space having more than one reverse mapping and add them to the xfarray. Any records owned byXFS_RMAP_OWN_COWare also added to the xfarray because these are extents allocated to stage a copy on write operation and are tracked in the refcount btree.Use any records owned by
XFS_RMAP_OWN_REFCto create a bitmap of old refcount btree blocks.Sort the records in physical extent order, putting the CoW staging extents at the end of the xfarray. This matches the sorting order of records in the refcount btree.
Use the
xfs_btree_bload_compute_geometryfunction to compute the number of blocks needed for the new tree.Allocate the number of blocks computed in the previous step.
Use
xfs_btree_bloadto write the xfarray records to btree blocks and generate the internal node blocks.Commit the location of new btree root block to the AGF.
Reap the old btree blocks using the bitmap created in step 1.
Details are as follows; the same algorithm is used by xfs_repair to
generate refcount information from reverse mapping records.
Until the reverse mapping btree runs out of records:
Retrieve the next record from the btree and put it in a bag.
Collect all records with the same starting block from the btree and put them in the bag.
While the bag isn’t empty:
Among the mappings in the bag, compute the lowest block number where the reference count changes. This position will be either the starting block number of the next unprocessed reverse mapping or the next block after the shortest mapping in the bag.
Remove all mappings from the bag that end at this position.
Collect all reverse mappings that start at this position from the btree and put them in the bag.
If the size of the bag changed and is greater than one, create a new refcount record associating the block number range that we just walked to the size of the bag.
The bag-like structure in this case is a type 2 xfarray as discussed in the
xfarray access patterns section.
Reverse mappings are added to the bag using xfarray_store_anywhere and
removed via xfarray_unset.
Bag members are examined through xfarray_iter loops.
The proposed patchset is the AG btree repair series.
4.5.6.3.3. Case Study: Rebuilding File Fork Mapping Indices¶
The high level process to rebuild a data/attr fork mapping btree is:
Walk the reverse mapping records to generate
struct xfs_bmbt_recrecords from the reverse mapping records for that inode and fork. Append these records to an xfarray. Compute the bitmap of the old bmap btree blocks from theBMBT_BLOCKrecords.Use the
xfs_btree_bload_compute_geometryfunction to compute the number of blocks needed for the new tree.Sort the records in file offset order.
If the extent records would fit in the inode fork immediate area, commit the records to that immediate area and skip to step 8.
Allocate the number of blocks computed in the previous step.
Use
xfs_btree_bloadto write the xfarray records to btree blocks and generate the internal node blocks.Commit the new btree root block to the inode fork immediate area.
Reap the old btree blocks using the bitmap created in step 1.
There are some complications here: First, it’s possible to move the fork offset to adjust the sizes of the immediate areas if the data and attr forks are not both in BMBT format. Second, if there are sufficiently few fork mappings, it may be possible to use EXTENTS format instead of BMBT, which may require a conversion. Third, the incore extent map must be reloaded carefully to avoid disturbing any delayed allocation extents.
The proposed patchset is the file mapping repair series.
4.5.7. Reaping Old Metadata Blocks¶
Whenever online fsck builds a new data structure to replace one that is suspect, there is a question of how to find and dispose of the blocks that belonged to the old structure. The laziest method of course is not to deal with them at all, but this slowly leads to service degradations as space leaks out of the filesystem. Hopefully, someone will schedule a rebuild of the free space information to plug all those leaks. Offline repair rebuilds all space metadata after recording the usage of the files and directories that it decides not to clear, hence it can build new structures in the discovered free space and avoid the question of reaping.
As part of a repair, online fsck relies heavily on the reverse mapping records to find space that is owned by the corresponding rmap owner yet truly free. Cross referencing rmap records with other rmap records is necessary because there may be other data structures that also think they own some of those blocks (e.g. crosslinked trees). Permitting the block allocator to hand them out again will not push the system towards consistency.
For space metadata, the process of finding extents to dispose of generally follows this format:
Create a bitmap of space used by data structures that must be preserved. The space reservations used to create the new metadata can be used here if the same rmap owner code is used to denote all of the objects being rebuilt.
Survey the reverse mapping data to create a bitmap of space owned by the same
XFS_RMAP_OWN_*number for the metadata that is being preserved.Use the bitmap disunion operator to subtract (1) from (2). The remaining set bits represent candidate extents that could be freed. The process moves on to step 4 below.
Repairs for file-based metadata such as extended attributes, directories, symbolic links, quota files and realtime bitmaps are performed by building a new structure attached to a temporary file and exchanging all mappings in the file forks. Afterward, the mappings in the old file fork are the candidate blocks for disposal.
The process for disposing of old extents is as follows:
For each candidate extent, count the number of reverse mapping records for the first block in that extent that do not have the same rmap owner for the data structure being repaired.
If zero, the block has a single owner and can be freed.
If not, the block is part of a crosslinked structure and must not be freed.
Starting with the next block in the extent, figure out how many more blocks have the same zero/nonzero other owner status as that first block.
If the region is crosslinked, delete the reverse mapping entry for the structure being repaired and move on to the next region.
If the region is to be freed, mark any corresponding buffers in the buffer cache as stale to prevent log writeback.
Free the region and move on.
However, there is one complication to this procedure. Transactions are of finite size, so the reaping process must be careful to roll the transactions to avoid overruns. Overruns come from two sources:
EFIs logged on behalf of space that is no longer occupied
Log items for buffer invalidations
This is also a window in which a crash during the reaping process can leak blocks. As stated earlier, online repair functions use very large transactions to minimize the chances of this occurring.
The proposed patchset is the preparation for bulk loading btrees series.
4.5.7.1. Case Study: Reaping After a Regular Btree Repair¶
Old reference count and inode btrees are the easiest to reap because they have
rmap records with special owner codes: XFS_RMAP_OWN_REFC for the refcount
btree, and XFS_RMAP_OWN_INOBT for the inode and free inode btrees.
Creating a list of extents to reap the old btree blocks is quite simple,
conceptually:
Lock the relevant AGI/AGF header buffers to prevent allocation and frees.
For each reverse mapping record with an rmap owner corresponding to the metadata structure being rebuilt, set the corresponding range in a bitmap.
Walk the current data structures that have the same rmap owner. For each block visited, clear that range in the above bitmap.
Each set bit in the bitmap represents a block that could be a block from the old data structures and hence is a candidate for reaping. In other words,
(rmap_records_owned_by & ~blocks_reachable_by_walk)are the blocks that might be freeable.
If it is possible to maintain the AGF lock throughout the repair (which is the common case), then step 2 can be performed at the same time as the reverse mapping record walk that creates the records for the new btree.
4.5.7.2. Case Study: Rebuilding the Free Space Indices¶
The high level process to rebuild the free space indices is:
Walk the reverse mapping records to generate
struct xfs_alloc_rec_incorerecords from the gaps in the reverse mapping btree.Append the records to an xfarray.
Use the
xfs_btree_bload_compute_geometryfunction to compute the number of blocks needed for each new tree.Allocate the number of blocks computed in the previous step from the free space information collected.
Use
xfs_btree_bloadto write the xfarray records to btree blocks and generate the internal node blocks for the free space by length index. Call it again for the free space by block number index.Commit the locations of the new btree root blocks to the AGF.
Reap the old btree blocks by looking for space that is not recorded by the reverse mapping btree, the new free space btrees, or the AGFL.
Repairing the free space btrees has three key complications over a regular btree repair:
First, free space is not explicitly tracked in the reverse mapping records. Hence, the new free space records must be inferred from gaps in the physical space component of the keyspace of the reverse mapping btree.
Second, free space repairs cannot use the common btree reservation code because new blocks are reserved out of the free space btrees. This is impossible when repairing the free space btrees themselves. However, repair holds the AGF buffer lock for the duration of the free space index reconstruction, so it can use the collected free space information to supply the blocks for the new free space btrees. It is not necessary to back each reserved extent with an EFI because the new free space btrees are constructed in what the ondisk filesystem thinks is unowned space. However, if reserving blocks for the new btrees from the collected free space information changes the number of free space records, repair must re-estimate the new free space btree geometry with the new record count until the reservation is sufficient. As part of committing the new btrees, repair must ensure that reverse mappings are created for the reserved blocks and that unused reserved blocks are inserted into the free space btrees. Deferrred rmap and freeing operations are used to ensure that this transition is atomic, similar to the other btree repair functions.
Third, finding the blocks to reap after the repair is not overly
straightforward.
Blocks for the free space btrees and the reverse mapping btrees are supplied by
the AGFL.
Blocks put onto the AGFL have reverse mapping records with the owner
XFS_RMAP_OWN_AG.
This ownership is retained when blocks move from the AGFL into the free space
btrees or the reverse mapping btrees.
When repair walks reverse mapping records to synthesize free space records, it
creates a bitmap (ag_owner_bitmap) of all the space claimed by
XFS_RMAP_OWN_AG records.
The repair context maintains a second bitmap corresponding to the rmap btree
blocks and the AGFL blocks (rmap_agfl_bitmap).
When the walk is complete, the bitmap disunion operation (ag_owner_bitmap &
~rmap_agfl_bitmap) computes the extents that are used by the old free space
btrees.
These blocks can then be reaped using the methods outlined above.
The proposed patchset is the AG btree repair series.
4.5.7.3. Case Study: Reaping After Repairing Reverse Mapping Btrees¶
Old reverse mapping btrees are less difficult to reap after a repair.
As mentioned in the previous section, blocks on the AGFL, the two free space
btree blocks, and the reverse mapping btree blocks all have reverse mapping
records with XFS_RMAP_OWN_AG as the owner.
The full process of gathering reverse mapping records and building a new btree
are described in the case study of
live rebuilds of rmap data, but a crucial point from that
discussion is that the new rmap btree will not contain any records for the old
rmap btree, nor will the old btree blocks be tracked in the free space btrees.
The list of candidate reaping blocks is computed by setting the bits
corresponding to the gaps in the new rmap btree records, and then clearing the
bits corresponding to extents in the free space btrees and the current AGFL
blocks.
The result (new_rmapbt_gaps & ~(agfl | bnobt_records)) are reaped using the
methods outlined above.
The rest of the process of rebuildng the reverse mapping btree is discussed in a separate case study.
The proposed patchset is the AG btree repair series.
4.5.7.4. Case Study: Rebuilding the AGFL¶
The allocation group free block list (AGFL) is repaired as follows:
Create a bitmap for all the space that the reverse mapping data claims is owned by
XFS_RMAP_OWN_AG.Subtract the space used by the two free space btrees and the rmap btree.
Subtract any space that the reverse mapping data claims is owned by any other owner, to avoid re-adding crosslinked blocks to the AGFL.
Once the AGFL is full, reap any blocks leftover.
