* [PATCH 6/7] locking: Move docs into Documentation/locking/
2014-07-28 5:18 [PATCH -tip/master 1/7] locking/mutex: Unify arguments in lock/unlock slowpaths Davidlohr Bueso
` (3 preceding siblings ...)
2014-07-28 5:18 ` [PATCH -tip/master 5/7] locking/mutex: Use MUTEX_SPIN_ON_OWNER when appropriate Davidlohr Bueso
@ 2014-07-28 5:18 ` Davidlohr Bueso
2014-07-28 5:18 ` [PATCH -tip/master 7/7] Documentation: Update locking/mutex-design.txt disadvantages Davidlohr Bueso
5 siblings, 0 replies; 28+ messages in thread
From: Davidlohr Bueso @ 2014-07-28 5:18 UTC (permalink / raw)
To: peterz, mingo; +Cc: jason.low2, davidlohr, aswin, linux-kernel
Specifically:
Documentation/locking/lockdep-design.txt
Documentation/locking/lockstat.txt
Documentation/locking/mutex-design.txt
Documentation/locking/rt-mutex-design.txt
Documentation/locking/rt-mutex.txt
Documentation/locking/spinlocks.txt
Documentation/locking/ww-mutex-design.txt
Signed-off-by: Davidlohr Bueso <davidlohr@hp.com>
---
Documentation/00-INDEX | 2 +
Documentation/DocBook/kernel-locking.tmpl | 2 +-
Documentation/lockdep-design.txt | 286 -----------
Documentation/locking/lockdep-design.txt | 286 +++++++++++
Documentation/locking/lockstat.txt | 178 +++++++
Documentation/locking/mutex-design.txt | 157 ++++++
Documentation/locking/rt-mutex-design.txt | 781 ++++++++++++++++++++++++++++++
Documentation/locking/rt-mutex.txt | 79 +++
Documentation/locking/spinlocks.txt | 167 +++++++
Documentation/locking/ww-mutex-design.txt | 344 +++++++++++++
Documentation/lockstat.txt | 178 -------
Documentation/mutex-design.txt | 157 ------
Documentation/rt-mutex-design.txt | 781 ------------------------------
Documentation/rt-mutex.txt | 79 ---
Documentation/spinlocks.txt | 167 -------
Documentation/ww-mutex-design.txt | 344 -------------
MAINTAINERS | 4 +-
drivers/gpu/drm/drm_modeset_lock.c | 2 +-
include/linux/lockdep.h | 2 +-
include/linux/mutex.h | 2 +-
include/linux/rwsem.h | 2 +-
kernel/locking/mutex.c | 2 +-
kernel/locking/rtmutex.c | 2 +-
lib/Kconfig.debug | 4 +-
24 files changed, 2005 insertions(+), 2003 deletions(-)
delete mode 100644 Documentation/lockdep-design.txt
create mode 100644 Documentation/locking/lockdep-design.txt
create mode 100644 Documentation/locking/lockstat.txt
create mode 100644 Documentation/locking/mutex-design.txt
create mode 100644 Documentation/locking/rt-mutex-design.txt
create mode 100644 Documentation/locking/rt-mutex.txt
create mode 100644 Documentation/locking/spinlocks.txt
create mode 100644 Documentation/locking/ww-mutex-design.txt
delete mode 100644 Documentation/lockstat.txt
delete mode 100644 Documentation/mutex-design.txt
delete mode 100644 Documentation/rt-mutex-design.txt
delete mode 100644 Documentation/rt-mutex.txt
delete mode 100644 Documentation/spinlocks.txt
delete mode 100644 Documentation/ww-mutex-design.txt
diff --git a/Documentation/00-INDEX b/Documentation/00-INDEX
index 27e67a9..1750fce 100644
--- a/Documentation/00-INDEX
+++ b/Documentation/00-INDEX
@@ -287,6 +287,8 @@ local_ops.txt
- semantics and behavior of local atomic operations.
lockdep-design.txt
- documentation on the runtime locking correctness validator.
+locking/
+ - directory with info about kernel locking primitives
lockstat.txt
- info on collecting statistics on locks (and contention).
lockup-watchdogs.txt
diff --git a/Documentation/DocBook/kernel-locking.tmpl b/Documentation/DocBook/kernel-locking.tmpl
index e584ee1..c70fd1b 100644
--- a/Documentation/DocBook/kernel-locking.tmpl
+++ b/Documentation/DocBook/kernel-locking.tmpl
@@ -1972,7 +1972,7 @@ machines due to caching.
<itemizedlist>
<listitem>
<para>
- <filename>Documentation/spinlocks.txt</filename>:
+ <filename>Documentation/locking/spinlocks.txt</filename>:
Linus Torvalds' spinlocking tutorial in the kernel sources.
</para>
</listitem>
diff --git a/Documentation/lockdep-design.txt b/Documentation/lockdep-design.txt
deleted file mode 100644
index 5dbc99c..0000000
--- a/Documentation/lockdep-design.txt
+++ /dev/null
@@ -1,286 +0,0 @@
-Runtime locking correctness validator
-=====================================
-
-started by Ingo Molnar <mingo@redhat.com>
-additions by Arjan van de Ven <arjan@linux.intel.com>
-
-Lock-class
-----------
-
-The basic object the validator operates upon is a 'class' of locks.
-
-A class of locks is a group of locks that are logically the same with
-respect to locking rules, even if the locks may have multiple (possibly
-tens of thousands of) instantiations. For example a lock in the inode
-struct is one class, while each inode has its own instantiation of that
-lock class.
-
-The validator tracks the 'state' of lock-classes, and it tracks
-dependencies between different lock-classes. The validator maintains a
-rolling proof that the state and the dependencies are correct.
-
-Unlike an lock instantiation, the lock-class itself never goes away: when
-a lock-class is used for the first time after bootup it gets registered,
-and all subsequent uses of that lock-class will be attached to this
-lock-class.
-
-State
------
-
-The validator tracks lock-class usage history into 4n + 1 separate state bits:
-
-- 'ever held in STATE context'
-- 'ever held as readlock in STATE context'
-- 'ever held with STATE enabled'
-- 'ever held as readlock with STATE enabled'
-
-Where STATE can be either one of (kernel/lockdep_states.h)
- - hardirq
- - softirq
- - reclaim_fs
-
-- 'ever used' [ == !unused ]
-
-When locking rules are violated, these state bits are presented in the
-locking error messages, inside curlies. A contrived example:
-
- modprobe/2287 is trying to acquire lock:
- (&sio_locks[i].lock){-.-...}, at: [<c02867fd>] mutex_lock+0x21/0x24
-
- but task is already holding lock:
- (&sio_locks[i].lock){-.-...}, at: [<c02867fd>] mutex_lock+0x21/0x24
-
-
-The bit position indicates STATE, STATE-read, for each of the states listed
-above, and the character displayed in each indicates:
-
- '.' acquired while irqs disabled and not in irq context
- '-' acquired in irq context
- '+' acquired with irqs enabled
- '?' acquired in irq context with irqs enabled.
-
-Unused mutexes cannot be part of the cause of an error.
-
-
-Single-lock state rules:
-------------------------
-
-A softirq-unsafe lock-class is automatically hardirq-unsafe as well. The
-following states are exclusive, and only one of them is allowed to be
-set for any lock-class:
-
- <hardirq-safe> and <hardirq-unsafe>
- <softirq-safe> and <softirq-unsafe>
-
-The validator detects and reports lock usage that violate these
-single-lock state rules.
-
-Multi-lock dependency rules:
-----------------------------
-
-The same lock-class must not be acquired twice, because this could lead
-to lock recursion deadlocks.
-
-Furthermore, two locks may not be taken in different order:
-
- <L1> -> <L2>
- <L2> -> <L1>
-
-because this could lead to lock inversion deadlocks. (The validator
-finds such dependencies in arbitrary complexity, i.e. there can be any
-other locking sequence between the acquire-lock operations, the
-validator will still track all dependencies between locks.)
-
-Furthermore, the following usage based lock dependencies are not allowed
-between any two lock-classes:
-
- <hardirq-safe> -> <hardirq-unsafe>
- <softirq-safe> -> <softirq-unsafe>
-
-The first rule comes from the fact the a hardirq-safe lock could be
-taken by a hardirq context, interrupting a hardirq-unsafe lock - and
-thus could result in a lock inversion deadlock. Likewise, a softirq-safe
-lock could be taken by an softirq context, interrupting a softirq-unsafe
-lock.
-
-The above rules are enforced for any locking sequence that occurs in the
-kernel: when acquiring a new lock, the validator checks whether there is
-any rule violation between the new lock and any of the held locks.
-
-When a lock-class changes its state, the following aspects of the above
-dependency rules are enforced:
-
-- if a new hardirq-safe lock is discovered, we check whether it
- took any hardirq-unsafe lock in the past.
-
-- if a new softirq-safe lock is discovered, we check whether it took
- any softirq-unsafe lock in the past.
-
-- if a new hardirq-unsafe lock is discovered, we check whether any
- hardirq-safe lock took it in the past.
-
-- if a new softirq-unsafe lock is discovered, we check whether any
- softirq-safe lock took it in the past.
-
-(Again, we do these checks too on the basis that an interrupt context
-could interrupt _any_ of the irq-unsafe or hardirq-unsafe locks, which
-could lead to a lock inversion deadlock - even if that lock scenario did
-not trigger in practice yet.)
-
-Exception: Nested data dependencies leading to nested locking
--------------------------------------------------------------
-
-There are a few cases where the Linux kernel acquires more than one
-instance of the same lock-class. Such cases typically happen when there
-is some sort of hierarchy within objects of the same type. In these
-cases there is an inherent "natural" ordering between the two objects
-(defined by the properties of the hierarchy), and the kernel grabs the
-locks in this fixed order on each of the objects.
-
-An example of such an object hierarchy that results in "nested locking"
-is that of a "whole disk" block-dev object and a "partition" block-dev
-object; the partition is "part of" the whole device and as long as one
-always takes the whole disk lock as a higher lock than the partition
-lock, the lock ordering is fully correct. The validator does not
-automatically detect this natural ordering, as the locking rule behind
-the ordering is not static.
-
-In order to teach the validator about this correct usage model, new
-versions of the various locking primitives were added that allow you to
-specify a "nesting level". An example call, for the block device mutex,
-looks like this:
-
-enum bdev_bd_mutex_lock_class
-{
- BD_MUTEX_NORMAL,
- BD_MUTEX_WHOLE,
- BD_MUTEX_PARTITION
-};
-
- mutex_lock_nested(&bdev->bd_contains->bd_mutex, BD_MUTEX_PARTITION);
-
-In this case the locking is done on a bdev object that is known to be a
-partition.
-
-The validator treats a lock that is taken in such a nested fashion as a
-separate (sub)class for the purposes of validation.
-
-Note: When changing code to use the _nested() primitives, be careful and
-check really thoroughly that the hierarchy is correctly mapped; otherwise
-you can get false positives or false negatives.
-
-Proof of 100% correctness:
---------------------------
-
-The validator achieves perfect, mathematical 'closure' (proof of locking
-correctness) in the sense that for every simple, standalone single-task
-locking sequence that occurred at least once during the lifetime of the
-kernel, the validator proves it with a 100% certainty that no
-combination and timing of these locking sequences can cause any class of
-lock related deadlock. [*]
-
-I.e. complex multi-CPU and multi-task locking scenarios do not have to
-occur in practice to prove a deadlock: only the simple 'component'
-locking chains have to occur at least once (anytime, in any
-task/context) for the validator to be able to prove correctness. (For
-example, complex deadlocks that would normally need more than 3 CPUs and
-a very unlikely constellation of tasks, irq-contexts and timings to
-occur, can be detected on a plain, lightly loaded single-CPU system as
-well!)
-
-This radically decreases the complexity of locking related QA of the
-kernel: what has to be done during QA is to trigger as many "simple"
-single-task locking dependencies in the kernel as possible, at least
-once, to prove locking correctness - instead of having to trigger every
-possible combination of locking interaction between CPUs, combined with
-every possible hardirq and softirq nesting scenario (which is impossible
-to do in practice).
-
-[*] assuming that the validator itself is 100% correct, and no other
- part of the system corrupts the state of the validator in any way.
- We also assume that all NMI/SMM paths [which could interrupt
- even hardirq-disabled codepaths] are correct and do not interfere
- with the validator. We also assume that the 64-bit 'chain hash'
- value is unique for every lock-chain in the system. Also, lock
- recursion must not be higher than 20.
-
-Performance:
-------------
-
-The above rules require _massive_ amounts of runtime checking. If we did
-that for every lock taken and for every irqs-enable event, it would
-render the system practically unusably slow. The complexity of checking
-is O(N^2), so even with just a few hundred lock-classes we'd have to do
-tens of thousands of checks for every event.
-
-This problem is solved by checking any given 'locking scenario' (unique
-sequence of locks taken after each other) only once. A simple stack of
-held locks is maintained, and a lightweight 64-bit hash value is
-calculated, which hash is unique for every lock chain. The hash value,
-when the chain is validated for the first time, is then put into a hash
-table, which hash-table can be checked in a lockfree manner. If the
-locking chain occurs again later on, the hash table tells us that we
-dont have to validate the chain again.
-
-Troubleshooting:
-----------------
-
-The validator tracks a maximum of MAX_LOCKDEP_KEYS number of lock classes.
-Exceeding this number will trigger the following lockdep warning:
-
- (DEBUG_LOCKS_WARN_ON(id >= MAX_LOCKDEP_KEYS))
-
-By default, MAX_LOCKDEP_KEYS is currently set to 8191, and typical
-desktop systems have less than 1,000 lock classes, so this warning
-normally results from lock-class leakage or failure to properly
-initialize locks. These two problems are illustrated below:
-
-1. Repeated module loading and unloading while running the validator
- will result in lock-class leakage. The issue here is that each
- load of the module will create a new set of lock classes for
- that module's locks, but module unloading does not remove old
- classes (see below discussion of reuse of lock classes for why).
- Therefore, if that module is loaded and unloaded repeatedly,
- the number of lock classes will eventually reach the maximum.
-
-2. Using structures such as arrays that have large numbers of
- locks that are not explicitly initialized. For example,
- a hash table with 8192 buckets where each bucket has its own
- spinlock_t will consume 8192 lock classes -unless- each spinlock
- is explicitly initialized at runtime, for example, using the
- run-time spin_lock_init() as opposed to compile-time initializers
- such as __SPIN_LOCK_UNLOCKED(). Failure to properly initialize
- the per-bucket spinlocks would guarantee lock-class overflow.
- In contrast, a loop that called spin_lock_init() on each lock
- would place all 8192 locks into a single lock class.
-
- The moral of this story is that you should always explicitly
- initialize your locks.
-
-One might argue that the validator should be modified to allow
-lock classes to be reused. However, if you are tempted to make this
-argument, first review the code and think through the changes that would
-be required, keeping in mind that the lock classes to be removed are
-likely to be linked into the lock-dependency graph. This turns out to
-be harder to do than to say.
-
-Of course, if you do run out of lock classes, the next thing to do is
-to find the offending lock classes. First, the following command gives
-you the number of lock classes currently in use along with the maximum:
-
- grep "lock-classes" /proc/lockdep_stats
-
-This command produces the following output on a modest system:
-
- lock-classes: 748 [max: 8191]
-
-If the number allocated (748 above) increases continually over time,
-then there is likely a leak. The following command can be used to
-identify the leaking lock classes:
-
- grep "BD" /proc/lockdep
-
-Run the command and save the output, then compare against the output from
-a later run of this command to identify the leakers. This same output
-can also help you find situations where runtime lock initialization has
-been omitted.
diff --git a/Documentation/locking/lockdep-design.txt b/Documentation/locking/lockdep-design.txt
new file mode 100644
index 0000000..5dbc99c
--- /dev/null
+++ b/Documentation/locking/lockdep-design.txt
@@ -0,0 +1,286 @@
+Runtime locking correctness validator
+=====================================
+
+started by Ingo Molnar <mingo@redhat.com>
+additions by Arjan van de Ven <arjan@linux.intel.com>
+
+Lock-class
+----------
+
+The basic object the validator operates upon is a 'class' of locks.
+
+A class of locks is a group of locks that are logically the same with
+respect to locking rules, even if the locks may have multiple (possibly
+tens of thousands of) instantiations. For example a lock in the inode
+struct is one class, while each inode has its own instantiation of that
+lock class.
+
+The validator tracks the 'state' of lock-classes, and it tracks
+dependencies between different lock-classes. The validator maintains a
+rolling proof that the state and the dependencies are correct.
+
+Unlike an lock instantiation, the lock-class itself never goes away: when
+a lock-class is used for the first time after bootup it gets registered,
+and all subsequent uses of that lock-class will be attached to this
+lock-class.
+
+State
+-----
+
+The validator tracks lock-class usage history into 4n + 1 separate state bits:
+
+- 'ever held in STATE context'
+- 'ever held as readlock in STATE context'
+- 'ever held with STATE enabled'
+- 'ever held as readlock with STATE enabled'
+
+Where STATE can be either one of (kernel/lockdep_states.h)
+ - hardirq
+ - softirq
+ - reclaim_fs
+
+- 'ever used' [ == !unused ]
+
+When locking rules are violated, these state bits are presented in the
+locking error messages, inside curlies. A contrived example:
+
+ modprobe/2287 is trying to acquire lock:
+ (&sio_locks[i].lock){-.-...}, at: [<c02867fd>] mutex_lock+0x21/0x24
+
+ but task is already holding lock:
+ (&sio_locks[i].lock){-.-...}, at: [<c02867fd>] mutex_lock+0x21/0x24
+
+
+The bit position indicates STATE, STATE-read, for each of the states listed
+above, and the character displayed in each indicates:
+
+ '.' acquired while irqs disabled and not in irq context
+ '-' acquired in irq context
+ '+' acquired with irqs enabled
+ '?' acquired in irq context with irqs enabled.
+
+Unused mutexes cannot be part of the cause of an error.
+
+
+Single-lock state rules:
+------------------------
+
+A softirq-unsafe lock-class is automatically hardirq-unsafe as well. The
+following states are exclusive, and only one of them is allowed to be
+set for any lock-class:
+
+ <hardirq-safe> and <hardirq-unsafe>
+ <softirq-safe> and <softirq-unsafe>
+
+The validator detects and reports lock usage that violate these
+single-lock state rules.
+
+Multi-lock dependency rules:
+----------------------------
+
+The same lock-class must not be acquired twice, because this could lead
+to lock recursion deadlocks.
+
+Furthermore, two locks may not be taken in different order:
+
+ <L1> -> <L2>
+ <L2> -> <L1>
+
+because this could lead to lock inversion deadlocks. (The validator
+finds such dependencies in arbitrary complexity, i.e. there can be any
+other locking sequence between the acquire-lock operations, the
+validator will still track all dependencies between locks.)
+
+Furthermore, the following usage based lock dependencies are not allowed
+between any two lock-classes:
+
+ <hardirq-safe> -> <hardirq-unsafe>
+ <softirq-safe> -> <softirq-unsafe>
+
+The first rule comes from the fact the a hardirq-safe lock could be
+taken by a hardirq context, interrupting a hardirq-unsafe lock - and
+thus could result in a lock inversion deadlock. Likewise, a softirq-safe
+lock could be taken by an softirq context, interrupting a softirq-unsafe
+lock.
+
+The above rules are enforced for any locking sequence that occurs in the
+kernel: when acquiring a new lock, the validator checks whether there is
+any rule violation between the new lock and any of the held locks.
+
+When a lock-class changes its state, the following aspects of the above
+dependency rules are enforced:
+
+- if a new hardirq-safe lock is discovered, we check whether it
+ took any hardirq-unsafe lock in the past.
+
+- if a new softirq-safe lock is discovered, we check whether it took
+ any softirq-unsafe lock in the past.
+
+- if a new hardirq-unsafe lock is discovered, we check whether any
+ hardirq-safe lock took it in the past.
+
+- if a new softirq-unsafe lock is discovered, we check whether any
+ softirq-safe lock took it in the past.
+
+(Again, we do these checks too on the basis that an interrupt context
+could interrupt _any_ of the irq-unsafe or hardirq-unsafe locks, which
+could lead to a lock inversion deadlock - even if that lock scenario did
+not trigger in practice yet.)
+
+Exception: Nested data dependencies leading to nested locking
+-------------------------------------------------------------
+
+There are a few cases where the Linux kernel acquires more than one
+instance of the same lock-class. Such cases typically happen when there
+is some sort of hierarchy within objects of the same type. In these
+cases there is an inherent "natural" ordering between the two objects
+(defined by the properties of the hierarchy), and the kernel grabs the
+locks in this fixed order on each of the objects.
+
+An example of such an object hierarchy that results in "nested locking"
+is that of a "whole disk" block-dev object and a "partition" block-dev
+object; the partition is "part of" the whole device and as long as one
+always takes the whole disk lock as a higher lock than the partition
+lock, the lock ordering is fully correct. The validator does not
+automatically detect this natural ordering, as the locking rule behind
+the ordering is not static.
