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-rw-r--r--Documentation/locking/futex-requeue-pi.rst132
-rw-r--r--Documentation/locking/hwspinlock.rst485
-rw-r--r--Documentation/locking/index.rst33
-rw-r--r--Documentation/locking/lockdep-design.rst663
-rw-r--r--Documentation/locking/lockstat.rst204
-rw-r--r--Documentation/locking/locktorture.rst169
-rw-r--r--Documentation/locking/locktypes.rst534
-rw-r--r--Documentation/locking/mutex-design.rst152
-rw-r--r--Documentation/locking/percpu-rw-semaphore.rst28
-rw-r--r--Documentation/locking/pi-futex.rst122
-rw-r--r--Documentation/locking/preempt-locking.rst144
-rw-r--r--Documentation/locking/robust-futex-ABI.rst184
-rw-r--r--Documentation/locking/robust-futexes.rst221
-rw-r--r--Documentation/locking/rt-mutex-design.rst574
-rw-r--r--Documentation/locking/rt-mutex.rst77
-rw-r--r--Documentation/locking/seqlock.rst239
-rw-r--r--Documentation/locking/spinlocks.rst165
-rw-r--r--Documentation/locking/ww-mutex-design.rst393
18 files changed, 4519 insertions, 0 deletions
diff --git a/Documentation/locking/futex-requeue-pi.rst b/Documentation/locking/futex-requeue-pi.rst
new file mode 100644
index 000000000..dd4ecf452
--- /dev/null
+++ b/Documentation/locking/futex-requeue-pi.rst
@@ -0,0 +1,132 @@
+================
+Futex Requeue PI
+================
+
+Requeueing of tasks from a non-PI futex to a PI futex requires
+special handling in order to ensure the underlying rt_mutex is never
+left without an owner if it has waiters; doing so would break the PI
+boosting logic [see rt-mutex-design.rst] For the purposes of
+brevity, this action will be referred to as "requeue_pi" throughout
+this document. Priority inheritance is abbreviated throughout as
+"PI".
+
+Motivation
+----------
+
+Without requeue_pi, the glibc implementation of
+pthread_cond_broadcast() must resort to waking all the tasks waiting
+on a pthread_condvar and letting them try to sort out which task
+gets to run first in classic thundering-herd formation. An ideal
+implementation would wake the highest-priority waiter, and leave the
+rest to the natural wakeup inherent in unlocking the mutex
+associated with the condvar.
+
+Consider the simplified glibc calls::
+
+ /* caller must lock mutex */
+ pthread_cond_wait(cond, mutex)
+ {
+ lock(cond->__data.__lock);
+ unlock(mutex);
+ do {
+ unlock(cond->__data.__lock);
+ futex_wait(cond->__data.__futex);
+ lock(cond->__data.__lock);
+ } while(...)
+ unlock(cond->__data.__lock);
+ lock(mutex);
+ }
+
+ pthread_cond_broadcast(cond)
+ {
+ lock(cond->__data.__lock);
+ unlock(cond->__data.__lock);
+ futex_requeue(cond->data.__futex, cond->mutex);
+ }
+
+Once pthread_cond_broadcast() requeues the tasks, the cond->mutex
+has waiters. Note that pthread_cond_wait() attempts to lock the
+mutex only after it has returned to user space. This will leave the
+underlying rt_mutex with waiters, and no owner, breaking the
+previously mentioned PI-boosting algorithms.
+
+In order to support PI-aware pthread_condvar's, the kernel needs to
+be able to requeue tasks to PI futexes. This support implies that
+upon a successful futex_wait system call, the caller would return to
+user space already holding the PI futex. The glibc implementation
+would be modified as follows::
+
+
+ /* caller must lock mutex */
+ pthread_cond_wait_pi(cond, mutex)
+ {
+ lock(cond->__data.__lock);
+ unlock(mutex);
+ do {
+ unlock(cond->__data.__lock);
+ futex_wait_requeue_pi(cond->__data.__futex);
+ lock(cond->__data.__lock);
+ } while(...)
+ unlock(cond->__data.__lock);
+ /* the kernel acquired the mutex for us */
+ }
+
+ pthread_cond_broadcast_pi(cond)
+ {
+ lock(cond->__data.__lock);
+ unlock(cond->__data.__lock);
+ futex_requeue_pi(cond->data.__futex, cond->mutex);
+ }
+
+The actual glibc implementation will likely test for PI and make the
+necessary changes inside the existing calls rather than creating new
+calls for the PI cases. Similar changes are needed for
+pthread_cond_timedwait() and pthread_cond_signal().
+
+Implementation
+--------------
+
+In order to ensure the rt_mutex has an owner if it has waiters, it
+is necessary for both the requeue code, as well as the waiting code,
+to be able to acquire the rt_mutex before returning to user space.
+The requeue code cannot simply wake the waiter and leave it to
+acquire the rt_mutex as it would open a race window between the
+requeue call returning to user space and the waiter waking and
+starting to run. This is especially true in the uncontended case.
+
+The solution involves two new rt_mutex helper routines,
+rt_mutex_start_proxy_lock() and rt_mutex_finish_proxy_lock(), which
+allow the requeue code to acquire an uncontended rt_mutex on behalf
+of the waiter and to enqueue the waiter on a contended rt_mutex.
+Two new system calls provide the kernel<->user interface to
+requeue_pi: FUTEX_WAIT_REQUEUE_PI and FUTEX_CMP_REQUEUE_PI.
+
+FUTEX_WAIT_REQUEUE_PI is called by the waiter (pthread_cond_wait()
+and pthread_cond_timedwait()) to block on the initial futex and wait
+to be requeued to a PI-aware futex. The implementation is the
+result of a high-speed collision between futex_wait() and
+futex_lock_pi(), with some extra logic to check for the additional
+wake-up scenarios.
+
+FUTEX_CMP_REQUEUE_PI is called by the waker
+(pthread_cond_broadcast() and pthread_cond_signal()) to requeue and
+possibly wake the waiting tasks. Internally, this system call is
+still handled by futex_requeue (by passing requeue_pi=1). Before
+requeueing, futex_requeue() attempts to acquire the requeue target
+PI futex on behalf of the top waiter. If it can, this waiter is
+woken. futex_requeue() then proceeds to requeue the remaining
+nr_wake+nr_requeue tasks to the PI futex, calling
+rt_mutex_start_proxy_lock() prior to each requeue to prepare the
+task as a waiter on the underlying rt_mutex. It is possible that
+the lock can be acquired at this stage as well, if so, the next
+waiter is woken to finish the acquisition of the lock.