The next operation to fix the freelist will right-size the list.
See fs/xfs/scrub/agheader_repair.c for more details.
4.5.8. Inode Record Repairs¶
Inode records must be handled carefully, because they have both ondisk records (“dinodes”) and an in-memory (“cached”) representation. There is a very high potential for cache coherency issues if online fsck is not careful to access the ondisk metadata only when the ondisk metadata is so badly damaged that the filesystem cannot load the in-memory representation. When online fsck wants to open a damaged file for scrubbing, it must use specialized resource acquisition functions that return either the in-memory representation or a lock on whichever object is necessary to prevent any update to the ondisk location.
The only repairs that should be made to the ondisk inode buffers are whatever
is necessary to get the in-core structure loaded.
This means fixing whatever is caught by the inode cluster buffer and inode fork
verifiers, and retrying the iget operation.
If the second iget fails, the repair has failed.
Once the in-memory representation is loaded, repair can lock the inode and can subject it to comprehensive checks, repairs, and optimizations. Most inode attributes are easy to check and constrain, or are user-controlled arbitrary bit patterns; these are both easy to fix. Dealing with the data and attr fork extent counts and the file block counts is more complicated, because computing the correct value requires traversing the forks, or if that fails, leaving the fields invalid and waiting for the fork fsck functions to run.
The proposed patchset is the inode repair series.
4.5.9. Quota Record Repairs¶
Similar to inodes, quota records (“dquots”) also have both ondisk records and an in-memory representation, and hence are subject to the same cache coherency issues. Somewhat confusingly, both are known as dquots in the XFS codebase.
The only repairs that should be made to the ondisk quota record buffers are whatever is necessary to get the in-core structure loaded. Once the in-memory representation is loaded, the only attributes needing checking are obviously bad limits and timer values.
Quota usage counters are checked, repaired, and discussed separately in the section about live quotacheck.
The proposed patchset is the quota repair series.
4.5.10. Freezing to Fix Summary Counters¶
Filesystem summary counters track availability of filesystem resources such as free blocks, free inodes, and allocated inodes. This information could be compiled by walking the free space and inode indexes, but this is a slow process, so XFS maintains a copy in the ondisk superblock that should reflect the ondisk metadata, at least when the filesystem has been unmounted cleanly. For performance reasons, XFS also maintains incore copies of those counters, which are key to enabling resource reservations for active transactions. Writer threads reserve the worst-case quantities of resources from the incore counter and give back whatever they don’t use at commit time. It is therefore only necessary to serialize on the superblock when the superblock is being committed to disk.
The lazy superblock counter feature introduced in XFS v5 took this even further by training log recovery to recompute the summary counters from the AG headers, which eliminated the need for most transactions even to touch the superblock. The only time XFS commits the summary counters is at filesystem unmount. To reduce contention even further, the incore counter is implemented as a percpu counter, which means that each CPU is allocated a batch of blocks from a global incore counter and can satisfy small allocations from the local batch.
The high-performance nature of the summary counters makes it difficult for online fsck to check them, since there is no way to quiesce a percpu counter while the system is running. Although online fsck can read the filesystem metadata to compute the correct values of the summary counters, there’s no way to hold the value of a percpu counter stable, so it’s quite possible that the counter will be out of date by the time the walk is complete. Earlier versions of online scrub would return to userspace with an incomplete scan flag, but this is not a satisfying outcome for a system administrator. For repairs, the in-memory counters must be stabilized while walking the filesystem metadata to get an accurate reading and install it in the percpu counter.
To satisfy this requirement, online fsck must prevent other programs in the system from initiating new writes to the filesystem, it must disable background garbage collection threads, and it must wait for existing writer programs to exit the kernel. Once that has been established, scrub can walk the AG free space indexes, the inode btrees, and the realtime bitmap to compute the correct value of all four summary counters. This is very similar to a filesystem freeze, though not all of the pieces are necessary:
The final freeze state is set one higher than
SB_FREEZE_COMPLETEto prevent other threads from thawing the filesystem, or other scrub threads from initiating another fscounters freeze.It does not quiesce the log.
With this code in place, it is now possible to pause the filesystem for just long enough to check and correct the summary counters.
Historical Sidebar: |
The initial implementation used the actual VFS filesystem freeze mechanism to quiesce filesystem activity. With the filesystem frozen, it is possible to resolve the counter values with exact precision, but there are many problems with calling the VFS methods directly:
|
The proposed patchset is the summary counter cleanup series.
4.5.11. Full Filesystem Scans¶
Certain types of metadata can only be checked by walking every file in the entire filesystem to record observations and comparing the observations against what’s recorded on disk. Like every other type of online repair, repairs are made by writing those observations to disk in a replacement structure and committing it atomically. However, it is not practical to shut down the entire filesystem to examine hundreds of billions of files because the downtime would be excessive. Therefore, online fsck must build the infrastructure to manage a live scan of all the files in the filesystem. There are two questions that need to be solved to perform a live walk:
How does scrub manage the scan while it is collecting data?
How does the scan keep abreast of changes being made to the system by other threads?
4.5.11.1. Coordinated Inode Scans¶
In the original Unix filesystems of the 1970s, each directory entry contained an index number (inumber) which was used as an index into on ondisk array (itable) of fixed-size records (inodes) describing a file’s attributes and its data block mapping. This system is described by J. Lions, “inode (5659)” in Lions’ Commentary on UNIX, 6th Edition, (Dept. of Computer Science, the University of New South Wales, November 1977), pp. 18-2; and later by D. Ritchie and K. Thompson, “Implementation of the File System”, from The UNIX Time-Sharing System, (The Bell System Technical Journal, July 1978), pp. 1913-4.
XFS retains most of this design, except now inumbers are search keys over all
the space in the data section filesystem.
They form a continuous keyspace that can be expressed as a 64-bit integer,
though the inodes themselves are sparsely distributed within the keyspace.
Scans proceed in a linear fashion across the inumber keyspace, starting from
0x0 and ending at 0xFFFFFFFFFFFFFFFF.
Naturally, a scan through a keyspace requires a scan cursor object to track the
scan progress.
Because this keyspace is sparse, this cursor contains two parts.
The first part of this scan cursor object tracks the inode that will be
examined next; call this the examination cursor.
Somewhat less obviously, the scan cursor object must also track which parts of
the keyspace have already been visited, which is critical for deciding if a
concurrent filesystem update needs to be incorporated into the scan data.
Call this the visited inode cursor.
Advancing the scan cursor is a multi-step process encapsulated in
xchk_iscan_iter:
Lock the AGI buffer of the AG containing the inode pointed to by the visited inode cursor. This guarantee that inodes in this AG cannot be allocated or freed while advancing the cursor.
Use the per-AG inode btree to look up the next inumber after the one that was just visited, since it may not be keyspace adjacent.
If there are no more inodes left in this AG:
Move the examination cursor to the point of the inumber keyspace that corresponds to the start of the next AG.
Adjust the visited inode cursor to indicate that it has “visited” the last possible inode in the current AG’s inode keyspace. XFS inumbers are segmented, so the cursor needs to be marked as having visited the entire keyspace up to just before the start of the next AG’s inode keyspace.
Unlock the AGI and return to step 1 if there are unexamined AGs in the filesystem.
If there are no more AGs to examine, set both cursors to the end of the inumber keyspace. The scan is now complete.
Otherwise, there is at least one more inode to scan in this AG:
Move the examination cursor ahead to the next inode marked as allocated by the inode btree.
Adjust the visited inode cursor to point to the inode just prior to where the examination cursor is now. Because the scanner holds the AGI buffer lock, no inodes could have been created in the part of the inode keyspace that the visited inode cursor just advanced.
Get the incore inode for the inumber of the examination cursor. By maintaining the AGI buffer lock until this point, the scanner knows that it was safe to advance the examination cursor across the entire keyspace, and that it has stabilized this next inode so that it cannot disappear from the filesystem until the scan releases the incore inode.
Drop the AGI lock and return the incore inode to the caller.
Online fsck functions scan all files in the filesystem as follows:
Start a scan by calling
xchk_iscan_start.Advance the scan cursor (
xchk_iscan_iter) to get the next inode. If one is provided:Lock the inode to prevent updates during the scan.
Scan the inode.
While still holding the inode lock, adjust the visited inode cursor (
xchk_iscan_mark_visited) to point to this inode.Unlock and release the inode.
Call
xchk_iscan_teardownto complete the scan.
There are subtleties with the inode cache that complicate grabbing the incore inode for the caller. Obviously, it is an absolute requirement that the inode metadata be consistent enough to load it into the inode cache. Second, if the incore inode is stuck in some intermediate state, the scan coordinator must release the AGI and push the main filesystem to get the inode back into a loadable state.
The proposed patches are the inode scanner series. The first user of the new functionality is the online quotacheck series.
4.5.11.2. Inode Management¶
In regular filesystem code, references to allocated XFS incore inodes are
always obtained (xfs_iget) outside of transaction context because the
creation of the incore context for an existing file does not require metadata
updates.
However, it is important to note that references to incore inodes obtained as
part of file creation must be performed in transaction context because the
filesystem must ensure the atomicity of the ondisk inode btree index updates
and the initialization of the actual ondisk inode.
References to incore inodes are always released (xfs_irele) outside of
transaction context because there are a handful of activities that might
require ondisk updates:
The VFS may decide to kick off writeback as part of a
DONTCACHEinode release.Speculative preallocations need to be unreserved.
An unlinked file may have lost its last reference, in which case the entire file must be inactivated, which involves releasing all of its resources in the ondisk metadata and freeing the inode.
These activities are collectively called inode inactivation. Inactivation has two parts -- the VFS part, which initiates writeback on all dirty file pages, and the XFS part, which cleans up XFS-specific information and frees the inode if it was unlinked. If the inode is unlinked (or unconnected after a file handle operation), the kernel drops the inode into the inactivation machinery immediately.
During normal operation, resource acquisition for an update follows this order to avoid deadlocks:
Inode reference (
iget).Filesystem freeze protection, if repairing (
mnt_want_write_file).Inode
IOLOCK(VFSi_rwsem) lock to control file IO.Inode
MMAPLOCK(page cacheinvalidate_lock) lock for operations that can update page cache mappings.Log feature enablement.
Transaction log space grant.
Space on the data and realtime devices for the transaction.
Incore dquot references, if a file is being repaired. Note that they are not locked, merely acquired.
Inode
ILOCKfor file metadata updates.AG header buffer locks / Realtime metadata inode ILOCK.
Realtime metadata buffer locks, if applicable.
Extent mapping btree blocks, if applicable.
