code:
.. SPDX-License-Identifier: GPL-2.0

Written by: Neil Brown
Please see MAINTAINERS file for where to send questions.

Overlay Filesystem
==================

This document describes a prototype for a new approach to providing
overlay-filesystem functionality in Linux (sometimes referred to as
union-filesystems).  An overlay-filesystem tries to present a
filesystem which is the result over overlaying one filesystem on top
of the other.

Overlay objects
---------------

The overlay filesystem approach is 'hybrid', because the objects that
appear in the filesystem do not always appear to belong to that filesystem.
In many cases, an object accessed in the union will be indistinguishable
from accessing the corresponding object from the original filesystem.
This is most obvious from the 'st_dev' field returned by stat(2).

While directories will report an st_dev from the overlay-filesystem,
non-directory objects may report an st_dev from the lower filesystem or
upper filesystem that is providing the object.  Similarly st_ino will
only be unique when combined with st_dev, and both of these can change
over the lifetime of a non-directory object.  Many applications and
tools ignore these values and will not be affected.

In the special case of all overlay layers on the same underlying
filesystem, all objects will report an st_dev from the overlay
filesystem and st_ino from the underlying filesystem.  This will
make the overlay mount more compliant with filesystem scanners and
overlay objects will be distinguishable from the corresponding
objects in the original filesystem.

On 64bit systems, even if all overlay layers are not on the same
underlying filesystem, the same compliant behavior could be achieved
with the "xino" feature.  The "xino" feature composes a unique object
identifier from the real object st_ino and an underlying fsid index.
The "xino" feature uses the high inode number bits for fsid, because the
underlying filesystems rarely use the high inode number bits.  In case
the underlying inode number does overflow into the high xino bits, overlay
filesystem will fall back to the non xino behavior for that inode.

The "xino" feature can be enabled with the "-o xino=on" overlay mount option.
If all underlying filesystems support NFS file handles, the value of st_ino
for overlay filesystem objects is not only unique, but also persistent over
the lifetime of the filesystem.  The "-o xino=auto" overlay mount option
enables the "xino" feature only if the persistent st_ino requirement is met.

The following table summarizes what can be expected in different overlay
configurations.