+
+FUTEX_CMP_REQUEUE_PI accepts nr_wake and nr_requeue as arguments, but
+their sum is all that really matters. futex_requeue() will wake or
+requeue up to nr_wake + nr_requeue tasks. It will wake only as many
+tasks as it can acquire the lock for, which in the majority of cases
+should be 0 as good programming practice dictates that the caller of
+either pthread_cond_broadcast() or pthread_cond_signal() acquire the
+mutex prior to making the call. FUTEX_CMP_REQUEUE_PI requires that
+nr_wake=1. nr_requeue should be INT_MAX for broadcast and 0 for
+signal.
diff --git a/Documentation/locking/hwspinlock.rst b/Documentation/locking/hwspinlock.rst
new file mode 100644
index 000000000..6f03713b7
--- /dev/null
+++ b/Documentation/locking/hwspinlock.rst
@@ -0,0 +1,485 @@
+===========================
+Hardware Spinlock Framework
+===========================
+
+Introduction
+============
+
+Hardware spinlock modules provide hardware assistance for synchronization
+and mutual exclusion between heterogeneous processors and those not operating
+under a single, shared operating system.
+
+For example, OMAP4 has dual Cortex-A9, dual Cortex-M3 and a C64x+ DSP,
+each of which is running a different Operating System (the master, A9,
+is usually running Linux and the slave processors, the M3 and the DSP,
+are running some flavor of RTOS).
+
+A generic hwspinlock framework allows platform-independent drivers to use
+the hwspinlock device in order to access data structures that are shared
+between remote processors, that otherwise have no alternative mechanism
+to accomplish synchronization and mutual exclusion operations.
+
+This is necessary, for example, for Inter-processor communications:
+on OMAP4, cpu-intensive multimedia tasks are offloaded by the host to the
+remote M3 and/or C64x+ slave processors (by an IPC subsystem called Syslink).
+
+To achieve fast message-based communications, a minimal kernel support
+is needed to deliver messages arriving from a remote processor to the
+appropriate user process.
+
+This communication is based on simple data structures that is shared between
+the remote processors, and access to it is synchronized using the hwspinlock
+module (remote processor directly places new messages in this shared data
+structure).
+
+A common hwspinlock interface makes it possible to have generic, platform-
+independent, drivers.
+
+User API
+========
+
+::
+
+ struct hwspinlock *hwspin_lock_request(void);
+
+Dynamically assign an hwspinlock and return its address, or NULL
+in case an unused hwspinlock isn't available. Users of this
+API will usually want to communicate the lock's id to the remote core
+before it can be used to achieve synchronization.
+
+Should be called from a process context (might sleep).
+
+::
+
+ struct hwspinlock *hwspin_lock_request_specific(unsigned int id);
+
+Assign a specific hwspinlock id and return its address, or NULL
+if that hwspinlock is already in use. Usually board code will
+be calling this function in order to reserve specific hwspinlock
+ids for predefined purposes.
+
+Should be called from a process context (might sleep).
+
+::
+
+ int of_hwspin_lock_get_id(struct device_node *np, int index);
+
+Retrieve the global lock id for an OF phandle-based specific lock.
+This function provides a means for DT users of a hwspinlock module
+to get the global lock id of a specific hwspinlock, so that it can
+be requested using the normal hwspin_lock_request_specific() API.
+
+The function returns a lock id number on success, -EPROBE_DEFER if
+the hwspinlock device is not yet registered with the core, or other
+error values.
+
+Should be called from a process context (might sleep).
+
+::
+
+ int hwspin_lock_free(struct hwspinlock *hwlock);
+
+Free a previously-assigned hwspinlock; returns 0 on success, or an
+appropriate error code on failure (e.g. -EINVAL if the hwspinlock
+is already free).
+
+Should be called from a process context (might sleep).
+
+::
+
+ int hwspin_lock_timeout(struct hwspinlock *hwlock, unsigned int timeout);
+
+Lock a previously-assigned hwspinlock with a timeout limit (specified in
+msecs). If the hwspinlock is already taken, the function will busy loop
+waiting for it to be released, but give up when the timeout elapses.
+Upon a successful return from this function, preemption is disabled so
+the caller must not sleep, and is advised to release the hwspinlock as
+soon as possible, in order to minimize remote cores polling on the
+hardware interconnect.
+
+Returns 0 when successful and an appropriate error code otherwise (most
+notably -ETIMEDOUT if the hwspinlock is still busy after timeout msecs).
+The function will never sleep.
+
+::
+
+ int hwspin_lock_timeout_irq(struct hwspinlock *hwlock, unsigned int timeout);
+
+Lock a previously-assigned hwspinlock with a timeout limit (specified in
+msecs). If the hwspinlock is already taken, the function will busy loop
+waiting for it to be released, but give up when the timeout elapses.
+Upon a successful return from this function, preemption and the local
+interrupts are disabled, so the caller must not sleep, and is advised to
+release the hwspinlock as soon as possible.
+
+Returns 0 when successful and an appropriate error code otherwise (most
+notably -ETIMEDOUT if the hwspinlock is still busy after timeout msecs).
+The function will never sleep.
+
+::
+
+ int hwspin_lock_timeout_irqsave(struct hwspinlock *hwlock, unsigned int to,
+ unsigned long *flags);
+
+Lock a previously-assigned hwspinlock with a timeout limit (specified in
+msecs). If the hwspinlock is already taken, the function will busy loop
+waiting for it to be released, but give up when the timeout elapses.
+Upon a successful return from this function, preemption is disabled,
+local interrupts are disabled and their previous state is saved at the
+given flags placeholder. The caller must not sleep, and is advised to
+release the hwspinlock as soon as possible.
+
+Returns 0 when successful and an appropriate error code otherwise (most
+notably -ETIMEDOUT if the hwspinlock is still busy after timeout msecs).
+
+The function will never sleep.
+
+::
+
+ int hwspin_lock_timeout_raw(struct hwspinlock *hwlock, unsigned int timeout);
+
+Lock a previously-assigned hwspinlock with a timeout limit (specified in
+msecs). If the hwspinlock is already taken, the function will busy loop
+waiting for it to be released, but give up when the timeout elapses.
+
+Caution: User must protect the routine of getting hardware lock with mutex
+or spinlock to avoid dead-lock, that will let user can do some time-consuming
+or sleepable operations under the hardware lock.
+
+Returns 0 when successful and an appropriate error code otherwise (most
+notably -ETIMEDOUT if the hwspinlock is still busy after timeout msecs).