Resources are often released in the reverse order, though this is not required. However, online fsck differs from regular XFS operations because it may examine an object that normally is acquired in a later stage of the locking order, and then decide to cross-reference the object with an object that is acquired earlier in the order. The next few sections detail the specific ways in which online fsck takes care to avoid deadlocks.
4.5.11.2.1. iget and irele During a Scrub¶
An inode scan performed on behalf of a scrub operation runs in transaction
context, and possibly with resources already locked and bound to it.
This isn’t much of a problem for iget since it can operate in the context
of an existing transaction, as long as all of the bound resources are acquired
before the inode reference in the regular filesystem.
When the VFS iput function is given a linked inode with no other
references, it normally puts the inode on an LRU list in the hope that it can
save time if another process re-opens the file before the system runs out
of memory and frees it.
Filesystem callers can short-circuit the LRU process by setting a DONTCACHE
flag on the inode to cause the kernel to try to drop the inode into the
inactivation machinery immediately.
In the past, inactivation was always done from the process that dropped the
inode, which was a problem for scrub because scrub may already hold a
transaction, and XFS does not support nesting transactions.
On the other hand, if there is no scrub transaction, it is desirable to drop
otherwise unused inodes immediately to avoid polluting caches.
To capture these nuances, the online fsck code has a separate xchk_irele
function to set or clear the DONTCACHE flag to get the required release
behavior.
Proposed patchsets include fixing scrub iget usage and dir iget usage.
4.5.11.2.2. Locking Inodes¶
In regular filesystem code, the VFS and XFS will acquire multiple IOLOCK locks
in a well-known order: parent → child when updating the directory tree, and
in numerical order of the addresses of their struct inode object otherwise.
For regular files, the MMAPLOCK can be acquired after the IOLOCK to stop page
faults.
If two MMAPLOCKs must be acquired, they are acquired in numerical order of
the addresses of their struct address_space objects.
Due to the structure of existing filesystem code, IOLOCKs and MMAPLOCKs must be
acquired before transactions are allocated.
If two ILOCKs must be acquired, they are acquired in inumber order.
Inode lock acquisition must be done carefully during a coordinated inode scan.
Online fsck cannot abide these conventions, because for a directory tree
scanner, the scrub process holds the IOLOCK of the file being scanned and it
needs to take the IOLOCK of the file at the other end of the directory link.
If the directory tree is corrupt because it contains a cycle, xfs_scrub
cannot use the regular inode locking functions and avoid becoming trapped in an
ABBA deadlock.
Solving both of these problems is straightforward -- any time online fsck needs to take a second lock of the same class, it uses trylock to avoid an ABBA deadlock. If the trylock fails, scrub drops all inode locks and use trylock loops to (re)acquire all necessary resources. Trylock loops enable scrub to check for pending fatal signals, which is how scrub avoids deadlocking the filesystem or becoming an unresponsive process. However, trylock loops means that online fsck must be prepared to measure the resource being scrubbed before and after the lock cycle to detect changes and react accordingly.
4.5.11.2.3. Case Study: Finding a Directory Parent¶
Consider the directory parent pointer repair code as an example. Online fsck must verify that the dotdot dirent of a directory points up to a parent directory, and that the parent directory contains exactly one dirent pointing down to the child directory. Fully validating this relationship (and repairing it if possible) requires a walk of every directory on the filesystem while holding the child locked, and while updates to the directory tree are being made. The coordinated inode scan provides a way to walk the filesystem without the possibility of missing an inode. The child directory is kept locked to prevent updates to the dotdot dirent, but if the scanner fails to lock a parent, it can drop and relock both the child and the prospective parent. If the dotdot entry changes while the directory is unlocked, then a move or rename operation must have changed the child’s parentage, and the scan can exit early.
The proposed patchset is the directory repair series.
4.5.11.3. Filesystem Hooks¶
The second piece of support that online fsck functions need during a full filesystem scan is the ability to stay informed about updates being made by other threads in the filesystem, since comparisons against the past are useless in a dynamic environment. Two pieces of Linux kernel infrastructure enable online fsck to monitor regular filesystem operations: filesystem hooks and static keys.
Filesystem hooks convey information about an ongoing filesystem operation to a downstream consumer. In this case, the downstream consumer is always an online fsck function. Because multiple fsck functions can run in parallel, online fsck uses the Linux notifier call chain facility to dispatch updates to any number of interested fsck processes. Call chains are a dynamic list, which means that they can be configured at run time. Because these hooks are private to the XFS module, the information passed along contains exactly what the checking function needs to update its observations.
The current implementation of XFS hooks uses SRCU notifier chains to reduce the impact to highly threaded workloads. Regular blocking notifier chains use a rwsem and seem to have a much lower overhead for single-threaded applications. However, it may turn out that the combination of blocking chains and static keys are a more performant combination; more study is needed here.
The following pieces are necessary to hook a certain point in the filesystem:
A
struct xfs_hooksobject must be embedded in a convenient place such as a well-known incore filesystem object.Each hook must define an action code and a structure containing more context about the action.
Hook providers should provide appropriate wrapper functions and structs around the
xfs_hooksandxfs_hookobjects to take advantage of type checking to ensure correct usage.A callsite in the regular filesystem code must be chosen to call
xfs_hooks_callwith the action code and data structure. This place should be adjacent to (and not earlier than) the place where the filesystem update is committed to the transaction. In general, when the filesystem calls a hook chain, it should be able to handle sleeping and should not be vulnerable to memory reclaim or locking recursion. However, the exact requirements are very dependent on the context of the hook caller and the callee.The online fsck function should define a structure to hold scan data, a lock to coordinate access to the scan data, and a
struct xfs_hookobject. The scanner function and the regular filesystem code must acquire resources in the same order; see the next section for details.The online fsck code must contain a C function to catch the hook action code and data structure. If the object being updated has already been visited by the scan, then the hook information must be applied to the scan data.
Prior to unlocking inodes to start the scan, online fsck must call
xfs_hooks_setupto initialize thestruct xfs_hook, andxfs_hooks_addto enable the hook.Online fsck must call
xfs_hooks_delto disable the hook once the scan is complete.
The number of hooks should be kept to a minimum to reduce complexity. Static keys are used to reduce the overhead of filesystem hooks to nearly zero when online fsck is not running.
4.5.11.4. Live Updates During a Scan¶
The code paths of the online fsck scanning code and the hooked filesystem code look like this:
other program
↓
inode lock ←────────────────────┐
↓ │
AG header lock │
↓ │
filesystem function │
↓ │
notifier call chain │ same
↓ ├─── inode
scrub hook function │ lock
↓ │
scan data mutex ←──┐ same │
↓ ├─── scan │
update scan data │ lock │
↑ │ │
scan data mutex ←──┘ │
↑ │
inode lock ←────────────────────┘
↑
scrub function
↑
inode scanner
↑
xfs_scrub
These rules must be followed to ensure correct interactions between the checking code and the code making an update to the filesystem:
Prior to invoking the notifier call chain, the filesystem function being hooked must acquire the same lock that the scrub scanning function acquires to scan the inode.
The scanning function and the scrub hook function must coordinate access to the scan data by acquiring a lock on the scan data.
Scrub hook function must not add the live update information to the scan observations unless the inode being updated has already been scanned. The scan coordinator has a helper predicate (
xchk_iscan_want_live_update) for this.Scrub hook functions must not change the caller’s state, including the transaction that it is running. They must not acquire any resources that might conflict with the filesystem function being hooked.
The hook function can abort the inode scan to avoid breaking the other rules.
The inode scan APIs are pretty simple:
xchk_iscan_startstarts a scanxchk_iscan_itergrabs a reference to the next inode in the scan or returns zero if there is nothing left to scanxchk_iscan_want_live_updateto decide if an inode has already been visited in the scan. This is critical for hook functions to decide if they need to update the in-memory scan information.xchk_iscan_mark_visitedto mark an inode as having been visited in the scanxchk_iscan_teardownto finish the scan
This functionality is also a part of the inode scanner series.
4.5.11.4.1. Case Study: Quota Counter Checking¶
It is useful to compare the mount time quotacheck code to the online repair quotacheck code. Mount time quotacheck does not have to contend with concurrent operations, so it does the following:
Make sure the ondisk dquots are in good enough shape that all the incore dquots will actually load, and zero the resource usage counters in the ondisk buffer.
Walk every inode in the filesystem. Add each file’s resource usage to the incore dquot.
Walk each incore dquot. If the incore dquot is not being flushed, add the ondisk buffer backing the incore dquot to a delayed write (delwri) list.
Write the buffer list to disk.
Like most online fsck functions, online quotacheck can’t write to regular
filesystem objects until the newly collected metadata reflect all filesystem
state.
Therefore, online quotacheck records file resource usage to a shadow dquot
index implemented with a sparse xfarray, and only writes to the real dquots
once the scan is complete.
Handling transactional updates is tricky because quota resource usage updates
are handled in phases to minimize contention on dquots:
The inodes involved are joined and locked to a transaction.
For each dquot attached to the file:
The dquot is locked.
A quota reservation is added to the dquot’s resource usage. The reservation is recorded in the transaction.
The dquot is unlocked.
Changes in actual quota usage are tracked in the transaction.
At transaction commit time, each dquot is examined again:
The dquot is locked again.
Quota usage changes are logged and unused reservation is given back to the dquot.
The dquot is unlocked.
For online quotacheck, hooks are placed in steps 2 and 4.
The step 2 hook creates a shadow version of the transaction dquot context
(dqtrx) that operates in a similar manner to the regular code.
The step 4 hook commits the shadow dqtrx changes to the shadow dquots.
Notice that both hooks are called with the inode locked, which is how the
live update coordinates with the inode scanner.
The quotacheck scan looks like this:
Set up a coordinated inode scan.
For each inode returned by the inode scan iterator:
Grab and lock the inode.
Determine that inode’s resource usage (data blocks, inode counts, realtime blocks) and add that to the shadow dquots for the user, group, and project ids associated with the inode.
Unlock and release the inode.
For each dquot in the system:
Grab and lock the dquot.
Check the dquot against the shadow dquots created by the scan and updated by the live hooks.
Live updates are key to being able to walk every quota record without needing to hold any locks for a long duration. If repairs are desired, the real and shadow dquots are locked and their resource counts are set to the values in the shadow dquot.
The proposed patchset is the online quotacheck series.
4.5.11.4.2. Case Study: File Link Count Checking¶
File link count checking also uses live update hooks.
The coordinated inode scanner is used to visit all directories on the
filesystem, and per-file link count records are stored in a sparse xfarray
indexed by inumber.
During the scanning phase, each entry in a directory generates observation
data as follows:
If the entry is a dotdot (
'..') entry of the root directory, the directory’s parent link count is bumped because the root directory’s dotdot entry is self referential.If the entry is a dotdot entry of a subdirectory, the parent’s backref count is bumped.