Inode properties
````````````````

+--------------+------------+------------+-----------------+----------------+
|Configuration | Persistent | Uniform    | st_ino == d_ino | d_ino == i_ino |
|              | st_ino     | st_dev     |                 | [*]            |
+==============+=====+======+=====+======+========+========+========+=======+
|              | dir | !dir | dir | !dir |  dir   +  !dir  |  dir   | !dir  |
+--------------+-----+------+-----+------+--------+--------+--------+-------+
| All layers   |  Y  |  Y   |  Y  |  Y   |  Y     |   Y    |  Y     |  Y    |
| on same fs   |     |      |     |      |        |        |        |       |
+--------------+-----+------+-----+------+--------+--------+--------+-------+
| Layers not   |  N  |  N   |  Y  |  N   |  N     |   Y    |  N     |  Y    |
| on same fs,  |     |      |     |      |        |        |        |       |
| xino=off     |     |      |     |      |        |        |        |       |
+--------------+-----+------+-----+------+--------+--------+--------+-------+
| xino=on/auto |  Y  |  Y   |  Y  |  Y   |  Y     |   Y    |  Y     |  Y    |
+--------------+-----+------+-----+------+--------+--------+--------+-------+
| xino=on/auto,|  N  |  N   |  Y  |  N   |  N     |   Y    |  N     |  Y    |
| ino overflow |     |      |     |      |        |        |        |       |
+--------------+-----+------+-----+------+--------+--------+--------+-------+

[*] nfsd v3 readdirplus verifies d_ino == i_ino. i_ino is exposed via several
/proc files, such as /proc/locks and /proc/self/fdinfo/<fd> of an inotify
file descriptor.

Upper and Lower
---------------

An overlay filesystem combines two filesystems - an 'upper' filesystem
and a 'lower' filesystem.  When a name exists in both filesystems, the
object in the 'upper' filesystem is visible while the object in the
'lower' filesystem is either hidden or, in the case of directories,
merged with the 'upper' object.

It would be more correct to refer to an upper and lower 'directory
tree' rather than 'filesystem' as it is quite possible for both
directory trees to be in the same filesystem and there is no
requirement that the root of a filesystem be given for either upper or
lower.

A wide range of filesystems supported by Linux can be the lower filesystem,
but not all filesystems that are mountable by Linux have the features
needed for OverlayFS to work.  The lower filesystem does not need to be
writable.  The lower filesystem can even be another overlayfs.  The upper
filesystem will normally be writable and if it is it must support the
creation of trusted.* and/or user.* extended attributes, and must provide
valid d_type in readdir responses, so NFS is not suitable.

A read-only overlay of two read-only filesystems may use any
filesystem type.

Directories
-----------

Overlaying mainly involves directories.  If a given name appears in both
upper and lower filesystems and refers to a non-directory in either,
then the lower object is hidden - the name refers only to the upper
object.

Where both upper and lower objects are directories, a merged directory
is formed.

At mount time, the two directories given as mount options "lowerdir" and
"upperdir" are combined into a merged directory:

  mount -t overlay overlay -olowerdir=/lower,upperdir=/upper,\
  workdir=/work /merged

The "workdir" needs to be an empty directory on the same filesystem
as upperdir.

Then whenever a lookup is requested in such a merged directory, the
lookup is performed in each actual directory and the combined result
is cached in the dentry belonging to the overlay filesystem.  If both
actual lookups find directories, both are stored and a merged
directory is created, otherwise only one is stored: the upper if it
exists, else the lower.

Only the lists of names from directories are merged.  Other content
such as metadata and extended attributes are reported for the upper
directory only.  These attributes of the lower directory are hidden.

whiteouts and opaque directories
--------------------------------

In order to support rm and rmdir without changing the lower
filesystem, an overlay filesystem needs to record in the upper filesystem
that files have been removed.  This is done using whiteouts and opaque
directories (non-directories are always opaque).

A whiteout is created as a character device with 0/0 device number.
When a whiteout is found in the upper level of a merged directory, any
matching name in the lower level is ignored, and the whiteout itself
is also hidden.

A directory is made opaque by setting the xattr "trusted.overlay.opaque"
to "y".  Where the upper filesystem contains an opaque directory, any
directory in the lower filesystem with the same name is ignored.

readdir
-------

When a 'readdir' request is made on a merged directory, the upper and
lower directories are each read and the name lists merged in the
obvious way (upper is read first, then lower - entries that already
exist are not re-added).  This merged name list is cached in the
'struct file' and so remains as long as the file is kept open.  If the
directory is opened and read by two processes at the same time, they
will each have separate caches.  A seekdir to the start of the
directory (offset 0) followed by a readdir will cause the cache to be
discarded and rebuilt.

This means that changes to the merged directory do not appear while a
directory is being read.  This is unlikely to be noticed by many
programs.

seek offsets are assigned sequentially when the directories are read.
Thus if

  - read part of a directory
  - remember an offset, and close the directory
  - re-open the directory some time later
  - seek to the remembered offset

there may be little correlation between the old and new locations in
the list of filenames, particularly if anything has changed in the
directory.

Readdir on directories that are not merged is simply handled by the
underlying directory (upper or lower).

renaming directories
--------------------

When renaming a directory that is on the lower layer or merged (i.e. the
directory was not created on the upper layer to start with) overlayfs can
handle it in two different ways:

1. return EXDEV error: this error is returned by rename(2) when trying to
   move a file or directory across filesystem boundaries.  Hence
   applications are usually prepared to hande this error (mv(1) for example
   recursively copies the directory tree).  This is the default behavior.

2. If the "redirect_dir" feature is enabled, then the directory will be
   copied up (but not the contents).  Then the "trusted.overlay.redirect"
   extended attribute is set to the path of the original location from the
   root of the overlay.  Finally the directory is moved to the new
   location.

There are several ways to tune the "redirect_dir" feature.

Kernel config options:

- OVERLAY_FS_REDIRECT_DIR:
    If this is enabled, then redirect_dir is turned on by  default.
- OVERLAY_FS_REDIRECT_ALWAYS_FOLLOW:
    If this is enabled, then redirects are always followed by default. Enabling
    this results in a less secure configuration.  Enable this option only when
    worried about backward compatibility with kernels that have the redirect_dir
    feature and follow redirects even if turned off.

Module options (can also be changed through /sys/module/overlay/parameters/):

- "redirect_dir=BOOL":
    See OVERLAY_FS_REDIRECT_DIR kernel config option above.
- "redirect_always_follow=BOOL":
    See OVERLAY_FS_REDIRECT_ALWAYS_FOLLOW kernel config option above.
- "redirect_max=NUM":
    The maximum number of bytes in an absolute redirect (default is 256).

Mount options:

- "redirect_dir=on":
    Redirects are enabled.
- "redirect_dir=follow":
    Redirects are not created, but followed.
- "redirect_dir=off":
    Redirects are not created and only followed if "redirect_always_follow"
    feature is enabled in the kernel/module config.
- "redirect_dir=nofollow":
    Redirects are not created and not followed (equivalent to "redirect_dir=off"
    if "redirect_always_follow" feature is not enabled).

When the NFS export feature is enabled, every copied up directory is
indexed by the file handle of the lower inode and a file handle of the
upper directory is stored in a "trusted.overlay.upper" extended attribute
on the index entry.  On lookup of a merged directory, if the upper
directory does not match the file handle stores in the index, that is an
indication that multiple upper directories may be redirected to the same
lower directory.  In that case, lookup returns an error and warns about
a possible inconsistency.

Because lower layer redirects cannot be verified with the index, enabling
NFS export support on an overlay filesystem with no upper layer requires
turning off redirect follow (e.g. "redirect_dir=nofollow").

Non-directories
---------------

Objects that are not directories (files, symlinks, device-special
files etc.) are presented either from the upper or lower filesystem as
appropriate.  When a file in the lower filesystem is accessed in a way
the requires write-access, such as opening for write access, changing
some metadata etc., the file is first copied from the lower filesystem
to the upper filesystem (copy_up).  Note that creating a hard-link
also requires copy_up, though of course creation of a symlink does
not.

The copy_up may turn out to be unnecessary, for example if the file is
opened for read-write but the data is not modified.

The copy_up process first makes sure that the containing directory
exists in the upper filesystem - creating it and any parents as
necessary.  It then creates the object with the same metadata (owner,
mode, mtime, symlink-target etc.) and then if the object is a file, the
data is copied from the lower to the upper filesystem.  Finally any
extended attributes are copied up.

Once the copy_up is complete, the overlay filesystem simply
provides direct access to the newly created file in the upper
filesystem - future operations on the file are barely noticed by the
overlay filesystem (though an operation on the name of the file such as
rename or unlink will of course be noticed and handled).

Permission model
----------------

Permission checking in the overlay filesystem follows these principles:

 1) permission check SHOULD return the same result before and after copy up

 2) task creating the overlay mount MUST NOT gain additional privileges

 3) non-mounting task MAY gain additional privileges through the overlay,
 compared to direct access on underlying lower or upper filesystems

This is achieved by performing two permission checks on each access

 a) check if current task is allowed access based on local DAC (owner,
    group, mode and posix acl), as well as MAC checks

 b) check if mounting task would be allowed real operation on lower or
    upper layer based on underlying filesystem permissions, again including
    MAC checks

Check (a) ensures consistency (1) since owner, group, mode and posix acls
are copied up.  On the other hand it can result in server enforced
permissions (used by NFS, for example) being ignored (3).

Check (b) ensures that no task gains permissions to underlying layers that
the mounting task does not have (2).  This also means that it is possible
to create setups where the consistency rule (1) does not hold; normally,
however, the mounting task will have sufficient privileges to perform all
operations.

Another way to demonstrate this model is drawing parallels between

  mount -t overlay overlay -olowerdir=/lower,upperdir=/upper,... /merged

and

  cp -a /lower /upper
  mount --bind /upper /merged

The resulting access permissions should be the same.  The difference is in
the time of copy (on-demand vs. up-front).

Multiple lower layers
---------------------

Multiple lower layers can now be given using the colon (":") as a
separator character between the directory names.  For example:

  mount -t overlay overlay -olowerdir=/lower1:/lower2:/lower3 /merged

As the example shows, "upperdir=" and "workdir=" may be omitted.  In
that case the overlay will be read-only.

The specified lower directories will be stacked beginning from the
rightmost one and going left.  In the above example lower1 will be the
top, lower2 the middle and lower3 the bottom layer.

Metadata only copy up
---------------------

When metadata only copy up feature is enabled, overlayfs will only copy
up metadata (as opposed to whole file), when a metadata specific operation
like chown/chmod is performed. Full file will be copied up later when
file is opened for WRITE operation.

In other words, this is delayed data copy up operation and data is copied
up when there is a need to actually modify data.

There are multiple ways to enable/disable this feature. A config option
CONFIG_OVERLAY_FS_METACOPY can be set/unset to enable/disable this feature
by default. Or one can enable/disable it at module load time with module
parameter metacopy=on/off. Lastly, there is also a per mount option
metacopy=on/off to enable/disable this feature per mount.

Do not use metacopy=on with untrusted upper/lower directories. Otherwise
it is possible that an attacker can create a handcrafted file with
appropriate REDIRECT and METACOPY xattrs, and gain access to file on lower
pointed by REDIRECT. This should not be possible on local system as setting
"trusted." xattrs will require CAP_SYS_ADMIN. But it should be possible
for untrusted layers like from a pen drive.

Note: redirect_dir={off|nofollow|follow[*]} and nfs_export=on mount options
conflict with metacopy=on, and will result in an error.

[*] redirect_dir=follow only conflicts with metacopy=on if upperdir=... is
given.

Sharing and copying layers
--------------------------

Lower layers may be shared among several overlay mounts and that is indeed
a very common practice.  An overlay mount may use the same lower layer
path as another overlay mount and it may use a lower layer path that is
beneath or above the path of another overlay lower layer path.

Using an upper layer path and/or a workdir path that are already used by
another overlay mount is not allowed and may fail with EBUSY.  Using
partially overlapping paths is not allowed and may fail with EBUSY.
If files are accessed from two overlayfs mounts which share or overlap the
upper layer and/or workdir path the behavior of the overlay is undefined,
though it will not result in a crash or deadlock.

Mounting an overlay using an upper layer path, where the upper layer path
was previously used by another mounted overlay in combination with a
different lower layer path, is allowed, unless the "inodes index" feature
or "metadata only copy up" feature is enabled.

With the "inodes index" feature, on the first time mount, an NFS file
handle of the lower layer root directory, along with the UUID of the lower
filesystem, are encoded and stored in the "trusted.overlay.origin" extended
attribute on the upper layer root directory.  On subsequent mount attempts,
the lower root directory file handle and lower filesystem UUID are compared
to the stored origin in upper root directory.  On failure to verify the
lower root origin, mount will fail with ESTALE.  An overlayfs mount with
"inodes index" enabled will fail with EOPNOTSUPP if the lower filesystem
does not support NFS export, lower filesystem does not have a valid UUID or
if the upper filesystem does not support extended attributes.

For "metadata only copy up" feature there is no verification mechanism at
mount time. So if same upper is mounted with different set of lower, mount
probably will succeed but expect the unexpected later on. So don't do it.

It is quite a common practice to copy overlay layers to a different
directory tree on the same or different underlying filesystem, and even
to a different machine.  With the "inodes index" feature, trying to mount
the copied layers will fail the verification of the lower root file handle.

Non-standard behavior
---------------------

Current version of overlayfs can act as a mostly POSIX compliant
filesystem.

This is the list of cases that overlayfs doesn't currently handle:

a) POSIX mandates updating st_atime for reads.  This is currently not
done in the case when the file resides on a lower layer.

b) If a file residing on a lower layer is opened for read-only and then
memory mapped with MAP_SHARED, then subsequent changes to the file are not
reflected in the memory mapping.

c) If a file residing on a lower layer is being executed, then opening that
file for write or truncating the file will not be denied with ETXTBSY.

The following options allow overlayfs to act more like a standards
compliant filesystem:

1) "redirect_dir"

Enabled with the mount option or module option: "redirect_dir=on" or with
the kernel config option CONFIG_OVERLAY_FS_REDIRECT_DIR=y.

If this feature is disabled, then rename(2) on a lower or merged directory
will fail with EXDEV ("Invalid cross-device link").

2) "inode index"

Enabled with the mount option or module option "index=on" or with the
kernel config option CONFIG_OVERLAY_FS_INDEX=y.

If this feature is disabled and a file with multiple hard links is copied
up, then this will "break" the link.  Changes will not be propagated to
other names referring to the same inode.

3) "xino"

Enabled with the mount option "xino=auto" or "xino=on", with the module
option "xino_auto=on" or with the kernel config option
CONFIG_OVERLAY_FS_XINO_AUTO=y.  Also implicitly enabled by using the same
underlying filesystem for all layers making up the overlay.

If this feature is disabled or the underlying filesystem doesn't have
enough free bits in the inode number, then overlayfs will not be able to
guarantee that the values of st_ino and st_dev returned by stat(2) and the
value of d_ino returned by readdir(3) will act like on a normal filesystem.
E.g. the value of st_dev may be different for two objects in the same
overlay filesystem and the value of st_ino for filesystem objects may not be
persistent and could change even while the overlay filesystem is mounted, as
summarized in the `Inode properties`_ table above.

Changes to underlying filesystems
---------------------------------

Changes to the underlying filesystems while part of a mounted overlay
filesystem are not allowed.  If the underlying filesystem is changed,
the behavior of the overlay is undefined, though it will not result in
a crash or deadlock.

Offline changes, when the overlay is not mounted, are allowed to the
upper tree.  Offline changes to the lower tree are only allowed if the
"metadata only copy up", "inode index", "xino" and "redirect_dir" features
have not been used.  If the lower tree is modified and any of these
features has been used, the behavior of the overlay is undefined,
though it will not result in a crash or deadlock.

When the overlay NFS export feature is enabled, overlay filesystems
behavior on offline changes of the underlying lower layer is different
than the behavior when NFS export is disabled.

On every copy_up, an NFS file handle of the lower inode, along with the
UUID of the lower filesystem, are encoded and stored in an extended
attribute "trusted.overlay.origin" on the upper inode.

When the NFS export feature is enabled, a lookup of a merged directory,
that found a lower directory at the lookup path or at the path pointed
to by the "trusted.overlay.redirect" extended attribute, will verify
that the found lower directory file handle and lower filesystem UUID
match the origin file handle that was stored at copy_up time.  If a
found lower directory does not match the stored origin, that directory
will not be merged with the upper directory.

NFS export
----------

When the underlying filesystems supports NFS export and the "nfs_export"
feature is enabled, an overlay filesystem may be exported to NFS.

With the "nfs_export" feature, on copy_up of any lower object, an index
entry is created under the index directory.  The index entry name is the
hexadecimal representation of the copy up origin file handle.  For a
non-directory object, the index entry is a hard link to the upper inode.
For a directory object, the index entry has an extended attribute
"trusted.overlay.upper" with an encoded file handle of the upper
directory inode.

When encoding a file handle from an overlay filesystem object, the
following rules apply:

1. For a non-upper object, encode a lower file handle from lower inode
2. For an indexed object, encode a lower file handle from copy_up origin
3. For a pure-upper object and for an existing non-indexed upper object,
   encode an upper file handle from upper inode

The encoded overlay file handle includes:
 - Header including path type information (e.g. lower/upper)
 - UUID of the underlying filesystem
 - Underlying filesystem encoding of underlying inode

This encoding format is identical to the encoding format file handles that
are stored in extended attribute "trusted.overlay.origin".

When decoding an overlay file handle, the following steps are followed:

1. Find underlying layer by UUID and path type information.
2. Decode the underlying filesystem file handle to underlying dentry.
3. For a lower file handle, lookup the handle in index directory by name.
4. If a whiteout is found in index, return ESTALE. This represents an
   overlay object that was deleted after its file handle was encoded.
5. For a non-directory, instantiate a disconnected overlay dentry from the
   decoded underlying dentry, the path type and index inode, if found.
6. For a directory, use the connected underlying decoded dentry, path type
   and index, to lookup a connected overlay dentry.

Decoding a non-directory file handle may return a disconnected dentry.
copy_up of that disconnected dentry will create an upper index entry with
no upper alias.

When overlay filesystem has multiple lower layers, a middle layer
directory may have a "redirect" to lower directory.  Because middle layer
"redirects" are not indexed, a lower file handle that was encoded from the
"redirect" origin directory, cannot be used to find the middle or upper
layer directory.  Similarly, a lower file handle that was encoded from a
descendant of the "redirect" origin directory, cannot be used to
reconstruct a connected overlay path.  To mitigate the cases of
directories that cannot be decoded from a lower file handle, these
directories are copied up on encode and encoded as an upper file handle.
On an overlay filesystem with no upper layer this mitigation cannot be
used NFS export in this setup requires turning off redirect follow (e.g.
"redirect_dir=nofollow").

The overlay filesystem does not support non-directory connectable file
handles, so exporting with the 'subtree_check' exportfs configuration will
cause failures to lookup files over NFS.

When the NFS export feature is enabled, all directory index entries are
verified on mount time to check that upper file handles are not stale.
This verification may cause significant overhead in some cases.

Note: the mount options index=off,nfs_export=on are conflicting for a
read-write mount and will result in an error.

Note: the mount option uuid=off can be used to replace UUID of the underlying
filesystem in file handles with null, and effectively disable UUID checks. This
can be useful in case the underlying disk is copied and the UUID of this copy
is changed. This is only applicable if all lower/upper/work directories are on
the same filesystem, otherwise it will fallback to normal behaviour.

Volatile mount
--------------

This is enabled with the "volatile" mount option.  Volatile mounts are not
guaranteed to survive a crash.  It is strongly recommended that volatile
mounts are only used if data written to the overlay can be recreated
without significant effort.

The advantage of mounting with the "volatile" option is that all forms of
sync calls to the upper filesystem are omitted.

In order to avoid a giving a false sense of safety, the syncfs (and fsync)
semantics of volatile mounts are slightly different than that of the rest of
VFS.  If any writeback error occurs on the upperdir's filesystem after a
volatile mount takes place, all sync functions will return an error.  Once this
condition is reached, the filesystem will not recover, and every subsequent sync
call will return an error, even if the upperdir has not experience a new error
since the last sync call.

When overlay is mounted with "volatile" option, the directory
"$workdir/work/incompat/volatile" is created.  During next mount, overlay
checks for this directory and refuses to mount if present. This is a strong
indicator that user should throw away upper and work directories and create
fresh one. In very limited cases where the user knows that the system has
not crashed and contents of upperdir are intact, The "volatile" directory
can be removed.

User xattr
----------

The "-o userxattr" mount option forces overlayfs to use the
"user.overlay." xattr namespace instead of "trusted.overlay.".  This is
useful for unprivileged mounting of overlayfs.

Testsuite
---------

There's a testsuite originally developed by David Howells and currently
maintained by Amir Goldstein at:

  https://github.com/amir73il/unionmount-testsuite.git

Run as root:

  # cd unionmount-testsuite
  # ./run --ov --verify


code:
===============
Pathname lookup
===============

This write-up is based on three articles published at lwn.net:

- <https://lwn.net/Articles/649115/> Pathname lookup in Linux
- <https://lwn.net/Articles/649729/> RCU-walk: faster pathname lookup in Linux
- <https://lwn.net/Articles/650786/> A walk among the symlinks

Written by Neil Brown with help from Al Viro and Jon Corbet.
It has subsequently been updated to reflect changes in the kernel
including:

- per-directory parallel name lookup.
- ``openat2()`` resolution restriction flags.

Introduction to pathname lookup
===============================

The most obvious aspect of pathname lookup, which very little
exploration is needed to discover, is that it is complex.  There are
many rules, special cases, and implementation alternatives that all
combine to confuse the unwary reader.  Computer science has long been
acquainted with such complexity and has tools to help manage it.  One
tool that we will make extensive use of is "divide and conquer".  For
the early parts of the analysis we will divide off symlinks - leaving
them until the final part.  Well before we get to symlinks we have
another major division based on the VFS's approach to locking which
will allow us to review "REF-walk" and "RCU-walk" separately.  But we
are getting ahead of ourselves.  There are some important low level
distinctions we need to clarify first.

There are two sorts of ...
--------------------------

.. _openat: http://man7.org/linux/man-pages/man2/openat.2.html

Pathnames (sometimes "file names"), used to identify objects in the
filesystem, will be familiar to most readers.  They contain two sorts
of elements: "slashes" that are sequences of one or more "``/``"
characters, and "components" that are sequences of one or more
non-"``/``" characters.  These form two kinds of paths.  Those that
start with slashes are "absolute" and start from the filesystem root.
The others are "relative" and start from the current directory, or
from some other location specified by a file descriptor given to
"``*at()``" system calls such as `openat() <openat_>`_.

.. _execveat: http://man7.org/linux/man-pages/man2/execveat.2.html

It is tempting to describe the second kind as starting with a
component, but that isn't always accurate: a pathname can lack both
slashes and components, it can be empty, in other words.  This is
generally forbidden in POSIX, but some of those "``*at()``" system calls
in Linux permit it when the ``AT_EMPTY_PATH`` flag is given.  For
example, if you have an open file descriptor on an executable file you
can execute it by calling `execveat() <execveat_>`_ passing
the file descriptor, an empty path, and the ``AT_EMPTY_PATH`` flag.

These paths can be divided into two sections: the final component and
everything else.  The "everything else" is the easy bit.  In all cases
it must identify a directory that already exists, otherwise an error
such as ``ENOENT`` or ``ENOTDIR`` will be reported.

The final component is not so simple.  Not only do different system
calls interpret it quite differently (e.g. some create it, some do
not), but it might not even exist: neither the empty pathname nor the
pathname that is just slashes have a final component.  If it does
exist, it could be "``.``" or "``..``" which are handled quite differently
from other components.

.. _POSIX: https://pubs.opengroup.org/onlinepubs/9699919799/basedefs/V1_chap04.html#tag_04_12

If a pathname ends with a slash, such as "``/tmp/foo/``" it might be
tempting to consider that to have an empty final component.  In many
ways that would lead to correct results, but not always.  In
particular, ``mkdir()`` and ``rmdir()`` each create or remove a directory named
by the final component, and they are required to work with pathnames
ending in "``/``".  According to POSIX_:

  A pathname that contains at least one non-<slash> character and
  that ends with one or more trailing <slash> characters shall not
  be resolved successfully unless the last pathname component before
  the trailing <slash> characters names an existing directory or a
  directory entry that is to be created for a directory immediately
  after the pathname is resolved.

The Linux pathname walking code (mostly in ``fs/namei.c``) deals with
all of these issues: breaking the path into components, handling the
"everything else" quite separately from the final component, and
checking that the trailing slash is not used where it isn't
permitted.  It also addresses the important issue of concurrent
access.

While one process is looking up a pathname, another might be making
changes that affect that lookup.  One fairly extreme case is that if
"a/b" were renamed to "a/c/b" while another process were looking up
"a/b/..", that process might successfully resolve on "a/c".
Most races are much more subtle, and a big part of the task of
pathname lookup is to prevent them from having damaging effects.  Many
of the possible races are seen most clearly in the context of the
"dcache" and an understanding of that is central to understanding
pathname lookup.

More than just a cache
----------------------

The "dcache" caches information about names in each filesystem to
make them quickly available for lookup.  Each entry (known as a
"dentry") contains three significant fields: a component name, a
pointer to a parent dentry, and a pointer to the "inode" which
contains further information about the object in that parent with
the given name.  The inode pointer can be ``NULL`` indicating that the
name doesn't exist in the parent.  While there can be linkage in the
dentry of a directory to the dentries of the children, that linkage is
not used for pathname lookup, and so will not be considered here.

The dcache has a number of uses apart from accelerating lookup.  One
that will be particularly relevant is that it is closely integrated
with the mount table that records which filesystem is mounted where.
What the mount table actually stores is which dentry is mounted on top
of which other dentry.

When considering the dcache, we have another of our "two types"
distinctions: there are two types of filesystems.

Some filesystems ensure that the information in the dcache is always
completely accurate (though not necessarily complete).  This can allow
the VFS to determine if a particular file does or doesn't exist
without checking with the filesystem, and means that the VFS can
protect the filesystem against certain races and other problems.
These are typically "local" filesystems such as ext3, XFS, and Btrfs.

Other filesystems don't provide that guarantee because they cannot.
These are typically filesystems that are shared across a network,
whether remote filesystems like NFS and 9P, or cluster filesystems
like ocfs2 or cephfs.  These filesystems allow the VFS to revalidate
cached information, and must provide their own protection against
awkward races.  The VFS can detect these filesystems by the
``DCACHE_OP_REVALIDATE`` flag being set in the dentry.

REF-walk: simple concurrency management with refcounts and spinlocks
--------------------------------------------------------------------

With all of those divisions carefully classified, we can now start
looking at the actual process of walking along a path.  In particular
we will start with the handling of the "everything else" part of a
pathname, and focus on the "REF-walk" approach to concurrency
management.  This code is found in the ``link_path_walk()`` function, if
you ignore all the places that only run when "``LOOKUP_RCU``"
(indicating the use of RCU-walk) is set.

.. _Meet the Lockers: https://lwn.net/Articles/453685/

REF-walk is fairly heavy-handed with locks and reference counts.  Not
as heavy-handed as in the old "big kernel lock" days, but certainly not
afraid of taking a lock when one is needed.  It uses a variety of
different concurrency controls.  A background understanding of the
various primitives is assumed, or can be gleaned from elsewhere such
as in `Meet the Lockers`_.

The locking mechanisms used by REF-walk include:

dentry->d_lockref
~~~~~~~~~~~~~~~~~

