[Feature]add MT2731_MP2_MR2_SVN388 baseline version
Change-Id: Ief04314834b31e27effab435d3ca8ba33b499059
diff --git a/src/kernel/linux/v4.14/Documentation/locking/00-INDEX b/src/kernel/linux/v4.14/Documentation/locking/00-INDEX
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+++ b/src/kernel/linux/v4.14/Documentation/locking/00-INDEX
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+00-INDEX
+ - this file.
+lockdep-design.txt
+ - documentation on the runtime locking correctness validator.
+lockstat.txt
+ - info on collecting statistics on locks (and contention).
+mutex-design.txt
+ - info on the generic mutex subsystem.
+rt-mutex-design.txt
+ - description of the RealTime mutex implementation design.
+rt-mutex.txt
+ - desc. of RT-mutex subsystem with PI (Priority Inheritance) support.
+spinlocks.txt
+ - info on using spinlocks to provide exclusive access in kernel.
+ww-mutex-design.txt
+ - Intro to Mutex wait/would deadlock handling.s
diff --git a/src/kernel/linux/v4.14/Documentation/locking/crossrelease.txt b/src/kernel/linux/v4.14/Documentation/locking/crossrelease.txt
new file mode 100644
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+Crossrelease
+============
+
+Started by Byungchul Park <byungchul.park@lge.com>
+
+Contents:
+
+ (*) Background
+
+ - What causes deadlock
+ - How lockdep works
+
+ (*) Limitation
+
+ - Limit lockdep
+ - Pros from the limitation
+ - Cons from the limitation
+ - Relax the limitation
+
+ (*) Crossrelease
+
+ - Introduce crossrelease
+ - Introduce commit
+
+ (*) Implementation
+
+ - Data structures
+ - How crossrelease works
+
+ (*) Optimizations
+
+ - Avoid duplication
+ - Lockless for hot paths
+
+ (*) APPENDIX A: What lockdep does to work aggresively
+
+ (*) APPENDIX B: How to avoid adding false dependencies
+
+
+==========
+Background
+==========
+
+What causes deadlock
+--------------------
+
+A deadlock occurs when a context is waiting for an event to happen,
+which is impossible because another (or the) context who can trigger the
+event is also waiting for another (or the) event to happen, which is
+also impossible due to the same reason.
+
+For example:
+
+ A context going to trigger event C is waiting for event A to happen.
+ A context going to trigger event A is waiting for event B to happen.
+ A context going to trigger event B is waiting for event C to happen.
+
+A deadlock occurs when these three wait operations run at the same time,
+because event C cannot be triggered if event A does not happen, which in
+turn cannot be triggered if event B does not happen, which in turn
+cannot be triggered if event C does not happen. After all, no event can
+be triggered since any of them never meets its condition to wake up.
+
+A dependency might exist between two waiters and a deadlock might happen
+due to an incorrect releationship between dependencies. Thus, we must
+define what a dependency is first. A dependency exists between them if:
+
+ 1. There are two waiters waiting for each event at a given time.
+ 2. The only way to wake up each waiter is to trigger its event.
+ 3. Whether one can be woken up depends on whether the other can.
+
+Each wait in the example creates its dependency like:
+
+ Event C depends on event A.
+ Event A depends on event B.
+ Event B depends on event C.
+
+ NOTE: Precisely speaking, a dependency is one between whether a
+ waiter for an event can be woken up and whether another waiter for
+ another event can be woken up. However from now on, we will describe
+ a dependency as if it's one between an event and another event for
+ simplicity.
+
+And they form circular dependencies like:
+
+ -> C -> A -> B -
+ / \
+ \ /
+ ----------------
+
+ where 'A -> B' means that event A depends on event B.
+
+Such circular dependencies lead to a deadlock since no waiter can meet
+its condition to wake up as described.
+
+CONCLUSION
+
+Circular dependencies cause a deadlock.
+
+
+How lockdep works
+-----------------
+
+Lockdep tries to detect a deadlock by checking dependencies created by
+lock operations, acquire and release. Waiting for a lock corresponds to
+waiting for an event, and releasing a lock corresponds to triggering an
+event in the previous section.
+
+In short, lockdep does:
+
+ 1. Detect a new dependency.
+ 2. Add the dependency into a global graph.
+ 3. Check if that makes dependencies circular.
+ 4. Report a deadlock or its possibility if so.
+
+For example, consider a graph built by lockdep that looks like:
+
+ A -> B -
+ \
+ -> E
+ /
+ C -> D -
+
+ where A, B,..., E are different lock classes.
+
+Lockdep will add a dependency into the graph on detection of a new
+dependency. For example, it will add a dependency 'E -> C' when a new
+dependency between lock E and lock C is detected. Then the graph will be:
+
+ A -> B -
+ \
+ -> E -
+ / \
+ -> C -> D - \
+ / /
+ \ /
+ ------------------
+
+ where A, B,..., E are different lock classes.
+
+This graph contains a subgraph which demonstrates circular dependencies:
+
+ -> E -
+ / \
+ -> C -> D - \
+ / /
+ \ /
+ ------------------
+
+ where C, D and E are different lock classes.
+
+This is the condition under which a deadlock might occur. Lockdep
+reports it on detection after adding a new dependency. This is the way
+how lockdep works.