+
+In order to teach the validator about this correct usage model, new
+versions of the various locking primitives were added that allow you to
+specify a "nesting level". An example call, for the block device mutex,
+looks like this:
+
+enum bdev_bd_mutex_lock_class
+{
+ BD_MUTEX_NORMAL,
+ BD_MUTEX_WHOLE,
+ BD_MUTEX_PARTITION
+};
+
+ mutex_lock_nested(&bdev->bd_contains->bd_mutex, BD_MUTEX_PARTITION);
+
+In this case the locking is done on a bdev object that is known to be a
+partition.
+
+The validator treats a lock that is taken in such a nested fashion as a
+separate (sub)class for the purposes of validation.
+
+Note: When changing code to use the _nested() primitives, be careful and
+check really thoroughly that the hierarchy is correctly mapped; otherwise
+you can get false positives or false negatives.
+
+Proof of 100% correctness:
+--------------------------
+
+The validator achieves perfect, mathematical 'closure' (proof of locking
+correctness) in the sense that for every simple, standalone single-task
+locking sequence that occurred at least once during the lifetime of the
+kernel, the validator proves it with a 100% certainty that no
+combination and timing of these locking sequences can cause any class of
+lock related deadlock. [*]
+
+I.e. complex multi-CPU and multi-task locking scenarios do not have to
+occur in practice to prove a deadlock: only the simple 'component'
+locking chains have to occur at least once (anytime, in any
+task/context) for the validator to be able to prove correctness. (For
+example, complex deadlocks that would normally need more than 3 CPUs and
+a very unlikely constellation of tasks, irq-contexts and timings to
+occur, can be detected on a plain, lightly loaded single-CPU system as
+well!)
+
+This radically decreases the complexity of locking related QA of the
+kernel: what has to be done during QA is to trigger as many "simple"
+single-task locking dependencies in the kernel as possible, at least
+once, to prove locking correctness - instead of having to trigger every
+possible combination of locking interaction between CPUs, combined with
+every possible hardirq and softirq nesting scenario (which is impossible
+to do in practice).
+
+[*] assuming that the validator itself is 100% correct, and no other
+ part of the system corrupts the state of the validator in any way.
+ We also assume that all NMI/SMM paths [which could interrupt
+ even hardirq-disabled codepaths] are correct and do not interfere
+ with the validator. We also assume that the 64-bit 'chain hash'
+ value is unique for every lock-chain in the system. Also, lock
+ recursion must not be higher than 20.
+
+Performance:
+------------
+
+The above rules require _massive_ amounts of runtime checking. If we did
+that for every lock taken and for every irqs-enable event, it would
+render the system practically unusably slow. The complexity of checking
+is O(N^2), so even with just a few hundred lock-classes we'd have to do
+tens of thousands of checks for every event.
+
+This problem is solved by checking any given 'locking scenario' (unique
+sequence of locks taken after each other) only once. A simple stack of
+held locks is maintained, and a lightweight 64-bit hash value is
+calculated, which hash is unique for every lock chain. The hash value,
+when the chain is validated for the first time, is then put into a hash
+table, which hash-table can be checked in a lockfree manner. If the
+locking chain occurs again later on, the hash table tells us that we
+dont have to validate the chain again.
+
+Troubleshooting:
+----------------
+
+The validator tracks a maximum of MAX_LOCKDEP_KEYS number of lock classes.
+Exceeding this number will trigger the following lockdep warning:
+
+ (DEBUG_LOCKS_WARN_ON(id >= MAX_LOCKDEP_KEYS))
+
+By default, MAX_LOCKDEP_KEYS is currently set to 8191, and typical
+desktop systems have less than 1,000 lock classes, so this warning
+normally results from lock-class leakage or failure to properly
+initialize locks. These two problems are illustrated below:
+
+1. Repeated module loading and unloading while running the validator
+ will result in lock-class leakage. The issue here is that each
+ load of the module will create a new set of lock classes for
+ that module's locks, but module unloading does not remove old
+ classes (see below discussion of reuse of lock classes for why).
+ Therefore, if that module is loaded and unloaded repeatedly,
+ the number of lock classes will eventually reach the maximum.
+
+2. Using structures such as arrays that have large numbers of
+ locks that are not explicitly initialized. For example,
+ a hash table with 8192 buckets where each bucket has its own
+ spinlock_t will consume 8192 lock classes -unless- each spinlock
+ is explicitly initialized at runtime, for example, using the
+ run-time spin_lock_init() as opposed to compile-time initializers
+ such as __SPIN_LOCK_UNLOCKED(). Failure to properly initialize
+ the per-bucket spinlocks would guarantee lock-class overflow.
+ In contrast, a loop that called spin_lock_init() on each lock
+ would place all 8192 locks into a single lock class.
+
+ The moral of this story is that you should always explicitly
+ initialize your locks.
+
+One might argue that the validator should be modified to allow
+lock classes to be reused. However, if you are tempted to make this
+argument, first review the code and think through the changes that would
+be required, keeping in mind that the lock classes to be removed are
+likely to be linked into the lock-dependency graph. This turns out to
+be harder to do than to say.
+
+Of course, if you do run out of lock classes, the next thing to do is
+to find the offending lock classes. First, the following command gives
+you the number of lock classes currently in use along with the maximum:
+
+ grep "lock-classes" /proc/lockdep_stats
+
+This command produces the following output on a modest system:
+
+ lock-classes: 748 [max: 8191]
+
+If the number allocated (748 above) increases continually over time,
+then there is likely a leak. The following command can be used to
+identify the leaking lock classes:
+
+ grep "BD" /proc/lockdep
+
+Run the command and save the output, then compare against the output from
+a later run of this command to identify the leakers. This same output
+can also help you find situations where runtime lock initialization has
+been omitted.
diff --git a/Documentation/locking/lockstat.txt b/Documentation/locking/lockstat.txt
new file mode 100644
index 0000000..7428773
--- /dev/null
+++ b/Documentation/locking/lockstat.txt
@@ -0,0 +1,178 @@
+
+LOCK STATISTICS
+
+- WHAT
+
+As the name suggests, it provides statistics on locks.
+
+- WHY
+
+Because things like lock contention can severely impact performance.
+
+- HOW
+
+Lockdep already has hooks in the lock functions and maps lock instances to
+lock classes. We build on that (see Documentation/lokcing/lockdep-design.txt).
+The graph below shows the relation between the lock functions and the various
+hooks therein.
+
+ __acquire
+ |
+ lock _____
+ | \
+ | __contended
+ | |
+ | <wait>
+ | _______/
+ |/
+ |
+ __acquired
+ |
+ .
+ <hold>
+ .
+ |
+ __release
+ |
+ unlock
+
+lock, unlock - the regular lock functions
+__* - the hooks
+<> - states
+
+With these hooks we provide the following statistics:
+
+ con-bounces - number of lock contention that involved x-cpu data
+ contentions - number of lock acquisitions that had to wait
+ wait time min - shortest (non-0) time we ever had to wait for a lock
+ max - longest time we ever had to wait for a lock
+ total - total time we spend waiting on this lock
+ avg - average time spent waiting on this lock
+ acq-bounces - number of lock acquisitions that involved x-cpu data
+ acquisitions - number of times we took the lock
+ hold time min - shortest (non-0) time we ever held the lock
+ max - longest time we ever held the lock
+ total - total time this lock was held
+ avg - average time this lock was held
+
+These numbers are gathered per lock class, per read/write state (when
+applicable).
+
+It also tracks 4 contention points per class. A contention point is a call site
+that had to wait on lock acquisition.
+
+ - CONFIGURATION
+
+Lock statistics are enabled via CONFIG_LOCK_STAT.
+
+ - USAGE
+
+Enable collection of statistics:
+
+# echo 1 >/proc/sys/kernel/lock_stat
+
+Disable collection of statistics:
+
+# echo 0 >/proc/sys/kernel/lock_stat
+
+Look at the current lock statistics:
+
+( line numbers not part of actual output, done for clarity in the explanation
+ below )
+
+# less /proc/lock_stat
+
+01 lock_stat version 0.4
+02-----------------------------------------------------------------------------------------------------------------------------------------------------------------------------------------------------------------------------
+03 class name con-bounces contentions waittime-min waittime-max waittime-total waittime-avg acq-bounces acquisitions holdtime-min holdtime-max holdtime-total holdtime-avg
+04-----------------------------------------------------------------------------------------------------------------------------------------------------------------------------------------------------------------------------
+05
+06 &mm->mmap_sem-W: 46 84 0.26 939.10 16371.53 194.90 47291 2922365 0.16 2220301.69 17464026916.32 5975.99
+07 &mm->mmap_sem-R: 37 100 1.31 299502.61 325629.52 3256.30 212344 34316685 0.10 7744.91 95016910.20 2.77
+08 ---------------
+09 &mm->mmap_sem 1 [<ffffffff811502a7>] khugepaged_scan_mm_slot+0x57/0x280
+19 &mm->mmap_sem 96 [<ffffffff815351c4>] __do_page_fault+0x1d4/0x510
+11 &mm->mmap_sem 34 [<ffffffff81113d77>] vm_mmap_pgoff+0x87/0xd0
+12 &mm->mmap_sem 17 [<ffffffff81127e71>] vm_munmap+0x41/0x80
+13 ---------------
+14 &mm->mmap_sem 1 [<ffffffff81046fda>] dup_mmap+0x2a/0x3f0
+15 &mm->mmap_sem 60 [<ffffffff81129e29>] SyS_mprotect+0xe9/0x250
+16 &mm->mmap_sem 41 [<ffffffff815351c4>] __do_page_fault+0x1d4/0x510
+17 &mm->mmap_sem 68 [<ffffffff81113d77>] vm_mmap_pgoff+0x87/0xd0
+18
+19.............................................................................................................................................................................................................................
+20
+21 unix_table_lock: 110 112 0.21 49.24 163.91 1.46 21094 66312 0.12 624.42 31589.81 0.48
+22 ---------------
+23 unix_table_lock 45 [<ffffffff8150ad8e>] unix_create1+0x16e/0x1b0
+24 unix_table_lock 47 [<ffffffff8150b111>] unix_release_sock+0x31/0x250
+25 unix_table_lock 15 [<ffffffff8150ca37>] unix_find_other+0x117/0x230
+26 unix_table_lock 5 [<ffffffff8150a09f>] unix_autobind+0x11f/0x1b0
+27 ---------------
+28 unix_table_lock 39 [<ffffffff8150b111>] unix_release_sock+0x31/0x250
+29 unix_table_lock 49 [<ffffffff8150ad8e>] unix_create1+0x16e/0x1b0
+30 unix_table_lock 20 [<ffffffff8150ca37>] unix_find_other+0x117/0x230
+31 unix_table_lock 4 [<ffffffff8150a09f>] unix_autobind+0x11f/0x1b0
+
+
+This excerpt shows the first two lock class statistics. Line 01 shows the
+output version - each time the format changes this will be updated. Line 02-04
+show the header with column descriptions. Lines 05-18 and 20-31 show the actual
+statistics. These statistics come in two parts; the actual stats separated by a
+short separator (line 08, 13) from the contention points.
+
+The first lock (05-18) is a read/write lock, and shows two lines above the
+short separator. The contention points don't match the column descriptors,
+they have two: contentions and [<IP>] symbol. The second set of contention
+points are the points we're contending with.
+
+The integer part of the time values is in us.
+
+Dealing with nested locks, subclasses may appear:
+
+32...........................................................................................................................................................................................................................
+33
+34 &rq->lock: 13128 13128 0.43 190.53 103881.26 7.91 97454 3453404 0.00 401.11 13224683.11 3.82
+35 ---------
+36 &rq->lock 645 [<ffffffff8103bfc4>] task_rq_lock+0x43/0x75
+37 &rq->lock 297 [<ffffffff8104ba65>] try_to_wake_up+0x127/0x25a
+38 &rq->lock 360 [<ffffffff8103c4c5>] select_task_rq_fair+0x1f0/0x74a
+39 &rq->lock 428 [<ffffffff81045f98>] scheduler_tick+0x46/0x1fb
+40 ---------
+41 &rq->lock 77 [<ffffffff8103bfc4>] task_rq_lock+0x43/0x75
+42 &rq->lock 174 [<ffffffff8104ba65>] try_to_wake_up+0x127/0x25a
+43 &rq->lock 4715 [<ffffffff8103ed4b>] double_rq_lock+0x42/0x54
+44 &rq->lock 893 [<ffffffff81340524>] schedule+0x157/0x7b8
+45
+46...........................................................................................................................................................................................................................
+47
+48 &rq->lock/1: 1526 11488 0.33 388.73 136294.31 11.86 21461 38404 0.00 37.93 109388.53 2.84
+49 -----------
+50 &rq->lock/1 11526 [<ffffffff8103ed58>] double_rq_lock+0x4f/0x54
+51 -----------
+52 &rq->lock/1 5645 [<ffffffff8103ed4b>] double_rq_lock+0x42/0x54
+53 &rq->lock/1 1224 [<ffffffff81340524>] schedule+0x157/0x7b8
+54 &rq->lock/1 4336 [<ffffffff8103ed58>] double_rq_lock+0x4f/0x54
+55 &rq->lock/1 181 [<ffffffff8104ba65>] try_to_wake_up+0x127/0x25a
+
+Line 48 shows statistics for the second subclass (/1) of &rq->lock class
+(subclass starts from 0), since in this case, as line 50 suggests,
+double_rq_lock actually acquires a nested lock of two spinlocks.
+
+View the top contending locks:
+
+# grep : /proc/lock_stat | head
+ clockevents_lock: 2926159 2947636 0.15 46882.81 1784540466.34 605.41 3381345 3879161 0.00 2260.97 53178395.68 13.71
+ tick_broadcast_lock: 346460 346717 0.18 2257.43 39364622.71 113.54 3642919 4242696 0.00 2263.79 49173646.60 11.59
+ &mapping->i_mmap_mutex: 203896 203899 3.36 645530.05 31767507988.39 155800.21 3361776 8893984 0.17 2254.15 14110121.02 1.59
+ &rq->lock: 135014 136909 0.18 606.09 842160.68 6.15 1540728 10436146 0.00 728.72 17606683.41 1.69
+ &(&zone->lru_lock)->rlock: 93000 94934 0.16 59.18 188253.78 1.98 1199912 3809894 0.15 391.40 3559518.81 0.93
+ tasklist_lock-W: 40667 41130 0.23 1189.42 428980.51 10.43 270278 510106 0.16 653.51 3939674.91 7.72
+ tasklist_lock-R: 21298 21305 0.20 1310.05 215511.12 10.12 186204 241258 0.14 1162.33 1179779.23 4.89
+ rcu_node_1: 47656 49022 0.16 635.41 193616.41 3.95 844888 1865423 0.00 764.26 1656226.96 0.89
+ &(&dentry->d_lockref.lock)->rlock: 39791 40179 0.15 1302.08 88851.96 2.21 2790851 12527025 0.10 1910.75 3379714.27 0.27
+ rcu_node_0: 29203 30064 0.16 786.55 1555573.00 51.74 88963 244254 0.00 398.87 428872.51 1.76
+
+Clear the statistics:
+
+# echo 0 > /proc/lock_stat
diff --git a/Documentation/locking/mutex-design.txt b/Documentation/locking/mutex-design.txt
new file mode 100644
index 0000000..ee231ed
--- /dev/null
+++ b/Documentation/locking/mutex-design.txt
@@ -0,0 +1,157 @@
+Generic Mutex Subsystem
+
+started by Ingo Molnar <mingo@redhat.com>
+updated by Davidlohr Bueso <davidlohr@hp.com>
+
+What are mutexes?
+-----------------
+
+In the Linux kernel, mutexes refer to a particular locking primitive
+that enforces serialization on shared memory systems, and not only to
+the generic term referring to 'mutual exclusion' found in academia
+or similar theoretical text books. Mutexes are sleeping locks which
+behave similarly to binary semaphores, and were introduced in 2006[1]
+as an alternative to these. This new data structure provided a number
+of advantages, including simpler interfaces, and at that time smaller
+code (see Disadvantages).
+
+[1] http://lwn.net/Articles/164802/
+
+Implementation
+--------------
+
+Mutexes are represented by 'struct mutex', defined in include/linux/mutex.h
+and implemented in kernel/locking/mutex.c. These locks use a three
+state atomic counter (->count) to represent the different possible
+transitions that can occur during the lifetime of a lock:
+
+ 1: unlocked
+ 0: locked, no waiters
+ negative: locked, with potential waiters
+
+In its most basic form it also includes a wait-queue and a spinlock
+that serializes access to it. CONFIG_SMP systems can also include
+a pointer to the lock task owner (->owner) as well as a spinner MCS
+lock (->osq), both described below in (ii).
+
+When acquiring a mutex, there are three possible paths that can be
+taken, depending on the state of the lock:
+
+(i) fastpath: tries to atomically acquire the lock by decrementing the
+ counter. If it was already taken by another task it goes to the next
+ possible path. This logic is architecture specific. On x86-64, the
+ locking fastpath is 2 instructions:
+
+ 0000000000000e10 <mutex_lock>:
+ e21: f0 ff 0b lock decl (%rbx)
+ e24: 79 08 jns e2e <mutex_lock+0x1e>
+
+ the unlocking fastpath is equally tight:
+
+ 0000000000000bc0 <mutex_unlock>:
+ bc8: f0 ff 07 lock incl (%rdi)
+ bcb: 7f 0a jg bd7 <mutex_unlock+0x17>
+
+
+(ii) midpath: aka optimistic spinning, tries to spin for acquisition
+ while the lock owner is running and there are no other tasks ready
+ to run that have higher priority (need_resched). The rationale is
+ that if the lock owner is running, it is likely to release the lock
+ soon. The mutex spinners are queued up using MCS lock so that only
+ one spinner can compete for the mutex.
+
+ The MCS lock (proposed by Mellor-Crummey and Scott) is a simple spinlock
+ with the desirable properties of being fair and with each cpu trying
+ to acquire the lock spinning on a local variable. It avoids expensive
+ cacheline bouncing that common test-and-set spinlock implementations
+ incur. An MCS-like lock is specially tailored for optimistic spinning
+ for sleeping lock implementation. An important feature of the customized
+ MCS lock is that it has the extra property that spinners are able to exit
+ the MCS spinlock queue when they need to reschedule. This further helps
+ avoid situations where MCS spinners that need to reschedule would continue
+ waiting to spin on mutex owner, only to go directly to slowpath upon
+ obtaining the MCS lock.
+
+
+(iii) slowpath: last resort, if the lock is still unable to be acquired,
+ the task is added to the wait-queue and sleeps until woken up by the
+ unlock path. Under normal circumstances it blocks as TASK_UNINTERRUPTIBLE.
+
+While formally kernel mutexes are sleepable locks, it is path (ii) that
+makes them more practically a hybrid type. By simply not interrupting a
+task and busy-waiting for a few cycles instead of immediately sleeping,
+the performance of this lock has been seen to significantly improve a
+number of workloads. Note that this technique is also used for rw-semaphores.
+
+Semantics
+---------
+
+The mutex subsystem checks and enforces the following rules:
+
+ - Only one task can hold the mutex at a time.
+ - Only the owner can unlock the mutex.
+ - Multiple unlocks are not permitted.
+ - Recursive locking/unlocking is not permitted.
+ - A mutex must only be initialized via the API (see below).
+ - A task may not exit with a mutex held.
+ - Memory areas where held locks reside must not be freed.
+ - Held mutexes must not be reinitialized.
+ - Mutexes may not be used in hardware or software interrupt
+ contexts such as tasklets and timers.
+
+These semantics are fully enforced when CONFIG DEBUG_MUTEXES is enabled.
+In addition, the mutex debugging code also implements a number of other
+features that make lock debugging easier and faster:
+
+ - Uses symbolic names of mutexes, whenever they are printed
+ in debug output.
+ - Point-of-acquire tracking, symbolic lookup of function names,
+ list of all locks held in the system, printout of them.
+ - Owner tracking.
+ - Detects self-recursing locks and prints out all relevant info.
+ - Detects multi-task circular deadlocks and prints out all affected
+ locks and tasks (and only those tasks).
+
+
+Interfaces
+----------
+Statically define the mutex:
+ DEFINE_MUTEX(name);
+
+Dynamically initialize the mutex:
+ mutex_init(mutex);
+
+Acquire the mutex, uninterruptible:
+ void mutex_lock(struct mutex *lock);
+ void mutex_lock_nested(struct mutex *lock, unsigned int subclass);
+ int mutex_trylock(struct mutex *lock);
+
+Acquire the mutex, interruptible:
+ int mutex_lock_interruptible_nested(struct mutex *lock,
+ unsigned int subclass);
+ int mutex_lock_interruptible(struct mutex *lock);
+
+Acquire the mutex, interruptible, if dec to 0:
+ int atomic_dec_and_mutex_lock(atomic_t *cnt, struct mutex *lock);
+
+Unlock the mutex:
+ void mutex_unlock(struct mutex *lock);
+
+Test if the mutex is taken:
+ int mutex_is_locked(struct mutex *lock);
+
+Disadvantages
+-------------
+
+Unlike its original design and purpose, 'struct mutex' is larger than
+most locks in the kernel. E.g: on x86-64 it is 40 bytes, almost twice
+as large as 'struct semaphore' (24 bytes) and 8 bytes shy of the
+'struct rw_semaphore' variant. Larger structure sizes mean more CPU
+cache and memory footprint.