This uses the lockref primitive to provide both a spinlock and a
reference count.  The special-sauce of this primitive is that the
conceptual sequence "lock; inc_ref; unlock;" can often be performed
with a single atomic memory operation.

Holding a reference on a dentry ensures that the dentry won't suddenly
be freed and used for something else, so the values in various fields
will behave as expected.  It also protects the ``->d_inode`` reference
to the inode to some extent.

The association between a dentry and its inode is fairly permanent.
For example, when a file is renamed, the dentry and inode move
together to the new location.  When a file is created the dentry will
initially be negative (i.e. ``d_inode`` is ``NULL``), and will be assigned
to the new inode as part of the act of creation.

When a file is deleted, this can be reflected in the cache either by
setting ``d_inode`` to ``NULL``, or by removing it from the hash table
(described shortly) used to look up the name in the parent directory.
If the dentry is still in use the second option is used as it is
perfectly legal to keep using an open file after it has been deleted
and having the dentry around helps.  If the dentry is not otherwise in
use (i.e. if the refcount in ``d_lockref`` is one), only then will
``d_inode`` be set to ``NULL``.  Doing it this way is more efficient for a
very common case.

So as long as a counted reference is held to a dentry, a non-``NULL`` ``->d_inode``
value will never be changed.

dentry->d_lock
~~~~~~~~~~~~~~

``d_lock`` is a synonym for the spinlock that is part of ``d_lockref`` above.
For our purposes, holding this lock protects against the dentry being
renamed or unlinked.  In particular, its parent (``d_parent``), and its
name (``d_name``) cannot be changed, and it cannot be removed from the
dentry hash table.

When looking for a name in a directory, REF-walk takes ``d_lock`` on
each candidate dentry that it finds in the hash table and then checks
that the parent and name are correct.  So it doesn't lock the parent
while searching in the cache; it only locks children.

When looking for the parent for a given name (to handle "``..``"),
REF-walk can take ``d_lock`` to get a stable reference to ``d_parent``,
but it first tries a more lightweight approach.  As seen in
``dget_parent()``, if a reference can be claimed on the parent, and if
subsequently ``d_parent`` can be seen to have not changed, then there is
no need to actually take the lock on the child.

rename_lock
~~~~~~~~~~~

Looking up a given name in a given directory involves computing a hash
from the two values (the name and the dentry of the directory),
accessing that slot in a hash table, and searching the linked list
that is found there.