+
+CONCLUSION
+
+Lockdep detects a deadlock or its possibility by checking if circular
+dependencies were created after adding each new dependency.
+
+
+==========
+Limitation
+==========
+
+Limit lockdep
+-------------
+
+Limiting lockdep to work on only typical locks e.g. spin locks and
+mutexes, which are released within the acquire context, the
+implementation becomes simple but its capacity for detection becomes
+limited. Let's check pros and cons in next section.
+
+
+Pros from the limitation
+------------------------
+
+Given the limitation, when acquiring a lock, locks in a held_locks
+cannot be released if the context cannot acquire it so has to wait to
+acquire it, which means all waiters for the locks in the held_locks are
+stuck. It's an exact case to create dependencies between each lock in
+the held_locks and the lock to acquire.
+
+For example:
+
+ CONTEXT X
+ ---------
+ acquire A
+ acquire B /* Add a dependency 'A -> B' */
+ release B
+ release A
+
+ where A and B are different lock classes.
+
+When acquiring lock A, the held_locks of CONTEXT X is empty thus no
+dependency is added. But when acquiring lock B, lockdep detects and adds
+a new dependency 'A -> B' between lock A in the held_locks and lock B.
+They can be simply added whenever acquiring each lock.
+
+And data required by lockdep exists in a local structure, held_locks
+embedded in task_struct. Forcing to access the data within the context,
+lockdep can avoid racy problems without explicit locks while handling
+the local data.
+
+Lastly, lockdep only needs to keep locks currently being held, to build
+a dependency graph. However, relaxing the limitation, it needs to keep
+even locks already released, because a decision whether they created
+dependencies might be long-deferred.
+
+To sum up, we can expect several advantages from the limitation:
+
+ 1. Lockdep can easily identify a dependency when acquiring a lock.
+ 2. Races are avoidable while accessing local locks in a held_locks.
+ 3. Lockdep only needs to keep locks currently being held.
+
+CONCLUSION
+
+Given the limitation, the implementation becomes simple and efficient.
+
+
+Cons from the limitation
+------------------------
+
+Given the limitation, lockdep is applicable only to typical locks. For
+example, page locks for page access or completions for synchronization
+cannot work with lockdep.
+
+Can we detect deadlocks below, under the limitation?
+
+Example 1:
+
+ CONTEXT X CONTEXT Y CONTEXT Z
+ --------- --------- ----------
+ mutex_lock A
+ lock_page B
+ lock_page B
+ mutex_lock A /* DEADLOCK */
+ unlock_page B held by X
+ unlock_page B
+ mutex_unlock A
+ mutex_unlock A
+
+ where A and B are different lock classes.
+
+No, we cannot.
+
+Example 2:
+
+ CONTEXT X CONTEXT Y
+ --------- ---------
+ mutex_lock A
+ mutex_lock A
+ wait_for_complete B /* DEADLOCK */
+ complete B
+ mutex_unlock A
+ mutex_unlock A
+
+ where A is a lock class and B is a completion variable.
+
+No, we cannot.
+
+CONCLUSION
+
+Given the limitation, lockdep cannot detect a deadlock or its
+possibility caused by page locks or completions.
+
+
+Relax the limitation
+--------------------
+
+Under the limitation, things to create dependencies are limited to
+typical locks. However, synchronization primitives like page locks and
+completions, which are allowed to be released in any context, also
+create dependencies and can cause a deadlock. So lockdep should track
+these locks to do a better job. We have to relax the limitation for
+these locks to work with lockdep.
+
+Detecting dependencies is very important for lockdep to work because
+adding a dependency means adding an opportunity to check whether it
+causes a deadlock. The more lockdep adds dependencies, the more it
+thoroughly works. Thus Lockdep has to do its best to detect and add as
+many true dependencies into a graph as possible.
+
+For example, considering only typical locks, lockdep builds a graph like:
+
+ A -> B -
+ \
+ -> E
+ /
+ C -> D -
+
+ where A, B,..., E are different lock classes.
+
+On the other hand, under the relaxation, additional dependencies might
+be created and added. Assuming additional 'FX -> C' and 'E -> GX' are
+added thanks to the relaxation, the graph will be:
+
+ A -> B -
+ \
+ -> E -> GX
+ /
+ FX -> C -> D -
+
+ where A, B,..., E, FX and GX are different lock classes, and a suffix
+ 'X' is added on non-typical locks.
+
+The latter graph gives us more chances to check circular dependencies
+than the former. However, it might suffer performance degradation since
+relaxing the limitation, with which design and implementation of lockdep
+can be efficient, might introduce inefficiency inevitably. So lockdep
+should provide two options, strong detection and efficient detection.
+
+Choosing efficient detection:
+
+ Lockdep works with only locks restricted to be released within the
+ acquire context. However, lockdep works efficiently.
+
+Choosing strong detection:
+
+ Lockdep works with all synchronization primitives. However, lockdep
+ suffers performance degradation.
+
+CONCLUSION
+
+Relaxing the limitation, lockdep can add additional dependencies giving
+additional opportunities to check circular dependencies.
+
+
+============
+Crossrelease
+============
+
+Introduce crossrelease
+----------------------
+
+In order to allow lockdep to handle additional dependencies by what
+might be released in any context, namely 'crosslock', we have to be able
+to identify those created by crosslocks. The proposed 'crossrelease'
+feature provoides a way to do that.
