This document captures the design of the online filesystem check feature for
XFS.
The purpose of this document is threefold:
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.
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. |
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.
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 “xfs_repair”.
The userspace driver program for the new online fsck tool can be
referred to as “xfs_scrub”.
The kernel portion of online fsck that validates metadata is called
“online scrub”, and portion of the kernel that fixes metadata is called
“online repair”. |
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.
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.
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.
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 fsstress and xfs_scrub -n to ensure that checking the whole
filesystem doesn’t cause problems.
Race fsstress and xfs_scrub in force-rebuild mode to ensure that
force-repairing the whole filesystem doesn’t cause problems.
Race xfs_scrub in check and force-repair mode against fsstress while
freezing and thawing the filesystem.
Race xfs_scrub in check and force-repair mode against fsstress while
remounting the filesystem read-only and read-write.
The same, but running fsx instead of fsstress. (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.
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.
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.
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:
The decision to enable the background scan is left to the system administrator.
This can be done by enabling either of the following services:
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.
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.
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.
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.
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.
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.
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)
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.
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.
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?
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.
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.
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.
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_pending objects, each of which maintains a list of struct
xfs_extent_free_item objects.
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 the struct xfs_extent_free_item object 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.
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_bump to increase the counter.
When the deferred item manager wants to finish the deferred work item, it
calls ->finish_item to complete it.
The ->finish_item implementation logs some changes and calls
xfs_defer_drain_drop to 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_busy returns 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.
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_FALSE macro takes care of this.
The static key itself should be declared as a static variable.
When deciding to invoke code that’s only used by scrub, the regular
filesystem should call the static_branch_unlikely predicate 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_inc to enable and static_branch_dec to 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_enable from 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; the TRY_HARDER flag 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.
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.
|
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.
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.
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.
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.
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 */
}
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.
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.
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.
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.
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.
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.
An online fsck function that wants to create an xfbtree should proceed as
follows:
Call xfile_create to create an xfile.
Call xfs_alloc_memory_buftarg to create a buffer cache target structure
pointing to the xfile.
Pass the buffer cache target, buffer ops, and other information to
xfbtree_init to initialize the passed in struct xfbtree and 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 define xfs_rmapbt_mem_create to 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_cursor takes 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.
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.
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.
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:
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 + 1 and 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.
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.
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_finish to 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.
The high level process to rebuild the inode index btree is:
Walk the reverse mapping records to generate struct xfs_inobt_rec
records 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_geometry function 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_bload to 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.
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_irec
records for any space having more than one reverse mapping and add them to
the xfarray.
Any records owned by XFS_RMAP_OWN_COW are 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_REFC to 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_geometry function to compute the number
of blocks needed for the new tree.
Allocate the number of blocks computed in the previous step.
Use xfs_btree_bload to 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.
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.
The high level process to rebuild a data/attr fork mapping btree is:
Walk the reverse mapping records to generate struct xfs_bmbt_rec
records 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 the BMBT_BLOCK
records.
Use the xfs_btree_bload_compute_geometry function 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_bload to 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.
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.
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.
The high level process to rebuild the free space indices is:
Walk the reverse mapping records to generate struct xfs_alloc_rec_incore
records from the gaps in the reverse mapping btree.
Append the records to an xfarray.
Use the xfs_btree_bload_compute_geometry function 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_bload to 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.
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.
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.
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.
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.
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_COMPLETE to
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:
Other programs can unfreeze the filesystem without our knowledge.
This leads to incorrect scan results and incorrect repairs. Adding an extra lock to prevent others from thawing the filesystem
required the addition of a ->freeze_super function to wrap
freeze_fs().
This in turn caused other subtle problems because it turns out that
the VFS freeze_super and thaw_super functions can drop the
last reference to the VFS superblock, and any subsequent access
becomes a UAF bug!
This can happen if the filesystem is unmounted while the underlying
block device has frozen the filesystem.
This problem could be solved by grabbing extra references to the
superblock, but it felt suboptimal given the other inadequacies of
this approach. The log need not be quiesced to check the summary counters, but a VFS
freeze initiates one anyway.