When a dentry is renamed, the name and the parent dentry can both
change so the hash will almost certainly change too.  This would move the
dentry to a different chain in the hash table.  If a filename search
happened to be looking at a dentry that was moved in this way,
it might end up continuing the search down the wrong chain,
and so miss out on part of the correct chain.

The name-lookup process (``d_lookup()``) does *not* try to prevent this
from happening, but only to detect when it happens.
``rename_lock`` is a seqlock that is updated whenever any dentry is
renamed.  If ``d_lookup`` finds that a rename happened while it
unsuccessfully scanned a chain in the hash table, it simply tries
again.

``rename_lock`` is also used to detect and defend against potential attacks
against ``LOOKUP_BENEATH`` and ``LOOKUP_IN_ROOT`` when resolving ".." (where
the parent directory is moved outside the root, bypassing the ``path_equal()``
check). If ``rename_lock`` is updated during the lookup and the path encounters
a "..", a potential attack occurred and ``handle_dots()`` will bail out with
``-EAGAIN``.

inode->i_rwsem
~~~~~~~~~~~~~~

``i_rwsem`` is a read/write semaphore that serializes all changes to a particular
directory.  This ensures that, for example, an ``unlink()`` and a ``rename()``
cannot both happen at the same time.  It also keeps the directory
stable while the filesystem is asked to look up a name that is not
currently in the dcache or, optionally, when the list of entries in a
directory is being retrieved with ``readdir()``.

This has a complementary role to that of ``d_lock``: ``i_rwsem`` on a
directory protects all of the names in that directory, while ``d_lock``
on a name protects just one name in a directory.  Most changes to the
dcache hold ``i_rwsem`` on the relevant directory inode and briefly take
``d_lock`` on one or more the dentries while the change happens.  One
exception is when idle dentries are removed from the dcache due to
memory pressure.  This uses ``d_lock``, but ``i_rwsem`` plays no role.

The semaphore affects pathname lookup in two distinct ways.  Firstly it
prevents changes during lookup of a name in a directory.  ``walk_component()`` uses
``lookup_fast()`` first which, in turn, checks to see if the name is in the cache,
using only ``d_lock`` locking.  If the name isn't found, then ``walk_component()``
falls back to ``lookup_slow()`` which takes a shared lock on ``i_rwsem``, checks again that
the name isn't in the cache, and then calls in to the filesystem to get a
definitive answer.  A new dentry will be added to the cache regardless of
the result.

Secondly, when pathname lookup reaches the final component, it will
sometimes need to take an exclusive lock on ``i_rwsem`` before performing the last lookup so
that the required exclusion can be achieved.  How path lookup chooses
to take, or not take, ``i_rwsem`` is one of the
issues addressed in a subsequent section.

If two threads attempt to look up the same name at the same time - a
name that is not yet in the dcache - the shared lock on ``i_rwsem`` will
not prevent them both adding new dentries with the same name.  As this
would result in confusion an extra level of interlocking is used,
based around a secondary hash table (``in_lookup_hashtable``) and a
per-dentry flag bit (``DCACHE_PAR_LOOKUP``).

To add a new dentry to the cache while only holding a shared lock on
``i_rwsem``, a thread must call ``d_alloc_parallel()``.  This allocates a
dentry, stores the required name and parent in it, checks if there
is already a matching dentry in the primary or secondary hash
tables, and if not, stores the newly allocated dentry in the secondary
hash table, with ``DCACHE_PAR_LOOKUP`` set.

If a matching dentry was found in the primary hash table then that is
returned and the caller can know that it lost a race with some other
thread adding the entry.  If no matching dentry is found in either
cache, the newly allocated dentry is returned and the caller can
detect this from the presence of ``DCACHE_PAR_LOOKUP``.  In this case it
knows that it has won any race and now is responsible for asking the
filesystem to perform the lookup and find the matching inode.  When
the lookup is complete, it must call ``d_lookup_done()`` which clears
the flag and does some other house keeping, including removing the
dentry from the secondary hash table - it will normally have been
added to the primary hash table already.  Note that a ``struct
waitqueue_head`` is passed to ``d_alloc_parallel()``, and
``d_lookup_done()`` must be called while this ``waitqueue_head`` is still
in scope.

If a matching dentry is found in the secondary hash table,
``d_alloc_parallel()`` has a little more work to do. It first waits for
``DCACHE_PAR_LOOKUP`` to be cleared, using a wait_queue that was passed
to the instance of ``d_alloc_parallel()`` that won the race and that
will be woken by the call to ``d_lookup_done()``.  It then checks to see
if the dentry has now been added to the primary hash table.  If it
has, the dentry is returned and the caller just sees that it lost any
race.  If it hasn't been added to the primary hash table, the most
likely explanation is that some other dentry was added instead using
``d_splice_alias()``.  In any case, ``d_alloc_parallel()`` repeats all the
look ups from the start and will normally return something from the
primary hash table.

mnt->mnt_count
~~~~~~~~~~~~~~

``mnt_count`` is a per-CPU reference counter on "``mount``" structures.
Per-CPU here means that incrementing the count is cheap as it only
uses CPU-local memory, but checking if the count is zero is expensive as
it needs to check with every CPU.  Taking a ``mnt_count`` reference
prevents the mount structure from disappearing as the result of regular
unmount operations, but does not prevent a "lazy" unmount.  So holding
``mnt_count`` doesn't ensure that the mount remains in the namespace and,
in particular, doesn't stabilize the link to the mounted-on dentry.  It
does, however, ensure that the ``mount`` data structure remains coherent,
and it provides a reference to the root dentry of the mounted
filesystem.  So a reference through ``->mnt_count`` provides a stable
reference to the mounted dentry, but not the mounted-on dentry.

mount_lock
~~~~~~~~~~

``mount_lock`` is a global seqlock, a bit like ``rename_lock``.  It can be used to
check if any change has been made to any mount points.

While walking down the tree (away from the root) this lock is used when
crossing a mount point to check that the crossing was safe.  That is,
the value in the seqlock is read, then the code finds the mount that
is mounted on the current directory, if there is one, and increments
the ``mnt_count``.  Finally the value in ``mount_lock`` is checked against
the old value.  If there is no change, then the crossing was safe.  If there
was a change, the ``mnt_count`` is decremented and the whole process is
retried.

When walking up the tree (towards the root) by following a ".." link,
a little more care is needed.  In this case the seqlock (which
contains both a counter and a spinlock) is fully locked to prevent
any changes to any mount points while stepping up.  This locking is
needed to stabilize the link to the mounted-on dentry, which the
refcount on the mount itself doesn't ensure.

``mount_lock`` is also used to detect and defend against potential attacks
against ``LOOKUP_BENEATH`` and ``LOOKUP_IN_ROOT`` when resolving ".." (where
the parent directory is moved outside the root, bypassing the ``path_equal()``
check). If ``mount_lock`` is updated during the lookup and the path encounters
a "..", a potential attack occurred and ``handle_dots()`` will bail out with
``-EAGAIN``.

RCU
~~~

Finally the global (but extremely lightweight) RCU read lock is held
from time to time to ensure certain data structures don't get freed
unexpectedly.

In particular it is held while scanning chains in the dcache hash
table, and the mount point hash table.

Bringing it together with ``struct nameidata``
----------------------------------------------

.. _First edition Unix: https://minnie.tuhs.org/cgi-bin/utree.pl?file=V1/u2.s

Throughout the process of walking a path, the current status is stored
in a ``struct nameidata``, "namei" being the traditional name - dating
all the way back to `First Edition Unix`_ - of the function that
converts a "name" to an "inode".  ``struct nameidata`` contains (among
other fields):

``struct path path``
~~~~~~~~~~~~~~~~~~~~

A ``path`` contains a ``struct vfsmount`` (which is
embedded in a ``struct mount``) and a ``struct dentry``.  Together these
record the current status of the walk.  They start out referring to the
starting point (the current working directory, the root directory, or some other
directory identified by a file descriptor), and are updated on each
step.  A reference through ``d_lockref`` and ``mnt_count`` is always
held.

``struct qstr last``
~~~~~~~~~~~~~~~~~~~~

This is a string together with a length (i.e. *not* ``nul`` terminated)
that is the "next" component in the pathname.

``int last_type``
~~~~~~~~~~~~~~~~~

This is one of ``LAST_NORM``, ``LAST_ROOT``, ``LAST_DOT`` or ``LAST_DOTDOT``.
The ``last`` field is only valid if the type is ``LAST_NORM``.

``struct path root``
~~~~~~~~~~~~~~~~~~~~

This is used to hold a reference to the effective root of the
filesystem.  Often that reference won't be needed, so this field is
only assigned the first time it is used, or when a non-standard root
is requested.  Keeping a reference in the ``nameidata`` ensures that
only one root is in effect for the entire path walk, even if it races
with a ``chroot()`` system call.

It should be noted that in the case of ``LOOKUP_IN_ROOT`` or
``LOOKUP_BENEATH``, the effective root becomes the directory file descriptor
passed to ``openat2()`` (which exposes these ``LOOKUP_`` flags).

The root is needed when either of two conditions holds: (1) either the
pathname or a symbolic link starts with a "'/'", or (2) a "``..``"
component is being handled, since "``..``" from the root must always stay
at the root.  The value used is usually the current root directory of
the calling process.  An alternate root can be provided as when
``sysctl()`` calls ``file_open_root()``, and when NFSv4 or Btrfs call
``mount_subtree()``.  In each case a pathname is being looked up in a very
specific part of the filesystem, and the lookup must not be allowed to
escape that subtree.  It works a bit like a local ``chroot()``.

Ignoring the handling of symbolic links, we can now describe the
"``link_path_walk()``" function, which handles the lookup of everything
except the final component as:

   Given a path (``name``) and a nameidata structure (``nd``), check that the
   current directory has execute permission and then advance ``name``
   over one component while updating ``last_type`` and ``last``.  If that
   was the final component, then return, otherwise call
   ``walk_component()`` and repeat from the top.