+
+Crossrelease feature has to do:
+
+ 1. Identify dependencies created by crosslocks.
+ 2. Add the dependencies into a dependency graph.
+
+That's all. Once a meaningful dependency is added into graph, then
+lockdep would work with the graph as it did. The most important thing
+crossrelease feature has to do is to correctly identify and add true
+dependencies into the global graph.
+
+A dependency e.g. 'A -> B' can be identified only in the A's release
+context because a decision required to identify the dependency can be
+made only in the release context. That is to decide whether A can be
+released so that a waiter for A can be woken up. It cannot be made in
+other than the A's release context.
+
+It's no matter for typical locks because each acquire context is same as
+its release context, thus lockdep can decide whether a lock can be
+released in the acquire context. However for crosslocks, lockdep cannot
+make the decision in the acquire context but has to wait until the
+release context is identified.
+
+Therefore, deadlocks by crosslocks cannot be detected just when it
+happens, because those cannot be identified until the crosslocks are
+released. However, deadlock possibilities can be detected and it's very
+worth. See 'APPENDIX A' section to check why.
+
+CONCLUSION
+
+Using crossrelease feature, lockdep can work with what might be released
+in any context, namely crosslock.
+
+
+Introduce commit
+----------------
+
+Since crossrelease defers the work adding true dependencies of
+crosslocks until they are actually released, crossrelease has to queue
+all acquisitions which might create dependencies with the crosslocks.
+Then it identifies dependencies using the queued data in batches at a
+proper time. We call it 'commit'.
+
+There are four types of dependencies:
+
+1. TT type: 'typical lock A -> typical lock B'
+
+ Just when acquiring B, lockdep can see it's in the A's release
+ context. So the dependency between A and B can be identified
+ immediately. Commit is unnecessary.
+
+2. TC type: 'typical lock A -> crosslock BX'
+
+ Just when acquiring BX, lockdep can see it's in the A's release
+ context. So the dependency between A and BX can be identified
+ immediately. Commit is unnecessary, too.
+
+3. CT type: 'crosslock AX -> typical lock B'
+
+ When acquiring B, lockdep cannot identify the dependency because
+ there's no way to know if it's in the AX's release context. It has
+ to wait until the decision can be made. Commit is necessary.
+
+4. CC type: 'crosslock AX -> crosslock BX'
+
+ When acquiring BX, lockdep cannot identify the dependency because
+ there's no way to know if it's in the AX's release context. It has
+ to wait until the decision can be made. Commit is necessary.
+ But, handling CC type is not implemented yet. It's a future work.
+
+Lockdep can work without commit for typical locks, but commit step is
+necessary once crosslocks are involved. Introducing commit, lockdep
+performs three steps. What lockdep does in each step is:
+
+1. Acquisition: For typical locks, lockdep does what it originally did
+ and queues the lock so that CT type dependencies can be checked using
+ it at the commit step. For crosslocks, it saves data which will be
+ used at the commit step and increases a reference count for it.
+
+2. Commit: No action is reauired for typical locks. For crosslocks,
+ lockdep adds CT type dependencies using the data saved at the
+ acquisition step.
+
+3. Release: No changes are required for typical locks. When a crosslock
+ is released, it decreases a reference count for it.
+
+CONCLUSION
+
+Crossrelease introduces commit step to handle dependencies of crosslocks
+in batches at a proper time.
+
+
+==============
+Implementation
+==============
+
+Data structures
+---------------
+
+Crossrelease introduces two main data structures.
+
+1. hist_lock
+
+ This is an array embedded in task_struct, for keeping lock history so
+ that dependencies can be added using them at the commit step. Since
+ it's local data, it can be accessed locklessly in the owner context.
+ The array is filled at the acquisition step and consumed at the
+ commit step. And it's managed in circular manner.
+
+2. cross_lock
+
+ One per lockdep_map exists. This is for keeping data of crosslocks
+ and used at the commit step.
+
+
+How crossrelease works
+----------------------
+
+It's the key of how crossrelease works, to defer necessary works to an
+appropriate point in time and perform in at once at the commit step.
+Let's take a look with examples step by step, starting from how lockdep
+works without crossrelease for typical locks.
+
+ acquire A /* Push A onto held_locks */
+ acquire B /* Push B onto held_locks and add 'A -> B' */
+ acquire C /* Push C onto held_locks and add 'B -> C' */
+ release C /* Pop C from held_locks */
+ release B /* Pop B from held_locks */
+ release A /* Pop A from held_locks */
+
+ where A, B and C are different lock classes.
+
+ NOTE: This document assumes that readers already understand how
+ lockdep works without crossrelease thus omits details. But there's
+ one thing to note. Lockdep pretends to pop a lock from held_locks
+ when releasing it. But it's subtly different from the original pop
+ operation because lockdep allows other than the top to be poped.
+
+In this case, lockdep adds 'the top of held_locks -> the lock to acquire'
+dependency every time acquiring a lock.
+
+After adding 'A -> B', a dependency graph will be:
+
+ A -> B
+
+ where A and B are different lock classes.
+
+And after adding 'B -> C', the graph will be:
+
+ A -> B -> C
+
+ where A, B and C are different lock classes.
+
+Let's performs commit step even for typical locks to add dependencies.