If the entry is neither a dot nor a dotdot entry, the target file’s parent count is bumped.
If the target is a subdirectory, the parent’s child link count is bumped.
A crucial point to understand about how the link count inode scanner interacts
with the live update hooks is that the scan cursor tracks which parent
directories have been scanned.
In other words, the live updates ignore any update about A → B when A has
not been scanned, even if B has been scanned.
Furthermore, a subdirectory A with a dotdot entry pointing back to B is
accounted as a backref counter in the shadow data for A, since child dotdot
entries affect the parent’s link count.
Live update hooks are carefully placed in all parts of the filesystem that
create, change, or remove directory entries, since those operations involve
bumplink and droplink.
For any file, the correct link count is the number of parents plus the number of child subdirectories. Non-directories never have children of any kind. The backref information is used to detect inconsistencies in the number of links pointing to child subdirectories and the number of dotdot entries pointing back.
After the scan completes, the link count of each file can be checked by locking both the inode and the shadow data, and comparing the link counts. A second coordinated inode scan cursor is used for comparisons. Live updates are key to being able to walk every inode without needing to hold any locks between inodes. If repairs are desired, the inode’s link count is set to the value in the shadow information. If no parents are found, the file must be reparented to the orphanage to prevent the file from being lost forever.
The proposed patchset is the file link count repair series.
4.5.11.4.3. Case Study: Rebuilding Reverse Mapping Records¶
Most repair functions follow the same pattern: lock filesystem resources, walk the surviving ondisk metadata looking for replacement metadata records, and use an in-memory array to store the gathered observations. The primary advantage of this approach is the simplicity and modularity of the repair code -- code and data are entirely contained within the scrub module, do not require hooks in the main filesystem, and are usually the most efficient in memory use. A secondary advantage of this repair approach is atomicity -- once the kernel decides a structure is corrupt, no other threads can access the metadata until the kernel finishes repairing and revalidating the metadata.
For repairs going on within a shard of the filesystem, these advantages outweigh the delays inherent in locking the shard while repairing parts of the shard. Unfortunately, repairs to the reverse mapping btree cannot use the “standard” btree repair strategy because it must scan every space mapping of every fork of every file in the filesystem, and the filesystem cannot stop. Therefore, rmap repair foregoes atomicity between scrub and repair. It combines a coordinated inode scanner, live update hooks, and an in-memory rmap btree to complete the scan for reverse mapping records.
Set up an xfbtree to stage rmap records.
While holding the locks on the AGI and AGF buffers acquired during the scrub, generate reverse mappings for all AG metadata: inodes, btrees, CoW staging extents, and the internal log.
Set up an inode scanner.
Hook into rmap updates for the AG being repaired so that the live scan data can receive updates to the rmap btree from the rest of the filesystem during the file scan.
For each space mapping found in either fork of each file scanned, decide if the mapping matches the AG of interest. If so:
Create a btree cursor for the in-memory btree.
Use the rmap code to add the record to the in-memory btree.
Use the special commit function to write the xfbtree changes to the xfile.
For each live update received via the hook, decide if the owner has already been scanned. If so, apply the live update into the scan data:
Create a btree cursor for the in-memory btree.
Replay the operation into the in-memory btree.
Use the special commit function to write the xfbtree changes to the xfile. This is performed with an empty transaction to avoid changing the caller’s state.
When the inode scan finishes, create a new scrub transaction and relock the two AG headers.
Compute the new btree geometry using the number of rmap records in the shadow btree, like all other btree rebuilding functions.
Allocate the number of blocks computed in the previous step.
Perform the usual btree bulk loading and commit to install the new rmap btree.
Reap the old rmap btree blocks as discussed in the case study about how to reap after rmap btree repair.
Free the xfbtree now that it not needed.
The proposed patchset is the rmap repair series.
4.5.12. Staging Repairs with Temporary Files on Disk¶
XFS stores a substantial amount of metadata in file forks: directories, extended attributes, symbolic link targets, free space bitmaps and summary information for the realtime volume, and quota records. File forks map 64-bit logical file fork space extents to physical storage space extents, similar to how a memory management unit maps 64-bit virtual addresses to physical memory addresses. Therefore, file-based tree structures (such as directories and extended attributes) use blocks mapped in the file fork offset address space that point to other blocks mapped within that same address space, and file-based linear structures (such as bitmaps and quota records) compute array element offsets in the file fork offset address space.
Because file forks can consume as much space as the entire filesystem, repairs cannot be staged in memory, even when a paging scheme is available. Therefore, online repair of file-based metadata createas a temporary file in the XFS filesystem, writes a new structure at the correct offsets into the temporary file, and atomically exchanges all file fork mappings (and hence the fork contents) to commit the repair. Once the repair is complete, the old fork can be reaped as necessary; if the system goes down during the reap, the iunlink code will delete the blocks during log recovery.
Note: All space usage and inode indices in the filesystem must be consistent to use a temporary file safely! This dependency is the reason why online repair can only use pageable kernel memory to stage ondisk space usage information.
Exchanging metadata file mappings with a temporary file requires the owner field of the block headers to match the file being repaired and not the temporary file. The directory, extended attribute, and symbolic link functions were all modified to allow callers to specify owner numbers explicitly.
There is a downside to the reaping process -- if the system crashes during the reap phase and the fork extents are crosslinked, the iunlink processing will fail because freeing space will find the extra reverse mappings and abort.
Temporary files created for repair are similar to O_TMPFILE files created
by userspace.
They are not linked into a directory and the entire file will be reaped when
the last reference to the file is lost.
The key differences are that these files must have no access permission outside
the kernel at all, they must be specially marked to prevent them from being
opened by handle, and they must never be linked into the directory tree.
Historical Sidebar: |
In the initial iteration of file metadata repair, the damaged metadata blocks would be scanned for salvageable data; the extents in the file fork would be reaped; and then a new structure would be built in its place. This strategy did not survive the introduction of the atomic repair requirement expressed earlier in this document. The second iteration explored building a second structure at a high
offset in the fork from the salvage data, reaping the old extents, and
using a This had many drawbacks:
This lead to the introduction of temporary file staging. |
4.5.12.1. Using a Temporary File¶
Online repair code should use the xrep_tempfile_create function to create a
temporary file inside the filesystem.
This allocates an inode, marks the in-core inode private, and attaches it to
the scrub context.
These files are hidden from userspace, may not be added to the directory tree,
and must be kept private.
Temporary files only use two inode locks: the IOLOCK and the ILOCK. The MMAPLOCK is not needed here, because there must not be page faults from userspace for data fork blocks. The usage patterns of these two locks are the same as for any other XFS file -- access to file data are controlled via the IOLOCK, and access to file metadata are controlled via the ILOCK. Locking helpers are provided so that the temporary file and its lock state can be cleaned up by the scrub context. To comply with the nested locking strategy laid out in the inode locking section, it is recommended that scrub functions use the xrep_tempfile_ilock*_nowait lock helpers.
Data can be written to a temporary file by two means:
xrep_tempfile_copyincan be used to set the contents of a regular temporary file from an xfile.The regular directory, symbolic link, and extended attribute functions can be used to write to the temporary file.
Once a good copy of a data file has been constructed in a temporary file, it must be conveyed to the file being repaired, which is the topic of the next section.
The proposed patches are in the repair temporary files series.
4.5.13. Logged File Content Exchanges¶
Once repair builds a temporary file with a new data structure written into
it, it must commit the new changes into the existing file.
It is not possible to swap the inumbers of two files, so instead the new
metadata must replace the old.
This suggests the need for the ability to swap extents, but the existing extent
swapping code used by the file defragmenting tool xfs_fsr is not sufficient
for online repair because:
When the reverse-mapping btree is enabled, the swap code must keep the reverse mapping information up to date with every exchange of mappings. Therefore, it can only exchange one mapping per transaction, and each transaction is independent.
Reverse-mapping is critical for the operation of online fsck, so the old defragmentation code (which swapped entire extent forks in a single operation) is not useful here.
Defragmentation is assumed to occur between two files with identical contents. For this use case, an incomplete exchange will not result in a user-visible change in file contents, even if the operation is interrupted.
Online repair needs to swap the contents of two files that are by definition not identical. For directory and xattr repairs, the user-visible contents might be the same, but the contents of individual blocks may be very different.
Old blocks in the file may be cross-linked with another structure and must not reappear if the system goes down mid-repair.
These problems are overcome by creating a new deferred operation and a new type
of log intent item to track the progress of an operation to exchange two file
ranges.
The new exchange operation type chains together the same transactions used by
the reverse-mapping extent swap code, but records intermedia progress in the
log so that operations can be restarted after a crash.
This new functionality is called the file contents exchange (xfs_exchrange)
code.
The underlying implementation exchanges file fork mappings (xfs_exchmaps).
The new log item records the progress of the exchange to ensure that once an
exchange begins, it will always run to completion, even there are
interruptions.
The new XFS_SB_FEAT_INCOMPAT_EXCHRANGE incompatible feature flag
in the superblock protects these new log item records from being replayed on
old kernels.
The proposed patchset is the file contents exchange series.
Sidebar: Using Log-Incompatible Feature Flags |
Starting with XFS v5, the superblock contains a
The log coordinates access to incompatible features through the use of
one Log-assisted extended attribute updates and file content exchanges bothe use log incompat features and provide convenience wrappers around the functionality. |
4.5.13.1. Mechanics of a Logged File Content Exchange¶
Exchanging contents between file forks is a complex task. The goal is to exchange all file fork mappings between two file fork offset ranges. There are likely to be many extent mappings in each fork, and the edges of the mappings aren’t necessarily aligned. Furthermore, there may be other updates that need to happen after the exchange, such as exchanging file sizes, inode flags, or conversion of fork data to local format. This is roughly the format of the new deferred exchange-mapping work item:
struct xfs_exchmaps_intent {
/* Inodes participating in the operation. */
struct xfs_inode *xmi_ip1;
struct xfs_inode *xmi_ip2;
/* File offset range information. */
xfs_fileoff_t xmi_startoff1;
xfs_fileoff_t xmi_startoff2;
xfs_filblks_t xmi_blockcount;
/* Set these file sizes after the operation, unless negative. */
xfs_fsize_t xmi_isize1;
xfs_fsize_t xmi_isize2;
/* XFS_EXCHMAPS_* log operation flags */
uint64_t xmi_flags;
};
The new log intent item contains enough information to track two logical fork
offset ranges: (inode1, startoff1, blockcount) and (inode2, startoff2,
blockcount).
Each step of an exchange operation exchanges the largest file range mapping
possible from one file to the other.
After each step in the exchange operation, the two startoff fields are
incremented and the blockcount field is decremented to reflect the progress
made.