+
+When to use mutexes
+-------------------
+
+Unless the strict semantics of mutexes are unsuitable and/or the critical
+region prevents the lock from being shared, always prefer them to any other
+locking primitive.
diff --git a/Documentation/locking/rt-mutex-design.txt b/Documentation/locking/rt-mutex-design.txt
new file mode 100644
index 0000000..8666070
--- /dev/null
+++ b/Documentation/locking/rt-mutex-design.txt
@@ -0,0 +1,781 @@
+#
+# Copyright (c) 2006 Steven Rostedt
+# Licensed under the GNU Free Documentation License, Version 1.2
+#
+
+RT-mutex implementation design
+------------------------------
+
+This document tries to describe the design of the rtmutex.c implementation.
+It doesn't describe the reasons why rtmutex.c exists. For that please see
+Documentation/rt-mutex.txt. Although this document does explain problems
+that happen without this code, but that is in the concept to understand
+what the code actually is doing.
+
+The goal of this document is to help others understand the priority
+inheritance (PI) algorithm that is used, as well as reasons for the
+decisions that were made to implement PI in the manner that was done.
+
+
+Unbounded Priority Inversion
+----------------------------
+
+Priority inversion is when a lower priority process executes while a higher
+priority process wants to run. This happens for several reasons, and
+most of the time it can't be helped. Anytime a high priority process wants
+to use a resource that a lower priority process has (a mutex for example),
+the high priority process must wait until the lower priority process is done
+with the resource. This is a priority inversion. What we want to prevent
+is something called unbounded priority inversion. That is when the high
+priority process is prevented from running by a lower priority process for
+an undetermined amount of time.
+
+The classic example of unbounded priority inversion is where you have three
+processes, let's call them processes A, B, and C, where A is the highest
+priority process, C is the lowest, and B is in between. A tries to grab a lock
+that C owns and must wait and lets C run to release the lock. But in the
+meantime, B executes, and since B is of a higher priority than C, it preempts C,
+but by doing so, it is in fact preempting A which is a higher priority process.
+Now there's no way of knowing how long A will be sleeping waiting for C
+to release the lock, because for all we know, B is a CPU hog and will
+never give C a chance to release the lock. This is called unbounded priority
+inversion.
+
+Here's a little ASCII art to show the problem.
+
+ grab lock L1 (owned by C)
+ |
+A ---+
+ C preempted by B
+ |
+C +----+
+
+B +-------->
+ B now keeps A from running.
+
+
+Priority Inheritance (PI)
+-------------------------
+
+There are several ways to solve this issue, but other ways are out of scope
+for this document. Here we only discuss PI.
+
+PI is where a process inherits the priority of another process if the other
+process blocks on a lock owned by the current process. To make this easier
+to understand, let's use the previous example, with processes A, B, and C again.
+
+This time, when A blocks on the lock owned by C, C would inherit the priority
+of A. So now if B becomes runnable, it would not preempt C, since C now has
+the high priority of A. As soon as C releases the lock, it loses its
+inherited priority, and A then can continue with the resource that C had.
+
+Terminology
+-----------
+
+Here I explain some terminology that is used in this document to help describe
+the design that is used to implement PI.
+
+PI chain - The PI chain is an ordered series of locks and processes that cause
+ processes to inherit priorities from a previous process that is
+ blocked on one of its locks. This is described in more detail
+ later in this document.
+
+mutex - In this document, to differentiate from locks that implement
+ PI and spin locks that are used in the PI code, from now on
+ the PI locks will be called a mutex.
+
+lock - In this document from now on, I will use the term lock when
+ referring to spin locks that are used to protect parts of the PI
+ algorithm. These locks disable preemption for UP (when
+ CONFIG_PREEMPT is enabled) and on SMP prevents multiple CPUs from
+ entering critical sections simultaneously.
+
+spin lock - Same as lock above.
+
+waiter - A waiter is a struct that is stored on the stack of a blocked
+ process. Since the scope of the waiter is within the code for
+ a process being blocked on the mutex, it is fine to allocate
+ the waiter on the process's stack (local variable). This
+ structure holds a pointer to the task, as well as the mutex that
+ the task is blocked on. It also has the plist node structures to
+ place the task in the waiter_list of a mutex as well as the
+ pi_list of a mutex owner task (described below).
+
+ waiter is sometimes used in reference to the task that is waiting
+ on a mutex. This is the same as waiter->task.
+
+waiters - A list of processes that are blocked on a mutex.
+
+top waiter - The highest priority process waiting on a specific mutex.
+
+top pi waiter - The highest priority process waiting on one of the mutexes
+ that a specific process owns.
+
+Note: task and process are used interchangeably in this document, mostly to
+ differentiate between two processes that are being described together.
+
+
+PI chain
+--------
+
+The PI chain is a list of processes and mutexes that may cause priority
+inheritance to take place. Multiple chains may converge, but a chain
+would never diverge, since a process can't be blocked on more than one
+mutex at a time.
+
+Example:
+
+ Process: A, B, C, D, E
+ Mutexes: L1, L2, L3, L4
+
+ A owns: L1
+ B blocked on L1
+ B owns L2
+ C blocked on L2
+ C owns L3
+ D blocked on L3
+ D owns L4
+ E blocked on L4
+
+The chain would be:
+
+ E->L4->D->L3->C->L2->B->L1->A
+
+To show where two chains merge, we could add another process F and
+another mutex L5 where B owns L5 and F is blocked on mutex L5.
+
+The chain for F would be:
+
+ F->L5->B->L1->A
+
+Since a process may own more than one mutex, but never be blocked on more than
+one, the chains merge.
+
+Here we show both chains:
+
+ E->L4->D->L3->C->L2-+
+ |
+ +->B->L1->A
+ |
+ F->L5-+
+
+For PI to work, the processes at the right end of these chains (or we may
+also call it the Top of the chain) must be equal to or higher in priority
+than the processes to the left or below in the chain.
+
+Also since a mutex may have more than one process blocked on it, we can
+have multiple chains merge at mutexes. If we add another process G that is
+blocked on mutex L2:
+
+ G->L2->B->L1->A
+
+And once again, to show how this can grow I will show the merging chains
+again.
+
+ E->L4->D->L3->C-+
+ +->L2-+
+ | |
+ G-+ +->B->L1->A
+ |
+ F->L5-+
+
+
+Plist
+-----
+
+Before I go further and talk about how the PI chain is stored through lists
+on both mutexes and processes, I'll explain the plist. This is similar to
+the struct list_head functionality that is already in the kernel.
+The implementation of plist is out of scope for this document, but it is
+very important to understand what it does.
+
+There are a few differences between plist and list, the most important one
+being that plist is a priority sorted linked list. This means that the
+priorities of the plist are sorted, such that it takes O(1) to retrieve the
+highest priority item in the list. Obviously this is useful to store processes
+based on their priorities.
+
+Another difference, which is important for implementation, is that, unlike
+list, the head of the list is a different element than the nodes of a list.
+So the head of the list is declared as struct plist_head and nodes that will
+be added to the list are declared as struct plist_node.
+
+
+Mutex Waiter List
+-----------------
+
+Every mutex keeps track of all the waiters that are blocked on itself. The mutex
+has a plist to store these waiters by priority. This list is protected by
+a spin lock that is located in the struct of the mutex. This lock is called
+wait_lock. Since the modification of the waiter list is never done in
+interrupt context, the wait_lock can be taken without disabling interrupts.
+
+
+Task PI List
+------------
+
+To keep track of the PI chains, each process has its own PI list. This is
+a list of all top waiters of the mutexes that are owned by the process.
+Note that this list only holds the top waiters and not all waiters that are
+blocked on mutexes owned by the process.
+
+The top of the task's PI list is always the highest priority task that
+is waiting on a mutex that is owned by the task. So if the task has
+inherited a priority, it will always be the priority of the task that is
+at the top of this list.
+
+This list is stored in the task structure of a process as a plist called
+pi_list. This list is protected by a spin lock also in the task structure,
+called pi_lock. This lock may also be taken in interrupt context, so when
+locking the pi_lock, interrupts must be disabled.
+
+
+Depth of the PI Chain
+---------------------
+
+The maximum depth of the PI chain is not dynamic, and could actually be
+defined. But is very complex to figure it out, since it depends on all
+the nesting of mutexes. Let's look at the example where we have 3 mutexes,
+L1, L2, and L3, and four separate functions func1, func2, func3 and func4.
+The following shows a locking order of L1->L2->L3, but may not actually
+be directly nested that way.
+
+void func1(void)
+{
+ mutex_lock(L1);
+
+ /* do anything */
+
+ mutex_unlock(L1);
+}
+
+void func2(void)
+{
+ mutex_lock(L1);
+ mutex_lock(L2);
+
+ /* do something */
+
+ mutex_unlock(L2);
+ mutex_unlock(L1);
+}
+
+void func3(void)
+{
+ mutex_lock(L2);
+ mutex_lock(L3);
+
+ /* do something else */
+
+ mutex_unlock(L3);
+ mutex_unlock(L2);
+}
+
+void func4(void)
+{
+ mutex_lock(L3);
+
+ /* do something again */
+
+ mutex_unlock(L3);
+}
+
+Now we add 4 processes that run each of these functions separately.
+Processes A, B, C, and D which run functions func1, func2, func3 and func4
+respectively, and such that D runs first and A last. With D being preempted
+in func4 in the "do something again" area, we have a locking that follows:
+
+D owns L3
+ C blocked on L3
+ C owns L2
+ B blocked on L2
+ B owns L1
+ A blocked on L1
+
+And thus we have the chain A->L1->B->L2->C->L3->D.
+
+This gives us a PI depth of 4 (four processes), but looking at any of the
+functions individually, it seems as though they only have at most a locking
+depth of two. So, although the locking depth is defined at compile time,
+it still is very difficult to find the possibilities of that depth.
+
+Now since mutexes can be defined by user-land applications, we don't want a DOS
+type of application that nests large amounts of mutexes to create a large
+PI chain, and have the code holding spin locks while looking at a large
+amount of data. So to prevent this, the implementation not only implements
+a maximum lock depth, but also only holds at most two different locks at a
+time, as it walks the PI chain. More about this below.
+
+
+Mutex owner and flags
+---------------------
+
+The mutex structure contains a pointer to the owner of the mutex. If the
+mutex is not owned, this owner is set to NULL. Since all architectures
+have the task structure on at least a four byte alignment (and if this is
+not true, the rtmutex.c code will be broken!), this allows for the two
+least significant bits to be used as flags. This part is also described
+in Documentation/rt-mutex.txt, but will also be briefly described here.
+
+Bit 0 is used as the "Pending Owner" flag. This is described later.
+Bit 1 is used as the "Has Waiters" flags. This is also described later
+ in more detail, but is set whenever there are waiters on a mutex.
+
+
+cmpxchg Tricks
+--------------
+
+Some architectures implement an atomic cmpxchg (Compare and Exchange). This
+is used (when applicable) to keep the fast path of grabbing and releasing
+mutexes short.
+
+cmpxchg is basically the following function performed atomically:
+
+unsigned long _cmpxchg(unsigned long *A, unsigned long *B, unsigned long *C)
+{
+ unsigned long T = *A;
+ if (*A == *B) {
+ *A = *C;
+ }
+ return T;
+}
+#define cmpxchg(a,b,c) _cmpxchg(&a,&b,&c)
+
+This is really nice to have, since it allows you to only update a variable
+if the variable is what you expect it to be. You know if it succeeded if
+the return value (the old value of A) is equal to B.
+
+The macro rt_mutex_cmpxchg is used to try to lock and unlock mutexes. If
+the architecture does not support CMPXCHG, then this macro is simply set
+to fail every time. But if CMPXCHG is supported, then this will
+help out extremely to keep the fast path short.
+
+The use of rt_mutex_cmpxchg with the flags in the owner field help optimize
+the system for architectures that support it. This will also be explained
+later in this document.
+
+
+Priority adjustments
+--------------------
+
+The implementation of the PI code in rtmutex.c has several places that a
+process must adjust its priority. With the help of the pi_list of a
+process this is rather easy to know what needs to be adjusted.
+
+The functions implementing the task adjustments are rt_mutex_adjust_prio,
+__rt_mutex_adjust_prio (same as the former, but expects the task pi_lock
+to already be taken), rt_mutex_getprio, and rt_mutex_setprio.
+
+rt_mutex_getprio and rt_mutex_setprio are only used in __rt_mutex_adjust_prio.
+
+rt_mutex_getprio returns the priority that the task should have. Either the
+task's own normal priority, or if a process of a higher priority is waiting on
+a mutex owned by the task, then that higher priority should be returned.
+Since the pi_list of a task holds an order by priority list of all the top
+waiters of all the mutexes that the task owns, rt_mutex_getprio simply needs
+to compare the top pi waiter to its own normal priority, and return the higher
+priority back.
+
+(Note: if looking at the code, you will notice that the lower number of
+ prio is returned. This is because the prio field in the task structure
+ is an inverse order of the actual priority. So a "prio" of 5 is
+ of higher priority than a "prio" of 10.)
+
+__rt_mutex_adjust_prio examines the result of rt_mutex_getprio, and if the
+result does not equal the task's current priority, then rt_mutex_setprio
+is called to adjust the priority of the task to the new priority.
+Note that rt_mutex_setprio is defined in kernel/sched/core.c to implement the
+actual change in priority.
+
+It is interesting to note that __rt_mutex_adjust_prio can either increase
+or decrease the priority of the task. In the case that a higher priority
+process has just blocked on a mutex owned by the task, __rt_mutex_adjust_prio
+would increase/boost the task's priority. But if a higher priority task
+were for some reason to leave the mutex (timeout or signal), this same function
+would decrease/unboost the priority of the task. That is because the pi_list
+always contains the highest priority task that is waiting on a mutex owned
+by the task, so we only need to compare the priority of that top pi waiter
+to the normal priority of the given task.
+
+
+High level overview of the PI chain walk
+----------------------------------------
+
+The PI chain walk is implemented by the function rt_mutex_adjust_prio_chain.
+
+The implementation has gone through several iterations, and has ended up
+with what we believe is the best. It walks the PI chain by only grabbing
+at most two locks at a time, and is very efficient.
+
+The rt_mutex_adjust_prio_chain can be used either to boost or lower process
+priorities.
+
+rt_mutex_adjust_prio_chain is called with a task to be checked for PI
+(de)boosting (the owner of a mutex that a process is blocking on), a flag to
+check for deadlocking, the mutex that the task owns, and a pointer to a waiter
+that is the process's waiter struct that is blocked on the mutex (although this
+parameter may be NULL for deboosting).
+
+For this explanation, I will not mention deadlock detection. This explanation
+will try to stay at a high level.
+
+When this function is called, there are no locks held. That also means
+that the state of the owner and lock can change when entered into this function.
+
+Before this function is called, the task has already had rt_mutex_adjust_prio
+performed on it. This means that the task is set to the priority that it
+should be at, but the plist nodes of the task's waiter have not been updated
+with the new priorities, and that this task may not be in the proper locations
+in the pi_lists and wait_lists that the task is blocked on. This function
+solves all that.
+
+A loop is entered, where task is the owner to be checked for PI changes that
+was passed by parameter (for the first iteration). The pi_lock of this task is
+taken to prevent any more changes to the pi_list of the task. This also
+prevents new tasks from completing the blocking on a mutex that is owned by this
+task.
+
+If the task is not blocked on a mutex then the loop is exited. We are at
+the top of the PI chain.
+
+A check is now done to see if the original waiter (the process that is blocked
+on the current mutex) is the top pi waiter of the task. That is, is this
+waiter on the top of the task's pi_list. If it is not, it either means that
+there is another process higher in priority that is blocked on one of the
+mutexes that the task owns, or that the waiter has just woken up via a signal
+or timeout and has left the PI chain. In either case, the loop is exited, since
+we don't need to do any more changes to the priority of the current task, or any
+task that owns a mutex that this current task is waiting on. A priority chain
+walk is only needed when a new top pi waiter is made to a task.
+
+The next check sees if the task's waiter plist node has the priority equal to
+the priority the task is set at. If they are equal, then we are done with
+the loop. Remember that the function started with the priority of the
+task adjusted, but the plist nodes that hold the task in other processes
+pi_lists have not been adjusted.
+
+Next, we look at the mutex that the task is blocked on. The mutex's wait_lock
+is taken. This is done by a spin_trylock, because the locking order of the
+pi_lock and wait_lock goes in the opposite direction. If we fail to grab the
+lock, the pi_lock is released, and we restart the loop.
+
+Now that we have both the pi_lock of the task as well as the wait_lock of
+the mutex the task is blocked on, we update the task's waiter's plist node
+that is located on the mutex's wait_list.
+
+Now we release the pi_lock of the task.
+
+Next the owner of the mutex has its pi_lock taken, so we can update the
+task's entry in the owner's pi_list. If the task is the highest priority
+process on the mutex's wait_list, then we remove the previous top waiter
+from the owner's pi_list, and replace it with the task.
+
+Note: It is possible that the task was the current top waiter on the mutex,
+ in which case the task is not yet on the pi_list of the waiter. This
+ is OK, since plist_del does nothing if the plist node is not on any
+ list.
+
+If the task was not the top waiter of the mutex, but it was before we
+did the priority updates, that means we are deboosting/lowering the
+task. In this case, the task is removed from the pi_list of the owner,
+and the new top waiter is added.
+
+Lastly, we unlock both the pi_lock of the task, as well as the mutex's
+wait_lock, and continue the loop again. On the next iteration of the
+loop, the previous owner of the mutex will be the task that will be
+processed.
+
+Note: One might think that the owner of this mutex might have changed
+ since we just grab the mutex's wait_lock. And one could be right.
+ The important thing to remember is that the owner could not have
+ become the task that is being processed in the PI chain, since
+ we have taken that task's pi_lock at the beginning of the loop.
+ So as long as there is an owner of this mutex that is not the same
+ process as the tasked being worked on, we are OK.
+
+ Looking closely at the code, one might be confused. The check for the
+ end of the PI chain is when the task isn't blocked on anything or the
+ task's waiter structure "task" element is NULL. This check is
+ protected only by the task's pi_lock. But the code to unlock the mutex
+ sets the task's waiter structure "task" element to NULL with only
+ the protection of the mutex's wait_lock, which was not taken yet.
+ Isn't this a race condition if the task becomes the new owner?
+
+ The answer is No! The trick is the spin_trylock of the mutex's
+ wait_lock. If we fail that lock, we release the pi_lock of the
+ task and continue the loop, doing the end of PI chain check again.
+
+ In the code to release the lock, the wait_lock of the mutex is held
+ the entire time, and it is not let go when we grab the pi_lock of the
+ new owner of the mutex. So if the switch of a new owner were to happen
+ after the check for end of the PI chain and the grabbing of the
+ wait_lock, the unlocking code would spin on the new owner's pi_lock
+ but never give up the wait_lock. So the PI chain loop is guaranteed to
+ fail the spin_trylock on the wait_lock, release the pi_lock, and
+ try again.
+
+ If you don't quite understand the above, that's OK. You don't have to,
+ unless you really want to make a proof out of it ;)
+
+
+Pending Owners and Lock stealing
+--------------------------------
+
+One of the flags in the owner field of the mutex structure is "Pending Owner".
+What this means is that an owner was chosen by the process releasing the
+mutex, but that owner has yet to wake up and actually take the mutex.
+
+Why is this important? Why can't we just give the mutex to another process
+and be done with it?
+
+The PI code is to help with real-time processes, and to let the highest
+priority process run as long as possible with little latencies and delays.
+If a high priority process owns a mutex that a lower priority process is
+blocked on, when the mutex is released it would be given to the lower priority
+process. What if the higher priority process wants to take that mutex again.
+The high priority process would fail to take that mutex that it just gave up
+and it would need to boost the lower priority process to run with full
+latency of that critical section (since the low priority process just entered
+it).
+
+There's no reason a high priority process that gives up a mutex should be
+penalized if it tries to take that mutex again. If the new owner of the
+mutex has not woken up yet, there's no reason that the higher priority process
+could not take that mutex away.
+
+To solve this, we introduced Pending Ownership and Lock Stealing. When a
+new process is given a mutex that it was blocked on, it is only given
+pending ownership. This means that it's the new owner, unless a higher
+priority process comes in and tries to grab that mutex. If a higher priority
+process does come along and wants that mutex, we let the higher priority
+process "steal" the mutex from the pending owner (only if it is still pending)
+and continue with the mutex.
+
+
+Taking of a mutex (The walk through)
+------------------------------------
+
+OK, now let's take a look at the detailed walk through of what happens when
+taking a mutex.
+
+The first thing that is tried is the fast taking of the mutex. This is
+done when we have CMPXCHG enabled (otherwise the fast taking automatically
+fails). Only when the owner field of the mutex is NULL can the lock be
+taken with the CMPXCHG and nothing else needs to be done.