+Of course, commit step is not necessary for them, however, it would work
+well because this is a more general way.
+
+ acquire A
+ /*
+ * Queue A into hist_locks
+ *
+ * In hist_locks: A
+ * In graph: Empty
+ */
+
+ acquire B
+ /*
+ * Queue B into hist_locks
+ *
+ * In hist_locks: A, B
+ * In graph: Empty
+ */
+
+ acquire C
+ /*
+ * Queue C into hist_locks
+ *
+ * In hist_locks: A, B, C
+ * In graph: Empty
+ */
+
+ commit C
+ /*
+ * Add 'C -> ?'
+ * Answer the following to decide '?'
+ * What has been queued since acquire C: Nothing
+ *
+ * In hist_locks: A, B, C
+ * In graph: Empty
+ */
+
+ release C
+
+ commit B
+ /*
+ * Add 'B -> ?'
+ * Answer the following to decide '?'
+ * What has been queued since acquire B: C
+ *
+ * In hist_locks: A, B, C
+ * In graph: 'B -> C'
+ */
+
+ release B
+
+ commit A
+ /*
+ * Add 'A -> ?'
+ * Answer the following to decide '?'
+ * What has been queued since acquire A: B, C
+ *
+ * In hist_locks: A, B, C
+ * In graph: 'B -> C', 'A -> B', 'A -> C'
+ */
+
+ release A
+
+ where A, B and C are different lock classes.
+
+In this case, dependencies are added at the commit step as described.
+
+After commits for A, B and C, the graph will be:
+
+ A -> B -> C
+
+ where A, B and C are different lock classes.
+
+ NOTE: A dependency 'A -> C' is optimized out.
+
+We can see the former graph built without commit step is same as the
+latter graph built using commit steps. Of course the former way leads to
+earlier finish for building the graph, which means we can detect a
+deadlock or its possibility sooner. So the former way would be prefered
+when possible. But we cannot avoid using the latter way for crosslocks.
+
+Let's look at how commit steps work for crosslocks. In this case, the
+commit step is performed only on crosslock AX as real. And it assumes
+that the AX release context is different from the AX acquire context.
+
+ BX RELEASE CONTEXT BX ACQUIRE CONTEXT
+ ------------------ ------------------
+ acquire A
+ /*
+ * Push A onto held_locks
+ * Queue A into hist_locks
+ *
+ * In held_locks: A
+ * In hist_locks: A
+ * In graph: Empty
+ */
+
+ acquire BX
+ /*
+ * Add 'the top of held_locks -> BX'
+ *
+ * In held_locks: A
+ * In hist_locks: A
+ * In graph: 'A -> BX'
+ */
+
+ ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~
+ It must be guaranteed that the following operations are seen after
+ acquiring BX globally. It can be done by things like barrier.
+ ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~
+
+ acquire C
+ /*
+ * Push C onto held_locks
+ * Queue C into hist_locks
+ *
+ * In held_locks: C
+ * In hist_locks: C
+ * In graph: 'A -> BX'
+ */
+
+ release C
+ /*
+ * Pop C from held_locks
+ *
+ * In held_locks: Empty
+ * In hist_locks: C
+ * In graph: 'A -> BX'
+ */
+ acquire D
+ /*
+ * Push D onto held_locks
+ * Queue D into hist_locks
+ * Add 'the top of held_locks -> D'
+ *
+ * In held_locks: A, D
+ * In hist_locks: A, D
+ * In graph: 'A -> BX', 'A -> D'
+ */
+ acquire E
+ /*
+ * Push E onto held_locks
+ * Queue E into hist_locks
+ *
+ * In held_locks: E
+ * In hist_locks: C, E
+ * In graph: 'A -> BX', 'A -> D'
+ */
+
+ release E
+ /*
+ * Pop E from held_locks
+ *
+ * In held_locks: Empty
+ * In hist_locks: D, E
+ * In graph: 'A -> BX', 'A -> D'
+ */
+ release D
+ /*
+ * Pop D from held_locks
+ *
+ * In held_locks: A
+ * In hist_locks: A, D
+ * In graph: 'A -> BX', 'A -> D'
+ */
+ commit BX
+ /*
+ * Add 'BX -> ?'
+ * What has been queued since acquire BX: C, E
+ *
+ * In held_locks: Empty
+ * In hist_locks: D, E
+ * In graph: 'A -> BX', 'A -> D',
+ * 'BX -> C', 'BX -> E'
+ */
+
+ release BX
+ /*
+ * In held_locks: Empty
+ * In hist_locks: D, E
+ * In graph: 'A -> BX', 'A -> D',
+ * 'BX -> C', 'BX -> E'
+ */
+ release A
+ /*
+ * Pop A from held_locks
+ *
+ * In held_locks: Empty
+ * In hist_locks: A, D
+ * In graph: 'A -> BX', 'A -> D',
+ * 'BX -> C', 'BX -> E'
+ */
+
+ where A, BX, C,..., E are different lock classes, and a suffix 'X' is
+ added on crosslocks.
+
+Crossrelease considers all acquisitions after acqiuring BX are
+candidates which might create dependencies with BX. True dependencies
+will be determined when identifying the release context of BX. Meanwhile,
+all typical locks are queued so that they can be used at the commit step.