The flags field captures behavioral parameters such as exchanging attr fork
mappings instead of the data fork and other work to be done after the exchange.
The two isize fields are used to exchange the file sizes at the end of the
operation if the file data fork is the target of the operation.
When the exchange is initiated, the sequence of operations is as follows:
Create a deferred work item for the file mapping exchange. At the start, it should contain the entirety of the file block ranges to be exchanged.
Call
xfs_defer_finishto process the exchange. This is encapsulated inxrep_tempexch_contentsfor scrub operations. This will log an extent swap intent item to the transaction for the deferred mapping exchange work item.Until
xmi_blockcountof the deferred mapping exchange work item is zero,Read the block maps of both file ranges starting at
xmi_startoff1andxmi_startoff2, respectively, and compute the longest extent that can be exchanged in a single step. This is the minimum of the twobr_blockcounts in the mappings. Keep advancing through the file forks until at least one of the mappings contains written blocks. Mutual holes, unwritten extents, and extent mappings to the same physical space are not exchanged.For the next few steps, this document will refer to the mapping that came from file 1 as “map1”, and the mapping that came from file 2 as “map2”.
Create a deferred block mapping update to unmap map1 from file 1.
Create a deferred block mapping update to unmap map2 from file 2.
Create a deferred block mapping update to map map1 into file 2.
Create a deferred block mapping update to map map2 into file 1.
Log the block, quota, and extent count updates for both files.
Extend the ondisk size of either file if necessary.
Log a mapping exchange done log item for th mapping exchange intent log item that was read at the start of step 3.
Compute the amount of file range that has just been covered. This quantity is
(map1.br_startoff + map1.br_blockcount - xmi_startoff1), because step 3a could have skipped holes.Increase the starting offsets of
xmi_startoff1andxmi_startoff2by the number of blocks computed in the previous step, and decreasexmi_blockcountby the same quantity. This advances the cursor.Log a new mapping exchange intent log item reflecting the advanced state of the work item.
Return the proper error code (EAGAIN) to the deferred operation manager to inform it that there is more work to be done. The operation manager completes the deferred work in steps 3b-3e before moving back to the start of step 3.
Perform any post-processing. This will be discussed in more detail in subsequent sections.
If the filesystem goes down in the middle of an operation, log recovery will find the most recent unfinished maping exchange log intent item and restart from there. This is how atomic file mapping exchanges guarantees that an outside observer will either see the old broken structure or the new one, and never a mismash of both.
4.5.13.2. Preparation for File Content Exchanges¶
There are a few things that need to be taken care of before initiating an atomic file mapping exchange operation. First, regular files require the page cache to be flushed to disk before the operation begins, and directio writes to be quiesced. Like any filesystem operation, file mapping exchanges must determine the maximum amount of disk space and quota that can be consumed on behalf of both files in the operation, and reserve that quantity of resources to avoid an unrecoverable out of space failure once it starts dirtying metadata. The preparation step scans the ranges of both files to estimate:
Data device blocks needed to handle the repeated updates to the fork mappings.
Change in data and realtime block counts for both files.
Increase in quota usage for both files, if the two files do not share the same set of quota ids.
The number of extent mappings that will be added to each file.
Whether or not there are partially written realtime extents. User programs must never be able to access a realtime file extent that maps to different extents on the realtime volume, which could happen if the operation fails to run to completion.
The need for precise estimation increases the run time of the exchange operation, but it is very important to maintain correct accounting. The filesystem must not run completely out of free space, nor can the mapping exchange ever add more extent mappings to a fork than it can support. Regular users are required to abide the quota limits, though metadata repairs may exceed quota to resolve inconsistent metadata elsewhere.
4.5.13.3. Special Features for Exchanging Metadata File Contents¶
Extended attributes, symbolic links, and directories can set the fork format to “local” and treat the fork as a literal area for data storage. Metadata repairs must take extra steps to support these cases:
If both forks are in local format and the fork areas are large enough, the exchange is performed by copying the incore fork contents, logging both forks, and committing. The atomic file mapping exchange mechanism is not necessary, since this can be done with a single transaction.
If both forks map blocks, then the regular atomic file mapping exchange is used.
Otherwise, only one fork is in local format. The contents of the local format fork are converted to a block to perform the exchange. The conversion to block format must be done in the same transaction that logs the initial mapping exchange intent log item. The regular atomic mapping exchange is used to exchange the metadata file mappings. Special flags are set on the exchange operation so that the transaction can be rolled one more time to convert the second file’s fork back to local format so that the second file will be ready to go as soon as the ILOCK is dropped.
Extended attributes and directories stamp the owning inode into every block, but the buffer verifiers do not actually check the inode number! Although there is no verification, it is still important to maintain referential integrity, so prior to performing the mapping exchange, online repair builds every block in the new data structure with the owner field of the file being repaired.
After a successful exchange operation, the repair operation must reap the old fork blocks by processing each fork mapping through the standard file extent reaping mechanism that is done post-repair. If the filesystem should go down during the reap part of the repair, the iunlink processing at the end of recovery will free both the temporary file and whatever blocks were not reaped. However, this iunlink processing omits the cross-link detection of online repair, and is not completely foolproof.
4.5.13.4. Exchanging Temporary File Contents¶
To repair a metadata file, online repair proceeds as follows:
Create a temporary repair file.
Use the staging data to write out new contents into the temporary repair file. The same fork must be written to as is being repaired.
Commit the scrub transaction, since the exchange resource estimation step must be completed before transaction reservations are made.
Call
xrep_tempexch_trans_allocto allocate a new scrub transaction with the appropriate resource reservations, locks, and fill out astruct xfs_exchmaps_reqwith the details of the exchange operation.Call
xrep_tempexch_contentsto exchange the contents.Commit the transaction to complete the repair.
4.5.13.4.1. Case Study: Repairing the Realtime Summary File¶
In the “realtime” section of an XFS filesystem, free space is tracked via a bitmap, similar to Unix FFS. Each bit in the bitmap represents one realtime extent, which is a multiple of the filesystem block size between 4KiB and 1GiB in size. The realtime summary file indexes the number of free extents of a given size to the offset of the block within the realtime free space bitmap where those free extents begin. In other words, the summary file helps the allocator find free extents by length, similar to what the free space by count (cntbt) btree does for the data section.
The summary file itself is a flat file (with no block headers or checksums!)
partitioned into log2(total rt extents) sections containing enough 32-bit
counters to match the number of blocks in the rt bitmap.
Each counter records the number of free extents that start in that bitmap block
and can satisfy a power-of-two allocation request.
To check the summary file against the bitmap:
Take the ILOCK of both the realtime bitmap and summary files.
For each free space extent recorded in the bitmap:
Compute the position in the summary file that contains a counter that represents this free extent.
Read the counter from the xfile.
Increment it, and write it back to the xfile.
Compare the contents of the xfile against the ondisk file.
To repair the summary file, write the xfile contents into the temporary file and use atomic mapping exchange to commit the new contents. The temporary file is then reaped.
The proposed patchset is the realtime summary repair series.
4.5.13.4.2. Case Study: Salvaging Extended Attributes¶
In XFS, extended attributes are implemented as a namespaced name-value store.
Values are limited in size to 64KiB, but there is no limit in the number of
names.
The attribute fork is unpartitioned, which means that the root of the attribute
structure is always in logical block zero, but attribute leaf blocks, dabtree
index blocks, and remote value blocks are intermixed.
Attribute leaf blocks contain variable-sized records that associate
user-provided names with the user-provided values.
Values larger than a block are allocated separate extents and written there.
If the leaf information expands beyond a single block, a directory/attribute
btree (dabtree) is created to map hashes of attribute names to entries
for fast lookup.
Salvaging extended attributes is done as follows:
Walk the attr fork mappings of the file being repaired to find the attribute leaf blocks. When one is found,
Walk the attr leaf block to find candidate keys. When one is found,
Check the name for problems, and ignore the name if there are.
Retrieve the value. If that succeeds, add the name and value to the staging xfarray and xfblob.
If the memory usage of the xfarray and xfblob exceed a certain amount of memory or there are no more attr fork blocks to examine, unlock the file and add the staged extended attributes to the temporary file.
Use atomic file mapping exchange to exchange the new and old extended attribute structures. The old attribute blocks are now attached to the temporary file.
Reap the temporary file.
The proposed patchset is the extended attribute repair series.
4.5.14. Fixing Directories¶
Fixing directories is difficult with currently available filesystem features,
since directory entries are not redundant.
The offline repair tool scans all inodes to find files with nonzero link count,
and then it scans all directories to establish parentage of those linked files.
Damaged files and directories are zapped, and files with no parent are
moved to the /lost+found directory.
It does not try to salvage anything.
The best that online repair can do at this time is to read directory data
blocks and salvage any dirents that look plausible, correct link counts, and
move orphans back into the directory tree.
The salvage process is discussed in the case study at the end of this section.
The file link count fsck code takes care of fixing link counts
and moving orphans to the /lost+found directory.
4.5.14.1. Case Study: Salvaging Directories¶
Unlike extended attributes, directory blocks are all the same size, so salvaging directories is straightforward:
Find the parent of the directory. If the dotdot entry is not unreadable, try to confirm that the alleged parent has a child entry pointing back to the directory being repaired. Otherwise, walk the filesystem to find it.
Walk the first partition of data fork of the directory to find the directory entry data blocks. When one is found,
Walk the directory data block to find candidate entries. When an entry is found:
Check the name for problems, and ignore the name if there are.
Retrieve the inumber and grab the inode. If that succeeds, add the name, inode number, and file type to the staging xfarray and xblob.
If the memory usage of the xfarray and xfblob exceed a certain amount of memory or there are no more directory data blocks to examine, unlock the directory and add the staged dirents into the temporary directory. Truncate the staging files.
Use atomic file mapping exchange to exchange the new and old directory structures. The old directory blocks are now attached to the temporary file.
Reap the temporary file.
Future Work Question: Should repair revalidate the dentry cache when rebuilding a directory?
Answer: Yes, it should.
In theory it is necessary to scan all dentry cache entries for a directory to ensure that one of the following apply:
The cached dentry reflects an ondisk dirent in the new directory.
The cached dentry no longer has a corresponding ondisk dirent in the new directory and the dentry can be purged from the cache.
The cached dentry no longer has an ondisk dirent but the dentry cannot be purged. This is the problem case.
Unfortunately, the current dentry cache design doesn’t provide a means to walk every child dentry of a specific directory, which makes this a hard problem. There is no known solution.
The proposed patchset is the directory repair series.