+
+If there is contention on the lock, whether it is owned or pending owner
+we go about the slow path (rt_mutex_slowlock).
+
+The slow path function is where the task's waiter structure is created on
+the stack. This is because the waiter structure is only needed for the
+scope of this function. The waiter structure holds the nodes to store
+the task on the wait_list of the mutex, and if need be, the pi_list of
+the owner.
+
+The wait_lock of the mutex is taken since the slow path of unlocking the
+mutex also takes this lock.
+
+We then call try_to_take_rt_mutex. This is where the architecture that
+does not implement CMPXCHG would always grab the lock (if there's no
+contention).
+
+try_to_take_rt_mutex is used every time the task tries to grab a mutex in the
+slow path. The first thing that is done here is an atomic setting of
+the "Has Waiters" flag of the mutex's owner field. Yes, this could really
+be false, because if the mutex has no owner, there are no waiters and
+the current task also won't have any waiters. But we don't have the lock
+yet, so we assume we are going to be a waiter. The reason for this is to
+play nice for those architectures that do have CMPXCHG. By setting this flag
+now, the owner of the mutex can't release the mutex without going into the
+slow unlock path, and it would then need to grab the wait_lock, which this
+code currently holds. So setting the "Has Waiters" flag forces the owner
+to synchronize with this code.
+
+Now that we know that we can't have any races with the owner releasing the
+mutex, we check to see if we can take the ownership. This is done if the
+mutex doesn't have a owner, or if we can steal the mutex from a pending
+owner. Let's look at the situations we have here.
+
+ 1) Has owner that is pending
+ ----------------------------
+
+ The mutex has a owner, but it hasn't woken up and the mutex flag
+ "Pending Owner" is set. The first check is to see if the owner isn't the
+ current task. This is because this function is also used for the pending
+ owner to grab the mutex. When a pending owner wakes up, it checks to see
+ if it can take the mutex, and this is done if the owner is already set to
+ itself. If so, we succeed and leave the function, clearing the "Pending
+ Owner" bit.
+
+ If the pending owner is not current, we check to see if the current priority is
+ higher than the pending owner. If not, we fail the function and return.
+
+ There's also something special about a pending owner. That is a pending owner
+ is never blocked on a mutex. So there is no PI chain to worry about. It also
+ means that if the mutex doesn't have any waiters, there's no accounting needed
+ to update the pending owner's pi_list, since we only worry about processes
+ blocked on the current mutex.
+
+ If there are waiters on this mutex, and we just stole the ownership, we need
+ to take the top waiter, remove it from the pi_list of the pending owner, and
+ add it to the current pi_list. Note that at this moment, the pending owner
+ is no longer on the list of waiters. This is fine, since the pending owner
+ would add itself back when it realizes that it had the ownership stolen
+ from itself. When the pending owner tries to grab the mutex, it will fail
+ in try_to_take_rt_mutex if the owner field points to another process.
+
+ 2) No owner
+ -----------
+
+ If there is no owner (or we successfully stole the lock), we set the owner
+ of the mutex to current, and set the flag of "Has Waiters" if the current
+ mutex actually has waiters, or we clear the flag if it doesn't. See, it was
+ OK that we set that flag early, since now it is cleared.
+
+ 3) Failed to grab ownership
+ ---------------------------
+
+ The most interesting case is when we fail to take ownership. This means that
+ there exists an owner, or there's a pending owner with equal or higher
+ priority than the current task.
+
+We'll continue on the failed case.
+
+If the mutex has a timeout, we set up a timer to go off to break us out
+of this mutex if we failed to get it after a specified amount of time.
+
+Now we enter a loop that will continue to try to take ownership of the mutex, or
+fail from a timeout or signal.
+
+Once again we try to take the mutex. This will usually fail the first time
+in the loop, since it had just failed to get the mutex. But the second time
+in the loop, this would likely succeed, since the task would likely be
+the pending owner.
+
+If the mutex is TASK_INTERRUPTIBLE a check for signals and timeout is done
+here.
+
+The waiter structure has a "task" field that points to the task that is blocked
+on the mutex. This field can be NULL the first time it goes through the loop
+or if the task is a pending owner and had its mutex stolen. If the "task"
+field is NULL then we need to set up the accounting for it.
+
+Task blocks on mutex
+--------------------
+
+The accounting of a mutex and process is done with the waiter structure of
+the process. The "task" field is set to the process, and the "lock" field
+to the mutex. The plist nodes are initialized to the processes current
+priority.
+
+Since the wait_lock was taken at the entry of the slow lock, we can safely
+add the waiter to the wait_list. If the current process is the highest
+priority process currently waiting on this mutex, then we remove the
+previous top waiter process (if it exists) from the pi_list of the owner,
+and add the current process to that list. Since the pi_list of the owner
+has changed, we call rt_mutex_adjust_prio on the owner to see if the owner
+should adjust its priority accordingly.
+
+If the owner is also blocked on a lock, and had its pi_list changed
+(or deadlock checking is on), we unlock the wait_lock of the mutex and go ahead
+and run rt_mutex_adjust_prio_chain on the owner, as described earlier.
+
+Now all locks are released, and if the current process is still blocked on a
+mutex (waiter "task" field is not NULL), then we go to sleep (call schedule).
+
+Waking up in the loop
+---------------------
+
+The schedule can then wake up for a few reasons.
+ 1) we were given pending ownership of the mutex.
+ 2) we received a signal and was TASK_INTERRUPTIBLE
+ 3) we had a timeout and was TASK_INTERRUPTIBLE
+
+In any of these cases, we continue the loop and once again try to grab the
+ownership of the mutex. If we succeed, we exit the loop, otherwise we continue
+and on signal and timeout, will exit the loop, or if we had the mutex stolen
+we just simply add ourselves back on the lists and go back to sleep.
+
+Note: For various reasons, because of timeout and signals, the steal mutex
+ algorithm needs to be careful. This is because the current process is
+ still on the wait_list. And because of dynamic changing of priorities,
+ especially on SCHED_OTHER tasks, the current process can be the
+ highest priority task on the wait_list.
+
+Failed to get mutex on Timeout or Signal
+----------------------------------------
+
+If a timeout or signal occurred, the waiter's "task" field would not be
+NULL and the task needs to be taken off the wait_list of the mutex and perhaps
+pi_list of the owner. If this process was a high priority process, then
+the rt_mutex_adjust_prio_chain needs to be executed again on the owner,
+but this time it will be lowering the priorities.
+
+
+Unlocking the Mutex
+-------------------
+
+The unlocking of a mutex also has a fast path for those architectures with
+CMPXCHG. Since the taking of a mutex on contention always sets the
+"Has Waiters" flag of the mutex's owner, we use this to know if we need to
+take the slow path when unlocking the mutex. If the mutex doesn't have any
+waiters, the owner field of the mutex would equal the current process and
+the mutex can be unlocked by just replacing the owner field with NULL.
+
+If the owner field has the "Has Waiters" bit set (or CMPXCHG is not available),
+the slow unlock path is taken.
+
+The first thing done in the slow unlock path is to take the wait_lock of the
+mutex. This synchronizes the locking and unlocking of the mutex.
+
+A check is made to see if the mutex has waiters or not. On architectures that
+do not have CMPXCHG, this is the location that the owner of the mutex will
+determine if a waiter needs to be awoken or not. On architectures that
+do have CMPXCHG, that check is done in the fast path, but it is still needed
+in the slow path too. If a waiter of a mutex woke up because of a signal
+or timeout between the time the owner failed the fast path CMPXCHG check and
+the grabbing of the wait_lock, the mutex may not have any waiters, thus the
+owner still needs to make this check. If there are no waiters then the mutex
+owner field is set to NULL, the wait_lock is released and nothing more is
+needed.
+
+If there are waiters, then we need to wake one up and give that waiter
+pending ownership.
+
+On the wake up code, the pi_lock of the current owner is taken. The top
+waiter of the lock is found and removed from the wait_list of the mutex
+as well as the pi_list of the current owner. The task field of the new
+pending owner's waiter structure is set to NULL, and the owner field of the
+mutex is set to the new owner with the "Pending Owner" bit set, as well
+as the "Has Waiters" bit if there still are other processes blocked on the
+mutex.
+
+The pi_lock of the previous owner is released, and the new pending owner's
+pi_lock is taken. Remember that this is the trick to prevent the race
+condition in rt_mutex_adjust_prio_chain from adding itself as a waiter
+on the mutex.
+
+We now clear the "pi_blocked_on" field of the new pending owner, and if
+the mutex still has waiters pending, we add the new top waiter to the pi_list
+of the pending owner.
+
+Finally we unlock the pi_lock of the pending owner and wake it up.
+
+
+Contact
+-------
+
+For updates on this document, please email Steven Rostedt <rostedt@goodmis.org>
+
+
+Credits
+-------
+
+Author: Steven Rostedt <rostedt@goodmis.org>
+
+Reviewers: Ingo Molnar, Thomas Gleixner, Thomas Duetsch, and Randy Dunlap
+
+Updates
+-------
+
+This document was originally written for 2.6.17-rc3-mm1
diff --git a/Documentation/locking/rt-mutex.txt b/Documentation/locking/rt-mutex.txt
new file mode 100644
index 0000000..243393d
--- /dev/null
+++ b/Documentation/locking/rt-mutex.txt
@@ -0,0 +1,79 @@
+RT-mutex subsystem with PI support
+----------------------------------
+
+RT-mutexes with priority inheritance are used to support PI-futexes,
+which enable pthread_mutex_t priority inheritance attributes
+(PTHREAD_PRIO_INHERIT). [See Documentation/pi-futex.txt for more details
+about PI-futexes.]
+
+This technology was developed in the -rt tree and streamlined for
+pthread_mutex support.
+
+Basic principles:
+-----------------
+
+RT-mutexes extend the semantics of simple mutexes by the priority
+inheritance protocol.
+
+A low priority owner of a rt-mutex inherits the priority of a higher
+priority waiter until the rt-mutex is released. If the temporarily
+boosted owner blocks on a rt-mutex itself it propagates the priority
+boosting to the owner of the other rt_mutex it gets blocked on. The
+priority boosting is immediately removed once the rt_mutex has been
+unlocked.
+
+This approach allows us to shorten the block of high-prio tasks on
+mutexes which protect shared resources. Priority inheritance is not a
+magic bullet for poorly designed applications, but it allows
+well-designed applications to use userspace locks in critical parts of
+an high priority thread, without losing determinism.
+
+The enqueueing of the waiters into the rtmutex waiter list is done in
+priority order. For same priorities FIFO order is chosen. For each
+rtmutex, only the top priority waiter is enqueued into the owner's
+priority waiters list. This list too queues in priority order. Whenever
+the top priority waiter of a task changes (for example it timed out or
+got a signal), the priority of the owner task is readjusted. [The
+priority enqueueing is handled by "plists", see include/linux/plist.h
+for more details.]
+
+RT-mutexes are optimized for fastpath operations and have no internal
+locking overhead when locking an uncontended mutex or unlocking a mutex
+without waiters. The optimized fastpath operations require cmpxchg
+support. [If that is not available then the rt-mutex internal spinlock
+is used]
+
+The state of the rt-mutex is tracked via the owner field of the rt-mutex
+structure:
+
+rt_mutex->owner holds the task_struct pointer of the owner. Bit 0 and 1
+are used to keep track of the "owner is pending" and "rtmutex has
+waiters" state.
+
+ owner bit1 bit0
+ NULL 0 0 mutex is free (fast acquire possible)
+ NULL 0 1 invalid state
+ NULL 1 0 Transitional state*
+ NULL 1 1 invalid state
+ taskpointer 0 0 mutex is held (fast release possible)
+ taskpointer 0 1 task is pending owner
+ taskpointer 1 0 mutex is held and has waiters
+ taskpointer 1 1 task is pending owner and mutex has waiters
+
+Pending-ownership handling is a performance optimization:
+pending-ownership is assigned to the first (highest priority) waiter of
+the mutex, when the mutex is released. The thread is woken up and once
+it starts executing it can acquire the mutex. Until the mutex is taken
+by it (bit 0 is cleared) a competing higher priority thread can "steal"
+the mutex which puts the woken up thread back on the waiters list.
+
+The pending-ownership optimization is especially important for the
+uninterrupted workflow of high-prio tasks which repeatedly
+takes/releases locks that have lower-prio waiters. Without this
+optimization the higher-prio thread would ping-pong to the lower-prio
+task [because at unlock time we always assign a new owner].
+
+(*) The "mutex has waiters" bit gets set to take the lock. If the lock
+doesn't already have an owner, this bit is quickly cleared if there are
+no waiters. So this is a transitional state to synchronize with looking
+at the owner field of the mutex and the mutex owner releasing the lock.
diff --git a/Documentation/locking/spinlocks.txt b/Documentation/locking/spinlocks.txt
new file mode 100644
index 0000000..97eaf57
--- /dev/null
+++ b/Documentation/locking/spinlocks.txt
@@ -0,0 +1,167 @@
+Lesson 1: Spin locks
+
+The most basic primitive for locking is spinlock.
+
+static DEFINE_SPINLOCK(xxx_lock);
+
+ unsigned long flags;
+
+ spin_lock_irqsave(&xxx_lock, flags);
+ ... critical section here ..
+ spin_unlock_irqrestore(&xxx_lock, flags);
+
+The above is always safe. It will disable interrupts _locally_, but the
+spinlock itself will guarantee the global lock, so it will guarantee that
+there is only one thread-of-control within the region(s) protected by that
+lock. This works well even under UP also, so the code does _not_ need to
+worry about UP vs SMP issues: the spinlocks work correctly under both.
+
+ NOTE! Implications of spin_locks for memory are further described in:
+
+ Documentation/memory-barriers.txt
+ (5) LOCK operations.
+ (6) UNLOCK operations.
+
+The above is usually pretty simple (you usually need and want only one
+spinlock for most things - using more than one spinlock can make things a
+lot more complex and even slower and is usually worth it only for
+sequences that you _know_ need to be split up: avoid it at all cost if you
+aren't sure).
+
+This is really the only really hard part about spinlocks: once you start
+using spinlocks they tend to expand to areas you might not have noticed
+before, because you have to make sure the spinlocks correctly protect the
+shared data structures _everywhere_ they are used. The spinlocks are most
+easily added to places that are completely independent of other code (for
+example, internal driver data structures that nobody else ever touches).
+
+ NOTE! The spin-lock is safe only when you _also_ use the lock itself
+ to do locking across CPU's, which implies that EVERYTHING that
+ touches a shared variable has to agree about the spinlock they want
+ to use.
+
+----
+
+Lesson 2: reader-writer spinlocks.
+
+If your data accesses have a very natural pattern where you usually tend
+to mostly read from the shared variables, the reader-writer locks
+(rw_lock) versions of the spinlocks are sometimes useful. They allow multiple
+readers to be in the same critical region at once, but if somebody wants
+to change the variables it has to get an exclusive write lock.
+
+ NOTE! reader-writer locks require more atomic memory operations than
+ simple spinlocks. Unless the reader critical section is long, you
+ are better off just using spinlocks.
+
+The routines look the same as above:
+
+ rwlock_t xxx_lock = __RW_LOCK_UNLOCKED(xxx_lock);
+
+ unsigned long flags;
+
+ read_lock_irqsave(&xxx_lock, flags);
+ .. critical section that only reads the info ...
+ read_unlock_irqrestore(&xxx_lock, flags);
+
+ write_lock_irqsave(&xxx_lock, flags);
+ .. read and write exclusive access to the info ...
+ write_unlock_irqrestore(&xxx_lock, flags);
+
+The above kind of lock may be useful for complex data structures like
+linked lists, especially searching for entries without changing the list
+itself. The read lock allows many concurrent readers. Anything that
+_changes_ the list will have to get the write lock.
+
+ NOTE! RCU is better for list traversal, but requires careful
+ attention to design detail (see Documentation/RCU/listRCU.txt).
+
+Also, you cannot "upgrade" a read-lock to a write-lock, so if you at _any_
+time need to do any changes (even if you don't do it every time), you have
+to get the write-lock at the very beginning.
+
+ NOTE! We are working hard to remove reader-writer spinlocks in most
+ cases, so please don't add a new one without consensus. (Instead, see
+ Documentation/RCU/rcu.txt for complete information.)
+
+----
+
+Lesson 3: spinlocks revisited.
+
+The single spin-lock primitives above are by no means the only ones. They
+are the most safe ones, and the ones that work under all circumstances,
+but partly _because_ they are safe they are also fairly slow. They are slower
+than they'd need to be, because they do have to disable interrupts
+(which is just a single instruction on a x86, but it's an expensive one -
+and on other architectures it can be worse).
+
+If you have a case where you have to protect a data structure across
+several CPU's and you want to use spinlocks you can potentially use
+cheaper versions of the spinlocks. IFF you know that the spinlocks are
+never used in interrupt handlers, you can use the non-irq versions:
+
+ spin_lock(&lock);
+ ...
+ spin_unlock(&lock);
+
+(and the equivalent read-write versions too, of course). The spinlock will
+guarantee the same kind of exclusive access, and it will be much faster.
+This is useful if you know that the data in question is only ever
+manipulated from a "process context", ie no interrupts involved.
+
+The reasons you mustn't use these versions if you have interrupts that
+play with the spinlock is that you can get deadlocks:
+
+ spin_lock(&lock);
+ ...
+ <- interrupt comes in:
+ spin_lock(&lock);
+
+where an interrupt tries to lock an already locked variable. This is ok if
+the other interrupt happens on another CPU, but it is _not_ ok if the
+interrupt happens on the same CPU that already holds the lock, because the
+lock will obviously never be released (because the interrupt is waiting
+for the lock, and the lock-holder is interrupted by the interrupt and will
+not continue until the interrupt has been processed).
+
+(This is also the reason why the irq-versions of the spinlocks only need
+to disable the _local_ interrupts - it's ok to use spinlocks in interrupts
+on other CPU's, because an interrupt on another CPU doesn't interrupt the
+CPU that holds the lock, so the lock-holder can continue and eventually
+releases the lock).
+
+Note that you can be clever with read-write locks and interrupts. For
+example, if you know that the interrupt only ever gets a read-lock, then
+you can use a non-irq version of read locks everywhere - because they
+don't block on each other (and thus there is no dead-lock wrt interrupts.
+But when you do the write-lock, you have to use the irq-safe version.
+
+For an example of being clever with rw-locks, see the "waitqueue_lock"
+handling in kernel/sched/core.c - nothing ever _changes_ a wait-queue from
+within an interrupt, they only read the queue in order to know whom to
+wake up. So read-locks are safe (which is good: they are very common
+indeed), while write-locks need to protect themselves against interrupts.
+
+ Linus
+
+----
+
+Reference information:
+
+For dynamic initialization, use spin_lock_init() or rwlock_init() as
+appropriate:
+
+ spinlock_t xxx_lock;
+ rwlock_t xxx_rw_lock;
+
+ static int __init xxx_init(void)
+ {
+ spin_lock_init(&xxx_lock);
+ rwlock_init(&xxx_rw_lock);
+ ...
+ }
+
+ module_init(xxx_init);
+
+For static initialization, use DEFINE_SPINLOCK() / DEFINE_RWLOCK() or
+__SPIN_LOCK_UNLOCKED() / __RW_LOCK_UNLOCKED() as appropriate.
diff --git a/Documentation/locking/ww-mutex-design.txt b/Documentation/locking/ww-mutex-design.txt
new file mode 100644
index 0000000..8a112dc
--- /dev/null
+++ b/Documentation/locking/ww-mutex-design.txt
@@ -0,0 +1,344 @@
+Wait/Wound Deadlock-Proof Mutex Design
+======================================
+
+Please read mutex-design.txt first, as it applies to wait/wound mutexes too.
+
+Motivation for WW-Mutexes
+-------------------------
+
+GPU's do operations that commonly involve many buffers. Those buffers
+can be shared across contexts/processes, exist in different memory
+domains (for example VRAM vs system memory), and so on. And with
+PRIME / dmabuf, they can even be shared across devices. So there are
+a handful of situations where the driver needs to wait for buffers to
+become ready. If you think about this in terms of waiting on a buffer
+mutex for it to become available, this presents a problem because
+there is no way to guarantee that buffers appear in a execbuf/batch in
+the same order in all contexts. That is directly under control of
+userspace, and a result of the sequence of GL calls that an application
+makes. Which results in the potential for deadlock. The problem gets
+more complex when you consider that the kernel may need to migrate the
+buffer(s) into VRAM before the GPU operates on the buffer(s), which
+may in turn require evicting some other buffers (and you don't want to
+evict other buffers which are already queued up to the GPU), but for a
+simplified understanding of the problem you can ignore this.
+
+The algorithm that the TTM graphics subsystem came up with for dealing with
+this problem is quite simple. For each group of buffers (execbuf) that need
+to be locked, the caller would be assigned a unique reservation id/ticket,
+from a global counter. In case of deadlock while locking all the buffers
+associated with a execbuf, the one with the lowest reservation ticket (i.e.