+And then two dependencies 'BX -> C' and 'BX -> E' are added at the
+commit step when identifying the release context.
+
+The final graph will be, with crossrelease:
+
+ -> C
+ /
+ -> BX -
+ / \
+ A - -> E
+ \
+ -> D
+
+ where A, BX, C,..., E are different lock classes, and a suffix 'X' is
+ added on crosslocks.
+
+However, the final graph will be, without crossrelease:
+
+ A -> D
+
+ where A and D are different lock classes.
+
+The former graph has three more dependencies, 'A -> BX', 'BX -> C' and
+'BX -> E' giving additional opportunities to check if they cause
+deadlocks. This way lockdep can detect a deadlock or its possibility
+caused by crosslocks.
+
+CONCLUSION
+
+We checked how crossrelease works with several examples.
+
+
+=============
+Optimizations
+=============
+
+Avoid duplication
+-----------------
+
+Crossrelease feature uses a cache like what lockdep already uses for
+dependency chains, but this time it's for caching CT type dependencies.
+Once that dependency is cached, the same will never be added again.
+
+
+Lockless for hot paths
+----------------------
+
+To keep all locks for later use at the commit step, crossrelease adopts
+a local array embedded in task_struct, which makes access to the data
+lockless by forcing it to happen only within the owner context. It's
+like how lockdep handles held_locks. Lockless implmentation is important
+since typical locks are very frequently acquired and released.
+
+
+=================================================
+APPENDIX A: What lockdep does to work aggresively
+=================================================
+
+A deadlock actually occurs when all wait operations creating circular
+dependencies run at the same time. Even though they don't, a potential
+deadlock exists if the problematic dependencies exist. Thus it's
+meaningful to detect not only an actual deadlock but also its potential
+possibility. The latter is rather valuable. When a deadlock occurs
+actually, we can identify what happens in the system by some means or
+other even without lockdep. However, there's no way to detect possiblity
+without lockdep unless the whole code is parsed in head. It's terrible.
+Lockdep does the both, and crossrelease only focuses on the latter.
+
+Whether or not a deadlock actually occurs depends on several factors.
+For example, what order contexts are switched in is a factor. Assuming
+circular dependencies exist, a deadlock would occur when contexts are
+switched so that all wait operations creating the dependencies run
+simultaneously. Thus to detect a deadlock possibility even in the case
+that it has not occured yet, lockdep should consider all possible
+combinations of dependencies, trying to:
+
+1. Use a global dependency graph.
+
+ Lockdep combines all dependencies into one global graph and uses them,
+ regardless of which context generates them or what order contexts are
+ switched in. Aggregated dependencies are only considered so they are
+ prone to be circular if a problem exists.
+
+2. Check dependencies between classes instead of instances.
+
+ What actually causes a deadlock are instances of lock. However,
+ lockdep checks dependencies between classes instead of instances.
+ This way lockdep can detect a deadlock which has not happened but
+ might happen in future by others but the same class.
+
+3. Assume all acquisitions lead to waiting.
+
+ Although locks might be acquired without waiting which is essential
+ to create dependencies, lockdep assumes all acquisitions lead to
+ waiting since it might be true some time or another.
+
+CONCLUSION
+
+Lockdep detects not only an actual deadlock but also its possibility,
+and the latter is more valuable.
+
+
+==================================================
+APPENDIX B: How to avoid adding false dependencies
+==================================================
+
+Remind what a dependency is. A dependency exists if:
+
+ 1. There are two waiters waiting for each event at a given time.
+ 2. The only way to wake up each waiter is to trigger its event.
+ 3. Whether one can be woken up depends on whether the other can.
+
+For example:
+
+ acquire A
+ acquire B /* A dependency 'A -> B' exists */
+ release B
+ release A
+
+ where A and B are different lock classes.
+
+A depedency 'A -> B' exists since:
+
+ 1. A waiter for A and a waiter for B might exist when acquiring B.
+ 2. Only way to wake up each is to release what it waits for.
+ 3. Whether the waiter for A can be woken up depends on whether the
+ other can. IOW, TASK X cannot release A if it fails to acquire B.
+
+For another example:
+
+ TASK X TASK Y
+ ------ ------
+ acquire AX
+ acquire B /* A dependency 'AX -> B' exists */
+ release B
+ release AX held by Y
+
+ where AX and B are different lock classes, and a suffix 'X' is added
+ on crosslocks.
+
+Even in this case involving crosslocks, the same rule can be applied. A
+depedency 'AX -> B' exists since:
+
+ 1. A waiter for AX and a waiter for B might exist when acquiring B.
+ 2. Only way to wake up each is to release what it waits for.
+ 3. Whether the waiter for AX can be woken up depends on whether the
+ other can. IOW, TASK X cannot release AX if it fails to acquire B.
+
+Let's take a look at more complicated example:
+
+ TASK X TASK Y
+ ------ ------
+ acquire B
+ release B
+ fork Y
+ acquire AX
+ acquire C /* A dependency 'AX -> C' exists */
+ release C
+ release AX held by Y
+
+ where AX, B and C are different lock classes, and a suffix 'X' is
+ added on crosslocks.
+
+Does a dependency 'AX -> B' exist? Nope.
+
+Two waiters are essential to create a dependency. However, waiters for
+AX and B to create 'AX -> B' cannot exist at the same time in this
+example. Thus the dependency 'AX -> B' cannot be created.