4.5.14.2. Parent Pointers¶
A parent pointer is a piece of file metadata that enables a user to locate the file’s parent directory without having to traverse the directory tree from the root. Without them, reconstruction of directory trees is hindered in much the same way that the historic lack of reverse space mapping information once hindered reconstruction of filesystem space metadata. The parent pointer feature, however, makes total directory reconstruction possible.
XFS parent pointers contain the information needed to identify the
corresponding directory entry in the parent directory.
In other words, child files use extended attributes to store pointers to
parents in the form (dirent_name) → (parent_inum, parent_gen).
The directory checking process can be strengthened to ensure that the target of
each dirent also contains a parent pointer pointing back to the dirent.
Likewise, each parent pointer can be checked by ensuring that the target of
each parent pointer is a directory and that it contains a dirent matching
the parent pointer.
Both online and offline repair can use this strategy.
Historical Sidebar: |
Directory parent pointers were first proposed as an XFS feature more than a decade ago by SGI. Each link from a parent directory to a child file is mirrored with an extended attribute in the child that could be used to identify the parent directory. Unfortunately, this early implementation had major shortcomings and was never merged into Linux XFS:
The original parent pointer design was too unstable for something like a file system repair to depend on. Allison Henderson, Chandan Babu, and Catherine Hoang are working on a second implementation that solves all shortcomings of the first. During 2022, Allison introduced log intent items to track physical manipulations of the extended attribute structures. This solves the referential integrity problem by making it possible to commit a dirent update and a parent pointer update in the same transaction. Chandan increased the maximum extent counts of both data and attribute forks, thereby ensuring that the extended attribute structure can grow to handle the maximum hardlink count of any file. For this second effort, the ondisk parent pointer format as originally
proposed was There were a few other ways to have solved that problem:
In the end, it was decided that solution #6 was the most compact and the most performant. A new hash function was designed for parent pointers. |
4.5.14.2.1. Case Study: Repairing Directories with Parent Pointers¶
Directory rebuilding uses a coordinated inode scan and a directory entry live update hook as follows:
Set up a temporary directory for generating the new directory structure, an xfblob for storing entry names, and an xfarray for stashing the fixed size fields involved in a directory update:
(child inumber, add vs. remove, name cookie, ftype).Set up an inode scanner and hook into the directory entry code to receive updates on directory operations.
For each parent pointer found in each file scanned, decide if the parent pointer references the directory of interest. If so:
Stash the parent pointer name and an addname entry for this dirent in the xfblob and xfarray, respectively.
When finished scanning that file or the kernel memory consumption exceeds a threshold, flush the stashed updates to the temporary directory.
For each live directory update received via the hook, decide if the child has already been scanned. If so:
Stash the parent pointer name an addname or removename entry for this dirent update in the xfblob and xfarray for later. We cannot write directly to the temporary directory because hook functions are not allowed to modify filesystem metadata. Instead, we stash updates in the xfarray and rely on the scanner thread to apply the stashed updates to the temporary directory.
When the scan is complete, replay any stashed entries in the xfarray.
When the scan is complete, atomically exchange the contents of the temporary directory and the directory being repaired. The temporary directory now contains the damaged directory structure.
Reap the temporary directory.
The proposed patchset is the parent pointers directory repair series.
4.5.14.2.2. Case Study: Repairing Parent Pointers¶
Online reconstruction of a file’s parent pointer information works similarly to directory reconstruction:
Set up a temporary file for generating a new extended attribute structure, an xfblob for storing parent pointer names, and an xfarray for stashing the fixed size fields involved in a parent pointer update:
(parent inumber, parent generation, add vs. remove, name cookie).Set up an inode scanner and hook into the directory entry code to receive updates on directory operations.
For each directory entry found in each directory scanned, decide if the dirent references the file of interest. If so:
Stash the dirent name and an addpptr entry for this parent pointer in the xfblob and xfarray, respectively.
When finished scanning the directory or the kernel memory consumption exceeds a threshold, flush the stashed updates to the temporary file.
For each live directory update received via the hook, decide if the parent has already been scanned. If so:
Stash the dirent name and an addpptr or removepptr entry for this dirent update in the xfblob and xfarray for later. We cannot write parent pointers directly to the temporary file because hook functions are not allowed to modify filesystem metadata. Instead, we stash updates in the xfarray and rely on the scanner thread to apply the stashed parent pointer updates to the temporary file.
When the scan is complete, replay any stashed entries in the xfarray.
Copy all non-parent pointer extended attributes to the temporary file.
When the scan is complete, atomically exchange the mappings of the attribute forks of the temporary file and the file being repaired. The temporary file now contains the damaged extended attribute structure.
Reap the temporary file.
The proposed patchset is the parent pointers repair series.
4.5.14.2.3. Digression: Offline Checking of Parent Pointers¶
Examining parent pointers in offline repair works differently because corrupt files are erased long before directory tree connectivity checks are performed. Parent pointer checks are therefore a second pass to be added to the existing connectivity checks:
After the set of surviving files has been established (phase 6), walk the surviving directories of each AG in the filesystem. This is already performed as part of the connectivity checks.
For each directory entry found,
If the name has already been stored in the xfblob, then use that cookie and skip the next step.
Otherwise, record the name in an xfblob, and remember the xfblob cookie. Unique mappings are critical for
Deduplicating names to reduce memory usage, and
Creating a stable sort key for the parent pointer indexes so that the parent pointer validation described below will work.
Store
(child_ag_inum, parent_inum, parent_gen, name_hash, name_len, name_cookie)tuples in a per-AG in-memory slab. Thename_hashreferenced in this section is the regular directory entry name hash, not the specialized one used for parent pointer xattrs.
For each AG in the filesystem,
Sort the per-AG tuple set in order of
child_ag_inum,parent_inum,name_hash, andname_cookie. Having a singlename_cookiefor eachnameis critical for handling the uncommon case of a directory containing multiple hardlinks to the same file where all the names hash to the same value.For each inode in the AG,
Scan the inode for parent pointers. For each parent pointer found,
Validate the ondisk parent pointer. If validation fails, move on to the next parent pointer in the file.
If the name has already been stored in the xfblob, then use that cookie and skip the next step.
Record the name in a per-file xfblob, and remember the xfblob cookie.
Store
(parent_inum, parent_gen, name_hash, name_len, name_cookie)tuples in a per-file slab.
Sort the per-file tuples in order of
parent_inum,name_hash, andname_cookie.Position one slab cursor at the start of the inode’s records in the per-AG tuple slab. This should be trivial since the per-AG tuples are in child inumber order.
Position a second slab cursor at the start of the per-file tuple slab.
Iterate the two cursors in lockstep, comparing the
parent_ino,name_hash, andname_cookiefields of the records under each cursor:If the per-AG cursor is at a lower point in the keyspace than the per-file cursor, then the per-AG cursor points to a missing parent pointer. Add the parent pointer to the inode and advance the per-AG cursor.
If the per-file cursor is at a lower point in the keyspace than the per-AG cursor, then the per-file cursor points to a dangling parent pointer. Remove the parent pointer from the inode and advance the per-file cursor.
Otherwise, both cursors point at the same parent pointer. Update the parent_gen component if necessary. Advance both cursors.
Move on to examining link counts, as we do today.
The proposed patchset is the offline parent pointers repair series.
Rebuilding directories from parent pointers in offline repair would be very challenging because xfs_repair currently uses two single-pass scans of the filesystem during phases 3 and 4 to decide which files are corrupt enough to be zapped. This scan would have to be converted into a multi-pass scan:
The first pass of the scan zaps corrupt inodes, forks, and attributes much as it does now. Corrupt directories are noted but not zapped.
The next pass records parent pointers pointing to the directories noted as being corrupt in the first pass. This second pass may have to happen after the phase 4 scan for duplicate blocks, if phase 4 is also capable of zapping directories.
The third pass resets corrupt directories to an empty shortform directory. Free space metadata has not been ensured yet, so repair cannot yet use the directory building code in libxfs.
At the start of phase 6, space metadata have been rebuilt. Use the parent pointer information recorded during step 2 to reconstruct the dirents and add them to the now-empty directories.
This code has not yet been constructed.
4.5.14.2.4. Case Study: Directory Tree Structure¶
As mentioned earlier, the filesystem directory tree is supposed to be a
directed acylic graph structure.
However, each node in this graph is a separate xfs_inode object with its
own locks, which makes validating the tree qualities difficult.
Fortunately, non-directories are allowed to have multiple parents and cannot
have children, so only directories need to be scanned.
Directories typically constitute 5-10% of the files in a filesystem, which
reduces the amount of work dramatically.
If the directory tree could be frozen, it would be easy to discover cycles and disconnected regions by running a depth (or breadth) first search downwards from the root directory and marking a bitmap for each directory found. At any point in the walk, trying to set an already set bit means there is a cycle. After the scan completes, XORing the marked inode bitmap with the inode allocation bitmap reveals disconnected inodes. However, one of online repair’s design goals is to avoid locking the entire filesystem unless it’s absolutely necessary. Directory tree updates can move subtrees across the scanner wavefront on a live filesystem, so the bitmap algorithm cannot be applied.
Directory parent pointers enable an incremental approach to validation of the tree structure. Instead of using one thread to scan the entire filesystem, multiple threads can walk from individual subdirectories upwards towards the root. For this to work, all directory entries and parent pointers must be internally consistent, each directory entry must have a parent pointer, and the link counts of all directories must be correct. Each scanner thread must be able to take the IOLOCK of an alleged parent directory while holding the IOLOCK of the child directory to prevent either directory from being moved within the tree. This is not possible since the VFS does not take the IOLOCK of a child subdirectory when moving that subdirectory, so instead the scanner stabilizes the parent -> child relationship by taking the ILOCKs and installing a dirent update hook to detect changes.
The scanning process uses a dirent hook to detect changes to the directories mentioned in the scan data. The scan works as follows:
For each subdirectory in the filesystem,
For each parent pointer of that subdirectory,
Create a path object for that parent pointer, and mark the subdirectory inode number in the path object’s bitmap.
Record the parent pointer name and inode number in a path structure.
If the alleged parent is the subdirectory being scrubbed, the path is a cycle. Mark the path for deletion and repeat step 1a with the next subdirectory parent pointer.
Try to mark the alleged parent inode number in a bitmap in the path object. If the bit is already set, then there is a cycle in the directory tree. Mark the path as a cycle and repeat step 1a with the next subdirectory parent pointer.
Load the alleged parent. If the alleged parent is not a linked directory, abort the scan because the parent pointer information is inconsistent.
For each parent pointer of this alleged ancestor directory,
Record the parent pointer name and inode number in the path object if no parent has been set for that level.
If an ancestor has more than one parent, mark the path as corrupt. Repeat step 1a with the next subdirectory parent pointer.