+
+The function will never sleep.
+
+::
+
+ int hwspin_lock_timeout_in_atomic(struct hwspinlock *hwlock, unsigned int to);
+
+Lock a previously-assigned hwspinlock with a timeout limit (specified in
+msecs). If the hwspinlock is already taken, the function will busy loop
+waiting for it to be released, but give up when the timeout elapses.
+
+This function shall be called only from an atomic context and the timeout
+value shall not exceed a few msecs.
+
+Returns 0 when successful and an appropriate error code otherwise (most
+notably -ETIMEDOUT if the hwspinlock is still busy after timeout msecs).
+
+The function will never sleep.
+
+::
+
+ int hwspin_trylock(struct hwspinlock *hwlock);
+
+
+Attempt to lock a previously-assigned hwspinlock, but immediately fail if
+it is already taken.
+
+Upon a successful return from this function, preemption is disabled so
+caller must not sleep, and is advised to release the hwspinlock as soon as
+possible, in order to minimize remote cores polling on the hardware
+interconnect.
+
+Returns 0 on success and an appropriate error code otherwise (most
+notably -EBUSY if the hwspinlock was already taken).
+The function will never sleep.
+
+::
+
+ int hwspin_trylock_irq(struct hwspinlock *hwlock);
+
+
+Attempt to lock a previously-assigned hwspinlock, but immediately fail if
+it is already taken.
+
+Upon a successful return from this function, preemption and the local
+interrupts are disabled so caller must not sleep, and is advised to
+release the hwspinlock as soon as possible.
+
+Returns 0 on success and an appropriate error code otherwise (most
+notably -EBUSY if the hwspinlock was already taken).
+
+The function will never sleep.
+
+::
+
+ int hwspin_trylock_irqsave(struct hwspinlock *hwlock, unsigned long *flags);
+
+Attempt to lock a previously-assigned hwspinlock, but immediately fail if
+it is already taken.
+
+Upon a successful return from this function, preemption is disabled,
+the local interrupts are disabled and their previous state is saved
+at the given flags placeholder. The caller must not sleep, and is advised
+to release the hwspinlock as soon as possible.
+
+Returns 0 on success and an appropriate error code otherwise (most
+notably -EBUSY if the hwspinlock was already taken).
+The function will never sleep.
+
+::
+
+ int hwspin_trylock_raw(struct hwspinlock *hwlock);
+
+Attempt to lock a previously-assigned hwspinlock, but immediately fail if
+it is already taken.
+
+Caution: User must protect the routine of getting hardware lock with mutex
+or spinlock to avoid dead-lock, that will let user can do some time-consuming
+or sleepable operations under the hardware lock.
+
+Returns 0 on success and an appropriate error code otherwise (most
+notably -EBUSY if the hwspinlock was already taken).
+The function will never sleep.
+
+::
+
+ int hwspin_trylock_in_atomic(struct hwspinlock *hwlock);
+
+Attempt to lock a previously-assigned hwspinlock, but immediately fail if
+it is already taken.
+
+This function shall be called only from an atomic context.
+
+Returns 0 on success and an appropriate error code otherwise (most
+notably -EBUSY if the hwspinlock was already taken).
+The function will never sleep.
+
+::
+
+ void hwspin_unlock(struct hwspinlock *hwlock);
+
+Unlock a previously-locked hwspinlock. Always succeed, and can be called
+from any context (the function never sleeps).
+
+.. note::
+
+ code should **never** unlock an hwspinlock which is already unlocked
+ (there is no protection against this).
+
+::
+
+ void hwspin_unlock_irq(struct hwspinlock *hwlock);
+
+Unlock a previously-locked hwspinlock and enable local interrupts.
+The caller should **never** unlock an hwspinlock which is already unlocked.
+
+Doing so is considered a bug (there is no protection against this).
+Upon a successful return from this function, preemption and local
+interrupts are enabled. This function will never sleep.
+
+::
+
+ void
+ hwspin_unlock_irqrestore(struct hwspinlock *hwlock, unsigned long *flags);
+
+Unlock a previously-locked hwspinlock.
+
+The caller should **never** unlock an hwspinlock which is already unlocked.
+Doing so is considered a bug (there is no protection against this).
+Upon a successful return from this function, preemption is reenabled,
+and the state of the local interrupts is restored to the state saved at
+the given flags. This function will never sleep.
+
+::
+
+ void hwspin_unlock_raw(struct hwspinlock *hwlock);
+
+Unlock a previously-locked hwspinlock.
+
+The caller should **never** unlock an hwspinlock which is already unlocked.
+Doing so is considered a bug (there is no protection against this).
+This function will never sleep.
+
+::
+
+ void hwspin_unlock_in_atomic(struct hwspinlock *hwlock);
+
+Unlock a previously-locked hwspinlock.
+
+The caller should **never** unlock an hwspinlock which is already unlocked.
+Doing so is considered a bug (there is no protection against this).
+This function will never sleep.
+
+::
+
+ int hwspin_lock_get_id(struct hwspinlock *hwlock);
+
+Retrieve id number of a given hwspinlock. This is needed when an
+hwspinlock is dynamically assigned: before it can be used to achieve
+mutual exclusion with a remote cpu, the id number should be communicated
+to the remote task with which we want to synchronize.
+
+Returns the hwspinlock id number, or -EINVAL if hwlock is null.
+
+Typical usage
+=============
+
+::
+
+ #include <linux/hwspinlock.h>
+ #include <linux/err.h>
+
+ int hwspinlock_example1(void)
+ {
+ struct hwspinlock *hwlock;
+ int ret;
+
+ /* dynamically assign a hwspinlock */
+ hwlock = hwspin_lock_request();
+ if (!hwlock)
+ ...
+
+ id = hwspin_lock_get_id(hwlock);
+ /* probably need to communicate id to a remote processor now */
+
+ /* take the lock, spin for 1 sec if it's already taken */
+ ret = hwspin_lock_timeout(hwlock, 1000);
+ if (ret)
+ ...
+
+ /*
+ * we took the lock, do our thing now, but do NOT sleep
+ */
+
+ /* release the lock */
+ hwspin_unlock(hwlock);
+
+ /* free the lock */
+ ret = hwspin_lock_free(hwlock);
+ if (ret)
+ ...
+
+ return ret;
+ }
+
+ int hwspinlock_example2(void)
+ {
+ struct hwspinlock *hwlock;
+ int ret;
+
+ /*
+ * assign a specific hwspinlock id - this should be called early
+ * by board init code.