+
+It would be ideal if the full set of true ones can be considered. But
+we can ensure nothing but what actually happened. Relying on what
+actually happens at runtime, we can anyway add only true ones, though
+they might be a subset of true ones. It's similar to how lockdep works
+for typical locks. There might be more true dependencies than what
+lockdep has detected in runtime. Lockdep has no choice but to rely on
+what actually happens. Crossrelease also relies on it.
+
+CONCLUSION
+
+Relying on what actually happens, lockdep can avoid adding false
+dependencies.
diff --git a/src/kernel/linux/v4.14/Documentation/locking/lockdep-design.txt b/src/kernel/linux/v4.14/Documentation/locking/lockdep-design.txt
new file mode 100644
index 0000000..9de1c15
--- /dev/null
+++ b/src/kernel/linux/v4.14/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/locking/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 that 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
+don't 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/src/kernel/linux/v4.14/Documentation/locking/lockstat.txt b/src/kernel/linux/v4.14/Documentation/locking/lockstat.txt
new file mode 100644
index 0000000..5786ad2
--- /dev/null
+++ b/src/kernel/linux/v4.14/Documentation/locking/lockstat.txt
@@ -0,0 +1,183 @@
+
+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/locking/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.
+
+Lines 09-12 show the first 4 recorded contention points (the code
+which tries to get the lock) and lines 14-17 show the first 4 recorded
+contended points (the lock holder). It is possible that the max
+con-bounces point is missing in the statistics.
+
+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/src/kernel/linux/v4.14/Documentation/locking/locktorture.txt b/src/kernel/linux/v4.14/Documentation/locking/locktorture.txt
new file mode 100644
index 0000000..a2ef3a9
--- /dev/null
+++ b/src/kernel/linux/v4.14/Documentation/locking/locktorture.txt
@@ -0,0 +1,150 @@
+Kernel Lock Torture Test Operation
+
+CONFIG_LOCK_TORTURE_TEST
+
+The CONFIG LOCK_TORTURE_TEST config option provides a kernel module
+that runs torture tests on core kernel locking primitives. The kernel
+module, 'locktorture', may be built after the fact on the running
+kernel to be tested, if desired. The tests periodically output status
+messages via printk(), which can be examined via the dmesg (perhaps
+grepping for "torture"). The test is started when the module is loaded,
+and stops when the module is unloaded. This program is based on how RCU
+is tortured, via rcutorture.
+
+This torture test consists of creating a number of kernel threads which
+acquire the lock and hold it for specific amount of time, thus simulating
+different critical region behaviors. The amount of contention on the lock
+can be simulated by either enlarging this critical region hold time and/or
+creating more kthreads.
+
+
+MODULE PARAMETERS
+
+This module has the following parameters:
+
+
+ ** Locktorture-specific **
+
+nwriters_stress Number of kernel threads that will stress exclusive lock
+ ownership (writers). The default value is twice the number
+ of online CPUs.
+
+nreaders_stress Number of kernel threads that will stress shared lock
+ ownership (readers). The default is the same amount of writer
+ locks. If the user did not specify nwriters_stress, then
+ both readers and writers be the amount of online CPUs.
+
+torture_type Type of lock to torture. By default, only spinlocks will
+ be tortured. This module can torture the following locks,
+ with string values as follows:
+
+ o "lock_busted": Simulates a buggy lock implementation.
+
+ o "spin_lock": spin_lock() and spin_unlock() pairs.
+
+ o "spin_lock_irq": spin_lock_irq() and spin_unlock_irq()
+ pairs.
+
+ o "rw_lock": read/write lock() and unlock() rwlock pairs.
+
+ o "rw_lock_irq": read/write lock_irq() and unlock_irq()
+ rwlock pairs.
+
+ o "mutex_lock": mutex_lock() and mutex_unlock() pairs.
+
+ o "rtmutex_lock": rtmutex_lock() and rtmutex_unlock()
+ pairs. Kernel must have CONFIG_RT_MUTEX=y.
+
+ o "rwsem_lock": read/write down() and up() semaphore pairs.
+
+torture_runnable Start locktorture at boot time in the case where the
+ module is built into the kernel, otherwise wait for
+ torture_runnable to be set via sysfs before starting.
+ By default it will begin once the module is loaded.
+
+
+ ** Torture-framework (RCU + locking) **
+
+shutdown_secs The number of seconds to run the test before terminating
+ the test and powering off the system. The default is
+ zero, which disables test termination and system shutdown.
+ This capability is useful for automated testing.
+
+onoff_interval The number of seconds between each attempt to execute a
+ randomly selected CPU-hotplug operation. Defaults
+ to zero, which disables CPU hotplugging. In
+ CONFIG_HOTPLUG_CPU=n kernels, locktorture will silently
+ refuse to do any CPU-hotplug operations regardless of
+ what value is specified for onoff_interval.
+
+onoff_holdoff The number of seconds to wait until starting CPU-hotplug
+ operations. This would normally only be used when
+ locktorture was built into the kernel and started
+ automatically at boot time, in which case it is useful
+ in order to avoid confusing boot-time code with CPUs
+ coming and going. This parameter is only useful if
+ CONFIG_HOTPLUG_CPU is enabled.