Repeat steps 1a3-1a6 for the ancestor identified in step 1a6a. This repeats until the directory tree root is reached or no parents are found.
If the walk terminates at the root directory, mark the path as ok.
If the walk terminates without reaching the root, mark the path as disconnected.
If the directory entry update hook triggers, check all paths already found by the scan. If the entry matches part of a path, mark that path and the scan stale. When the scanner thread sees that the scan has been marked stale, it deletes all scan data and starts over.
Repairing the directory tree works as follows:
Walk each path of the target subdirectory.
Corrupt paths and cycle paths are counted as suspect.
Paths already marked for deletion are counted as bad.
Paths that reached the root are counted as good.
If the subdirectory is either the root directory or has zero link count, delete all incoming directory entries in the immediate parents. Repairs are complete.
If the subdirectory has exactly one path, set the dotdot entry to the parent and exit.
If the subdirectory has at least one good path, delete all the other incoming directory entries in the immediate parents.
If the subdirectory has no good paths and more than one suspect path, delete all the other incoming directory entries in the immediate parents.
If the subdirectory has zero paths, attach it to the lost and found.
The proposed patches are in the directory tree repair series.
4.5.15. The Orphanage¶
Filesystems present files as a directed, and hopefully acyclic, graph. In other words, a tree. The root of the filesystem is a directory, and each entry in a directory points downwards either to more subdirectories or to non-directory files. Unfortunately, a disruption in the directory graph pointers result in a disconnected graph, which makes files impossible to access via regular path resolution.
Without parent pointers, the directory parent pointer online scrub code can detect a dotdot entry pointing to a parent directory that doesn’t have a link back to the child directory and the file link count checker can detect a file that isn’t pointed to by any directory in the filesystem. If such a file has a positive link count, the file is an orphan.
With parent pointers, directories can be rebuilt by scanning parent pointers
and parent pointers can be rebuilt by scanning directories.
This should reduce the incidence of files ending up in /lost+found.
When orphans are found, they should be reconnected to the directory tree.
Offline fsck solves the problem by creating a directory /lost+found to
serve as an orphanage, and linking orphan files into the orphanage by using the
inumber as the name.
Reparenting a file to the orphanage does not reset any of its permissions or
ACLs.
This process is more involved in the kernel than it is in userspace. The directory and file link count repair setup functions must use the regular VFS mechanisms to create the orphanage directory with all the necessary security attributes and dentry cache entries, just like a regular directory tree modification.
Orphaned files are adopted by the orphanage as follows:
Call
xrep_orphanage_try_createat the start of the scrub setup function to try to ensure that the lost and found directory actually exists. This also attaches the orphanage directory to the scrub context.If the decision is made to reconnect a file, take the IOLOCK of both the orphanage and the file being reattached. The
xrep_orphanage_iolock_twofunction follows the inode locking strategy discussed earlier.Use
xrep_adoption_trans_allocto reserve resources to the repair transaction.Call
xrep_orphanage_compute_nameto compute the new name in the orphanage.If the adoption is going to happen, call
xrep_adoption_reparentto reparent the orphaned file into the lost and found and invalidate the dentry cache.Call
xrep_adoption_finishto commit any filesystem updates, release the orphanage ILOCK, and clean the scrub transaction. Callxrep_adoption_committo commit the updates and the scrub transaction.If a runtime error happens, call
xrep_adoption_cancelto release all resources.
The proposed patches are in the orphanage adoption series.
4.6. 6. Userspace Algorithms and Data Structures¶
This section discusses the key algorithms and data structures of the userspace
program, xfs_scrub, that provide the ability to drive metadata checks and
repairs in the kernel, verify file data, and look for other potential problems.
4.6.1. Checking Metadata¶
Recall the phases of fsck work outlined earlier. That structure follows naturally from the data dependencies designed into the filesystem from its beginnings in 1993. In XFS, there are several groups of metadata dependencies:
Filesystem summary counts depend on consistency within the inode indices, the allocation group space btrees, and the realtime volume space information.
Quota resource counts depend on consistency within the quota file data forks, inode indices, inode records, and the forks of every file on the system.
The naming hierarchy depends on consistency within the directory and extended attribute structures. This includes file link counts.
Directories, extended attributes, and file data depend on consistency within the file forks that map directory and extended attribute data to physical storage media.
The file forks depends on consistency within inode records and the space metadata indices of the allocation groups and the realtime volume. This includes quota and realtime metadata files.
Inode records depends on consistency within the inode metadata indices.
Realtime space metadata depend on the inode records and data forks of the realtime metadata inodes.
The allocation group metadata indices (free space, inodes, reference count, and reverse mapping btrees) depend on consistency within the AG headers and between all the AG metadata btrees.
xfs_scrubdepends on the filesystem being mounted and kernel support for online fsck functionality.
Therefore, a metadata dependency graph is a convenient way to schedule checking
operations in the xfs_scrub program:
Phase 1 checks that the provided path maps to an XFS filesystem and detect the kernel’s scrubbing abilities, which validates group (i).
Phase 2 scrubs groups (g) and (h) in parallel using a threaded workqueue.
Phase 3 scans inodes in parallel. For each inode, groups (f), (e), and (d) are checked, in that order.
Phase 4 repairs everything in groups (i) through (d) so that phases 5 and 6 may run reliably.
Phase 5 starts by checking groups (b) and (c) in parallel before moving on to checking names.
Phase 6 depends on groups (i) through (b) to find file data blocks to verify, to read them, and to report which blocks of which files are affected.
Phase 7 checks group (a), having validated everything else.
Notice that the data dependencies between groups are enforced by the structure of the program flow.
4.6.2. Parallel Inode Scans¶
An XFS filesystem can easily contain hundreds of millions of inodes. Given that XFS targets installations with large high-performance storage, it is desirable to scrub inodes in parallel to minimize runtime, particularly if the program has been invoked manually from a command line. This requires careful scheduling to keep the threads as evenly loaded as possible.
Early iterations of the xfs_scrub inode scanner naïvely created a single
workqueue and scheduled a single workqueue item per AG.
Each workqueue item walked the inode btree (with XFS_IOC_INUMBERS) to find
inode chunks and then called bulkstat (XFS_IOC_BULKSTAT) to gather enough
information to construct file handles.
The file handle was then passed to a function to generate scrub items for each
metadata object of each inode.
This simple algorithm leads to thread balancing problems in phase 3 if the
filesystem contains one AG with a few large sparse files and the rest of the
AGs contain many smaller files.
The inode scan dispatch function was not sufficiently granular; it should have
been dispatching at the level of individual inodes, or, to constrain memory
consumption, inode btree records.
Thanks to Dave Chinner, bounded workqueues in userspace enable xfs_scrub to
avoid this problem with ease by adding a second workqueue.
Just like before, the first workqueue is seeded with one workqueue item per AG,
and it uses INUMBERS to find inode btree chunks.
The second workqueue, however, is configured with an upper bound on the number
of items that can be waiting to be run.
Each inode btree chunk found by the first workqueue’s workers are queued to the
second workqueue, and it is this second workqueue that queries BULKSTAT,
creates a file handle, and passes it to a function to generate scrub items for
each metadata object of each inode.
If the second workqueue is too full, the workqueue add function blocks the
first workqueue’s workers until the backlog eases.
This doesn’t completely solve the balancing problem, but reduces it enough to
move on to more pressing issues.
The proposed patchsets are the scrub performance tweaks and the inode scan rebalance series.
4.6.3. Scheduling Repairs¶
During phase 2, corruptions and inconsistencies reported in any AGI header or inode btree are repaired immediately, because phase 3 relies on proper functioning of the inode indices to find inodes to scan. Failed repairs are rescheduled to phase 4. Problems reported in any other space metadata are deferred to phase 4. Optimization opportunities are always deferred to phase 4, no matter their origin.
During phase 3, corruptions and inconsistencies reported in any part of a file’s metadata are repaired immediately if all space metadata were validated during phase 2. Repairs that fail or cannot be repaired immediately are scheduled for phase 4.
In the original design of xfs_scrub, it was thought that repairs would be
so infrequent that the struct xfs_scrub_metadata objects used to
communicate with the kernel could also be used as the primary object to
schedule repairs.
With recent increases in the number of optimizations possible for a given
filesystem object, it became much more memory-efficient to track all eligible
repairs for a given filesystem object with a single repair item.
Each repair item represents a single lockable object -- AGs, metadata files,
individual inodes, or a class of summary information.
Phase 4 is responsible for scheduling a lot of repair work in as quick a
manner as is practical.
The data dependencies outlined earlier still apply, which
means that xfs_scrub must try to complete the repair work scheduled by
phase 2 before trying repair work scheduled by phase 3.
The repair process is as follows:
Start a round of repair with a workqueue and enough workers to keep the CPUs as busy as the user desires.
For each repair item queued by phase 2,
Ask the kernel to repair everything listed in the repair item for a given filesystem object.
Make a note if the kernel made any progress in reducing the number of repairs needed for this object.
If the object no longer requires repairs, revalidate all metadata associated with this object. If the revalidation succeeds, drop the repair item. If not, requeue the item for more repairs.
If any repairs were made, jump back to 1a to retry all the phase 2 items.
For each repair item queued by phase 3,
Ask the kernel to repair everything listed in the repair item for a given filesystem object.
Make a note if the kernel made any progress in reducing the number of repairs needed for this object.
If the object no longer requires repairs, revalidate all metadata associated with this object. If the revalidation succeeds, drop the repair item. If not, requeue the item for more repairs.
If any repairs were made, jump back to 1c to retry all the phase 3 items.
If step 1 made any repair progress of any kind, jump back to step 1 to start another round of repair.
If there are items left to repair, run them all serially one more time. Complain if the repairs were not successful, since this is the last chance to repair anything.
Corruptions and inconsistencies encountered during phases 5 and 7 are repaired immediately. Corrupt file data blocks reported by phase 6 cannot be recovered by the filesystem.
The proposed patchsets are the repair warning improvements, refactoring of the repair data dependency and object tracking, and the repair scheduling improvement series.
4.6.4. Checking Names for Confusable Unicode Sequences¶
If xfs_scrub succeeds in validating the filesystem metadata by the end of
phase 4, it moves on to phase 5, which checks for suspicious looking names in
the filesystem.
These names consist of the filesystem label, names in directory entries, and
the names of extended attributes.
Like most Unix filesystems, XFS imposes the sparest of constraints on the
contents of a name:
Slashes and null bytes are not allowed in directory entries.
Null bytes are not allowed in userspace-visible extended attributes.
Null bytes are not allowed in the filesystem label.
Directory entries and attribute keys store the length of the name explicitly ondisk, which means that nulls are not name terminators. For this section, the term “naming domain” refers to any place where names are presented together -- all the names in a directory, or all the attributes of a file.