+ */
+ hwlock = hwspin_lock_request_specific(PREDEFINED_LOCK_ID);
+ if (!hwlock)
+ ...
+
+ /* try to take it, but don't spin on it */
+ ret = hwspin_trylock(hwlock);
+ if (!ret) {
+ pr_info("lock is already taken\n");
+ return -EBUSY;
+ }
+
+ /*
+ * we took the lock, do our thing now, but do NOT sleep
+ */
+
+ /* release the lock */
+ hwspin_unlock(hwlock);
+
+ /* free the lock */
+ ret = hwspin_lock_free(hwlock);
+ if (ret)
+ ...
+
+ return ret;
+ }
+
+
+API for implementors
+====================
+
+::
+
+ int hwspin_lock_register(struct hwspinlock_device *bank, struct device *dev,
+ const struct hwspinlock_ops *ops, int base_id, int num_locks);
+
+To be called from the underlying platform-specific implementation, in
+order to register a new hwspinlock device (which is usually a bank of
+numerous locks). Should be called from a process context (this function
+might sleep).
+
+Returns 0 on success, or appropriate error code on failure.
+
+::
+
+ int hwspin_lock_unregister(struct hwspinlock_device *bank);
+
+To be called from the underlying vendor-specific implementation, in order
+to unregister an hwspinlock device (which is usually a bank of numerous
+locks).
+
+Should be called from a process context (this function might sleep).
+
+Returns the address of hwspinlock on success, or NULL on error (e.g.
+if the hwspinlock is still in use).
+
+Important structs
+=================
+
+struct hwspinlock_device is a device which usually contains a bank
+of hardware locks. It is registered by the underlying hwspinlock
+implementation using the hwspin_lock_register() API.
+
+::
+
+ /**
+ * struct hwspinlock_device - a device which usually spans numerous hwspinlocks
+ * @dev: underlying device, will be used to invoke runtime PM api
+ * @ops: platform-specific hwspinlock handlers
+ * @base_id: id index of the first lock in this device
+ * @num_locks: number of locks in this device
+ * @lock: dynamically allocated array of 'struct hwspinlock'
+ */
+ struct hwspinlock_device {
+ struct device *dev;
+ const struct hwspinlock_ops *ops;
+ int base_id;
+ int num_locks;
+ struct hwspinlock lock[0];
+ };
+
+struct hwspinlock_device contains an array of hwspinlock structs, each
+of which represents a single hardware lock::
+
+ /**
+ * struct hwspinlock - this struct represents a single hwspinlock instance
+ * @bank: the hwspinlock_device structure which owns this lock
+ * @lock: initialized and used by hwspinlock core
+ * @priv: private data, owned by the underlying platform-specific hwspinlock drv
+ */
+ struct hwspinlock {
+ struct hwspinlock_device *bank;
+ spinlock_t lock;
+ void *priv;
+ };
+
+When registering a bank of locks, the hwspinlock driver only needs to
+set the priv members of the locks. The rest of the members are set and
+initialized by the hwspinlock core itself.
+
+Implementation callbacks
+========================
+
+There are three possible callbacks defined in 'struct hwspinlock_ops'::
+
+ struct hwspinlock_ops {
+ int (*trylock)(struct hwspinlock *lock);
+ void (*unlock)(struct hwspinlock *lock);
+ void (*relax)(struct hwspinlock *lock);
+ };
+
+The first two callbacks are mandatory:
+
+The ->trylock() callback should make a single attempt to take the lock, and
+return 0 on failure and 1 on success. This callback may **not** sleep.
+
+The ->unlock() callback releases the lock. It always succeed, and it, too,
+may **not** sleep.
+
+The ->relax() callback is optional. It is called by hwspinlock core while
+spinning on a lock, and can be used by the underlying implementation to force
+a delay between two successive invocations of ->trylock(). It may **not** sleep.
diff --git a/Documentation/locking/index.rst b/Documentation/locking/index.rst
new file mode 100644
index 000000000..7003bd5ae
--- /dev/null
+++ b/Documentation/locking/index.rst
@@ -0,0 +1,33 @@
+.. SPDX-License-Identifier: GPL-2.0
+
+=======
+locking
+=======
+
+.. toctree::
+ :maxdepth: 1
+
+ locktypes
+ lockdep-design
+ lockstat
+ locktorture
+ mutex-design
+ rt-mutex-design
+ rt-mutex
+ seqlock
+ spinlocks
+ ww-mutex-design
+ preempt-locking
+ pi-futex
+ futex-requeue-pi
+ hwspinlock
+ percpu-rw-semaphore
+ robust-futexes
+ robust-futex-ABI
+
+.. only:: subproject and html
+
+ Indices
+ =======
+
+ * :ref:`genindex`
diff --git a/Documentation/locking/lockdep-design.rst b/Documentation/locking/lockdep-design.rst
new file mode 100644
index 000000000..82f36cab6
--- /dev/null
+++ b/Documentation/locking/lockdep-design.rst
@@ -0,0 +1,663 @@
+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 'usage state' of lock-classes, and it tracks
+the dependencies between different lock-classes. Lock usage indicates
+how a lock is used with regard to its IRQ contexts, while lock
+dependency can be understood as lock order, where L1 -> L2 suggests that
+a task is attempting to acquire L2 while holding L1. From lockdep's
+perspective, the two locks (L1 and L2) are not necessarily related; that
+dependency just means the order ever happened. The validator maintains a
+continuing effort to prove lock usages and dependencies are correct or
+the validator will shoot a splat if incorrect.
+
+A lock-class's behavior is constructed by its instances collectively:
+when the first instance of a lock-class is used after bootup the class
+gets registered, then all (subsequent) instances will be mapped to the
+class and hence their usages and dependecies will contribute to those of
+the class. A lock-class does not go away when a lock instance does, but
+it can be removed if the memory space of the lock class (static or
+dynamic) is reclaimed, this happens for example when a module is
+unloaded or a workqueue is destroyed.
+
+State
+-----
+
+The validator tracks lock-class usage history and divides the usage into
+(4 usages * n STATEs + 1) categories:
+
+where the 4 usages can be:
+
+- '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 the n STATEs are coded in kernel/locking/lockdep_states.h and as of
+now they include:
+
+- hardirq
+- softirq
+
+where the last 1 category is:
+
+- 'ever used' [ == !unused ]
+
+When locking rules are violated, these usage bits are presented in the
+locking error messages, inside curlies, with a total of 2 * n STATEs bits.