+
+stat_interval Number of seconds between statistics-related printk()s.
+ By default, locktorture will report stats every 60 seconds.
+ Setting the interval to zero causes the statistics to
+ be printed -only- when the module is unloaded, and this
+ is the default.
+
+stutter The length of time to run the test before pausing for this
+ same period of time. Defaults to "stutter=5", so as
+ to run and pause for (roughly) five-second intervals.
+ Specifying "stutter=0" causes the test to run continuously
+ without pausing, which is the old default behavior.
+
+shuffle_interval The number of seconds to keep the test threads affinitied
+ to a particular subset of the CPUs, defaults to 3 seconds.
+ Used in conjunction with test_no_idle_hz.
+
+verbose Enable verbose debugging printing, via printk(). Enabled
+ by default. This extra information is mostly related to
+ high-level errors and reports from the main 'torture'
+ framework.
+
+
+STATISTICS
+
+Statistics are printed in the following format:
+
+spin_lock-torture: Writes: Total: 93746064 Max/Min: 0/0 Fail: 0
+ (A) (B) (C) (D) (E)
+
+(A): Lock type that is being tortured -- torture_type parameter.
+
+(B): Number of writer lock acquisitions. If dealing with a read/write primitive
+ a second "Reads" statistics line is printed.
+
+(C): Number of times the lock was acquired.
+
+(D): Min and max number of times threads failed to acquire the lock.
+
+(E): true/false values if there were errors acquiring the lock. This should
+ -only- be positive if there is a bug in the locking primitive's
+ implementation. Otherwise a lock should never fail (i.e., spin_lock()).
+ Of course, the same applies for (C), above. A dummy example of this is
+ the "lock_busted" type.
+
+USAGE
+
+The following script may be used to torture locks:
+
+ #!/bin/sh
+
+ modprobe locktorture
+ sleep 3600
+ rmmod locktorture
+ dmesg | grep torture:
+
+The output can be manually inspected for the error flag of "!!!".
+One could of course create a more elaborate script that automatically
+checked for such errors. The "rmmod" command forces a "SUCCESS",
+"FAILURE", or "RCU_HOTPLUG" indication to be printk()ed. The first
+two are self-explanatory, while the last indicates that while there
+were no locking failures, CPU-hotplug problems were detected.
+
+Also see: Documentation/RCU/torture.txt
diff --git a/src/kernel/linux/v4.14/Documentation/locking/mutex-design.txt b/src/kernel/linux/v4.14/Documentation/locking/mutex-design.txt
new file mode 100644
index 0000000..60c482d
--- /dev/null
+++ b/src/kernel/linux/v4.14/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 tied, along with rwsems,
+for the largest lock in the kernel. 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/src/kernel/linux/v4.14/Documentation/locking/rt-mutex-design.txt b/src/kernel/linux/v4.14/Documentation/locking/rt-mutex-design.txt
new file mode 100644
index 0000000..6c6e8c2
--- /dev/null
+++ b/src/kernel/linux/v4.14/Documentation/locking/rt-mutex-design.txt
@@ -0,0 +1,559 @@
+#
+# 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 rbtree node structures to
+ place the task in the waiters rbtree of a mutex as well as the
+ pi_waiters rbtree 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-+
+
+If process G has the highest priority in the chain, then all the tasks up
+the chain (A and B in this example), must have their priorities increased
+to that of G.
+
+Mutex Waiters Tree
+-----------------
+
+Every mutex keeps track of all the waiters that are blocked on itself. The
+mutex has a rbtree to store these waiters by priority. This tree is protected
+by a spin lock that is located in the struct of the mutex. This lock is called
+wait_lock.
+
+
+Task PI Tree
+------------
+
+To keep track of the PI chains, each process has its own PI rbtree. This is
+a tree of all top waiters of the mutexes that are owned by the process.
+Note that this tree 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 tree 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 tree.
+
+This tree is stored in the task structure of a process as a rbtree called
+pi_waiters. It 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 two byte alignment (and if this is
+not true, the rtmutex.c code will be broken!), this allows for the least
+significant bit to be used as a flag. Bit 0 is used as the "Has Waiters"
+flag. It's set whenever there are waiters on a mutex.
+
+See Documentation/locking/rt-mutex.txt for further details.
+
+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_waiters 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
+and rt_mutex_setprio. rt_mutex_setprio is only used in rt_mutex_adjust_prio.
+
+rt_mutex_adjust_prio examines the priority of the task, and the highest
+priority process that is waiting any of mutexes owned by the task. Since
+the pi_waiters of a task holds an order by priority of all the top waiters
+of all the mutexes that the task owns, we simply need to compare the top
+pi waiter to its own normal/deadline priority and take the higher one.
+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.
+
+(Note: For the "prio" field in task_struct, the lower the number, the
+ higher the priority. A "prio" of 5 is of higher priority than a
+ "prio" of 10.)
+
+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_waiters
+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, 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), a pointer to the mutex on which the task
+is blocked, and a top_task as the top waiter of the mutex.
+
+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 rbtree nodes of the task's waiter have not been updated
+with the new priorities, and this task may not be in the proper locations
+in the pi_waiters and waiters trees that the task is blocked on. This function
+solves all that.