+the oldest task) wins, and the one with the higher reservation id (i.e. the
+younger task) unlocks all of the buffers that it has already locked, and then
+tries again.
+
+In the RDBMS literature this deadlock handling approach is called wait/wound:
+The older tasks waits until it can acquire the contended lock. The younger tasks
+needs to back off and drop all the locks it is currently holding, i.e. the
+younger task is wounded.
+
+Concepts
+--------
+
+Compared to normal mutexes two additional concepts/objects show up in the lock
+interface for w/w mutexes:
+
+Acquire context: To ensure eventual forward progress it is important the a task
+trying to acquire locks doesn't grab a new reservation id, but keeps the one it
+acquired when starting the lock acquisition. This ticket is stored in the
+acquire context. Furthermore the acquire context keeps track of debugging state
+to catch w/w mutex interface abuse.
+
+W/w class: In contrast to normal mutexes the lock class needs to be explicit for
+w/w mutexes, since it is required to initialize the acquire context.
+
+Furthermore there are three different class of w/w lock acquire functions:
+
+* Normal lock acquisition with a context, using ww_mutex_lock.
+
+* Slowpath lock acquisition on the contending lock, used by the wounded task
+ after having dropped all already acquired locks. These functions have the
+ _slow postfix.
+
+ From a simple semantics point-of-view the _slow functions are not strictly
+ required, since simply calling the normal ww_mutex_lock functions on the
+ contending lock (after having dropped all other already acquired locks) will
+ work correctly. After all if no other ww mutex has been acquired yet there's
+ no deadlock potential and hence the ww_mutex_lock call will block and not
+ prematurely return -EDEADLK. The advantage of the _slow functions is in
+ interface safety:
+ - ww_mutex_lock has a __must_check int return type, whereas ww_mutex_lock_slow
+ has a void return type. Note that since ww mutex code needs loops/retries
+ anyway the __must_check doesn't result in spurious warnings, even though the
+ very first lock operation can never fail.
+ - When full debugging is enabled ww_mutex_lock_slow checks that all acquired
+ ww mutex have been released (preventing deadlocks) and makes sure that we
+ block on the contending lock (preventing spinning through the -EDEADLK
+ slowpath until the contended lock can be acquired).
+
+* Functions to only acquire a single w/w mutex, which results in the exact same
+ semantics as a normal mutex. This is done by calling ww_mutex_lock with a NULL
+ context.
+
+ Again this is not strictly required. But often you only want to acquire a
+ single lock in which case it's pointless to set up an acquire context (and so
+ better to avoid grabbing a deadlock avoidance ticket).
+
+Of course, all the usual variants for handling wake-ups due to signals are also
+provided.
+
+Usage
+-----
+
+Three different ways to acquire locks within the same w/w class. Common
+definitions for methods #1 and #2:
+
+static DEFINE_WW_CLASS(ww_class);
+
+struct obj {
+ struct ww_mutex lock;
+ /* obj data */
+};
+
+struct obj_entry {
+ struct list_head head;
+ struct obj *obj;
+};
+
+Method 1, using a list in execbuf->buffers that's not allowed to be reordered.
+This is useful if a list of required objects is already tracked somewhere.
+Furthermore the lock helper can use propagate the -EALREADY return code back to
+the caller as a signal that an object is twice on the list. This is useful if
+the list is constructed from userspace input and the ABI requires userspace to
+not have duplicate entries (e.g. for a gpu commandbuffer submission ioctl).
+
+int lock_objs(struct list_head *list, struct ww_acquire_ctx *ctx)
+{
+ struct obj *res_obj = NULL;
+ struct obj_entry *contended_entry = NULL;
+ struct obj_entry *entry;
+
+ ww_acquire_init(ctx, &ww_class);
+
+retry:
+ list_for_each_entry (entry, list, head) {
+ if (entry->obj == res_obj) {
+ res_obj = NULL;
+ continue;
+ }
+ ret = ww_mutex_lock(&entry->obj->lock, ctx);
+ if (ret < 0) {
+ contended_entry = entry;
+ goto err;
+ }
+ }
+
+ ww_acquire_done(ctx);
+ return 0;
+
+err:
+ list_for_each_entry_continue_reverse (entry, list, head)
+ ww_mutex_unlock(&entry->obj->lock);
+
+ if (res_obj)
+ ww_mutex_unlock(&res_obj->lock);
+
+ if (ret == -EDEADLK) {
+ /* we lost out in a seqno race, lock and retry.. */
+ ww_mutex_lock_slow(&contended_entry->obj->lock, ctx);
+ res_obj = contended_entry->obj;
+ goto retry;
+ }
+ ww_acquire_fini(ctx);
+
+ return ret;
+}
+
+Method 2, using a list in execbuf->buffers that can be reordered. Same semantics
+of duplicate entry detection using -EALREADY as method 1 above. But the
+list-reordering allows for a bit more idiomatic code.
+
+int lock_objs(struct list_head *list, struct ww_acquire_ctx *ctx)
+{
+ struct obj_entry *entry, *entry2;
+
+ ww_acquire_init(ctx, &ww_class);
+
+ list_for_each_entry (entry, list, head) {
+ ret = ww_mutex_lock(&entry->obj->lock, ctx);
+ if (ret < 0) {
+ entry2 = entry;
+
+ list_for_each_entry_continue_reverse (entry2, list, head)
+ ww_mutex_unlock(&entry2->obj->lock);
+
+ if (ret != -EDEADLK) {
+ ww_acquire_fini(ctx);
+ return ret;
+ }
+
+ /* we lost out in a seqno race, lock and retry.. */
+ ww_mutex_lock_slow(&entry->obj->lock, ctx);
+
+ /*
+ * Move buf to head of the list, this will point
+ * buf->next to the first unlocked entry,
+ * restarting the for loop.
+ */
+ list_del(&entry->head);
+ list_add(&entry->head, list);
+ }
+ }
+
+ ww_acquire_done(ctx);
+ return 0;
+}
+
+Unlocking works the same way for both methods #1 and #2:
+
+void unlock_objs(struct list_head *list, struct ww_acquire_ctx *ctx)
+{
+ struct obj_entry *entry;
+
+ list_for_each_entry (entry, list, head)
+ ww_mutex_unlock(&entry->obj->lock);
+
+ ww_acquire_fini(ctx);
+}
+
+Method 3 is useful if the list of objects is constructed ad-hoc and not upfront,
+e.g. when adjusting edges in a graph where each node has its own ww_mutex lock,
+and edges can only be changed when holding the locks of all involved nodes. w/w
+mutexes are a natural fit for such a case for two reasons:
+- They can handle lock-acquisition in any order which allows us to start walking
+ a graph from a starting point and then iteratively discovering new edges and
+ locking down the nodes those edges connect to.
+- Due to the -EALREADY return code signalling that a given objects is already
+ held there's no need for additional book-keeping to break cycles in the graph
+ or keep track off which looks are already held (when using more than one node
+ as a starting point).
+
+Note that this approach differs in two important ways from the above methods:
+- Since the list of objects is dynamically constructed (and might very well be
+ different when retrying due to hitting the -EDEADLK wound condition) there's
+ no need to keep any object on a persistent list when it's not locked. We can
+ therefore move the list_head into the object itself.
+- On the other hand the dynamic object list construction also means that the -EALREADY return
+ code can't be propagated.
+
+Note also that methods #1 and #2 and method #3 can be combined, e.g. to first lock a
+list of starting nodes (passed in from userspace) using one of the above
+methods. And then lock any additional objects affected by the operations using
+method #3 below. The backoff/retry procedure will be a bit more involved, since
+when the dynamic locking step hits -EDEADLK we also need to unlock all the
+objects acquired with the fixed list. But the w/w mutex debug checks will catch
+any interface misuse for these cases.
+
+Also, method 3 can't fail the lock acquisition step since it doesn't return
+-EALREADY. Of course this would be different when using the _interruptible
+variants, but that's outside of the scope of these examples here.
+
+struct obj {
+ struct ww_mutex ww_mutex;
+ struct list_head locked_list;
+};
+
+static DEFINE_WW_CLASS(ww_class);
+
+void __unlock_objs(struct list_head *list)
+{
+ struct obj *entry, *temp;
+
+ list_for_each_entry_safe (entry, temp, list, locked_list) {
+ /* need to do that before unlocking, since only the current lock holder is
+ allowed to use object */
+ list_del(&entry->locked_list);
+ ww_mutex_unlock(entry->ww_mutex)
+ }
+}
+
+void lock_objs(struct list_head *list, struct ww_acquire_ctx *ctx)
+{
+ struct obj *obj;
+
+ ww_acquire_init(ctx, &ww_class);
+
+retry:
+ /* re-init loop start state */
+ loop {
+ /* magic code which walks over a graph and decides which objects
+ * to lock */
+
+ ret = ww_mutex_lock(obj->ww_mutex, ctx);
+ if (ret == -EALREADY) {
+ /* we have that one already, get to the next object */
+ continue;
+ }
+ if (ret == -EDEADLK) {
+ __unlock_objs(list);
+
+ ww_mutex_lock_slow(obj, ctx);
+ list_add(&entry->locked_list, list);
+ goto retry;
+ }
+
+ /* locked a new object, add it to the list */
+ list_add_tail(&entry->locked_list, list);
+ }
+
+ ww_acquire_done(ctx);
+ return 0;
+}
+
+void unlock_objs(struct list_head *list, struct ww_acquire_ctx *ctx)
+{
+ __unlock_objs(list);
+ ww_acquire_fini(ctx);
+}
+
+Method 4: Only lock one single objects. In that case deadlock detection and
+prevention is obviously overkill, since with grabbing just one lock you can't
+produce a deadlock within just one class. To simplify this case the w/w mutex
+api can be used with a NULL context.
+
+Implementation Details
+----------------------
+
+Design:
+ ww_mutex currently encapsulates a struct mutex, this means no extra overhead for
+ normal mutex locks, which are far more common. As such there is only a small
+ increase in code size if wait/wound mutexes are not used.
+
+ In general, not much contention is expected. The locks are typically used to
+ serialize access to resources for devices. The only way to make wakeups
+ smarter would be at the cost of adding a field to struct mutex_waiter. This
+ would add overhead to all cases where normal mutexes are used, and
+ ww_mutexes are generally less performance sensitive.
+
+Lockdep:
+ Special care has been taken to warn for as many cases of api abuse
+ as possible. Some common api abuses will be caught with
+ CONFIG_DEBUG_MUTEXES, but CONFIG_PROVE_LOCKING is recommended.
+
+ Some of the errors which will be warned about:
+ - Forgetting to call ww_acquire_fini or ww_acquire_init.
+ - Attempting to lock more mutexes after ww_acquire_done.
+ - Attempting to lock the wrong mutex after -EDEADLK and
+ unlocking all mutexes.
+ - Attempting to lock the right mutex after -EDEADLK,
+ before unlocking all mutexes.
+
+ - Calling ww_mutex_lock_slow before -EDEADLK was returned.
+
+ - Unlocking mutexes with the wrong unlock function.
+ - Calling one of the ww_acquire_* twice on the same context.
+ - Using a different ww_class for the mutex than for the ww_acquire_ctx.
+ - Normal lockdep errors that can result in deadlocks.
+
+ Some of the lockdep errors that can result in deadlocks:
+ - Calling ww_acquire_init to initialize a second ww_acquire_ctx before
+ having called ww_acquire_fini on the first.
+ - 'normal' deadlocks that can occur.
+
+FIXME: Update this section once we have the TASK_DEADLOCK task state flag magic
+implemented.
diff --git a/Documentation/lockstat.txt b/Documentation/lockstat.txt
deleted file mode 100644
index 72d0106..0000000
--- a/Documentation/lockstat.txt
+++ /dev/null
@@ -1,178 +0,0 @@
-
-LOCK STATISTICS
-
-- WHAT
-
-As the name suggests, it provides statistics on locks.
-
-- WHY
-
-Because things like lock contention can severely impact performance.
-
-- HOW
-
-Lockdep already has hooks in the lock functions and maps lock instances to
-lock classes. We build on that (see Documentation/lockdep-design.txt).
-The graph below shows the relation between the lock functions and the various
-hooks therein.
-
- __acquire
- |
- lock _____
- | \
- | __contended
- | |
- | <wait>
- | _______/
- |/
- |
- __acquired
- |
- .
- <hold>
- .
- |
- __release
- |
- unlock
-
-lock, unlock - the regular lock functions
-__* - the hooks
-<> - states
-
-With these hooks we provide the following statistics:
-
- con-bounces - number of lock contention that involved x-cpu data
- contentions - number of lock acquisitions that had to wait
- wait time min - shortest (non-0) time we ever had to wait for a lock
- max - longest time we ever had to wait for a lock
- total - total time we spend waiting on this lock
- avg - average time spent waiting on this lock
- acq-bounces - number of lock acquisitions that involved x-cpu data
- acquisitions - number of times we took the lock
- hold time min - shortest (non-0) time we ever held the lock
- max - longest time we ever held the lock
- total - total time this lock was held
- avg - average time this lock was held
-
-These numbers are gathered per lock class, per read/write state (when
-applicable).
-
-It also tracks 4 contention points per class. A contention point is a call site
-that had to wait on lock acquisition.
-
- - CONFIGURATION
-
-Lock statistics are enabled via CONFIG_LOCK_STAT.
-
- - USAGE
-
-Enable collection of statistics:
-
-# echo 1 >/proc/sys/kernel/lock_stat
-
-Disable collection of statistics:
-
-# echo 0 >/proc/sys/kernel/lock_stat
-
-Look at the current lock statistics:
-
-( line numbers not part of actual output, done for clarity in the explanation
- below )
-
-# less /proc/lock_stat
-
-01 lock_stat version 0.4
-02-----------------------------------------------------------------------------------------------------------------------------------------------------------------------------------------------------------------------------
-03 class name con-bounces contentions waittime-min waittime-max waittime-total waittime-avg acq-bounces acquisitions holdtime-min holdtime-max holdtime-total holdtime-avg
-04-----------------------------------------------------------------------------------------------------------------------------------------------------------------------------------------------------------------------------
-05
-06 &mm->mmap_sem-W: 46 84 0.26 939.10 16371.53 194.90 47291 2922365 0.16 2220301.69 17464026916.32 5975.99
-07 &mm->mmap_sem-R: 37 100 1.31 299502.61 325629.52 3256.30 212344 34316685 0.10 7744.91 95016910.20 2.77
-08 ---------------
-09 &mm->mmap_sem 1 [<ffffffff811502a7>] khugepaged_scan_mm_slot+0x57/0x280
-19 &mm->mmap_sem 96 [<ffffffff815351c4>] __do_page_fault+0x1d4/0x510
-11 &mm->mmap_sem 34 [<ffffffff81113d77>] vm_mmap_pgoff+0x87/0xd0
-12 &mm->mmap_sem 17 [<ffffffff81127e71>] vm_munmap+0x41/0x80
-13 ---------------
-14 &mm->mmap_sem 1 [<ffffffff81046fda>] dup_mmap+0x2a/0x3f0
-15 &mm->mmap_sem 60 [<ffffffff81129e29>] SyS_mprotect+0xe9/0x250
-16 &mm->mmap_sem 41 [<ffffffff815351c4>] __do_page_fault+0x1d4/0x510
-17 &mm->mmap_sem 68 [<ffffffff81113d77>] vm_mmap_pgoff+0x87/0xd0
-18
-19.............................................................................................................................................................................................................................
-20
-21 unix_table_lock: 110 112 0.21 49.24 163.91 1.46 21094 66312 0.12 624.42 31589.81 0.48
-22 ---------------
-23 unix_table_lock 45 [<ffffffff8150ad8e>] unix_create1+0x16e/0x1b0
-24 unix_table_lock 47 [<ffffffff8150b111>] unix_release_sock+0x31/0x250
-25 unix_table_lock 15 [<ffffffff8150ca37>] unix_find_other+0x117/0x230
-26 unix_table_lock 5 [<ffffffff8150a09f>] unix_autobind+0x11f/0x1b0
-27 ---------------
-28 unix_table_lock 39 [<ffffffff8150b111>] unix_release_sock+0x31/0x250
-29 unix_table_lock 49 [<ffffffff8150ad8e>] unix_create1+0x16e/0x1b0
-30 unix_table_lock 20 [<ffffffff8150ca37>] unix_find_other+0x117/0x230
-31 unix_table_lock 4 [<ffffffff8150a09f>] unix_autobind+0x11f/0x1b0
-
-
-This excerpt shows the first two lock class statistics. Line 01 shows the
-output version - each time the format changes this will be updated. Line 02-04
-show the header with column descriptions. Lines 05-18 and 20-31 show the actual
-statistics. These statistics come in two parts; the actual stats separated by a
-short separator (line 08, 13) from the contention points.
-
-The first lock (05-18) is a read/write lock, and shows two lines above the
-short separator. The contention points don't match the column descriptors,
-they have two: contentions and [<IP>] symbol. The second set of contention
-points are the points we're contending with.
-
-The integer part of the time values is in us.
-
-Dealing with nested locks, subclasses may appear:
-
-32...........................................................................................................................................................................................................................
-33
-34 &rq->lock: 13128 13128 0.43 190.53 103881.26 7.91 97454 3453404 0.00 401.11 13224683.11 3.82
-35 ---------
-36 &rq->lock 645 [<ffffffff8103bfc4>] task_rq_lock+0x43/0x75
-37 &rq->lock 297 [<ffffffff8104ba65>] try_to_wake_up+0x127/0x25a
-38 &rq->lock 360 [<ffffffff8103c4c5>] select_task_rq_fair+0x1f0/0x74a
-39 &rq->lock 428 [<ffffffff81045f98>] scheduler_tick+0x46/0x1fb
-40 ---------
-41 &rq->lock 77 [<ffffffff8103bfc4>] task_rq_lock+0x43/0x75
-42 &rq->lock 174 [<ffffffff8104ba65>] try_to_wake_up+0x127/0x25a
-43 &rq->lock 4715 [<ffffffff8103ed4b>] double_rq_lock+0x42/0x54
-44 &rq->lock 893 [<ffffffff81340524>] schedule+0x157/0x7b8
-45
-46...........................................................................................................................................................................................................................
-47
-48 &rq->lock/1: 1526 11488 0.33 388.73 136294.31 11.86 21461 38404 0.00 37.93 109388.53 2.84
-49 -----------
-50 &rq->lock/1 11526 [<ffffffff8103ed58>] double_rq_lock+0x4f/0x54
-51 -----------
-52 &rq->lock/1 5645 [<ffffffff8103ed4b>] double_rq_lock+0x42/0x54
-53 &rq->lock/1 1224 [<ffffffff81340524>] schedule+0x157/0x7b8
-54 &rq->lock/1 4336 [<ffffffff8103ed58>] double_rq_lock+0x4f/0x54
-55 &rq->lock/1 181 [<ffffffff8104ba65>] try_to_wake_up+0x127/0x25a
-
-Line 48 shows statistics for the second subclass (/1) of &rq->lock class
-(subclass starts from 0), since in this case, as line 50 suggests,
-double_rq_lock actually acquires a nested lock of two spinlocks.
-
-View the top contending locks:
-
-# grep : /proc/lock_stat | head
- clockevents_lock: 2926159 2947636 0.15 46882.81 1784540466.34 605.41 3381345 3879161 0.00 2260.97 53178395.68 13.71
- tick_broadcast_lock: 346460 346717 0.18 2257.43 39364622.71 113.54 3642919 4242696 0.00 2263.79 49173646.60 11.59
- &mapping->i_mmap_mutex: 203896 203899 3.36 645530.05 31767507988.39 155800.21 3361776 8893984 0.17 2254.15 14110121.02 1.59
- &rq->lock: 135014 136909 0.18 606.09 842160.68 6.15 1540728 10436146 0.00 728.72 17606683.41 1.69
- &(&zone->lru_lock)->rlock: 93000 94934 0.16 59.18 188253.78 1.98 1199912 3809894 0.15 391.40 3559518.81 0.93
- tasklist_lock-W: 40667 41130 0.23 1189.42 428980.51 10.43 270278 510106 0.16 653.51 3939674.91 7.72
- tasklist_lock-R: 21298 21305 0.20 1310.05 215511.12 10.12 186204 241258 0.14 1162.33 1179779.23 4.89
- rcu_node_1: 47656 49022 0.16 635.41 193616.41 3.95 844888 1865423 0.00 764.26 1656226.96 0.89
- &(&dentry->d_lockref.lock)->rlock: 39791 40179 0.15 1302.08 88851.96 2.21 2790851 12527025 0.10 1910.75 3379714.27 0.27
- rcu_node_0: 29203 30064 0.16 786.55 1555573.00 51.74 88963 244254 0.00 398.87 428872.51 1.76
-
-Clear the statistics:
-
-# echo 0 > /proc/lock_stat
diff --git a/Documentation/mutex-design.txt b/Documentation/mutex-design.txt
deleted file mode 100644
index ee231ed..0000000
--- a/Documentation/mutex-design.txt
+++ /dev/null
@@ -1,157 +0,0 @@
-Generic Mutex Subsystem
-
-started by Ingo Molnar <mingo@redhat.com>
-updated by Davidlohr Bueso <davidlohr@hp.com>
-
-What are mutexes?