+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
+
+
+For a given lock, the bit positions from left to right indicate the usage
+of the lock and readlock (if exists), for each of the n STATEs listed
+above respectively, and the character displayed at each bit position
+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.
+ === ===================================================
+
+The bits are illustrated with an example::
+
+ (&sio_locks[i].lock){-.-.}, at: [<c02867fd>] mutex_lock+0x21/0x24
+ ||||
+ ||| \-> softirq disabled and not in softirq context
+ || \--> acquired in softirq context
+ | \---> hardirq disabled and not in hardirq context
+ \----> acquired in hardirq context
+
+
+For a given STATE, whether the lock is ever acquired in that STATE
+context and whether that STATE is enabled yields four possible cases as
+shown in the table below. The bit character is able to indicate which
+exact case is for the lock as of the reporting time.
+
+ +--------------+-------------+--------------+
+ | | irq enabled | irq disabled |
+ +--------------+-------------+--------------+
+ | ever in irq | '?' | '-' |
+ +--------------+-------------+--------------+
+ | never in irq | '+' | '.' |
+ +--------------+-------------+--------------+
+
+The character '-' suggests irq is disabled because if otherwise the
+charactor '?' would have been shown instead. Similar deduction can be
+applied for '+' too.
+
+Unused locks (e.g., mutexes) cannot be part of the cause of an error.
+
+
+Single-lock state rules:
+------------------------
+
+A lock is irq-safe means it was ever used in an irq context, while a lock
+is irq-unsafe means it was ever acquired with irq enabled.
+
+A softirq-unsafe lock-class is automatically hardirq-unsafe as well. The
+following states must be exclusive: only one of them is allowed to be set
+for any lock-class based on its usage::
+
+ <hardirq-safe> or <hardirq-unsafe>
+ <softirq-safe> or <softirq-unsafe>
+
+This is because if a lock can be used in irq context (irq-safe) then it
+cannot be ever acquired with irq enabled (irq-unsafe). Otherwise, a
+deadlock may happen. For example, in the scenario that after this lock
+was acquired but before released, if the context is interrupted this
+lock will be attempted to acquire twice, which creates a deadlock,
+referred to as lock recursion deadlock.
+
+The validator detects and reports lock usage that violates 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 can not be taken in inverse order::
+
+ <L1> -> <L2>
+ <L2> -> <L1>
+
+because this could lead to a deadlock - referred to as lock inversion
+deadlock - as attempts to acquire the two locks form a circle which
+could lead to the two contexts waiting for each other permanently. The
+validator will find such dependency circle in arbitrary complexity,
+i.e., there can be any other locking sequence between the acquire-lock
+operations; the validator will still find whether these locks can be
+acquired in a circular fashion.
+
+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.
+
+Annotations
+-----------
+
+Two constructs can be used to annotate and check where and if certain locks
+must be held: lockdep_assert_held*(&lock) and lockdep_*pin_lock(&lock).
+
+As the name suggests, lockdep_assert_held* family of macros assert that a
+particular lock is held at a certain time (and generate a WARN() otherwise).
+This annotation is largely used all over the kernel, e.g. kernel/sched/
+core.c::
+
+ void update_rq_clock(struct rq *rq)
+ {
+ s64 delta;
+
+ lockdep_assert_held(&rq->lock);
+ [...]
+ }
+
+where holding rq->lock is required to safely update a rq's clock.
+
+The other family of macros is lockdep_*pin_lock(), which is admittedly only
+used for rq->lock ATM. Despite their limited adoption these annotations
+generate a WARN() if the lock of interest is "accidentally" unlocked. This turns
+out to be especially helpful to debug code with callbacks, where an upper
+layer assumes a lock remains taken, but a lower layer thinks it can maybe drop
+and reacquire the lock ("unwittingly" introducing races). lockdep_pin_lock()
+returns a 'struct pin_cookie' that is then used by lockdep_unpin_lock() to check
+that nobody tampered with the lock, e.g. kernel/sched/sched.h::
+
+ static inline void rq_pin_lock(struct rq *rq, struct rq_flags *rf)
+ {
+ rf->cookie = lockdep_pin_lock(&rq->lock);
+ [...]
+ }
+
+ static inline void rq_unpin_lock(struct rq *rq, struct rq_flags *rf)
+ {
+ [...]
+ lockdep_unpin_lock(&rq->lock, rf->cookie);
+ }
+
+While comments about locking requirements might provide useful information,
+the runtime checks performed by annotations are invaluable when debugging
+locking problems and they carry the same level of details when inspecting
+code. Always prefer annotations when in doubt!
+
+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. [1]_
+
+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).
+
+.. [1]
+
+ 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.
+
+Recursive read locks:
+---------------------
+The whole of the rest document tries to prove a certain type of cycle is equivalent
+to deadlock possibility.
+
+There are three types of lockers: writers (i.e. exclusive lockers, like
+spin_lock() or write_lock()), non-recursive readers (i.e. shared lockers, like
+down_read()) and recursive readers (recursive shared lockers, like rcu_read_lock()).
+And we use the following notations of those lockers in the rest of the document:
+
+ W or E: stands for writers (exclusive lockers).
+ r: stands for non-recursive readers.
+ R: stands for recursive readers.
+ S: stands for all readers (non-recursive + recursive), as both are shared lockers.
+ N: stands for writers and non-recursive readers, as both are not recursive.
+
+Obviously, N is "r or W" and S is "r or R".
+
+Recursive readers, as their name indicates, are the lockers allowed to acquire
+even inside the critical section of another reader of the same lock instance,
+in other words, allowing nested read-side critical sections of one lock instance.
+
+While non-recursive readers will cause a self deadlock if trying to acquire inside
+the critical section of another reader of the same lock instance.
+
+The difference between recursive readers and non-recursive readers is because:
+recursive readers get blocked only by a write lock *holder*, while non-recursive
+readers could get blocked by a write lock *waiter*. Considering the follow
+example::
+
+ TASK A: TASK B:
+
+ read_lock(X);
+ write_lock(X);
+ read_lock_2(X);
+
+Task A gets the reader (no matter whether recursive or non-recursive) on X via
+read_lock() first. And when task B tries to acquire writer on X, it will block
+and become a waiter for writer on X. Now if read_lock_2() is recursive readers,
+task A will make progress, because writer waiters don't block recursive readers,
+and there is no deadlock. However, if read_lock_2() is non-recursive readers,
+it will get blocked by writer waiter B, and cause a self deadlock.
+
+Block conditions on readers/writers of the same lock instance:
+--------------------------------------------------------------
+There are simply four block conditions:
+
+1. Writers block other writers.