This adds unnecessary runtime to live fscounter fsck operations. Quiescing the log means that XFS flushes the (possibly incorrect)
counters to disk as part of cleaning the log. A bug in the VFS meant that freeze could complete even when
sync_filesystem fails to flush the filesystem and returns an error.
This bug was fixed in Linux 5.17.
|
The proposed patchset is the
summary counter cleanup
series.
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:
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_teardown to 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.
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 DONTCACHE inode
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 (VFS i_rwsem) lock to control file IO.
Inode MMAPLOCK (page cache invalidate_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 ILOCK for 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.
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.
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.
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.
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_hooks object 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_hooks and xfs_hook objects to take advantage of type
checking to ensure correct usage.
A callsite in the regular filesystem code must be chosen to call
xfs_hooks_call with 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_hook object.
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_setup to initialize the struct xfs_hook, and
xfs_hooks_add to enable the hook.
Online fsck must call xfs_hooks_del to 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.
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_start starts a scan
xchk_iscan_iter grabs a reference to the next inode in the scan or
returns zero if there is nothing left to scan
xchk_iscan_want_live_update to 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_visited to mark an inode as having been visited in the
scan
xchk_iscan_teardown to finish the scan
This functionality is also a part of the
inode scanner
series.
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.
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.
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.
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 COLLAPSE_RANGE operation to slide the new extents into
place.
This had many drawbacks:
Array structures are linearly addressed, and the regular filesystem
codebase does not have the concept of a linear offset that could be
applied to the record offset computation to build an alternate copy. Extended attributes are allowed to use the entire attr fork offset
address space. Even if repair could build an alternate copy of a data structure in a
different part of the fork address space, the atomic repair commit
requirement means that online repair would have to be able to perform
a log assisted COLLAPSE_RANGE operation to ensure that the old
structure was completely replaced. A crash after construction of the secondary tree but before the range
collapse would leave unreachable blocks in the file fork.
This would likely confuse things further. Reaping blocks after a repair is not a simple operation, and
initiating a reap operation from a restarted range collapse operation
during log recovery is daunting. Directory entry blocks and quota records record the file fork offset
in the header area of each block.
An atomic range collapse operation would have to rewrite this part of
each block header.
Rewriting a single field in block headers is not a huge problem, but
it’s something to be aware of. Each block in a directory or extended attributes btree index contains
sibling and child block pointers.
Were the atomic commit to use a range collapse operation, each block
would have to be rewritten very carefully to preserve the graph
structure.
Doing this as part of a range collapse means rewriting a large number
of blocks repeatedly, which is not conducive to quick repairs.
This lead to the introduction of temporary file staging.
|
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_copyin can 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.
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
sb_features_log_incompat field to indicate that the log contains
records that might not readable by all kernels that could mount this
filesystem.
In short, log incompat features protect the log contents against kernels
that will not understand the contents.
Unlike the other superblock feature bits, log incompat bits are
ephemeral because an empty (clean) log does not need protection.
The log cleans itself after its contents have been committed into the
filesystem, either as part of an unmount or because the system is
otherwise idle.
Because upper level code can be working on a transaction at the same
time that the log cleans itself, it is necessary for upper level code to
communicate to the log when it is going to use a log incompatible
feature.
The log coordinates access to incompatible features through the use of
one struct rw_semaphore for each feature.
The log cleaning code tries to take this rwsem in exclusive mode to
clear the bit; if the lock attempt fails, the feature bit remains set.
The code supporting a log incompat feature should create wrapper
functions to obtain the log feature and call
xfs_add_incompat_log_feature to set the feature bits in the primary
superblock.
The superblock update is performed transactionally, so the wrapper to
obtain log assistance must be called just prior to the creation of the
transaction that uses the functionality.
For a file operation, this step must happen after taking the IOLOCK
and the MMAPLOCK, but before allocating the transaction.
When the transaction is complete, the xlog_drop_incompat_feat
function is called to release the feature.
The feature bit will not be cleared from the superblock until the log
becomes clean.
Log-assisted extended attribute updates and file content exchanges bothe
use log incompat features and provide convenience wrappers around the
functionality.
|
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_finish to process the exchange.