------------------
-
-In the Linux kernel, mutexes refer to a particular locking primitive
-that enforces serialization on shared memory systems, and not only to
-the generic term referring to 'mutual exclusion' found in academia
-or similar theoretical text books. Mutexes are sleeping locks which
-behave similarly to binary semaphores, and were introduced in 2006[1]
-as an alternative to these. This new data structure provided a number
-of advantages, including simpler interfaces, and at that time smaller
-code (see Disadvantages).
-
-[1] http://lwn.net/Articles/164802/
-
-Implementation
---------------
-
-Mutexes are represented by 'struct mutex', defined in include/linux/mutex.h
-and implemented in kernel/locking/mutex.c. These locks use a three
-state atomic counter (->count) to represent the different possible
-transitions that can occur during the lifetime of a lock:
-
- 1: unlocked
- 0: locked, no waiters
- negative: locked, with potential waiters
-
-In its most basic form it also includes a wait-queue and a spinlock
-that serializes access to it. CONFIG_SMP systems can also include
-a pointer to the lock task owner (->owner) as well as a spinner MCS
-lock (->osq), both described below in (ii).
-
-When acquiring a mutex, there are three possible paths that can be
-taken, depending on the state of the lock:
-
-(i) fastpath: tries to atomically acquire the lock by decrementing the
- counter. If it was already taken by another task it goes to the next
- possible path. This logic is architecture specific. On x86-64, the
- locking fastpath is 2 instructions:
-
- 0000000000000e10 <mutex_lock>:
- e21: f0 ff 0b lock decl (%rbx)
- e24: 79 08 jns e2e <mutex_lock+0x1e>
-
- the unlocking fastpath is equally tight:
-
- 0000000000000bc0 <mutex_unlock>:
- bc8: f0 ff 07 lock incl (%rdi)
- bcb: 7f 0a jg bd7 <mutex_unlock+0x17>
-
-
-(ii) midpath: aka optimistic spinning, tries to spin for acquisition
- while the lock owner is running and there are no other tasks ready
- to run that have higher priority (need_resched). The rationale is
- that if the lock owner is running, it is likely to release the lock
- soon. The mutex spinners are queued up using MCS lock so that only
- one spinner can compete for the mutex.
-
- The MCS lock (proposed by Mellor-Crummey and Scott) is a simple spinlock
- with the desirable properties of being fair and with each cpu trying
- to acquire the lock spinning on a local variable. It avoids expensive
- cacheline bouncing that common test-and-set spinlock implementations
- incur. An MCS-like lock is specially tailored for optimistic spinning
- for sleeping lock implementation. An important feature of the customized
- MCS lock is that it has the extra property that spinners are able to exit
- the MCS spinlock queue when they need to reschedule. This further helps
- avoid situations where MCS spinners that need to reschedule would continue
- waiting to spin on mutex owner, only to go directly to slowpath upon
- obtaining the MCS lock.
-
-
-(iii) slowpath: last resort, if the lock is still unable to be acquired,
- the task is added to the wait-queue and sleeps until woken up by the
- unlock path. Under normal circumstances it blocks as TASK_UNINTERRUPTIBLE.
-
-While formally kernel mutexes are sleepable locks, it is path (ii) that
-makes them more practically a hybrid type. By simply not interrupting a
-task and busy-waiting for a few cycles instead of immediately sleeping,
-the performance of this lock has been seen to significantly improve a
-number of workloads. Note that this technique is also used for rw-semaphores.
-
-Semantics
----------
-
-The mutex subsystem checks and enforces the following rules:
-
- - Only one task can hold the mutex at a time.
- - Only the owner can unlock the mutex.
- - Multiple unlocks are not permitted.
- - Recursive locking/unlocking is not permitted.
- - A mutex must only be initialized via the API (see below).
- - A task may not exit with a mutex held.
- - Memory areas where held locks reside must not be freed.
- - Held mutexes must not be reinitialized.
- - Mutexes may not be used in hardware or software interrupt
- contexts such as tasklets and timers.
-
-These semantics are fully enforced when CONFIG DEBUG_MUTEXES is enabled.
-In addition, the mutex debugging code also implements a number of other
-features that make lock debugging easier and faster:
-
- - Uses symbolic names of mutexes, whenever they are printed
- in debug output.
- - Point-of-acquire tracking, symbolic lookup of function names,
- list of all locks held in the system, printout of them.
- - Owner tracking.
- - Detects self-recursing locks and prints out all relevant info.
- - Detects multi-task circular deadlocks and prints out all affected
- locks and tasks (and only those tasks).
-
-
-Interfaces
-----------
-Statically define the mutex:
- DEFINE_MUTEX(name);
-
-Dynamically initialize the mutex:
- mutex_init(mutex);
-
-Acquire the mutex, uninterruptible:
- void mutex_lock(struct mutex *lock);
- void mutex_lock_nested(struct mutex *lock, unsigned int subclass);
- int mutex_trylock(struct mutex *lock);
-
-Acquire the mutex, interruptible:
- int mutex_lock_interruptible_nested(struct mutex *lock,
- unsigned int subclass);
- int mutex_lock_interruptible(struct mutex *lock);
-
-Acquire the mutex, interruptible, if dec to 0:
- int atomic_dec_and_mutex_lock(atomic_t *cnt, struct mutex *lock);
-
-Unlock the mutex:
- void mutex_unlock(struct mutex *lock);
-
-Test if the mutex is taken:
- int mutex_is_locked(struct mutex *lock);
-
-Disadvantages
--------------
-
-Unlike its original design and purpose, 'struct mutex' is larger than
-most locks in the kernel. E.g: on x86-64 it is 40 bytes, almost twice
-as large as 'struct semaphore' (24 bytes) and 8 bytes shy of the
-'struct rw_semaphore' variant. Larger structure sizes mean more CPU
-cache and memory footprint.
-
-When to use mutexes
--------------------
-
-Unless the strict semantics of mutexes are unsuitable and/or the critical
-region prevents the lock from being shared, always prefer them to any other
-locking primitive.
diff --git a/Documentation/rt-mutex-design.txt b/Documentation/rt-mutex-design.txt
deleted file mode 100644
index 8666070..0000000
--- a/Documentation/rt-mutex-design.txt
+++ /dev/null
@@ -1,781 +0,0 @@
-#
-# Copyright (c) 2006 Steven Rostedt
-# Licensed under the GNU Free Documentation License, Version 1.2
-#
-
-RT-mutex implementation design
-------------------------------
-
-This document tries to describe the design of the rtmutex.c implementation.
-It doesn't describe the reasons why rtmutex.c exists. For that please see
-Documentation/rt-mutex.txt. Although this document does explain problems
-that happen without this code, but that is in the concept to understand
-what the code actually is doing.
-
-The goal of this document is to help others understand the priority
-inheritance (PI) algorithm that is used, as well as reasons for the
-decisions that were made to implement PI in the manner that was done.
-
-
-Unbounded Priority Inversion
-----------------------------
-
-Priority inversion is when a lower priority process executes while a higher
-priority process wants to run. This happens for several reasons, and
-most of the time it can't be helped. Anytime a high priority process wants
-to use a resource that a lower priority process has (a mutex for example),
-the high priority process must wait until the lower priority process is done
-with the resource. This is a priority inversion. What we want to prevent
-is something called unbounded priority inversion. That is when the high
-priority process is prevented from running by a lower priority process for
-an undetermined amount of time.
-
-The classic example of unbounded priority inversion is where you have three
-processes, let's call them processes A, B, and C, where A is the highest
-priority process, C is the lowest, and B is in between. A tries to grab a lock
-that C owns and must wait and lets C run to release the lock. But in the
-meantime, B executes, and since B is of a higher priority than C, it preempts C,
-but by doing so, it is in fact preempting A which is a higher priority process.
-Now there's no way of knowing how long A will be sleeping waiting for C
-to release the lock, because for all we know, B is a CPU hog and will
-never give C a chance to release the lock. This is called unbounded priority
-inversion.
-
-Here's a little ASCII art to show the problem.
-
- grab lock L1 (owned by C)
- |
-A ---+
- C preempted by B
- |
-C +----+
-
-B +-------->
- B now keeps A from running.
-
-
-Priority Inheritance (PI)
--------------------------
-
-There are several ways to solve this issue, but other ways are out of scope
-for this document. Here we only discuss PI.
-
-PI is where a process inherits the priority of another process if the other
-process blocks on a lock owned by the current process. To make this easier
-to understand, let's use the previous example, with processes A, B, and C again.
-
-This time, when A blocks on the lock owned by C, C would inherit the priority
-of A. So now if B becomes runnable, it would not preempt C, since C now has
-the high priority of A. As soon as C releases the lock, it loses its
-inherited priority, and A then can continue with the resource that C had.
-
-Terminology
------------
-
-Here I explain some terminology that is used in this document to help describe
-the design that is used to implement PI.
-
-PI chain - The PI chain is an ordered series of locks and processes that cause
- processes to inherit priorities from a previous process that is
- blocked on one of its locks. This is described in more detail
- later in this document.
-
-mutex - In this document, to differentiate from locks that implement
- PI and spin locks that are used in the PI code, from now on
- the PI locks will be called a mutex.
-
-lock - In this document from now on, I will use the term lock when
- referring to spin locks that are used to protect parts of the PI
- algorithm. These locks disable preemption for UP (when
- CONFIG_PREEMPT is enabled) and on SMP prevents multiple CPUs from
- entering critical sections simultaneously.
-
-spin lock - Same as lock above.
-
-waiter - A waiter is a struct that is stored on the stack of a blocked
- process. Since the scope of the waiter is within the code for
- a process being blocked on the mutex, it is fine to allocate
- the waiter on the process's stack (local variable). This
- structure holds a pointer to the task, as well as the mutex that
- the task is blocked on. It also has the plist node structures to
- place the task in the waiter_list of a mutex as well as the
- pi_list of a mutex owner task (described below).
-
- waiter is sometimes used in reference to the task that is waiting
- on a mutex. This is the same as waiter->task.
-
-waiters - A list of processes that are blocked on a mutex.
-
-top waiter - The highest priority process waiting on a specific mutex.
-
-top pi waiter - The highest priority process waiting on one of the mutexes
- that a specific process owns.
-
-Note: task and process are used interchangeably in this document, mostly to
- differentiate between two processes that are being described together.
-
-
-PI chain
---------
-
-The PI chain is a list of processes and mutexes that may cause priority
-inheritance to take place. Multiple chains may converge, but a chain
-would never diverge, since a process can't be blocked on more than one
-mutex at a time.
-
-Example:
-
- Process: A, B, C, D, E
- Mutexes: L1, L2, L3, L4
-
- A owns: L1
- B blocked on L1
- B owns L2
- C blocked on L2
- C owns L3
- D blocked on L3
- D owns L4
- E blocked on L4
-
-The chain would be:
-
- E->L4->D->L3->C->L2->B->L1->A
-
-To show where two chains merge, we could add another process F and
-another mutex L5 where B owns L5 and F is blocked on mutex L5.
-
-The chain for F would be:
-
- F->L5->B->L1->A
-
-Since a process may own more than one mutex, but never be blocked on more than
-one, the chains merge.
-
-Here we show both chains:
-
- E->L4->D->L3->C->L2-+
- |
- +->B->L1->A
- |
- F->L5-+
-
-For PI to work, the processes at the right end of these chains (or we may
-also call it the Top of the chain) must be equal to or higher in priority
-than the processes to the left or below in the chain.
-
-Also since a mutex may have more than one process blocked on it, we can
-have multiple chains merge at mutexes. If we add another process G that is
-blocked on mutex L2:
-
- G->L2->B->L1->A
-
-And once again, to show how this can grow I will show the merging chains
-again.
-
- E->L4->D->L3->C-+
- +->L2-+
- | |
- G-+ +->B->L1->A
- |
- F->L5-+
-
-
-Plist
------
-
-Before I go further and talk about how the PI chain is stored through lists
-on both mutexes and processes, I'll explain the plist. This is similar to
-the struct list_head functionality that is already in the kernel.
-The implementation of plist is out of scope for this document, but it is
-very important to understand what it does.
-
-There are a few differences between plist and list, the most important one
-being that plist is a priority sorted linked list. This means that the
-priorities of the plist are sorted, such that it takes O(1) to retrieve the
-highest priority item in the list. Obviously this is useful to store processes
-based on their priorities.
-
-Another difference, which is important for implementation, is that, unlike
-list, the head of the list is a different element than the nodes of a list.
-So the head of the list is declared as struct plist_head and nodes that will
-be added to the list are declared as struct plist_node.
-
-
-Mutex Waiter List
------------------
-
-Every mutex keeps track of all the waiters that are blocked on itself. The mutex
-has a plist to store these waiters by priority. This list is protected by
-a spin lock that is located in the struct of the mutex. This lock is called
-wait_lock. Since the modification of the waiter list is never done in
-interrupt context, the wait_lock can be taken without disabling interrupts.
-
-
-Task PI List
-------------
-
-To keep track of the PI chains, each process has its own PI list. This is
-a list of all top waiters of the mutexes that are owned by the process.
-Note that this list only holds the top waiters and not all waiters that are
-blocked on mutexes owned by the process.
-
-The top of the task's PI list is always the highest priority task that
-is waiting on a mutex that is owned by the task. So if the task has
-inherited a priority, it will always be the priority of the task that is
-at the top of this list.
-
-This list is stored in the task structure of a process as a plist called
-pi_list. This list is protected by a spin lock also in the task structure,
-called pi_lock. This lock may also be taken in interrupt context, so when
-locking the pi_lock, interrupts must be disabled.
-
-
-Depth of the PI Chain
----------------------
-
-The maximum depth of the PI chain is not dynamic, and could actually be
-defined. But is very complex to figure it out, since it depends on all
-the nesting of mutexes. Let's look at the example where we have 3 mutexes,
-L1, L2, and L3, and four separate functions func1, func2, func3 and func4.
-The following shows a locking order of L1->L2->L3, but may not actually
-be directly nested that way.
-
-void func1(void)
-{
- mutex_lock(L1);
-
- /* do anything */
-
- mutex_unlock(L1);
-}
-
-void func2(void)
-{
- mutex_lock(L1);
- mutex_lock(L2);
-
- /* do something */
-
- mutex_unlock(L2);
- mutex_unlock(L1);
-}
-
-void func3(void)
-{
- mutex_lock(L2);
- mutex_lock(L3);
-
- /* do something else */
-
- mutex_unlock(L3);
- mutex_unlock(L2);
-}
-
-void func4(void)
-{
- mutex_lock(L3);
-
- /* do something again */
-
- mutex_unlock(L3);
-}
-
-Now we add 4 processes that run each of these functions separately.
-Processes A, B, C, and D which run functions func1, func2, func3 and func4
-respectively, and such that D runs first and A last. With D being preempted
-in func4 in the "do something again" area, we have a locking that follows:
-
-D owns L3
- C blocked on L3
- C owns L2
- B blocked on L2
- B owns L1
- A blocked on L1
-
-And thus we have the chain A->L1->B->L2->C->L3->D.
-
-This gives us a PI depth of 4 (four processes), but looking at any of the
-functions individually, it seems as though they only have at most a locking
-depth of two. So, although the locking depth is defined at compile time,
-it still is very difficult to find the possibilities of that depth.
-
-Now since mutexes can be defined by user-land applications, we don't want a DOS
-type of application that nests large amounts of mutexes to create a large
-PI chain, and have the code holding spin locks while looking at a large
-amount of data. So to prevent this, the implementation not only implements
-a maximum lock depth, but also only holds at most two different locks at a
-time, as it walks the PI chain. More about this below.
-
-
-Mutex owner and flags
----------------------
-
-The mutex structure contains a pointer to the owner of the mutex. If the
-mutex is not owned, this owner is set to NULL. Since all architectures
-have the task structure on at least a four byte alignment (and if this is
-not true, the rtmutex.c code will be broken!), this allows for the two
-least significant bits to be used as flags. This part is also described
-in Documentation/rt-mutex.txt, but will also be briefly described here.
-
-Bit 0 is used as the "Pending Owner" flag. This is described later.
-Bit 1 is used as the "Has Waiters" flags. This is also described later
- in more detail, but is set whenever there are waiters on a mutex.
-
-
-cmpxchg Tricks
---------------
-
-Some architectures implement an atomic cmpxchg (Compare and Exchange). This
-is used (when applicable) to keep the fast path of grabbing and releasing
-mutexes short.
-
-cmpxchg is basically the following function performed atomically:
-
-unsigned long _cmpxchg(unsigned long *A, unsigned long *B, unsigned long *C)
-{
- unsigned long T = *A;
- if (*A == *B) {
- *A = *C;
- }
- return T;
-}
-#define cmpxchg(a,b,c) _cmpxchg(&a,&b,&c)
-
-This is really nice to have, since it allows you to only update a variable
-if the variable is what you expect it to be. You know if it succeeded if
-the return value (the old value of A) is equal to B.
-
-The macro rt_mutex_cmpxchg is used to try to lock and unlock mutexes. If
-the architecture does not support CMPXCHG, then this macro is simply set
-to fail every time. But if CMPXCHG is supported, then this will
-help out extremely to keep the fast path short.
-
-The use of rt_mutex_cmpxchg with the flags in the owner field help optimize
-the system for architectures that support it. This will also be explained
-later in this document.
-
-
-Priority adjustments
---------------------
-
-The implementation of the PI code in rtmutex.c has several places that a
-process must adjust its priority. With the help of the pi_list of a
-process this is rather easy to know what needs to be adjusted.
-
-The functions implementing the task adjustments are rt_mutex_adjust_prio,
-__rt_mutex_adjust_prio (same as the former, but expects the task pi_lock
-to already be taken), rt_mutex_getprio, and rt_mutex_setprio.
-
-rt_mutex_getprio and rt_mutex_setprio are only used in __rt_mutex_adjust_prio.
-
-rt_mutex_getprio returns the priority that the task should have. Either the
-task's own normal priority, or if a process of a higher priority is waiting on
-a mutex owned by the task, then that higher priority should be returned.
-Since the pi_list of a task holds an order by priority list of all the top
-waiters of all the mutexes that the task owns, rt_mutex_getprio simply needs
-to compare the top pi waiter to its own normal priority, and return the higher
-priority back.
-
-(Note: if looking at the code, you will notice that the lower number of
- prio is returned. This is because the prio field in the task structure
- is an inverse order of the actual priority. So a "prio" of 5 is
- of higher priority than a "prio" of 10.)
-
-__rt_mutex_adjust_prio examines the result of rt_mutex_getprio, and if the
-result does not equal the task's current priority, then rt_mutex_setprio
-is called to adjust the priority of the task to the new priority.
-Note that rt_mutex_setprio is defined in kernel/sched/core.c to implement the
-actual change in priority.
-
-It is interesting to note that __rt_mutex_adjust_prio can either increase
-or decrease the priority of the task. In the case that a higher priority
-process has just blocked on a mutex owned by the task, __rt_mutex_adjust_prio
-would increase/boost the task's priority. But if a higher priority task
-were for some reason to leave the mutex (timeout or signal), this same function
-would decrease/unboost the priority of the task. That is because the pi_list
-always contains the highest priority task that is waiting on a mutex owned
-by the task, so we only need to compare the priority of that top pi waiter
-to the normal priority of the given task.
-
-
-High level overview of the PI chain walk
-----------------------------------------
-
-The PI chain walk is implemented by the function rt_mutex_adjust_prio_chain.
-
-The implementation has gone through several iterations, and has ended up
-with what we believe is the best. It walks the PI chain by only grabbing
-at most two locks at a time, and is very efficient.
-
-The rt_mutex_adjust_prio_chain can be used either to boost or lower process
-priorities.
-
-rt_mutex_adjust_prio_chain is called with a task to be checked for PI
-(de)boosting (the owner of a mutex that a process is blocking on), a flag to
-check for deadlocking, the mutex that the task owns, and a pointer to a waiter
-that is the process's waiter struct that is blocked on the mutex (although this
-parameter may be NULL for deboosting).
-
-For this explanation, I will not mention deadlock detection. This explanation
-will try to stay at a high level.
-
-When this function is called, there are no locks held. That also means
-that the state of the owner and lock can change when entered into this function.
-
-Before this function is called, the task has already had rt_mutex_adjust_prio
-performed on it. This means that the task is set to the priority that it
-should be at, but the plist nodes of the task's waiter have not been updated
-with the new priorities, and that this task may not be in the proper locations
-in the pi_lists and wait_lists that the task is blocked on. This function
-solves all that.
-
-A loop is entered, where task is the owner to be checked for PI changes that
-was passed by parameter (for the first iteration). The pi_lock of this task is
-taken to prevent any more changes to the pi_list of the task. This also
-prevents new tasks from completing the blocking on a mutex that is owned by this
-task.
-
-If the task is not blocked on a mutex then the loop is exited. We are at
-the top of the PI chain.