+
+The main operation of this function is summarized by Thomas Gleixner in
+rtmutex.c. See the 'Chain walk basics and protection scope' comment for further
+details.
+
+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, 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 waiters tree of the mutex, and if need be, the pi_waiters
+tree 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. By setting this flag
+now, the current owner of the mutex being contended for 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 current owner to synchronize with this code.
+
+The lock is taken if the following are true:
+ 1) The lock has no owner
+ 2) The current task is the highest priority against all other
+ waiters of the lock
+
+If the task succeeds to acquire the lock, then the task is set as the
+owner of the lock, and if the lock still has waiters, the top_waiter
+(highest priority task waiting on the lock) is added to this task's
+pi_waiters tree.
+
+If the lock is not taken by try_to_take_rt_mutex(), then the
+task_blocks_on_rt_mutex() function is called. This will add the task to
+the lock's waiter tree and propagate the pi chain of the lock as well
+as the lock's owner's pi_waiters tree. This is described in the next
+section.
+
+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 rbtree node of waiter 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 task waiter tree. 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_waiters of the owner,
+and add the current process to that tree. Since the pi_waiter 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_waiters 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 task can then wake up for a couple of reasons:
+ 1) The previous lock owner released the lock, and the task now is top_waiter
+ 2) we received a signal or timeout
+
+In both cases, the task will try again to acquire the lock. If it
+does, then it will take itself off the waiters tree and set itself back
+to the TASK_RUNNING state.
+
+In first case, if the lock was acquired by another task before this task
+could get the lock, then it will go back to sleep and wait to be woken again.
+
+The second case is only applicable for tasks that are grabbing a mutex
+that can wake up before getting the lock, either due to a signal or
+a timeout (i.e. rt_mutex_timed_futex_lock()). When woken, it will try to
+take the lock again, if it succeeds, then the task will return with the
+lock held, otherwise it will return with -EINTR if the task was woken
+by a signal, or -ETIMEDOUT if it timed out.
+
+
+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.
+
+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 waiters tree of the mutex
+as well as the pi_waiters tree of the current owner. The "Has Waiters" bit is
+marked to prevent lower priority tasks from stealing the lock.
+
+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>
+Updated: Alex Shi <alex.shi@linaro.org> - 7/6/2017
+
+Original Reviewers: Ingo Molnar, Thomas Gleixner, Thomas Duetsch, and
+ Randy Dunlap
+Update (7/6/2017) Reviewers: Steven Rostedt and Sebastian Siewior
+
+Updates
+-------
+
+This document was originally written for 2.6.17-rc3-mm1
+was updated on 4.12
diff --git a/src/kernel/linux/v4.14/Documentation/locking/rt-mutex.txt b/src/kernel/linux/v4.14/Documentation/locking/rt-mutex.txt
new file mode 100644
index 0000000..35793e0
--- /dev/null
+++ b/src/kernel/linux/v4.14/Documentation/locking/rt-mutex.txt
@@ -0,0 +1,73 @@
+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 tree 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 tree. This tree 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 "pi_waiters".
+
+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:
+
+lock->owner holds the task_struct pointer of the owner. Bit 0 is used to
+keep track of the "lock has waiters" state.
+
+ owner bit0
+ NULL 0 lock is free (fast acquire possible)
+ NULL 1 lock is free and has waiters and the top waiter
+ is going to take the lock*
+ taskpointer 0 lock is held (fast release possible)
+ taskpointer 1 lock is held and has waiters**
+
+The fast atomic compare exchange based acquire and release is only
+possible when bit 0 of lock->owner is 0.
+
+(*) It also can be a transitional state when grabbing the lock
+with ->wait_lock is held. To prevent any fast path cmpxchg to the lock,
+we need to set the bit0 before looking at the lock, and the owner may be
+NULL in this small time, hence this can be a transitional state.
+
+(**) There is a small time when bit 0 is set but there are no
+waiters. This can happen when grabbing the lock in the slow path.
+To prevent a cmpxchg of the owner releasing the lock, we need to
+set this bit before looking at the lock.
+
+BTW, there is still technically a "Pending Owner", it's just not called
+that anymore. The pending owner happens to be the top_waiter of a lock
+that has no owner and has been woken up to grab the lock.
diff --git a/src/kernel/linux/v4.14/Documentation/locking/spinlocks.txt b/src/kernel/linux/v4.14/Documentation/locking/spinlocks.txt
new file mode 100644
index 0000000..ff35e40
--- /dev/null
+++ b/src/kernel/linux/v4.14/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/src/kernel/linux/v4.14/Documentation/locking/ww-mutex-design.txt b/src/kernel/linux/v4.14/Documentation/locking/ww-mutex-design.txt
new file mode 100644
index 0000000..34c3a1b
--- /dev/null
+++ b/src/kernel/linux/v4.14/Documentation/locking/ww-mutex-design.txt
@@ -0,0 +1,348 @@
+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.
+
+ We maintain the following invariants for the wait list:
+ (1) Waiters with an acquire context are sorted by stamp order; waiters
+ without an acquire context are interspersed in FIFO order.
+ (2) Among waiters with contexts, only the first one can have other locks
+ acquired already (ctx->acquired > 0). Note that this waiter may come
+ after other waiters without contexts in the list.
+
+ In general, not much contention is expected. The locks are typically used to
+ serialize access to resources for devices.
+
+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.