``walk_component()`` is even easier.  If the component is ``LAST_DOTS``,
it calls ``handle_dots()`` which does the necessary locking as already
described.  If it finds a ``LAST_NORM`` component it first calls
"``lookup_fast()``" which only looks in the dcache, but will ask the
filesystem to revalidate the result if it is that sort of filesystem.
If that doesn't get a good result, it calls "``lookup_slow()``" which
takes ``i_rwsem``, rechecks the cache, and then asks the filesystem
to find a definitive answer.

As the last step of walk_component(), step_into() will be called either
directly from walk_component() or from handle_dots().  It calls
handle_mounts(), to check and handle mount points, in which a new
``struct path`` is created containing a counted reference to the new dentry and
a reference to the new ``vfsmount`` which is only counted if it is
different from the previous ``vfsmount``. Then if there is
a symbolic link, step_into() calls pick_link() to deal with it,
otherwise it installs the new ``struct path`` in the ``struct nameidata``, and
drops the unneeded references.

This "hand-over-hand" sequencing of getting a reference to the new
dentry before dropping the reference to the previous dentry may
seem obvious, but is worth pointing out so that we will recognize its
analogue in the "RCU-walk" version.

Handling the final component
----------------------------

``link_path_walk()`` only walks as far as setting ``nd->last`` and
``nd->last_type`` to refer to the final component of the path.  It does
not call ``walk_component()`` that last time.  Handling that final
component remains for the caller to sort out. Those callers are
path_lookupat(), path_parentat() and
path_openat() each of which handles the differing requirements of
different system calls.

``path_parentat()`` is clearly the simplest - it just wraps a little bit
of housekeeping around ``link_path_walk()`` and returns the parent
directory and final component to the caller.  The caller will be either
aiming to create a name (via ``filename_create()``) or remove or rename
a name (in which case ``user_path_parent()`` is used).  They will use
``i_rwsem`` to exclude other changes while they validate and then
perform their operation.

``path_lookupat()`` is nearly as simple - it is used when an existing
object is wanted such as by ``stat()`` or ``chmod()``.  It essentially just
calls ``walk_component()`` on the final component through a call to
``lookup_last()``.  ``path_lookupat()`` returns just the final dentry.
It is worth noting that when flag ``LOOKUP_MOUNTPOINT`` is set,
path_lookupat() will unset LOOKUP_JUMPED in nameidata so that in the
subsequent path traversal d_weak_revalidate() won't be called.
This is important when unmounting a filesystem that is inaccessible, such as
one provided by a dead NFS server.

Finally ``path_openat()`` is used for the ``open()`` system call; it
contains, in support functions starting with "open_last_lookups()", all the
complexity needed to handle the different subtleties of O_CREAT (with
or without O_EXCL), final "``/``" characters, and trailing symbolic
links.  We will revisit this in the final part of this series, which
focuses on those symbolic links.  "open_last_lookups()" will sometimes, but
not always, take ``i_rwsem``, depending on what it finds.

Each of these, or the functions which call them, need to be alert to
the possibility that the final component is not ``LAST_NORM``.  If the
goal of the lookup is to create something, then any value for
``last_type`` other than ``LAST_NORM`` will result in an error.  For
example if ``path_parentat()`` reports ``LAST_DOTDOT``, then the caller
won't try to create that name.  They also check for trailing slashes
by testing ``last.name[last.len]``.  If there is any character beyond
the final component, it must be a trailing slash.

Revalidation and automounts
---------------------------

Apart from symbolic links, there are only two parts of the "REF-walk"
process not yet covered.  One is the handling of stale cache entries
and the other is automounts.

On filesystems that require it, the lookup routines will call the
``->d_revalidate()`` dentry method to ensure that the cached information
is current.  This will often confirm validity or update a few details
from a server.  In some cases it may find that there has been change
further up the path and that something that was thought to be valid
previously isn't really.  When this happens the lookup of the whole
path is aborted and retried with the "``LOOKUP_REVAL``" flag set.  This
forces revalidation to be more thorough.  We will see more details of
this retry process in the next article.

Automount points are locations in the filesystem where an attempt to
lookup a name can trigger changes to how that lookup should be
handled, in particular by mounting a filesystem there.  These are
covered in greater detail in autofs.txt in the Linux documentation
tree, but a few notes specifically related to path lookup are in order
here.

The Linux VFS has a concept of "managed" dentries.  There are three
potentially interesting things about these dentries corresponding
to three different flags that might be set in ``dentry->d_flags``:

``DCACHE_MANAGE_TRANSIT``
~~~~~~~~~~~~~~~~~~~~~~~~~

If this flag has been set, then the filesystem has requested that the
``d_manage()`` dentry operation be called before handling any possible
mount point.  This can perform two particular services:

It can block to avoid races.  If an automount point is being
unmounted, the ``d_manage()`` function will usually wait for that
process to complete before letting the new lookup proceed and possibly
trigger a new automount.

It can selectively allow only some processes to transit through a
mount point.  When a server process is managing automounts, it may
need to access a directory without triggering normal automount
processing.  That server process can identify itself to the ``autofs``
filesystem, which will then give it a special pass through
``d_manage()`` by returning ``-EISDIR``.

``DCACHE_MOUNTED``
~~~~~~~~~~~~~~~~~~

This flag is set on every dentry that is mounted on.  As Linux
supports multiple filesystem namespaces, it is possible that the
dentry may not be mounted on in *this* namespace, just in some
other.  So this flag is seen as a hint, not a promise.

If this flag is set, and ``d_manage()`` didn't return ``-EISDIR``,
``lookup_mnt()`` is called to examine the mount hash table (honoring the
``mount_lock`` described earlier) and possibly return a new ``vfsmount``
and a new ``dentry`` (both with counted references).