-
-A check is now done to see if the original waiter (the process that is blocked
-on the current mutex) is the top pi waiter of the task. That is, is this
-waiter on the top of the task's pi_list. If it is not, it either means that
-there is another process higher in priority that is blocked on one of the
-mutexes that the task owns, or that the waiter has just woken up via a signal
-or timeout and has left the PI chain. In either case, the loop is exited, since
-we don't need to do any more changes to the priority of the current task, or any
-task that owns a mutex that this current task is waiting on. A priority chain
-walk is only needed when a new top pi waiter is made to a task.
-
-The next check sees if the task's waiter plist node has the priority equal to
-the priority the task is set at. If they are equal, then we are done with
-the loop. Remember that the function started with the priority of the
-task adjusted, but the plist nodes that hold the task in other processes
-pi_lists have not been adjusted.
-
-Next, we look at the mutex that the task is blocked on. The mutex's wait_lock
-is taken. This is done by a spin_trylock, because the locking order of the
-pi_lock and wait_lock goes in the opposite direction. If we fail to grab the
-lock, the pi_lock is released, and we restart the loop.
-
-Now that we have both the pi_lock of the task as well as the wait_lock of
-the mutex the task is blocked on, we update the task's waiter's plist node
-that is located on the mutex's wait_list.
-
-Now we release the pi_lock of the task.
-
-Next the owner of the mutex has its pi_lock taken, so we can update the
-task's entry in the owner's pi_list. If the task is the highest priority
-process on the mutex's wait_list, then we remove the previous top waiter
-from the owner's pi_list, and replace it with the task.
-
-Note: It is possible that the task was the current top waiter on the mutex,
- in which case the task is not yet on the pi_list of the waiter. This
- is OK, since plist_del does nothing if the plist node is not on any
- list.
-
-If the task was not the top waiter of the mutex, but it was before we
-did the priority updates, that means we are deboosting/lowering the
-task. In this case, the task is removed from the pi_list of the owner,
-and the new top waiter is added.
-
-Lastly, we unlock both the pi_lock of the task, as well as the mutex's
-wait_lock, and continue the loop again. On the next iteration of the
-loop, the previous owner of the mutex will be the task that will be
-processed.
-
-Note: One might think that the owner of this mutex might have changed
- since we just grab the mutex's wait_lock. And one could be right.
- The important thing to remember is that the owner could not have
- become the task that is being processed in the PI chain, since
- we have taken that task's pi_lock at the beginning of the loop.
- So as long as there is an owner of this mutex that is not the same
- process as the tasked being worked on, we are OK.
-
- Looking closely at the code, one might be confused. The check for the
- end of the PI chain is when the task isn't blocked on anything or the
- task's waiter structure "task" element is NULL. This check is
- protected only by the task's pi_lock. But the code to unlock the mutex
- sets the task's waiter structure "task" element to NULL with only
- the protection of the mutex's wait_lock, which was not taken yet.
- Isn't this a race condition if the task becomes the new owner?
-
- The answer is No! The trick is the spin_trylock of the mutex's
- wait_lock. If we fail that lock, we release the pi_lock of the
- task and continue the loop, doing the end of PI chain check again.
-
- In the code to release the lock, the wait_lock of the mutex is held
- the entire time, and it is not let go when we grab the pi_lock of the
- new owner of the mutex. So if the switch of a new owner were to happen
- after the check for end of the PI chain and the grabbing of the
- wait_lock, the unlocking code would spin on the new owner's pi_lock
- but never give up the wait_lock. So the PI chain loop is guaranteed to
- fail the spin_trylock on the wait_lock, release the pi_lock, and
- try again.
-
- If you don't quite understand the above, that's OK. You don't have to,
- unless you really want to make a proof out of it ;)
-
-
-Pending Owners and Lock stealing
---------------------------------
-
-One of the flags in the owner field of the mutex structure is "Pending Owner".
-What this means is that an owner was chosen by the process releasing the
-mutex, but that owner has yet to wake up and actually take the mutex.
-
-Why is this important? Why can't we just give the mutex to another process
-and be done with it?
-
-The PI code is to help with real-time processes, and to let the highest
-priority process run as long as possible with little latencies and delays.
-If a high priority process owns a mutex that a lower priority process is
-blocked on, when the mutex is released it would be given to the lower priority
-process. What if the higher priority process wants to take that mutex again.
-The high priority process would fail to take that mutex that it just gave up
-and it would need to boost the lower priority process to run with full
-latency of that critical section (since the low priority process just entered
-it).
-
-There's no reason a high priority process that gives up a mutex should be
-penalized if it tries to take that mutex again. If the new owner of the
-mutex has not woken up yet, there's no reason that the higher priority process
-could not take that mutex away.
-
-To solve this, we introduced Pending Ownership and Lock Stealing. When a
-new process is given a mutex that it was blocked on, it is only given
-pending ownership. This means that it's the new owner, unless a higher
-priority process comes in and tries to grab that mutex. If a higher priority
-process does come along and wants that mutex, we let the higher priority
-process "steal" the mutex from the pending owner (only if it is still pending)
-and continue with the mutex.
-
-
-Taking of a mutex (The walk through)
-------------------------------------
-
-OK, now let's take a look at the detailed walk through of what happens when
-taking a mutex.
-
-The first thing that is tried is the fast taking of the mutex. This is
-done when we have CMPXCHG enabled (otherwise the fast taking automatically
-fails). Only when the owner field of the mutex is NULL can the lock be
-taken with the CMPXCHG and nothing else needs to be done.
-
-If there is contention on the lock, whether it is owned or pending owner
-we go about the slow path (rt_mutex_slowlock).
-
-The slow path function is where the task's waiter structure is created on
-the stack. This is because the waiter structure is only needed for the
-scope of this function. The waiter structure holds the nodes to store
-the task on the wait_list of the mutex, and if need be, the pi_list of
-the owner.
-
-The wait_lock of the mutex is taken since the slow path of unlocking the
-mutex also takes this lock.
-
-We then call try_to_take_rt_mutex. This is where the architecture that
-does not implement CMPXCHG would always grab the lock (if there's no
-contention).
-
-try_to_take_rt_mutex is used every time the task tries to grab a mutex in the
-slow path. The first thing that is done here is an atomic setting of
-the "Has Waiters" flag of the mutex's owner field. Yes, this could really
-be false, because if the mutex has no owner, there are no waiters and
-the current task also won't have any waiters. But we don't have the lock
-yet, so we assume we are going to be a waiter. The reason for this is to
-play nice for those architectures that do have CMPXCHG. By setting this flag
-now, the owner of the mutex can't release the mutex without going into the
-slow unlock path, and it would then need to grab the wait_lock, which this
-code currently holds. So setting the "Has Waiters" flag forces the owner
-to synchronize with this code.
-
-Now that we know that we can't have any races with the owner releasing the
-mutex, we check to see if we can take the ownership. This is done if the
-mutex doesn't have a owner, or if we can steal the mutex from a pending
-owner. Let's look at the situations we have here.
-
- 1) Has owner that is pending
- ----------------------------
-
- The mutex has a owner, but it hasn't woken up and the mutex flag
- "Pending Owner" is set. The first check is to see if the owner isn't the
- current task. This is because this function is also used for the pending
- owner to grab the mutex. When a pending owner wakes up, it checks to see
- if it can take the mutex, and this is done if the owner is already set to
- itself. If so, we succeed and leave the function, clearing the "Pending
- Owner" bit.
-
- If the pending owner is not current, we check to see if the current priority is
- higher than the pending owner. If not, we fail the function and return.
-
- There's also something special about a pending owner. That is a pending owner
- is never blocked on a mutex. So there is no PI chain to worry about. It also
- means that if the mutex doesn't have any waiters, there's no accounting needed
- to update the pending owner's pi_list, since we only worry about processes
- blocked on the current mutex.
-
- If there are waiters on this mutex, and we just stole the ownership, we need
- to take the top waiter, remove it from the pi_list of the pending owner, and
- add it to the current pi_list. Note that at this moment, the pending owner
- is no longer on the list of waiters. This is fine, since the pending owner
- would add itself back when it realizes that it had the ownership stolen
- from itself. When the pending owner tries to grab the mutex, it will fail
- in try_to_take_rt_mutex if the owner field points to another process.
-
- 2) No owner
- -----------
-
- If there is no owner (or we successfully stole the lock), we set the owner
- of the mutex to current, and set the flag of "Has Waiters" if the current
- mutex actually has waiters, or we clear the flag if it doesn't. See, it was
- OK that we set that flag early, since now it is cleared.
-
- 3) Failed to grab ownership
- ---------------------------
-
- The most interesting case is when we fail to take ownership. This means that
- there exists an owner, or there's a pending owner with equal or higher
- priority than the current task.
-
-We'll continue on the failed case.
-
-If the mutex has a timeout, we set up a timer to go off to break us out
-of this mutex if we failed to get it after a specified amount of time.
-
-Now we enter a loop that will continue to try to take ownership of the mutex, or
-fail from a timeout or signal.
-
-Once again we try to take the mutex. This will usually fail the first time
-in the loop, since it had just failed to get the mutex. But the second time
-in the loop, this would likely succeed, since the task would likely be
-the pending owner.
-
-If the mutex is TASK_INTERRUPTIBLE a check for signals and timeout is done
-here.
-
-The waiter structure has a "task" field that points to the task that is blocked
-on the mutex. This field can be NULL the first time it goes through the loop
-or if the task is a pending owner and had its mutex stolen. If the "task"
-field is NULL then we need to set up the accounting for it.
-
-Task blocks on mutex
---------------------
-
-The accounting of a mutex and process is done with the waiter structure of
-the process. The "task" field is set to the process, and the "lock" field
-to the mutex. The plist nodes are initialized to the processes current
-priority.
-
-Since the wait_lock was taken at the entry of the slow lock, we can safely
-add the waiter to the wait_list. If the current process is the highest
-priority process currently waiting on this mutex, then we remove the
-previous top waiter process (if it exists) from the pi_list of the owner,
-and add the current process to that list. Since the pi_list of the owner
-has changed, we call rt_mutex_adjust_prio on the owner to see if the owner
-should adjust its priority accordingly.
-
-If the owner is also blocked on a lock, and had its pi_list changed
-(or deadlock checking is on), we unlock the wait_lock of the mutex and go ahead
-and run rt_mutex_adjust_prio_chain on the owner, as described earlier.
-
-Now all locks are released, and if the current process is still blocked on a
-mutex (waiter "task" field is not NULL), then we go to sleep (call schedule).
-
-Waking up in the loop
----------------------
-
-The schedule can then wake up for a few reasons.
- 1) we were given pending ownership of the mutex.
- 2) we received a signal and was TASK_INTERRUPTIBLE
- 3) we had a timeout and was TASK_INTERRUPTIBLE
-
-In any of these cases, we continue the loop and once again try to grab the
-ownership of the mutex. If we succeed, we exit the loop, otherwise we continue
-and on signal and timeout, will exit the loop, or if we had the mutex stolen
-we just simply add ourselves back on the lists and go back to sleep.
-
-Note: For various reasons, because of timeout and signals, the steal mutex
- algorithm needs to be careful. This is because the current process is
- still on the wait_list. And because of dynamic changing of priorities,
- especially on SCHED_OTHER tasks, the current process can be the
- highest priority task on the wait_list.
-
-Failed to get mutex on Timeout or Signal
-----------------------------------------
-
-If a timeout or signal occurred, the waiter's "task" field would not be
-NULL and the task needs to be taken off the wait_list of the mutex and perhaps
-pi_list of the owner. If this process was a high priority process, then
-the rt_mutex_adjust_prio_chain needs to be executed again on the owner,
-but this time it will be lowering the priorities.
-
-
-Unlocking the Mutex
--------------------
-
-The unlocking of a mutex also has a fast path for those architectures with
-CMPXCHG. Since the taking of a mutex on contention always sets the
-"Has Waiters" flag of the mutex's owner, we use this to know if we need to
-take the slow path when unlocking the mutex. If the mutex doesn't have any
-waiters, the owner field of the mutex would equal the current process and
-the mutex can be unlocked by just replacing the owner field with NULL.
-
-If the owner field has the "Has Waiters" bit set (or CMPXCHG is not available),
-the slow unlock path is taken.
-
-The first thing done in the slow unlock path is to take the wait_lock of the
-mutex. This synchronizes the locking and unlocking of the mutex.
-
-A check is made to see if the mutex has waiters or not. On architectures that
-do not have CMPXCHG, this is the location that the owner of the mutex will
-determine if a waiter needs to be awoken or not. On architectures that
-do have CMPXCHG, that check is done in the fast path, but it is still needed
-in the slow path too. If a waiter of a mutex woke up because of a signal
-or timeout between the time the owner failed the fast path CMPXCHG check and
-the grabbing of the wait_lock, the mutex may not have any waiters, thus the
-owner still needs to make this check. If there are no waiters then the mutex
-owner field is set to NULL, the wait_lock is released and nothing more is
-needed.
-
-If there are waiters, then we need to wake one up and give that waiter
-pending ownership.
-
-On the wake up code, the pi_lock of the current owner is taken. The top
-waiter of the lock is found and removed from the wait_list of the mutex
-as well as the pi_list of the current owner. The task field of the new
-pending owner's waiter structure is set to NULL, and the owner field of the
-mutex is set to the new owner with the "Pending Owner" bit set, as well
-as the "Has Waiters" bit if there still are other processes blocked on the
-mutex.
-
-The pi_lock of the previous owner is released, and the new pending owner's
-pi_lock is taken. Remember that this is the trick to prevent the race
-condition in rt_mutex_adjust_prio_chain from adding itself as a waiter
-on the mutex.
-
-We now clear the "pi_blocked_on" field of the new pending owner, and if
-the mutex still has waiters pending, we add the new top waiter to the pi_list
-of the pending owner.
-
-Finally we unlock the pi_lock of the pending owner and wake it up.
-
-
-Contact
--------
-
-For updates on this document, please email Steven Rostedt <rostedt@goodmis.org>
-
-
-Credits
--------
-
-Author: Steven Rostedt <rostedt@goodmis.org>
-
-Reviewers: Ingo Molnar, Thomas Gleixner, Thomas Duetsch, and Randy Dunlap
-
-Updates
--------
-
-This document was originally written for 2.6.17-rc3-mm1
diff --git a/Documentation/rt-mutex.txt b/Documentation/rt-mutex.txt
deleted file mode 100644
index 243393d..0000000
--- a/Documentation/rt-mutex.txt
+++ /dev/null
@@ -1,79 +0,0 @@
-RT-mutex subsystem with PI support
-----------------------------------
-
-RT-mutexes with priority inheritance are used to support PI-futexes,
-which enable pthread_mutex_t priority inheritance attributes
-(PTHREAD_PRIO_INHERIT). [See Documentation/pi-futex.txt for more details
-about PI-futexes.]
-
-This technology was developed in the -rt tree and streamlined for
-pthread_mutex support.
-
-Basic principles:
------------------
-
-RT-mutexes extend the semantics of simple mutexes by the priority
-inheritance protocol.
-
-A low priority owner of a rt-mutex inherits the priority of a higher
-priority waiter until the rt-mutex is released. If the temporarily
-boosted owner blocks on a rt-mutex itself it propagates the priority
-boosting to the owner of the other rt_mutex it gets blocked on. The
-priority boosting is immediately removed once the rt_mutex has been
-unlocked.
-
-This approach allows us to shorten the block of high-prio tasks on
-mutexes which protect shared resources. Priority inheritance is not a
-magic bullet for poorly designed applications, but it allows
-well-designed applications to use userspace locks in critical parts of
-an high priority thread, without losing determinism.
-
-The enqueueing of the waiters into the rtmutex waiter list is done in
-priority order. For same priorities FIFO order is chosen. For each
-rtmutex, only the top priority waiter is enqueued into the owner's
-priority waiters list. This list too queues in priority order. Whenever
-the top priority waiter of a task changes (for example it timed out or
-got a signal), the priority of the owner task is readjusted. [The
-priority enqueueing is handled by "plists", see include/linux/plist.h
-for more details.]
-
-RT-mutexes are optimized for fastpath operations and have no internal
-locking overhead when locking an uncontended mutex or unlocking a mutex
-without waiters. The optimized fastpath operations require cmpxchg
-support. [If that is not available then the rt-mutex internal spinlock
-is used]
-
-The state of the rt-mutex is tracked via the owner field of the rt-mutex
-structure:
-
-rt_mutex->owner holds the task_struct pointer of the owner. Bit 0 and 1
-are used to keep track of the "owner is pending" and "rtmutex has
-waiters" state.
-
- owner bit1 bit0
- NULL 0 0 mutex is free (fast acquire possible)
- NULL 0 1 invalid state
- NULL 1 0 Transitional state*
- NULL 1 1 invalid state
- taskpointer 0 0 mutex is held (fast release possible)
- taskpointer 0 1 task is pending owner
- taskpointer 1 0 mutex is held and has waiters
- taskpointer 1 1 task is pending owner and mutex has waiters
-
-Pending-ownership handling is a performance optimization:
-pending-ownership is assigned to the first (highest priority) waiter of
-the mutex, when the mutex is released. The thread is woken up and once
-it starts executing it can acquire the mutex. Until the mutex is taken
-by it (bit 0 is cleared) a competing higher priority thread can "steal"
-the mutex which puts the woken up thread back on the waiters list.
-
-The pending-ownership optimization is especially important for the
-uninterrupted workflow of high-prio tasks which repeatedly
-takes/releases locks that have lower-prio waiters. Without this
-optimization the higher-prio thread would ping-pong to the lower-prio
-task [because at unlock time we always assign a new owner].
-
-(*) The "mutex has waiters" bit gets set to take the lock. If the lock
-doesn't already have an owner, this bit is quickly cleared if there are
-no waiters. So this is a transitional state to synchronize with looking
-at the owner field of the mutex and the mutex owner releasing the lock.
diff --git a/Documentation/spinlocks.txt b/Documentation/spinlocks.txt
deleted file mode 100644
index 97eaf57..0000000
--- a/Documentation/spinlocks.txt
+++ /dev/null
@@ -1,167 +0,0 @@
-Lesson 1: Spin locks
-
-The most basic primitive for locking is spinlock.
-
-static DEFINE_SPINLOCK(xxx_lock);
-
- unsigned long flags;
-
- spin_lock_irqsave(&xxx_lock, flags);
- ... critical section here ..
- spin_unlock_irqrestore(&xxx_lock, flags);
-
-The above is always safe. It will disable interrupts _locally_, but the
-spinlock itself will guarantee the global lock, so it will guarantee that
-there is only one thread-of-control within the region(s) protected by that
-lock. This works well even under UP also, so the code does _not_ need to
-worry about UP vs SMP issues: the spinlocks work correctly under both.
-
- NOTE! Implications of spin_locks for memory are further described in:
-
- Documentation/memory-barriers.txt
- (5) LOCK operations.
- (6) UNLOCK operations.
-
-The above is usually pretty simple (you usually need and want only one
-spinlock for most things - using more than one spinlock can make things a
-lot more complex and even slower and is usually worth it only for
-sequences that you _know_ need to be split up: avoid it at all cost if you
-aren't sure).
-
-This is really the only really hard part about spinlocks: once you start
-using spinlocks they tend to expand to areas you might not have noticed
-before, because you have to make sure the spinlocks correctly protect the
-shared data structures _everywhere_ they are used. The spinlocks are most
-easily added to places that are completely independent of other code (for
-example, internal driver data structures that nobody else ever touches).
-
- NOTE! The spin-lock is safe only when you _also_ use the lock itself
- to do locking across CPU's, which implies that EVERYTHING that
- touches a shared variable has to agree about the spinlock they want
- to use.
-
-----
-
-Lesson 2: reader-writer spinlocks.
-
-If your data accesses have a very natural pattern where you usually tend
-to mostly read from the shared variables, the reader-writer locks
-(rw_lock) versions of the spinlocks are sometimes useful. They allow multiple
-readers to be in the same critical region at once, but if somebody wants
-to change the variables it has to get an exclusive write lock.
-
- NOTE! reader-writer locks require more atomic memory operations than
- simple spinlocks. Unless the reader critical section is long, you
- are better off just using spinlocks.
-
-The routines look the same as above:
-
- rwlock_t xxx_lock = __RW_LOCK_UNLOCKED(xxx_lock);
-
- unsigned long flags;
-
- read_lock_irqsave(&xxx_lock, flags);
- .. critical section that only reads the info ...
- read_unlock_irqrestore(&xxx_lock, flags);
-
- write_lock_irqsave(&xxx_lock, flags);
- .. read and write exclusive access to the info ...
- write_unlock_irqrestore(&xxx_lock, flags);
-
-The above kind of lock may be useful for complex data structures like
-linked lists, especially searching for entries without changing the list
-itself. The read lock allows many concurrent readers. Anything that
-_changes_ the list will have to get the write lock.
-
- NOTE! RCU is better for list traversal, but requires careful
- attention to design detail (see Documentation/RCU/listRCU.txt).
-
-Also, you cannot "upgrade" a read-lock to a write-lock, so if you at _any_
-time need to do any changes (even if you don't do it every time), you have
-to get the write-lock at the very beginning.
-
- NOTE! We are working hard to remove reader-writer spinlocks in most
- cases, so please don't add a new one without consensus. (Instead, see
- Documentation/RCU/rcu.txt for complete information.)