Although the Unix naming constraints are very permissive, the reality of most modern-day Linux systems is that programs work with Unicode character code points to support international languages. These programs typically encode those code points in UTF-8 when interfacing with the C library because the kernel expects null-terminated names. In the common case, therefore, names found in an XFS filesystem are actually UTF-8 encoded Unicode data.
To maximize its expressiveness, the Unicode standard defines separate control points for various characters that render similarly or identically in writing systems around the world. For example, the character “Cyrillic Small Letter A” U+0430 “а” often renders identically to “Latin Small Letter A” U+0061 “a”.
The standard also permits characters to be constructed in multiple ways -- either by using a defined code point, or by combining one code point with various combining marks. For example, the character “Angstrom Sign U+212B “Å” can also be expressed as “Latin Capital Letter A” U+0041 “A” followed by “Combining Ring Above” U+030A “◌̊”. Both sequences render identically.
Like the standards that preceded it, Unicode also defines various control characters to alter the presentation of text. For example, the character “Right-to-Left Override” U+202E can trick some programs into rendering “moo\xe2\x80\xaegnp.txt” as “mootxt.png”. A second category of rendering problems involves whitespace characters. If the character “Zero Width Space” U+200B is encountered in a file name, the name will render identically to a name that does not have the zero width space.
If two names within a naming domain have different byte sequences but render identically, a user may be confused by it. The kernel, in its indifference to upper level encoding schemes, permits this. Most filesystem drivers persist the byte sequence names that are given to them by the VFS.
Techniques for detecting confusable names are explained in great detail in
sections 4 and 5 of the
Unicode Security Mechanisms
document.
When xfs_scrub detects UTF-8 encoding in use on a system, it uses the
Unicode normalization form NFD in conjunction with the confusable name
detection component of
libicu
to identify names with a directory or within a file’s extended attributes that
could be confused for each other.
Names are also checked for control characters, non-rendering characters, and
mixing of bidirectional characters.
All of these potential issues are reported to the system administrator during
phase 5.
4.6.5. Media Verification of File Data Extents¶
The system administrator can elect to initiate a media scan of all file data
blocks.
This scan after validation of all filesystem metadata (except for the summary
counters) as phase 6.
The scan starts by calling FS_IOC_GETFSMAP to scan the filesystem space map
to find areas that are allocated to file data fork extents.
Gaps between data fork extents that are smaller than 64k are treated as if
they were data fork extents to reduce the command setup overhead.
When the space map scan accumulates a region larger than 32MB, a media
verification request is sent to the disk as a directio read of the raw block
device.
If the verification read fails, xfs_scrub retries with single-block reads
to narrow down the failure to the specific region of the media and recorded.
When it has finished issuing verification requests, it again uses the space
mapping ioctl to map the recorded media errors back to metadata structures
and report what has been lost.
For media errors in blocks owned by files, parent pointers can be used to
construct file paths from inode numbers for user-friendly reporting.
4.7. 7. Conclusion and Future Work¶
It is hoped that the reader of this document has followed the designs laid out in this document and now has some familiarity with how XFS performs online rebuilding of its metadata indices, and how filesystem users can interact with that functionality. Although the scope of this work is daunting, it is hoped that this guide will make it easier for code readers to understand what has been built, for whom it has been built, and why. Please feel free to contact the XFS mailing list with questions.
4.7.1. XFS_IOC_EXCHANGE_RANGE¶
As discussed earlier, a second frontend to the atomic file mapping exchange mechanism is a new ioctl call that userspace programs can use to commit updates to files atomically. This frontend has been out for review for several years now, though the necessary refinements to online repair and lack of customer demand mean that the proposal has not been pushed very hard.
4.7.1.1. File Content Exchanges with Regular User Files¶
As mentioned earlier, XFS has long had the ability to swap extents between
files, which is used almost exclusively by xfs_fsr to defragment files.
The earliest form of this was the fork swap mechanism, where the entire
contents of data forks could be exchanged between two files by exchanging the
raw bytes in each inode fork’s immediate area.
When XFS v5 came along with self-describing metadata, this old mechanism grew
some log support to continue rewriting the owner fields of BMBT blocks during
log recovery.
When the reverse mapping btree was later added to XFS, the only way to maintain
the consistency of the fork mappings with the reverse mapping index was to
develop an iterative mechanism that used deferred bmap and rmap operations to
swap mappings one at a time.
This mechanism is identical to steps 2-3 from the procedure above except for
the new tracking items, because the atomic file mapping exchange mechanism is
an iteration of an existing mechanism and not something totally novel.
For the narrow case of file defragmentation, the file contents must be
identical, so the recovery guarantees are not much of a gain.
Atomic file content exchanges are much more flexible than the existing swapext implementations because it can guarantee that the caller never sees a mix of old and new contents even after a crash, and it can operate on two arbitrary file fork ranges. The extra flexibility enables several new use cases:
Atomic commit of file writes: A userspace process opens a file that it wants to update. Next, it opens a temporary file and calls the file clone operation to reflink the first file’s contents into the temporary file. Writes to the original file should instead be written to the temporary file. Finally, the process calls the atomic file mapping exchange system call (
XFS_IOC_EXCHANGE_RANGE) to exchange the file contents, thereby committing all of the updates to the original file, or none of them.
Transactional file updates: The same mechanism as above, but the caller only wants the commit to occur if the original file’s contents have not changed. To make this happen, the calling process snapshots the file modification and change timestamps of the original file before reflinking its data to the temporary file. When the program is ready to commit the changes, it passes the timestamps into the kernel as arguments to the atomic file mapping exchange system call. The kernel only commits the changes if the provided timestamps match the original file. A new ioctl (
XFS_IOC_COMMIT_RANGE) is provided to perform this.Emulation of atomic block device writes: Export a block device with a logical sector size matching the filesystem block size to force all writes to be aligned to the filesystem block size. Stage all writes to a temporary file, and when that is complete, call the atomic file mapping exchange system call with a flag to indicate that holes in the temporary file should be ignored. This emulates an atomic device write in software, and can support arbitrary scattered writes.
4.7.2. Vectorized Scrub¶
As it turns out, the refactoring of repair items mentioned earlier was a catalyst for enabling a vectorized scrub system call. Since 2018, the cost of making a kernel call has increased considerably on some systems to mitigate the effects of speculative execution attacks. This incentivizes program authors to make as few system calls as possible to reduce the number of times an execution path crosses a security boundary.
With vectorized scrub, userspace pushes to the kernel the identity of a
filesystem object, a list of scrub types to run against that object, and a
simple representation of the data dependencies between the selected scrub
types.
The kernel executes as much of the caller’s plan as it can until it hits a
dependency that cannot be satisfied due to a corruption, and tells userspace
how much was accomplished.
It is hoped that io_uring will pick up enough of this functionality that
online fsck can use that instead of adding a separate vectored scrub system
call to XFS.
The relevant patchsets are the kernel vectorized scrub and userspace vectorized scrub series.
4.7.3. Quality of Service Targets for Scrub¶
One serious shortcoming of the online fsck code is that the amount of time that
it can spend in the kernel holding resource locks is basically unbounded.
Userspace is allowed to send a fatal signal to the process which will cause
xfs_scrub to exit when it reaches a good stopping point, but there’s no way
for userspace to provide a time budget to the kernel.
Given that the scrub codebase has helpers to detect fatal signals, it shouldn’t
be too much work to allow userspace to specify a timeout for a scrub/repair
operation and abort the operation if it exceeds budget.
However, most repair functions have the property that once they begin to touch
ondisk metadata, the operation cannot be cancelled cleanly, after which a QoS
timeout is no longer useful.
4.7.4. Defragmenting Free Space¶
Over the years, many XFS users have requested the creation of a program to
clear a portion of the physical storage underlying a filesystem so that it
becomes a contiguous chunk of free space.
Call this free space defragmenter clearspace for short.
The first piece the clearspace program needs is the ability to read the
reverse mapping index from userspace.
This already exists in the form of the FS_IOC_GETFSMAP ioctl.
The second piece it needs is a new fallocate mode
(FALLOC_FL_MAP_FREE_SPACE) that allocates the free space in a region and
maps it to a file.
Call this file the “space collector” file.
The third piece is the ability to force an online repair.
To clear all the metadata out of a portion of physical storage, clearspace
uses the new fallocate map-freespace call to map any free space in that region
to the space collector file.
Next, clearspace finds all metadata blocks in that region by way of
GETFSMAP and issues forced repair requests on the data structure.
This often results in the metadata being rebuilt somewhere that is not being
cleared.
After each relocation, clearspace calls the “map free space” function again to
collect any newly freed space in the region being cleared.
To clear all the file data out of a portion of the physical storage, clearspace
uses the FSMAP information to find relevant file data blocks.
Having identified a good target, it uses the FICLONERANGE call on that part
of the file to try to share the physical space with a dummy file.
Cloning the extent means that the original owners cannot overwrite the
contents; any changes will be written somewhere else via copy-on-write.
Clearspace makes its own copy of the frozen extent in an area that is not being
cleared, and uses FIEDEUPRANGE (or the atomic file content exchanges feature) to change the target file’s data extent
mapping away from the area being cleared.
When all other mappings have been moved, clearspace reflinks the space into the
space collector file so that it becomes unavailable.
There are further optimizations that could apply to the above algorithm.
To clear a piece of physical storage that has a high sharing factor, it is
strongly desirable to retain this sharing factor.
In fact, these extents should be moved first to maximize sharing factor after
the operation completes.
To make this work smoothly, clearspace needs a new ioctl
(FS_IOC_GETREFCOUNTS) to report reference count information to userspace.
With the refcount information exposed, clearspace can quickly find the longest,
most shared data extents in the filesystem, and target them first.
Future Work Question: How might the filesystem move inode chunks?
Answer: To move inode chunks, Dave Chinner constructed a prototype program that creates a new file with the old contents and then locklessly runs around the filesystem updating directory entries. The operation cannot complete if the filesystem goes down. That problem isn’t totally insurmountable: create an inode remapping table hidden behind a jump label, and a log item that tracks the kernel walking the filesystem to update directory entries. The trouble is, the kernel can’t do anything about open files, since it cannot revoke them.
Future Work Question: Can static keys be used to minimize the cost of
supporting revoke() on XFS files?
Answer: Yes. Until the first revocation, the bailout code need not be in the call path at all.
The relevant patchsets are the kernel freespace defrag and userspace freespace defrag series.
4.7.5. Shrinking Filesystems¶
Removing the end of the filesystem ought to be a simple matter of evacuating the data and metadata at the end of the filesystem, and handing the freed space to the shrink code. That requires an evacuation of the space at end of the filesystem, which is a use of free space defragmentation!