This is encapsulated in xrep_tempexch_contents for scrub operations.
This will log an extent swap intent item to the transaction for the deferred
mapping exchange work item.
Until xmi_blockcount of the deferred mapping exchange work item is zero,
Read the block maps of both file ranges starting at xmi_startoff1 and
xmi_startoff2, respectively, and compute the longest extent that can
be exchanged in a single step.
This is the minimum of the two br_blockcount s 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_startoff1 and xmi_startoff2
by the number of blocks computed in the previous step, and decrease
xmi_blockcount by 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.
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.
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.
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_alloc to allocate a new scrub transaction with
the appropriate resource reservations, locks, and fill out a struct
xfs_exchmaps_req with the details of the exchange operation.
Call xrep_tempexch_contents to exchange the contents.
Commit the transaction to complete the repair.
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.
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.
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.
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.
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 XFS codebase of the late 2000s did not have the infrastructure to
enforce strong referential integrity in the directory tree.
It did not guarantee that a change in a forward link would always be
followed up with the corresponding change to the reverse links. Referential integrity was not integrated into offline repair.
Checking and repairs were performed on mounted filesystems without
taking any kernel or inode locks to coordinate access.
It is not clear how this actually worked properly. The extended attribute did not record the name of the directory entry
in the parent, so the SGI parent pointer implementation cannot be
used to reconnect the directory tree. Extended attribute forks only support 65,536 extents, which means
that parent pointer attribute creation is likely to fail at some
point before the maximum file link count is achieved.
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 (parent_inum, parent_gen, dirent_pos) → (dirent_name).
The format was changed during development to eliminate the requirement
of repair tools needing to to ensure that the dirent_pos field
always matched when reconstructing a directory.
There were a few other ways to have solved that problem:
The field could be designated advisory, since the other three values
are sufficient to find the entry in the parent.
However, this makes indexed key lookup impossible while repairs are
ongoing. We could allow creating directory entries at specified offsets, which
solves the referential integrity problem but runs the risk that
dirent creation will fail due to conflicts with the free space in the
directory.
These conflicts could be resolved by appending the directory entry
and amending the xattr code to support updating an xattr key and
reindexing the dabtree, though this would have to be performed with
the parent directory still locked.
Same as above, but remove the old parent pointer entry and add a new
one atomically. Change the ondisk xattr format to
(parent_inum, name) → (parent_gen), which would provide the attr
name uniqueness that we require, without forcing repair code to
update the dirent position.
Unfortunately, this requires changes to the xattr code to support
attr names as long as 263 bytes. Change the ondisk xattr format to (parent_inum, hash(name)) →
(name, parent_gen).
If the hash is sufficiently resistant to collisions (e.g. sha256)
then this should provide the attr name uniqueness that we require.
Names shorter than 247 bytes could be stored directly. Change the ondisk xattr format to (dirent_name) → (parent_ino,
parent_gen). This format doesn’t require any of the complicated
nested name hashing of the previous suggestions. However, it was
discovered that multiple hardlinks to the same inode with the same
filename caused performance problems with hashed xattr lookups, so
the parent inumber is now xor’d into the hash index.
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.
|
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.
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.
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. The name_hash
referenced 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, and name_cookie.
Having a single name_cookie for each name is 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,
and name_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, and name_cookie fields 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.
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.
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_create at 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_two function follows the inode locking
strategy discussed earlier.
Use xrep_adoption_trans_alloc to reserve resources to the repair
transaction.
Call xrep_orphanage_compute_name to compute the new name in the
orphanage.
If the adoption is going to happen, call xrep_adoption_reparent to
reparent the orphaned file into the lost and found and invalidate the dentry
cache.
Call xrep_adoption_finish to commit any filesystem updates, release the
orphanage ILOCK, and clean the scrub transaction. Call
xrep_adoption_commit to commit the updates and the scrub transaction.
If a runtime error happens, call xrep_adoption_cancel to release all
resources.
The proposed patches are in the
orphanage adoption
series.
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.
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.
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.
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.
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.
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.
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.
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.
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.
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.
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!