-
-----
-
-Lesson 3: spinlocks revisited.
-
-The single spin-lock primitives above are by no means the only ones. They
-are the most safe ones, and the ones that work under all circumstances,
-but partly _because_ they are safe they are also fairly slow. They are slower
-than they'd need to be, because they do have to disable interrupts
-(which is just a single instruction on a x86, but it's an expensive one -
-and on other architectures it can be worse).
-
-If you have a case where you have to protect a data structure across
-several CPU's and you want to use spinlocks you can potentially use
-cheaper versions of the spinlocks. IFF you know that the spinlocks are
-never used in interrupt handlers, you can use the non-irq versions:
-
- spin_lock(&lock);
- ...
- spin_unlock(&lock);
-
-(and the equivalent read-write versions too, of course). The spinlock will
-guarantee the same kind of exclusive access, and it will be much faster.
-This is useful if you know that the data in question is only ever
-manipulated from a "process context", ie no interrupts involved.
-
-The reasons you mustn't use these versions if you have interrupts that
-play with the spinlock is that you can get deadlocks:
-
- spin_lock(&lock);
- ...
- <- interrupt comes in:
- spin_lock(&lock);
-
-where an interrupt tries to lock an already locked variable. This is ok if
-the other interrupt happens on another CPU, but it is _not_ ok if the
-interrupt happens on the same CPU that already holds the lock, because the
-lock will obviously never be released (because the interrupt is waiting
-for the lock, and the lock-holder is interrupted by the interrupt and will
-not continue until the interrupt has been processed).
-
-(This is also the reason why the irq-versions of the spinlocks only need
-to disable the _local_ interrupts - it's ok to use spinlocks in interrupts
-on other CPU's, because an interrupt on another CPU doesn't interrupt the
-CPU that holds the lock, so the lock-holder can continue and eventually
-releases the lock).
-
-Note that you can be clever with read-write locks and interrupts. For
-example, if you know that the interrupt only ever gets a read-lock, then
-you can use a non-irq version of read locks everywhere - because they
-don't block on each other (and thus there is no dead-lock wrt interrupts.
-But when you do the write-lock, you have to use the irq-safe version.
-
-For an example of being clever with rw-locks, see the "waitqueue_lock"
-handling in kernel/sched/core.c - nothing ever _changes_ a wait-queue from
-within an interrupt, they only read the queue in order to know whom to
-wake up. So read-locks are safe (which is good: they are very common
-indeed), while write-locks need to protect themselves against interrupts.
-
- Linus
-
-----
-
-Reference information:
-
-For dynamic initialization, use spin_lock_init() or rwlock_init() as
-appropriate:
-
- spinlock_t xxx_lock;
- rwlock_t xxx_rw_lock;
-
- static int __init xxx_init(void)
- {
- spin_lock_init(&xxx_lock);
- rwlock_init(&xxx_rw_lock);
- ...
- }
-
- module_init(xxx_init);
-
-For static initialization, use DEFINE_SPINLOCK() / DEFINE_RWLOCK() or
-__SPIN_LOCK_UNLOCKED() / __RW_LOCK_UNLOCKED() as appropriate.
diff --git a/Documentation/ww-mutex-design.txt b/Documentation/ww-mutex-design.txt
deleted file mode 100644
index 8a112dc..0000000
--- a/Documentation/ww-mutex-design.txt
+++ /dev/null
@@ -1,344 +0,0 @@
-Wait/Wound Deadlock-Proof Mutex Design
-======================================
-
-Please read mutex-design.txt first, as it applies to wait/wound mutexes too.
-
-Motivation for WW-Mutexes
--------------------------
-
-GPU's do operations that commonly involve many buffers. Those buffers
-can be shared across contexts/processes, exist in different memory
-domains (for example VRAM vs system memory), and so on. And with
-PRIME / dmabuf, they can even be shared across devices. So there are
-a handful of situations where the driver needs to wait for buffers to
-become ready. If you think about this in terms of waiting on a buffer
-mutex for it to become available, this presents a problem because
-there is no way to guarantee that buffers appear in a execbuf/batch in
-the same order in all contexts. That is directly under control of
-userspace, and a result of the sequence of GL calls that an application
-makes. Which results in the potential for deadlock. The problem gets
-more complex when you consider that the kernel may need to migrate the
-buffer(s) into VRAM before the GPU operates on the buffer(s), which
-may in turn require evicting some other buffers (and you don't want to
-evict other buffers which are already queued up to the GPU), but for a
-simplified understanding of the problem you can ignore this.
-
-The algorithm that the TTM graphics subsystem came up with for dealing with
-this problem is quite simple. For each group of buffers (execbuf) that need
-to be locked, the caller would be assigned a unique reservation id/ticket,
-from a global counter. In case of deadlock while locking all the buffers
-associated with a execbuf, the one with the lowest reservation ticket (i.e.
-the oldest task) wins, and the one with the higher reservation id (i.e. the
-younger task) unlocks all of the buffers that it has already locked, and then
-tries again.
-
-In the RDBMS literature this deadlock handling approach is called wait/wound:
-The older tasks waits until it can acquire the contended lock. The younger tasks
-needs to back off and drop all the locks it is currently holding, i.e. the
-younger task is wounded.
-
-Concepts
---------
-
-Compared to normal mutexes two additional concepts/objects show up in the lock
-interface for w/w mutexes:
-
-Acquire context: To ensure eventual forward progress it is important the a task
-trying to acquire locks doesn't grab a new reservation id, but keeps the one it
-acquired when starting the lock acquisition. This ticket is stored in the
-acquire context. Furthermore the acquire context keeps track of debugging state
-to catch w/w mutex interface abuse.
-
-W/w class: In contrast to normal mutexes the lock class needs to be explicit for
-w/w mutexes, since it is required to initialize the acquire context.
-
-Furthermore there are three different class of w/w lock acquire functions:
-
-* Normal lock acquisition with a context, using ww_mutex_lock.
-
-* Slowpath lock acquisition on the contending lock, used by the wounded task
- after having dropped all already acquired locks. These functions have the
- _slow postfix.
-
- From a simple semantics point-of-view the _slow functions are not strictly
- required, since simply calling the normal ww_mutex_lock functions on the
- contending lock (after having dropped all other already acquired locks) will
- work correctly. After all if no other ww mutex has been acquired yet there's
- no deadlock potential and hence the ww_mutex_lock call will block and not
- prematurely return -EDEADLK. The advantage of the _slow functions is in
- interface safety:
- - ww_mutex_lock has a __must_check int return type, whereas ww_mutex_lock_slow
- has a void return type. Note that since ww mutex code needs loops/retries
- anyway the __must_check doesn't result in spurious warnings, even though the
- very first lock operation can never fail.
- - When full debugging is enabled ww_mutex_lock_slow checks that all acquired
- ww mutex have been released (preventing deadlocks) and makes sure that we
- block on the contending lock (preventing spinning through the -EDEADLK
- slowpath until the contended lock can be acquired).
-
-* Functions to only acquire a single w/w mutex, which results in the exact same
- semantics as a normal mutex. This is done by calling ww_mutex_lock with a NULL
- context.
-
- Again this is not strictly required. But often you only want to acquire a
- single lock in which case it's pointless to set up an acquire context (and so
- better to avoid grabbing a deadlock avoidance ticket).
-
-Of course, all the usual variants for handling wake-ups due to signals are also
-provided.
-
-Usage
------
-
-Three different ways to acquire locks within the same w/w class. Common
-definitions for methods #1 and #2:
-
-static DEFINE_WW_CLASS(ww_class);
-
-struct obj {
- struct ww_mutex lock;
- /* obj data */
-};
-
-struct obj_entry {
- struct list_head head;
- struct obj *obj;
-};
-
-Method 1, using a list in execbuf->buffers that's not allowed to be reordered.
-This is useful if a list of required objects is already tracked somewhere.
-Furthermore the lock helper can use propagate the -EALREADY return code back to
-the caller as a signal that an object is twice on the list. This is useful if
-the list is constructed from userspace input and the ABI requires userspace to
-not have duplicate entries (e.g. for a gpu commandbuffer submission ioctl).
-
-int lock_objs(struct list_head *list, struct ww_acquire_ctx *ctx)
-{
- struct obj *res_obj = NULL;
- struct obj_entry *contended_entry = NULL;
- struct obj_entry *entry;
-
- ww_acquire_init(ctx, &ww_class);
-
-retry:
- list_for_each_entry (entry, list, head) {
- if (entry->obj == res_obj) {
- res_obj = NULL;
- continue;
- }
- ret = ww_mutex_lock(&entry->obj->lock, ctx);
- if (ret < 0) {
- contended_entry = entry;
- goto err;
- }
- }
-
- ww_acquire_done(ctx);
- return 0;
-
-err:
- list_for_each_entry_continue_reverse (entry, list, head)
- ww_mutex_unlock(&entry->obj->lock);
-
- if (res_obj)
- ww_mutex_unlock(&res_obj->lock);
-
- if (ret == -EDEADLK) {
- /* we lost out in a seqno race, lock and retry.. */
- ww_mutex_lock_slow(&contended_entry->obj->lock, ctx);
- res_obj = contended_entry->obj;
- goto retry;
- }
- ww_acquire_fini(ctx);
-
- return ret;
-}
-
-Method 2, using a list in execbuf->buffers that can be reordered. Same semantics
-of duplicate entry detection using -EALREADY as method 1 above. But the
-list-reordering allows for a bit more idiomatic code.
-
-int lock_objs(struct list_head *list, struct ww_acquire_ctx *ctx)
-{
- struct obj_entry *entry, *entry2;
-
- ww_acquire_init(ctx, &ww_class);
-
- list_for_each_entry (entry, list, head) {
- ret = ww_mutex_lock(&entry->obj->lock, ctx);
- if (ret < 0) {
- entry2 = entry;
-
- list_for_each_entry_continue_reverse (entry2, list, head)
- ww_mutex_unlock(&entry2->obj->lock);
-
- if (ret != -EDEADLK) {
- ww_acquire_fini(ctx);
- return ret;
- }
-
- /* we lost out in a seqno race, lock and retry.. */
- ww_mutex_lock_slow(&entry->obj->lock, ctx);
-
- /*
- * Move buf to head of the list, this will point
- * buf->next to the first unlocked entry,
- * restarting the for loop.
- */
- list_del(&entry->head);
- list_add(&entry->head, list);
- }
- }
-
- ww_acquire_done(ctx);
- return 0;
-}
-
-Unlocking works the same way for both methods #1 and #2:
-
-void unlock_objs(struct list_head *list, struct ww_acquire_ctx *ctx)
-{
- struct obj_entry *entry;
-
- list_for_each_entry (entry, list, head)
- ww_mutex_unlock(&entry->obj->lock);
-
- ww_acquire_fini(ctx);
-}
-
-Method 3 is useful if the list of objects is constructed ad-hoc and not upfront,
-e.g. when adjusting edges in a graph where each node has its own ww_mutex lock,
-and edges can only be changed when holding the locks of all involved nodes. w/w
-mutexes are a natural fit for such a case for two reasons:
-- They can handle lock-acquisition in any order which allows us to start walking
- a graph from a starting point and then iteratively discovering new edges and
- locking down the nodes those edges connect to.
-- Due to the -EALREADY return code signalling that a given objects is already
- held there's no need for additional book-keeping to break cycles in the graph
- or keep track off which looks are already held (when using more than one node
- as a starting point).
-
-Note that this approach differs in two important ways from the above methods:
-- Since the list of objects is dynamically constructed (and might very well be
- different when retrying due to hitting the -EDEADLK wound condition) there's
- no need to keep any object on a persistent list when it's not locked. We can
- therefore move the list_head into the object itself.
-- On the other hand the dynamic object list construction also means that the -EALREADY return
- code can't be propagated.
-
-Note also that methods #1 and #2 and method #3 can be combined, e.g. to first lock a
-list of starting nodes (passed in from userspace) using one of the above
-methods. And then lock any additional objects affected by the operations using
-method #3 below. The backoff/retry procedure will be a bit more involved, since
-when the dynamic locking step hits -EDEADLK we also need to unlock all the
-objects acquired with the fixed list. But the w/w mutex debug checks will catch
-any interface misuse for these cases.
-
-Also, method 3 can't fail the lock acquisition step since it doesn't return
--EALREADY. Of course this would be different when using the _interruptible
-variants, but that's outside of the scope of these examples here.
-
-struct obj {
- struct ww_mutex ww_mutex;
- struct list_head locked_list;
-};
-
-static DEFINE_WW_CLASS(ww_class);
-
-void __unlock_objs(struct list_head *list)
-{
- struct obj *entry, *temp;
-
- list_for_each_entry_safe (entry, temp, list, locked_list) {
- /* need to do that before unlocking, since only the current lock holder is
- allowed to use object */
- list_del(&entry->locked_list);
- ww_mutex_unlock(entry->ww_mutex)
- }
-}
-
-void lock_objs(struct list_head *list, struct ww_acquire_ctx *ctx)
-{
- struct obj *obj;
-
- ww_acquire_init(ctx, &ww_class);
-
-retry:
- /* re-init loop start state */
- loop {
- /* magic code which walks over a graph and decides which objects
- * to lock */
-
- ret = ww_mutex_lock(obj->ww_mutex, ctx);
- if (ret == -EALREADY) {
- /* we have that one already, get to the next object */
- continue;
- }
- if (ret == -EDEADLK) {
- __unlock_objs(list);
-
- ww_mutex_lock_slow(obj, ctx);
- list_add(&entry->locked_list, list);
- goto retry;
- }
-
- /* locked a new object, add it to the list */
- list_add_tail(&entry->locked_list, list);
- }
-
- ww_acquire_done(ctx);
- return 0;
-}
-
-void unlock_objs(struct list_head *list, struct ww_acquire_ctx *ctx)
-{
- __unlock_objs(list);
- ww_acquire_fini(ctx);
-}
-
-Method 4: Only lock one single objects. In that case deadlock detection and
-prevention is obviously overkill, since with grabbing just one lock you can't
-produce a deadlock within just one class. To simplify this case the w/w mutex
-api can be used with a NULL context.
-
-Implementation Details
-----------------------
-
-Design:
- ww_mutex currently encapsulates a struct mutex, this means no extra overhead for
- normal mutex locks, which are far more common. As such there is only a small
- increase in code size if wait/wound mutexes are not used.
-
- In general, not much contention is expected. The locks are typically used to
- serialize access to resources for devices. The only way to make wakeups
- smarter would be at the cost of adding a field to struct mutex_waiter. This
- would add overhead to all cases where normal mutexes are used, and
- ww_mutexes are generally less performance sensitive.
-
-Lockdep:
- Special care has been taken to warn for as many cases of api abuse
- as possible. Some common api abuses will be caught with
- CONFIG_DEBUG_MUTEXES, but CONFIG_PROVE_LOCKING is recommended.
-
- Some of the errors which will be warned about:
- - Forgetting to call ww_acquire_fini or ww_acquire_init.
- - Attempting to lock more mutexes after ww_acquire_done.
- - Attempting to lock the wrong mutex after -EDEADLK and
- unlocking all mutexes.
- - Attempting to lock the right mutex after -EDEADLK,
- before unlocking all mutexes.
-
- - Calling ww_mutex_lock_slow before -EDEADLK was returned.
-
- - Unlocking mutexes with the wrong unlock function.
- - Calling one of the ww_acquire_* twice on the same context.
- - Using a different ww_class for the mutex than for the ww_acquire_ctx.
- - Normal lockdep errors that can result in deadlocks.
-
- Some of the lockdep errors that can result in deadlocks:
- - Calling ww_acquire_init to initialize a second ww_acquire_ctx before
- having called ww_acquire_fini on the first.
- - 'normal' deadlocks that can occur.
-
-FIXME: Update this section once we have the TASK_DEADLOCK task state flag magic
-implemented.
diff --git a/MAINTAINERS b/MAINTAINERS
index b4534ee..053d8c2 100644
--- a/MAINTAINERS
+++ b/MAINTAINERS
@@ -5516,8 +5516,8 @@ M: Ingo Molnar <mingo@redhat.com>
L: linux-kernel@vger.kernel.org
T: git git://git.kernel.org/pub/scm/linux/kernel/git/tip/tip.git core/locking
S: Maintained
-F: Documentation/lockdep*.txt
-F: Documentation/lockstat.txt
+F: Documentation/locking/lockdep*.txt
+F: Documentation/locking/lockstat.txt
F: include/linux/lockdep.h
F: kernel/locking/
diff --git a/drivers/gpu/drm/drm_modeset_lock.c b/drivers/gpu/drm/drm_modeset_lock.c
index 0dc57d5..3a02e5e 100644
--- a/drivers/gpu/drm/drm_modeset_lock.c
+++ b/drivers/gpu/drm/drm_modeset_lock.c
@@ -35,7 +35,7 @@
* of extra utility/tracking out of our acquire-ctx. This is provided
* by drm_modeset_lock / drm_modeset_acquire_ctx.
*
- * For basic principles of ww_mutex, see: Documentation/ww-mutex-design.txt
+ * For basic principles of ww_mutex, see: Documentation/locking/ww-mutex-design.txt
*
* The basic usage pattern is to:
*
diff --git a/include/linux/lockdep.h b/include/linux/lockdep.h
index dadd6ba..b5a84b62 100644
--- a/include/linux/lockdep.h
+++ b/include/linux/lockdep.h
@@ -4,7 +4,7 @@
* Copyright (C) 2006,2007 Red Hat, Inc., Ingo Molnar <mingo@redhat.com>
* Copyright (C) 2007 Red Hat, Inc., Peter Zijlstra <pzijlstr@redhat.com>
*
- * see Documentation/lockdep-design.txt for more details.
+ * see Documentation/locking/lockdep-design.txt for more details.
*/
#ifndef __LINUX_LOCKDEP_H
#define __LINUX_LOCKDEP_H
diff --git a/include/linux/mutex.h b/include/linux/mutex.h
index e4c2941..cc31498 100644
--- a/include/linux/mutex.h
+++ b/include/linux/mutex.h
@@ -133,7 +133,7 @@ static inline int mutex_is_locked(struct mutex *lock)
/*
* See kernel/locking/mutex.c for detailed documentation of these APIs.
- * Also see Documentation/mutex-design.txt.
+ * Also see Documentation/locking/mutex-design.txt.
*/
#ifdef CONFIG_DEBUG_LOCK_ALLOC
extern void mutex_lock_nested(struct mutex *lock, unsigned int subclass);
diff --git a/include/linux/rwsem.h b/include/linux/rwsem.h
index 035d3c5..8f498cd 100644
--- a/include/linux/rwsem.h
+++ b/include/linux/rwsem.h
@@ -149,7 +149,7 @@ extern void downgrade_write(struct rw_semaphore *sem);
* static then another method for expressing nested locking is
* the explicit definition of lock class keys and the use of
* lockdep_set_class() at lock initialization time.
- * See Documentation/lockdep-design.txt for more details.)
+ * See Documentation/locking/lockdep-design.txt for more details.)
*/
extern void down_read_nested(struct rw_semaphore *sem, int subclass);
extern void down_write_nested(struct rw_semaphore *sem, int subclass);
diff --git a/kernel/locking/mutex.c b/kernel/locking/mutex.c
index a925597..50933f3 100644
--- a/kernel/locking/mutex.c
+++ b/kernel/locking/mutex.c
@@ -15,7 +15,7 @@
* by Steven Rostedt, based on work by Gregory Haskins, Peter Morreale
* and Sven Dietrich.
*
- * Also see Documentation/mutex-design.txt.
+ * Also see Documentation/locking/mutex-design.txt.
*/
#include <linux/mutex.h>
#include <linux/ww_mutex.h>
diff --git a/kernel/locking/rtmutex.c b/kernel/locking/rtmutex.c
index a0ea2a1..7c98873 100644
--- a/kernel/locking/rtmutex.c
+++ b/kernel/locking/rtmutex.c
@@ -8,7 +8,7 @@
* Copyright (C) 2005 Kihon Technologies Inc., Steven Rostedt
* Copyright (C) 2006 Esben Nielsen
*
- * See Documentation/rt-mutex-design.txt for details.
+ * See Documentation/locking/rt-mutex-design.txt for details.
*/
#include <linux/spinlock.h>
#include <linux/export.h>
diff --git a/lib/Kconfig.debug b/lib/Kconfig.debug
index 03f461f..269da61 100644
--- a/lib/Kconfig.debug
+++ b/lib/Kconfig.debug
@@ -924,7 +924,7 @@ config PROVE_LOCKING
the proof of observed correctness is also maintained for an
arbitrary combination of these separate locking variants.
- For more details, see Documentation/lockdep-design.txt.
+ For more details, see Documentation/locking/lockdep-design.txt.
config LOCKDEP
bool
@@ -945,7 +945,7 @@ config LOCK_STAT
help
This feature enables tracking lock contention points
- For more details, see Documentation/lockstat.txt
+ For more details, see Documentation/locking/lockstat.txt
This also enables lock events required by "perf lock",
subcommand of perf.
--
1.8.1.4
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