``DCACHE_NEED_AUTOMOUNT``
~~~~~~~~~~~~~~~~~~~~~~~~~

If ``d_manage()`` allowed us to get this far, and ``lookup_mnt()`` didn't
find a mount point, then this flag causes the ``d_automount()`` dentry
operation to be called.

The ``d_automount()`` operation can be arbitrarily complex and may
communicate with server processes etc. but it should ultimately either
report that there was an error, that there was nothing to mount, or
should provide an updated ``struct path`` with new ``dentry`` and ``vfsmount``.

In the latter case, ``finish_automount()`` will be called to safely
install the new mount point into the mount table.

There is no new locking of import here and it is important that no
locks (only counted references) are held over this processing due to
the very real possibility of extended delays.
This will become more important next time when we examine RCU-walk
which is particularly sensitive to delays.

RCU-walk - faster pathname lookup in Linux
==========================================

RCU-walk is another algorithm for performing pathname lookup in Linux.
It is in many ways similar to REF-walk and the two share quite a bit
of code.  The significant difference in RCU-walk is how it allows for
the possibility of concurrent access.

We noted that REF-walk is complex because there are numerous details
and special cases.  RCU-walk reduces this complexity by simply
refusing to handle a number of cases -- it instead falls back to
REF-walk.  The difficulty with RCU-walk comes from a different
direction: unfamiliarity.  The locking rules when depending on RCU are
quite different from traditional locking, so we will spend a little extra
time when we come to those.

Clear demarcation of roles
--------------------------

The easiest way to manage concurrency is to forcibly stop any other
thread from changing the data structures that a given thread is
looking at.  In cases where no other thread would even think of
changing the data and lots of different threads want to read at the
same time, this can be very costly.  Even when using locks that permit
multiple concurrent readers, the simple act of updating the count of
the number of current readers can impose an unwanted cost.  So the
goal when reading a shared data structure that no other process is
changing is to avoid writing anything to memory at all.  Take no
locks, increment no counts, leave no footprints.

The REF-walk mechanism already described certainly doesn't follow this
principle, but then it is really designed to work when there may well
be other threads modifying the data.  RCU-walk, in contrast, is
designed for the common situation where there are lots of frequent
readers and only occasional writers.  This may not be common in all
parts of the filesystem tree, but in many parts it will be.  For the
other parts it is important that RCU-walk can quickly fall back to
using REF-walk.

Pathname lookup always starts in RCU-walk mode but only remains there
as long as what it is looking for is in the cache and is stable.  It
dances lightly down the cached filesystem image, leaving no footprints
and carefully watching where it is, to be sure it doesn't trip.  If it
notices that something has changed or is changing, or if something
isn't in the cache, then it tries to stop gracefully and switch to
REF-walk.

This stopping requires getting a counted reference on the current
``vfsmount`` and ``dentry``, and ensuring that these are still valid -
that a path walk with REF-walk would have found the same entries.
This is an invariant that RCU-walk must guarantee.  It can only make
decisions, such as selecting the next step, that are decisions which
REF-walk could also have made if it were walking down the tree at the
same time.  If the graceful stop succeeds, the rest of the path is
processed with the reliable, if slightly sluggish, REF-walk.  If
RCU-walk finds it cannot stop gracefully, it simply gives up and
restarts from the top with REF-walk.

This pattern of "try RCU-walk, if that fails try REF-walk" can be
clearly seen in functions like filename_lookup(),
filename_parentat(),
do_filp_open(), and do_file_open_root().  These four
correspond roughly to the three ``path_*()`` functions we met earlier,
each of which calls ``link_path_walk()``.  The ``path_*()`` functions are
called using different mode flags until a mode is found which works.
They are first called with ``LOOKUP_RCU`` set to request "RCU-walk".  If
that fails with the error ``ECHILD`` they are called again with no
special flag to request "REF-walk".  If either of those report the
error ``ESTALE`` a final attempt is made with ``LOOKUP_REVAL`` set (and no
``LOOKUP_RCU``) to ensure that entries found in the cache are forcibly
revalidated - normally entries are only revalidated if the filesystem
determines that they are too old to trust.

The ``LOOKUP_RCU`` attempt may drop that flag internally and switch to
REF-walk, but will never then try to switch back to RCU-walk.  Places
that trip up RCU-walk are much more likely to be near the leaves and
so it is very unlikely that there will be much, if any, benefit from
switching back.

RCU and seqlocks: fast and light
--------------------------------

RCU is, unsurprisingly, critical to RCU-walk mode.  The
``rcu_read_lock()`` is held for the entire time that RCU-walk is walking
down a path.  The particular guarantee it provides is that the key
data structures - dentries, inodes, super_blocks, and mounts - will
not be freed while the lock is held.  They might be unlinked or
invalidated in one way or another, but the memory will not be
repurposed so values in various fields will still be meaningful.  This
is the only guarantee that RCU provides; everything else is done using
seqlocks.

As we saw above, REF-walk holds a counted reference to the current
dentry and the current vfsmount, and does not release those references
before taking references to the "next" dentry or vfsmount.  It also
sometimes takes the ``d_lock`` spinlock.  These references and locks are
taken to prevent certain changes from happening.  RCU-walk must not
take those references or locks and so cannot prevent such changes.
Instead, it checks to see if a change has been made, and aborts or
retries if it has.

To preserve the invariant mentioned above (that RCU-walk may only make
decisions that REF-walk could have made), it must make the checks at
or near the same places that REF-walk holds the references.  So, when
REF-walk increments a reference count or takes a spinlock, RCU-walk
samples the status of a seqlock using ``read_seqcount_begin()`` or a
similar function.  When REF-walk decrements the count or drops the
lock, RCU-walk checks if the sampled status is still valid using
``read_seqcount_retry()`` or similar.

However, there is a little bit more to seqlocks than that.  If
RCU-walk accesses two different fields in a seqlock-protected
structure, or accesses the same field twice, there is no a priori
guarantee of any consistency between those accesses.  When consistency
is needed - which it usually is - RCU-walk must take a copy and then
use ``read_seqcount_retry()`` to validate that copy.

``read_seqcount_retry()`` not only checks the sequence number, but also
imposes a memory barrier so that no memory-read instruction from
*before* the call can be delayed until *after* the call, either by the
CPU or by the compiler.  A simple example of this can be seen in
``slow_dentry_cmp()`` which, for filesystems which do not use simple
byte-wise name equality, calls into the filesystem to compare a name
against a dentry.  The length and name pointer are copied into local
variables, then ``read_seqcount_retry()`` is called to confirm the two
are consistent, and only then is ``->d_compare()`` called.  When
standard filename comparison is used, ``dentry_cmp()`` is called
instead.  Notably it does *not* use ``read_seqcount_retry()``, but
instead has a large comment explaining why the consistency guarantee
isn't necessary.  A subsequent ``read_seqcount_retry()`` will be
sufficient to catch any problem that could occur at this point.

With that little refresher on seqlocks out of the way we can look at
the bigger picture of how RCU-walk uses seqlocks.

``mount_lock`` and ``nd->m_seq``
~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~

We already met the ``mount_lock`` seqlock when REF-walk used it to
ensure that crossing a mount point is performed safely.  RCU-walk uses
it for that too, but for quite a bit more.

Instead of taking a counted reference to each ``vfsmount`` as it
descends the tree, RCU-walk samples the state of ``mount_lock`` at the
start of the walk and stores this initial sequence number in the
``struct nameidata`` in the ``m_seq`` field.  This one lock and one
sequence number are used to validate all accesses to all ``vfsmounts``,
and all mount point crossings.  As changes to the mount table are
relatively rare, it is reasonable to fall back on REF-walk any time
that any "mount" or "unmount" happens.

``m_seq`` is checked (using ``read_seqretry()``) at the end of an RCU-walk
sequence, whether switching to REF-walk for the rest of the path or
when the end of the path is reached.  It is also checked when stepping
down over a mount point (in ``__follow_mount_rcu()``) or up (in
``follow_dotdot_rcu()``).  If it is ever found to have changed, the
whole RCU-walk sequence is aborted and the path is processed again by
REF-walk.

If RCU-walk finds that ``mount_lock`` hasn't changed then it can be sure
that, had REF-walk taken counted references on each vfsmount, the
results would have been the same.  This ensures the invariant holds,
at least for vfsmount structures.

``dentry->d_seq`` and ``nd->seq``
~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~

In place of taking a count or lock on ``d_reflock``, RCU-walk samples
the per-dentry ``d_seq`` seqlock, and stores the sequence number in the
``seq`` field of the nameidata structure, so ``nd->seq`` should always be
the current sequence number of ``nd->dentry``.  This number needs to be
revalidated after copying, and before using, the name, parent, or
inode of the dentry.

The handling of the name we have already looked at, and the parent is
only accessed in ``follow_dotdot_rcu()`` which fairly trivially follows
the required pattern, though it does so for three different cases.

When not at a mount point, ``d_parent`` is followed and its ``d_seq`` is
collected.  When we are at a mount point, we instead follow the
``mnt->mnt_mountpoint`` link to get a new dentry and collect its
``d_seq``.  Then, after finally finding a ``d_parent`` to follow, we must
check if we have landed on a mount point and, if so, must find that
mount point and follow the ``mnt->mnt_root`` link.  This would imply a
somewhat unusual, but certainly possible, circumstance where the
starting point of the path lookup was in part of the filesystem that
was mounted on, and so not visible from the root.

The inode pointer, stored in ``->d_inode``, is a little more
interesting.  The inode will always need to be accessed at least
twice, once to determine if it is NULL and once to verify access
permissions.  Symlink handling requires a validated inode pointer too.
Rather than revalidating on each access, a copy is made on the first
access and it is stored in the ``inode`` field of ``nameidata`` from where
it can be safely accessed without further validation.

``lookup_fast()`` is the only lookup routine that is used in RCU-mode,
``lookup_slow()`` being too slow and requiring locks.  It is in
``lookup_fast()`` that we find the important "hand over hand" tracking
of the current dentry.

The current ``dentry`` and current ``seq`` number are passed to
``__d_lookup_rcu()`` which, on success, returns a new ``dentry`` and a
new ``seq`` number.  ``lookup_fast()`` then copies the inode pointer and
revalidates the new ``seq`` number.  It then validates the old ``dentry``
with the old ``seq`` number one last time and only then continues.  This
process of getting the ``seq`` number of the new dentry and then
checking the ``seq`` number of the old exactly mirrors the process of
getting a counted reference to the new dentry before dropping that for
the old dentry which we saw in REF-walk.

No ``inode->i_rwsem`` or even ``rename_lock``
~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~

A semaphore is a fairly heavyweight lock that can only be taken when it is
permissible to sleep.  As ``rcu_read_lock()`` forbids sleeping,
``inode->i_rwsem`` plays no role in RCU-walk.  If some other thread does
take ``i_rwsem`` and modifies the directory in a way that RCU-walk needs
to notice, the result will be either that RCU-walk fails to find the
dentry that it is looking for, or it will find a dentry which
``read_seqretry()`` won't validate.  In either case it will drop down to
REF-walk mode which can take whatever locks are needed.

Though ``rename_lock`` could be used by RCU-walk as it doesn't require
any sleeping, RCU-walk doesn't bother.  REF-walk uses ``rename_lock`` to
protect against the possibility of hash chains in the dcache changing
while they are being searched.  This can result in failing to find
something that actually is there.  When RCU-walk fails to find
something in the dentry cache, whether it is really there or not, it
already drops down to REF-walk and tries again with appropriate
locking.  This neatly handles all cases, so adding extra checks on
rename_lock would bring no significant value.

``unlazy walk()`` and ``complete_walk()``
-----------------------------------------

That "dropping down to REF-walk" typically involves a call to
``unlazy_walk()``, so named because "RCU-walk" is also sometimes
referred to as "lazy walk".  ``unlazy_walk()`` is called when
following the path down to the current vfsmount/dentry pair seems to
have proceeded successfully, but the next step is problematic.  This
can happen if the next name cannot be found in the dcache, if
permission checking or name revalidation couldn't be achieved while
the ``rcu_read_lock()`` is held (which forbids sleeping), if an
automount point is found, or in a couple of cases involving symlinks.
It is also called from ``complete_walk()`` when the lookup has reached
the final component, or the very end of the path, depending on which
particular flavor of lookup is used.

Other reasons for dropping out of RCU-walk that do not trigger a call
to ``unlazy_walk()`` are when some inconsistency is found that cannot be
handled immediately, such as ``mount_lock`` or one of the ``d_seq``
seqlocks reporting a change.  In these cases the relevant function
will return ``-ECHILD`` which will percolate up until it triggers a new
attempt from the top using REF-walk.

For those cases where ``unlazy_walk()`` is an option, it essentially
takes a reference on each of the pointers that it holds (vfsmount,
dentry, and possibly some symbolic links) and then verifies that the
relevant seqlocks have not been changed.  If there have been changes,
it, too, aborts with ``-ECHILD``, otherwise the transition to REF-walk
has been a success and the lookup process continues.

Taking a reference on those pointers is not quite as simple as just
incrementing a counter.  That works to take a second reference if you
already have one (often indirectly through another object), but it
isn't sufficient if you don't actually have a counted reference at
all.  For ``dentry->d_lockref``, it is safe to increment the reference
counter to get a reference unless it has been explicitly marked as
"dead" which involves setting the counter to ``-128``.
``lockref_get_not_dead()`` achieves this.

For ``mnt->mnt_count`` it is safe to take a reference as long as
``mount_lock`` is then used to validate the reference.  If that
validation fails, it may *not* be safe to just drop that reference in
the standard way of calling ``mnt_put()`` - an unmount may have
progressed too far.  So the code in ``legitimize_mnt()``, when it
finds that the reference it got might not be safe, checks the
``MNT_SYNC_UMOUNT`` flag to determine if a simple ``mnt_put()`` is
correct, or if it should just decrement the count and pretend none of
this ever happened.

Taking care in filesystems
--------------------------

RCU-walk depends almost entirely on cached information and often will
not call into the filesystem at all.  However there are two places,
besides the already-mentioned component-name comparison, where the
file system might be included in RCU-walk, and it must know to be
careful.

If the filesystem has non-standard permission-checking requirements -
such as a networked filesystem which may need to check with the server
- the ``i_op->permission`` interface might be called during RCU-walk.
In this case an extra "``MAY_NOT_BLOCK``" flag is passed so that it
knows not to sleep, but to return ``-ECHILD`` if it cannot complete
promptly.  ``i_op->permission`` is given the inode pointer, not the
dentry, so it doesn't need to worry about further consistency checks.
However if it accesses any other filesystem data structures, it must
ensure they are safe to be accessed with only the ``rcu_read_lock()``
held.  This typically means they must be freed using ``kfree_rcu()`` or
similar.

.. _READ_ONCE: https://lwn.net/Articles/624126/

If the filesystem may need to revalidate dcache entries, then
``d_op->d_revalidate`` may be called in RCU-walk too.  This interface
*is* passed the dentry but does not have access to the ``inode`` or the
``seq`` number from the ``nameidata``, so it needs to be extra careful
when accessing fields in the dentry.  This "extra care" typically
involves using  `READ_ONCE() <READ_ONCE_>`_ to access fields, and verifying the
result is not NULL before using it.  This pattern can be seen in
``nfs_lookup_revalidate()``.

A pair of patterns
------------------

In various places in the details of REF-walk and RCU-walk, and also in
the big picture, there are a couple of related patterns that are worth
being aware of.

The first is "try quickly and check, if that fails try slowly".  We
can see that in the high-level approach of first trying RCU-walk and
then trying REF-walk, and in places where ``unlazy_walk()`` is used to
switch to REF-walk for the rest of the path.  We also saw it earlier
in ``dget_parent()`` when following a "``..``" link.  It tries a quick way
to get a reference, then falls back to taking locks if needed.

The second pattern is "try quickly and check, if that fails try
again - repeatedly".  This is seen with the use of ``rename_lock`` and
``mount_lock`` in REF-walk.  RCU-walk doesn't make use of this pattern -
if anything goes wrong it is much safer to just abort and try a more
sedate approach.

The emphasis here is "try quickly and check".  It should probably be
"try quickly *and carefully*, then check".  The fact that checking is