+2. Readers block writers.
+3. Writers block both recursive readers and non-recursive readers.
+4. And readers (recursive or not) don't block other recursive readers but
+ may block non-recursive readers (because of the potential co-existing
+ writer waiters)
+
+Block condition matrix, Y means the row blocks the column, and N means otherwise.
+
+ +---+---+---+---+
+ | | W | r | R |
+ +---+---+---+---+
+ | W | Y | Y | Y |
+ +---+---+---+---+
+ | r | Y | Y | N |
+ +---+---+---+---+
+ | R | Y | Y | N |
+ +---+---+---+---+
+
+ (W: writers, r: non-recursive readers, R: recursive readers)
+
+
+acquired recursively. Unlike non-recursive read locks, recursive read locks
+only get blocked by current write lock *holders* other than write lock
+*waiters*, for example::
+
+ TASK A: TASK B:
+
+ read_lock(X);
+
+ write_lock(X);
+
+ read_lock(X);
+
+is not a deadlock for recursive read locks, as while the task B is waiting for
+the lock X, the second read_lock() doesn't need to wait because it's a recursive
+read lock. However if the read_lock() is non-recursive read lock, then the above
+case is a deadlock, because even if the write_lock() in TASK B cannot get the
+lock, but it can block the second read_lock() in TASK A.
+
+Note that a lock can be a write lock (exclusive lock), a non-recursive read
+lock (non-recursive shared lock) or a recursive read lock (recursive shared
+lock), depending on the lock operations used to acquire it (more specifically,
+the value of the 'read' parameter for lock_acquire()). In other words, a single
+lock instance has three types of acquisition depending on the acquisition
+functions: exclusive, non-recursive read, and recursive read.
+
+To be concise, we call that write locks and non-recursive read locks as
+"non-recursive" locks and recursive read locks as "recursive" locks.
+
+Recursive locks don't block each other, while non-recursive locks do (this is
+even true for two non-recursive read locks). A non-recursive lock can block the
+corresponding recursive lock, and vice versa.
+
+A deadlock case with recursive locks involved is as follow::
+
+ TASK A: TASK B:
+
+ read_lock(X);
+ read_lock(Y);
+ write_lock(Y);
+ write_lock(X);
+
+Task A is waiting for task B to read_unlock() Y and task B is waiting for task
+A to read_unlock() X.
+
+Dependency types and strong dependency paths:
+---------------------------------------------
+Lock dependencies record the orders of the acquisitions of a pair of locks, and
+because there are 3 types for lockers, there are, in theory, 9 types of lock
+dependencies, but we can show that 4 types of lock dependencies are enough for
+deadlock detection.
+
+For each lock dependency::
+
+ L1 -> L2
+
+, which means lockdep has seen L1 held before L2 held in the same context at runtime.
+And in deadlock detection, we care whether we could get blocked on L2 with L1 held,
+IOW, whether there is a locker L3 that L1 blocks L3 and L2 gets blocked by L3. So
+we only care about 1) what L1 blocks and 2) what blocks L2. As a result, we can combine
+recursive readers and non-recursive readers for L1 (as they block the same types) and
+we can combine writers and non-recursive readers for L2 (as they get blocked by the
+same types).
+
+With the above combination for simplification, there are 4 types of dependency edges
+in the lockdep graph:
+
+1) -(ER)->:
+ exclusive writer to recursive reader dependency, "X -(ER)-> Y" means
+ X -> Y and X is a writer and Y is a recursive reader.
+
+2) -(EN)->:
+ exclusive writer to non-recursive locker dependency, "X -(EN)-> Y" means
+ X -> Y and X is a writer and Y is either a writer or non-recursive reader.
+
+3) -(SR)->:
+ shared reader to recursive reader dependency, "X -(SR)-> Y" means
+ X -> Y and X is a reader (recursive or not) and Y is a recursive reader.
+
+4) -(SN)->:
+ shared reader to non-recursive locker dependency, "X -(SN)-> Y" means
+ X -> Y and X is a reader (recursive or not) and Y is either a writer or
+ non-recursive reader.
+
+Note that given two locks, they may have multiple dependencies between them,
+for example::
+
+ TASK A:
+
+ read_lock(X);
+ write_lock(Y);
+ ...
+
+ TASK B:
+
+ write_lock(X);
+ write_lock(Y);
+
+, we have both X -(SN)-> Y and X -(EN)-> Y in the dependency graph.
+
+We use -(xN)-> to represent edges that are either -(EN)-> or -(SN)->, the
+similar for -(Ex)->, -(xR)-> and -(Sx)->
+
+A "path" is a series of conjunct dependency edges in the graph. And we define a
+"strong" path, which indicates the strong dependency throughout each dependency
+in the path, as the path that doesn't have two conjunct edges (dependencies) as
+-(xR)-> and -(Sx)->. In other words, a "strong" path is a path from a lock
+walking to another through the lock dependencies, and if X -> Y -> Z is in the
+path (where X, Y, Z are locks), and the walk from X to Y is through a -(SR)-> or
+-(ER)-> dependency, the walk from Y to Z must not be through a -(SN)-> or
+-(SR)-> dependency.
+
+We will see why the path is called "strong" in next section.
+
+Recursive Read Deadlock Detection:
+----------------------------------
+
+We now prove two things:
+
+Lemma 1:
+
+If there is a closed strong path (i.e. a strong circle), then there is a
+combination of locking sequences that causes deadlock. I.e. a strong circle is
+sufficient for deadlock detection.
+
+Lemma 2:
+
+If there is no closed strong path (i.e. strong circle), then there is no
+combination of locking sequences that could cause deadlock. I.e. strong
+circles are necessary for deadlock detection.
+
+With these two Lemmas, we can easily say a closed strong path is both sufficient
+and necessary for deadlocks, therefore a closed strong path is equivalent to
+deadlock possibility. As a closed strong path stands for a dependency chain that
+could cause deadlocks, so we call it "strong", considering there are dependency
+circles that won't cause deadlocks.
+
+Proof for sufficiency (Lemma 1):
+
+Let's say we have a strong circle::
+
+ L1 -> L2 ... -> Ln -> L1
+
+, which means we have dependencies::
+
+ L1 -> L2
+ L2 -> L3
+ ...
+ Ln-1 -> Ln
+ Ln -> L1
+
+We now can construct a combination of locking sequences that cause deadlock:
+
+Firstly let's make one CPU/task get the L1 in L1 -> L2, and then another get
+the L2 in L2 -> L3, and so on. After this, all of the Lx in Lx -> Lx+1 are
+held by different CPU/tasks.