needed is a reminder that the system is dynamic and only a limited
number of things are safe at all.  The most likely cause of errors in
this whole process is assuming something is safe when in reality it
isn't.  Careful consideration of what exactly guarantees the safety of
each access is sometimes necessary.

A walk among the symlinks
=========================

There are several basic issues that we will examine to understand the
handling of symbolic links:  the symlink stack, together with cache
lifetimes, will help us understand the overall recursive handling of
symlinks and lead to the special care needed for the final component.
Then a consideration of access-time updates and summary of the various
flags controlling lookup will finish the story.

The symlink stack
-----------------

There are only two sorts of filesystem objects that can usefully
appear in a path prior to the final component: directories and symlinks.
Handling directories is quite straightforward: the new directory
simply becomes the starting point at which to interpret the next
component on the path.  Handling symbolic links requires a bit more
work.

Conceptually, symbolic links could be handled by editing the path.  If
a component name refers to a symbolic link, then that component is
replaced by the body of the link and, if that body starts with a '/',
then all preceding parts of the path are discarded.  This is what the
"``readlink -f``" command does, though it also edits out "``.``" and
"``..``" components.

Directly editing the path string is not really necessary when looking
up a path, and discarding early components is pointless as they aren't
looked at anyway.  Keeping track of all remaining components is
important, but they can of course be kept separately; there is no need
to concatenate them.  As one symlink may easily refer to another,
which in turn can refer to a third, we may need to keep the remaining
components of several paths, each to be processed when the preceding
ones are completed.  These path remnants are kept on a stack of
limited size.

There are two reasons for placing limits on how many symlinks can
occur in a single path lookup.  The most obvious is to avoid loops.
If a symlink referred to itself either directly or through
intermediaries, then following the symlink can never complete
successfully - the error ``ELOOP`` must be returned.  Loops can be
detected without imposing limits, but limits are the simplest solution
and, given the second reason for restriction, quite sufficient.

.. _outlined recently: http://thread.gmane.org/gmane.linux.kernel/1934390/focus=1934550

The second reason was `outlined recently`_ by Linus:

   Because it's a latency and DoS issue too. We need to react well to
   true loops, but also to "very deep" non-loops. It's not about memory
   use, it's about users triggering unreasonable CPU resources.

Linux imposes a limit on the length of any pathname: ``PATH_MAX``, which
is 4096.  There are a number of reasons for this limit; not letting the
kernel spend too much time on just one path is one of them.  With
symbolic links you can effectively generate much longer paths so some
sort of limit is needed for the same reason.  Linux imposes a limit of
at most 40 (MAXSYMLINKS) symlinks in any one path lookup.  It previously imposed
a further limit of eight on the maximum depth of recursion, but that was
raised to 40 when a separate stack was implemented, so there is now
just the one limit.

The ``nameidata`` structure that we met in an earlier article contains a
small stack that can be used to store the remaining part of up to two
symlinks.  In many cases this will be sufficient.  If it isn't, a
separate stack is allocated with room for 40 symlinks.  Pathname
lookup will never exceed that stack as, once the 40th symlink is
detected, an error is returned.

It might seem that the name remnants are all that needs to be stored on
this stack, but we need a bit more.  To see that, we need to move on to
cache lifetimes.

Storage and lifetime of cached symlinks
---------------------------------------

Like other filesystem resources, such as inodes and directory
entries, symlinks are cached by Linux to avoid repeated costly access
to external storage.  It is particularly important for RCU-walk to be
able to find and temporarily hold onto these cached entries, so that
it doesn't need to drop down into REF-walk.

.. _object-oriented design pattern: https://lwn.net/Articles/446317/

While each filesystem is free to make its own choice, symlinks are
typically stored in one of two places.  Short symlinks are often
stored directly in the inode.  When a filesystem allocates a ``struct
inode`` it typically allocates extra space to store private data (a
common `object-oriented design pattern`_ in the kernel).  This will
sometimes include space for a symlink.  The other common location is
in the page cache, which normally stores the content of files.  The
pathname in a symlink can be seen as the content of that symlink and
can easily be stored in the page cache just like file content.

When neither of these is suitable, the next most likely scenario is
that the filesystem will allocate some temporary memory and copy or
construct the symlink content into that memory whenever it is needed.

When the symlink is stored in the inode, it has the same lifetime as
the inode which, itself, is protected by RCU or by a counted reference
on the dentry.  This means that the mechanisms that pathname lookup
uses to access the dcache and icache (inode cache) safely are quite
sufficient for accessing some cached symlinks safely.  In these cases,
the ``i_link`` pointer in the inode is set to point to wherever the
symlink is stored and it can be accessed directly whenever needed.

When the symlink is stored in the page cache or elsewhere, the
situation is not so straightforward.  A reference on a dentry or even
on an inode does not imply any reference on cached pages of that
inode, and even an ``rcu_read_lock()`` is not sufficient to ensure that
a page will not disappear.  So for these symlinks the pathname lookup
code needs to ask the filesystem to provide a stable reference and,
significantly, needs to release that reference when it is finished
with it.

Taking a reference to a cache page is often possible even in RCU-walk
mode.  It does require making changes to memory, which is best avoided,
but that isn't necessarily a big cost and it is better than dropping
out of RCU-walk mode completely.  Even filesystems that allocate
space to copy the symlink into can use ``GFP_ATOMIC`` to often successfully
allocate memory without the need to drop out of RCU-walk.  If a
filesystem cannot successfully get a reference in RCU-walk mode, it
must return ``-ECHILD`` and ``unlazy_walk()`` will be called to return to
REF-walk mode in which the filesystem is allowed to sleep.

The place for all this to happen is the ``i_op->get_link()`` inode
method. This is called both in RCU-walk and REF-walk. In RCU-walk the
``dentry*`` argument is NULL, ``->get_link()`` can return -ECHILD to drop out of
RCU-walk.  Much like the ``i_op->permission()`` method we
looked at previously, ``->get_link()`` would need to be careful that
all the data structures it references are safe to be accessed while
holding no counted reference, only the RCU lock. A callback
``struct delayed_called`` will be passed to ``->get_link()``:
file systems can set their own put_link function and argument through
set_delayed_call(). Later on, when VFS wants to put link, it will call
do_delayed_call() to invoke that callback function with the argument.

In order for the reference to each symlink to be dropped when the walk completes,
whether in RCU-walk or REF-walk, the symlink stack needs to contain,
along with the path remnants:

- the ``struct path`` to provide a reference to the previous path
- the ``const char *`` to provide a reference to the to previous name
- the ``seq`` to allow the path to be safely switched from RCU-walk to REF-walk
- the ``struct delayed_call`` for later invocation.

This means that each entry in the symlink stack needs to hold five
pointers and an integer instead of just one pointer (the path
remnant).  On a 64-bit system, this is about 40 bytes per entry;
with 40 entries it adds up to 1600 bytes total, which is less than
half a page.  So it might seem like a lot, but is by no means
excessive.

Note that, in a given stack frame, the path remnant (``name``) is not
part of the symlink that the other fields refer to.  It is the remnant
to be followed once that symlink has been fully parsed.

Following the symlink
---------------------

The main loop in ``link_path_walk()`` iterates seamlessly over all
components in the path and all of the non-final symlinks.  As symlinks
are processed, the ``name`` pointer is adjusted to point to a new
symlink, or is restored from the stack, so that much of the loop
doesn't need to notice.  Getting this ``name`` variable on and off the
stack is very straightforward; pushing and popping the references is
a little more complex.

When a symlink is found, walk_component() calls pick_link() via step_into()
which returns the link from the filesystem.
Providing that operation is successful, the old path ``name`` is placed on the
stack, and the new value is used as the ``name`` for a while.  When the end of
the path is found (i.e. ``*name`` is ``'\0'``) the old ``name`` is restored
off the stack and path walking continues.

Pushing and popping the reference pointers (inode, cookie, etc.) is more
complex in part because of the desire to handle tail recursion.  When
the last component of a symlink itself points to a symlink, we
want to pop the symlink-just-completed off the stack before pushing
the symlink-just-found to avoid leaving empty path remnants that would
just get in the way.

It is most convenient to push the new symlink references onto the
stack in ``walk_component()`` immediately when the symlink is found;
``walk_component()`` is also the last piece of code that needs to look at the
old symlink as it walks that last component.  So it is quite
convenient for ``walk_component()`` to release the old symlink and pop
the references just before pushing the reference information for the
new symlink.  It is guided in this by three flags: ``WALK_NOFOLLOW`` which
forbids it from following a symlink if it finds one, ``WALK_MORE``
which indicates that it is yet too early to release the
current symlink, and ``WALK_TRAILING`` which indicates that it is on the final
component of the lookup, so we will check userspace flag ``LOOKUP_FOLLOW`` to
decide whether follow it when it is a symlink and call ``may_follow_link()`` to
check if we have privilege to follow it.

Symlinks with no final component
~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~

A pair of special-case symlinks deserve a little further explanation.
Both result in a new ``struct path`` (with mount and dentry) being set
up in the ``nameidata``, and result in pick_link() returning ``NULL``.

The more obvious case is a symlink to "``/``".  All symlinks starting
with "``/``" are detected in pick_link() which resets the ``nameidata``
to point to the effective filesystem root.  If the symlink only
contains "``/``" then there is nothing more to do, no components at all,
so ``NULL`` is returned to indicate that the symlink can be released and
the stack frame discarded.

The other case involves things in ``/proc`` that look like symlinks but
aren't really (and are therefore commonly referred to as "magic-links")::

     $ ls -l /proc/self/fd/1
     lrwx------ 1 neilb neilb 64 Jun 13 10:19 /proc/self/fd/1 -> /dev/pts/4

Every open file descriptor in any process is represented in ``/proc`` by
something that looks like a symlink.  It is really a reference to the
target file, not just the name of it.  When you ``readlink`` these
objects you get a name that might refer to the same file - unless it
has been unlinked or mounted over.  When ``walk_component()`` follows
one of these, the ``->get_link()`` method in "procfs" doesn't return
a string name, but instead calls nd_jump_link() which updates the
``nameidata`` in place to point to that target.  ``->get_link()`` then
returns ``NULL``.  Again there is no final component and pick_link()
returns ``NULL``.

Following the symlink in the final component
--------------------------------------------

All this leads to ``link_path_walk()`` walking down every component, and
following all symbolic links it finds, until it reaches the final
component.  This is just returned in the ``last`` field of ``nameidata``.
For some callers, this is all they need; they want to create that
``last`` name if it doesn't exist or give an error if it does.  Other
callers will want to follow a symlink if one is found, and possibly
apply special handling to the last component of that symlink, rather
than just the last component of the original file name.  These callers
potentially need to call ``link_path_walk()`` again and again on
successive symlinks until one is found that doesn't point to another
symlink.

This case is handled by relevant callers of link_path_walk(), such as
path_lookupat(), path_openat() using a loop that calls link_path_walk(),
and then handles the final component by calling open_last_lookups() or
lookup_last(). If it is a symlink that needs to be followed,
open_last_lookups() or lookup_last() will set things up properly and
return the path so that the loop repeats, calling
link_path_walk() again.  This could loop as many as 40 times if the last
component of each symlink is another symlink.

Of the various functions that examine the final component,
open_last_lookups() is the most interesting as it works in tandem
with do_open() for opening a file.  Part of open_last_lookups() runs
with ``i_rwsem`` held and this part is in a separate function: lookup_open().

Explaining open_last_lookups() and do_open() completely is beyond the scope
of this article, but a few highlights should help those interested in exploring
the code.

1. Rather than just finding the target file, do_open() is used after
   open_last_lookup() to open
   it.  If the file was found in the dcache, then ``vfs_open()`` is used for
   this.  If not, then ``lookup_open()`` will either call ``atomic_open()`` (if
   the filesystem provides it) to combine the final lookup with the open, or
   will perform the separate ``i_op->lookup()`` and ``i_op->create()`` steps
   directly.  In the later case the actual "open" of this newly found or
   created file will be performed by vfs_open(), just as if the name
   were found in the dcache.

2. vfs_open() can fail with ``-EOPENSTALE`` if the cached information
   wasn't quite current enough.  If it's in RCU-walk ``-ECHILD`` will be returned
   otherwise ``-ESTALE`` is returned.  When ``-ESTALE`` is returned, the caller may
   retry with ``LOOKUP_REVAL`` flag set.

3. An open with O_CREAT **does** follow a symlink in the final component,
   unlike other creation system calls (like ``mkdir``).  So the sequence::

          ln -s bar /tmp/foo
          echo hello > /tmp/foo

   will create a file called ``/tmp/bar``.  This is not permitted if
   ``O_EXCL`` is set but otherwise is handled for an O_CREAT open much
   like for a non-creating open: lookup_last() or open_last_lookup()
   returns a non ``NULL`` value, and link_path_walk() gets called and the
   open process continues on the symlink that was found.

Updating the access time
------------------------

We previously said of RCU-walk that it would "take no locks, increment
no counts, leave no footprints."  We have since seen that some
"footprints" can be needed when handling symlinks as a counted
reference (or even a memory allocation) may be needed.  But these
footprints are best kept to a minimum.

One other place where walking down a symlink can involve leaving
footprints in a way that doesn't affect directories is in updating access times.
In Unix (and Linux) every filesystem object has a "last accessed
time", or "``atime``".  Passing through a directory to access a file
within is not considered to be an access for the purposes of
``atime``; only listing the contents of a directory can update its ``atime``.
Symlinks are different it seems.  Both reading a symlink (with ``readlink()``)
and looking up a symlink on the way to some other destination can
update the atime on that symlink.

.. _clearest statement: https://pubs.opengroup.org/onlinepubs/9699919799/basedefs/V1_chap04.html#tag_04_08

It is not clear why this is the case; POSIX has little to say on the
subject.  The `clearest statement`_ is that, if a particular implementation
updates a timestamp in a place not specified by POSIX, this must be
documented "except that any changes caused by pathname resolution need
not be documented".  This seems to imply that POSIX doesn't really
care about access-time updates during pathname lookup.

.. _Linux 1.3.87: https://git.kernel.org/cgit/linux/kernel/git/history/history.git/diff/fs/ext2/symlink.c?id=f806c6db77b8eaa6e00dcfb6b567706feae8dbb8

An examination of history shows that prior to `Linux 1.3.87`_, the ext2
filesystem, at least, didn't update atime when following a link.
Unfortunately we have no record of why that behavior was changed.

In any case, access time must now be updated and that operation can be
quite complex.  Trying to stay in RCU-walk while doing it is best
avoided.  Fortunately it is often permitted to skip the ``atime``
update.  Because ``atime`` updates cause performance problems in various
areas, Linux supports the ``relatime`` mount option, which generally
limits the updates of ``atime`` to once per day on files that aren't
being changed (and symlinks never change once created).  Even without
``relatime``, many filesystems record ``atime`` with a one-second
granularity, so only one update per second is required.

It is easy to test if an ``atime`` update is needed while in RCU-walk
mode and, if it isn't, the update can be skipped and RCU-walk mode
continues.  Only when an ``atime`` update is actually required does the
path walk drop down to REF-walk.  All of this is handled in the
``get_link()`` function.

A few flags
-----------

A suitable way to wrap up this tour of pathname walking is to list
the various flags that can be stored in the ``nameidata`` to guide the
lookup process.  Many of these are only meaningful on the final
component, others reflect the current state of the pathname lookup, and some
apply restrictions to all path components encountered in the path lookup.

And then there is ``LOOKUP_EMPTY``, which doesn't fit conceptually with
the others.  If this is not set, an empty pathname causes an error
very early on.  If it is set, empty pathnames are not considered to be
an error.

Global state flags
~~~~~~~~~~~~~~~~~~

We have already met two global state flags: ``LOOKUP_RCU`` and
``LOOKUP_REVAL``.  These select between one of three overall approaches
to lookup: RCU-walk, REF-walk, and REF-walk with forced revalidation.

``LOOKUP_PARENT`` indicates that the final component hasn't been reached
yet.  This is primarily used to tell the audit subsystem the full
context of a particular access being audited.

``ND_ROOT_PRESET`` indicates that the ``root`` field in the ``nameidata`` was
provided by the caller, so it shouldn't be released when it is no
longer needed.

``ND_JUMPED`` means that the current dentry was chosen not because
it had the right name but for some other reason.  This happens when
following "``..``", following a symlink to ``/``, crossing a mount point
or accessing a "``/proc/$PID/fd/$FD``" symlink (also known as a "magic
link"). In this case the filesystem has not been asked to revalidate the
name (with ``d_revalidate()``).  In such cases the inode may still need
to be revalidated, so ``d_op->d_weak_revalidate()`` is called if
``ND_JUMPED`` is set when the look completes - which may be at the
final component or, when creating, unlinking, or renaming, at the penultimate component.

Resolution-restriction flags
~~~~~~~~~~~~~~~~~~~~~~~~~~~~

In order to allow userspace to protect itself against certain race conditions
and attack scenarios involving changing path components, a series of flags are
available which apply restrictions to all path components encountered during
path lookup. These flags are exposed through ``openat2()``'s ``resolve`` field.

``LOOKUP_NO_SYMLINKS`` blocks all symlink traversals (including magic-links).
This is distinctly different from ``LOOKUP_FOLLOW``, because the latter only
relates to restricting the following of trailing symlinks.

``LOOKUP_NO_MAGICLINKS`` blocks all magic-link traversals. Filesystems must
ensure that they return errors from ``nd_jump_link()``, because that is how
``LOOKUP_NO_MAGICLINKS`` and other magic-link restrictions are implemented.

``LOOKUP_NO_XDEV`` blocks all ``vfsmount`` traversals (this includes both
bind-mounts and ordinary mounts). Note that the ``vfsmount`` which contains the
lookup is determined by the first mountpoint the path lookup reaches --
absolute paths start with the ``vfsmount`` of ``/``, and relative paths start
with the ``dfd``'s ``vfsmount``. Magic-links are only permitted if the
``vfsmount`` of the path is unchanged.

``LOOKUP_BENEATH`` blocks any path components which resolve outside the
starting point of the resolution. This is done by blocking ``nd_jump_root()``
as well as blocking ".." if it would jump outside the starting point.
``rename_lock`` and ``mount_lock`` are used to detect attacks against the
resolution of "..". Magic-links are also blocked.

``LOOKUP_IN_ROOT`` resolves all path components as though the starting point
were the filesystem root. ``nd_jump_root()`` brings the resolution back to
the starting point, and ".." at the starting point will act as a no-op. As with
``LOOKUP_BENEATH``, ``rename_lock`` and ``mount_lock`` are used to detect
attacks against ".." resolution. Magic-links are also blocked.

Final-component flags
~~~~~~~~~~~~~~~~~~~~~

Some of these flags are only set when the final component is being
considered.  Others are only checked for when considering that final
component.

``LOOKUP_AUTOMOUNT`` ensures that, if the final component is an automount
point, then the mount is triggered.  Some operations would trigger it
anyway, but operations like ``stat()`` deliberately don't.  ``statfs()``
needs to trigger the mount but otherwise behaves a lot like ``stat()``, so
it sets ``LOOKUP_AUTOMOUNT``, as does "``quotactl()``" and the handling of
"``mount --bind``".

``LOOKUP_FOLLOW`` has a similar function to ``LOOKUP_AUTOMOUNT`` but for
symlinks.  Some system calls set or clear it implicitly, while
others have API flags such as ``AT_SYMLINK_FOLLOW`` and
``UMOUNT_NOFOLLOW`` to control it.  Its effect is similar to
``WALK_GET`` that we already met, but it is used in a different way.

``LOOKUP_DIRECTORY`` insists that the final component is a directory.
Various callers set this and it is also set when the final component
is found to be followed by a slash.

Finally ``LOOKUP_OPEN``, ``LOOKUP_CREATE``, ``LOOKUP_EXCL``, and
``LOOKUP_RENAME_TARGET`` are not used directly by the VFS but are made
available to the filesystem and particularly the ``->d_revalidate()``
method.  A filesystem can choose not to bother revalidating too hard
if it knows that it will be asked to open or create the file soon.
These flags were previously useful for ``->lookup()`` too but with the
introduction of ``->atomic_open()`` they are less relevant there.

End of the road
---------------

Despite its complexity, all this pathname lookup code appears to be
in good shape - various parts are certainly easier to understand now
than even a couple of releases ago.  But that doesn't mean it is
"finished".   As already mentioned, RCU-walk currently only follows
symlinks that are stored in the inode so, while it handles many ext4
symlinks, it doesn't help with NFS, XFS, or Btrfs.  That support
is not likely to be long delayed.

repo_name: angery/linux