+
+And then because we have L1 -> L2, so the holder of L1 is going to acquire L2
+in L1 -> L2, however since L2 is already held by another CPU/task, plus L1 ->
+L2 and L2 -> L3 are not -(xR)-> and -(Sx)-> (the definition of strong), which
+means either L2 in L1 -> L2 is a non-recursive locker (blocked by anyone) or
+the L2 in L2 -> L3, is writer (blocking anyone), therefore the holder of L1
+cannot get L2, it has to wait L2's holder to release.
+
+Moreover, we can have a similar conclusion for L2's holder: it has to wait L3's
+holder to release, and so on. We now can prove that Lx's holder has to wait for
+Lx+1's holder to release, and note that Ln+1 is L1, so we have a circular
+waiting scenario and nobody can get progress, therefore a deadlock.
+
+Proof for necessary (Lemma 2):
+
+Lemma 2 is equivalent to: If there is a deadlock scenario, then there must be a
+strong circle in the dependency graph.
+
+According to Wikipedia[1], if there is a deadlock, then there must be a circular
+waiting scenario, means there are N CPU/tasks, where CPU/task P1 is waiting for
+a lock held by P2, and P2 is waiting for a lock held by P3, ... and Pn is waiting
+for a lock held by P1. Let's name the lock Px is waiting as Lx, so since P1 is waiting
+for L1 and holding Ln, so we will have Ln -> L1 in the dependency graph. Similarly,
+we have L1 -> L2, L2 -> L3, ..., Ln-1 -> Ln in the dependency graph, which means we
+have a circle::
+
+ Ln -> L1 -> L2 -> ... -> Ln
+
+, and now let's prove the circle is strong:
+
+For a lock Lx, Px contributes the dependency Lx-1 -> Lx and Px+1 contributes
+the dependency Lx -> Lx+1, and since Px is waiting for Px+1 to release Lx,
+so it's impossible that Lx on Px+1 is a reader and Lx on Px is a recursive
+reader, because readers (no matter recursive or not) don't block recursive
+readers, therefore Lx-1 -> Lx and Lx -> Lx+1 cannot be a -(xR)-> -(Sx)-> pair,
+and this is true for any lock in the circle, therefore, the circle is strong.
+
+References:
+-----------
+[1]: https://en.wikipedia.org/wiki/Deadlock
+[2]: Shibu, K. (2009). Intro To Embedded Systems (1st ed.). Tata McGraw-Hill
diff --git a/Documentation/locking/lockstat.rst b/Documentation/locking/lockstat.rst
new file mode 100644
index 000000000..536eab8db
--- /dev/null
+++ b/Documentation/locking/lockstat.rst
@@ -0,0 +1,204 @@
+===============
+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.rst).
+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
+ 10 &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/Documentation/locking/locktorture.rst b/Documentation/locking/locktorture.rst
new file mode 100644
index 000000000..dfaf9fc88
--- /dev/null
+++ b/Documentation/locking/locktorture.rst
@@ -0,0 +1,169 @@
+==================================
+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:
+
+ - "lock_busted":
+ Simulates a buggy lock implementation.
+
+ - "spin_lock":
+ spin_lock() and spin_unlock() pairs.
+
+ - "spin_lock_irq":
+ spin_lock_irq() and spin_unlock_irq() pairs.
+
+ - "rw_lock":
+ read/write lock() and unlock() rwlock pairs.
+
+ - "rw_lock_irq":
+ read/write lock_irq() and unlock_irq()
+ rwlock pairs.
+
+ - "mutex_lock":
+ mutex_lock() and mutex_unlock() pairs.
+
+ - "rtmutex_lock":
+ rtmutex_lock() and rtmutex_unlock() pairs.
+ Kernel must have CONFIG_RT_MUTEX=y.
+
+ - "rwsem_lock":
+ read/write down() and up() semaphore pairs.
+
+
+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.
+
+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.
+
+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.rst
diff --git a/Documentation/locking/locktypes.rst b/Documentation/locking/locktypes.rst
new file mode 100644
index 000000000..9933faad4
--- /dev/null
+++ b/Documentation/locking/locktypes.rst
@@ -0,0 +1,534 @@
+.. SPDX-License-Identifier: GPL-2.0
+
+.. _kernel_hacking_locktypes:
+
+==========================
+Lock types and their rules
+==========================
+
+Introduction
+============
+
+The kernel provides a variety of locking primitives which can be divided
+into three categories:
+
+ - Sleeping locks
+ - CPU local locks
+ - Spinning locks
+
+This document conceptually describes these lock types and provides rules
+for their nesting, including the rules for use under PREEMPT_RT.
+
+
+Lock categories
+===============
+
+Sleeping locks
+--------------
+
+Sleeping locks can only be acquired in preemptible task context.
+
+Although implementations allow try_lock() from other contexts, it is
+necessary to carefully evaluate the safety of unlock() as well as of
+try_lock(). Furthermore, it is also necessary to evaluate the debugging
+versions of these primitives. In short, don't acquire sleeping locks from
+other contexts unless there is no other option.
+
+Sleeping lock types:
+
+ - mutex
+ - rt_mutex
+ - semaphore
+ - rw_semaphore
+ - ww_mutex
+ - percpu_rw_semaphore
+
+On PREEMPT_RT kernels, these lock types are converted to sleeping locks:
+
+ - local_lock
+ - spinlock_t
+ - rwlock_t
+
+
+CPU local locks
+---------------
+
+ - local_lock
+
+On non-PREEMPT_RT kernels, local_lock functions are wrappers around
+preemption and interrupt disabling primitives. Contrary to other locking
+mechanisms, disabling preemption or interrupts are pure CPU local
+concurrency control mechanisms and not suited for inter-CPU concurrency
+control.
+
+
+Spinning locks
+--------------
+
+ - raw_spinlock_t
+ - bit spinlocks
+
+On non-PREEMPT_RT kernels, these lock types are also spinning locks:
+
+ - spinlock_t
+ - rwlock_t
+
+Spinning locks implicitly disable preemption and the lock / unlock functions
+can have suffixes which apply further protections:
+
+ =================== ====================================================
+ _bh() Disable / enable bottom halves (soft interrupts)
+ _irq() Disable / enable interrupts
+ _irqsave/restore() Save and disable / restore interrupt disabled state
+ =================== ====================================================
+
+
+Owner semantics
+===============
+
+The aforementioned lock types except semaphores have strict owner
+semantics:
+
+ The context (task) that acquired the lock must release it.
+
+rw_semaphores have a special interface which allows non-owner release for
+readers.
+
+
+rtmutex
+=======
+
+RT-mutexes are mutexes with support for priority inheritance (PI).
+
+PI has limitations on non-PREEMPT_RT kernels due to preemption and
+interrupt disabled sections.
+
+PI clearly cannot preempt preemption-disabled or interrupt-disabled
+regions of code, even on PREEMPT_RT kernels. Instead, PREEMPT_RT kernels
+execute most such regions of code in preemptible task context, especially
+interrupt handlers and soft interrupts. This conversion allows spinlock_t
+and rwlock_t to be implemented via RT-mutexes.
+
+
+semaphore
+=========
+
+semaphore is a counting semaphore implementation.
+
+Semaphores are often used for both serialization and waiting, but new use
+cases should instead use separate serialization and wait mechanisms, such
+as mutexes and completions.
+
+semaphores and PREEMPT_RT
+----------------------------
+
+PREEMPT_RT does not change the semaphore implementation because counting
+semaphores have no concept of owners, thus preventing PREEMPT_RT from
+providing priority inheritance for semaphores. After all, an unknown
+owner cannot be boosted. As a consequence, blocking on semaphores can
+result in priority inversion.
+
+
+rw_semaphore
+============
+
+rw_semaphore is a multiple readers and single writer lock mechanism.
+
+On non-PREEMPT_RT kernels the implementation is fair, thus preventing
+writer starvation.
+
+rw_semaphore complies by default with the strict owner semantics, but there
+exist special-purpose interfaces that allow non-owner release for readers.
+These interfaces work independent of the kernel configuration.
+
+rw_semaphore and PREEMPT_RT
+---------------------------
+
+PREEMPT_RT kernels map rw_semaphore to a separate rt_mutex-based
+implementation, thus changing the fairness:
+
+ Because an rw_semaphore writer cannot grant its priority to multiple
+ readers, a preempted low-priority reader will continue holding its lock,
+ thus starving even high-priority writers. In contrast, because readers
+ can grant their priority to a writer, a preempted low-priority writer will
+ have its priority boosted until it releases the lock, thus preventing that
+ writer from starving readers.
+
+
+local_lock
+==========
+
+local_lock provides a named scope to critical sections which are protected
+by disabling preemption or interrupts.
+
+On non-PREEMPT_RT kernels local_lock operations map to the preemption and
+interrupt disabling and enabling primitives:
+
+ =============================== ======================
+ local_lock(&llock) preempt_disable()
+ local_unlock(&llock) preempt_enable()
+ local_lock_irq(&llock) local_irq_disable()
+ local_unlock_irq(&llock) local_irq_enable()
+ local_lock_irqsave(&llock) local_irq_save()
+ local_unlock_irqrestore(&llock) local_irq_restore()
+ =============================== ======================
+
+The named scope of local_lock has two advantages over the regular
+primitives:
+
+ - The lock name allows static analysis and is also a clear documentation
+ of the protection scope while the regular primitives are scopeless and
+ opaque.
+
+ - If lockdep is enabled the local_lock gains a lockmap which allows to
+ validate the correctness of the protection. This can detect cases where
+ e.g. a function using preempt_disable() as protection mechanism is
+ invoked from interrupt or soft-interrupt context. Aside of that
+ lockdep_assert_held(&llock) works as with any other locking primitive.
+
+local_lock and PREEMPT_RT
+-------------------------
+
+PREEMPT_RT kernels map local_lock to a per-CPU spinlock_t, thus changing
+semantics:
+
+ - All spinlock_t changes also apply to local_lock.
+
+local_lock usage
+----------------
+
+local_lock should be used in situations where disabling preemption or
+interrupts is the appropriate form of concurrency control to protect
+per-CPU data structures on a non PREEMPT_RT kernel.
+
+local_lock is not suitable to protect against preemption or interrupts on a
+PREEMPT_RT kernel due to the PREEMPT_RT specific spinlock_t semantics.
+
+
+raw_spinlock_t and spinlock_t
+=============================
+
+raw_spinlock_t
+--------------
+
+raw_spinlock_t is a strict spinning lock implementation in all kernels,
+including PREEMPT_RT kernels. Use raw_spinlock_t only in real critical
+core code, low-level interrupt handling and places where disabling
+preemption or interrupts is required, for example, to safely access
+hardware state. raw_spinlock_t can sometimes also be used when the
+critical section is tiny, thus avoiding RT-mutex overhead.
+
+spinlock_t
+----------
+
+The semantics of spinlock_t change with the state of PREEMPT_RT.
+
+On a non-PREEMPT_RT kernel spinlock_t is mapped to raw_spinlock_t and has
+exactly the same semantics.
+
+spinlock_t and PREEMPT_RT
+-------------------------
+
+On a PREEMPT_RT kernel spinlock_t is mapped to a separate implementation
+based on rt_mutex which changes the semantics:
+
+ - Preemption is not disabled.
+
+ - The hard interrupt related suffixes for spin_lock / spin_unlock
+ operations (_irq, _irqsave / _irqrestore) do not affect the CPU's
+ interrupt disabled state.
+
+ - The soft interrupt related suffix (_bh()) still disables softirq
+ handlers.
+
+ Non-PREEMPT_RT kernels disable preemption to get this effect.
+
+ PREEMPT_RT kernels use a per-CPU lock for serialization which keeps
+ preemption enabled. The lock disables softirq handlers and also
+ prevents reentrancy due to task preemption.
+
+PREEMPT_RT kernels preserve all other spinlock_t semantics:
+
+ - Tasks holding a spinlock_t do not migrate. Non-PREEMPT_RT kernels
+ avoid migration by disabling preemption. PREEMPT_RT kernels instead
+ disable migration, which ensures that pointers to per-CPU variables
+ remain valid even if the task is preempted.
+
+ - Task state is preserved across spinlock acquisition, ensuring that the
+ task-state rules apply to all kernel configurations. Non-PREEMPT_RT
+ kernels leave task state untouched. However, PREEMPT_RT must change
+ task state if the task blocks during acquisition. Therefore, it saves
+ the current task state before blocking and the corresponding lock wakeup
+ restores it, as shown below::
+
+ task->state = TASK_INTERRUPTIBLE
+ lock()
+ block()
+ task->saved_state = task->state
+ task->state = TASK_UNINTERRUPTIBLE
+ schedule()
+ lock wakeup
+ task->state = task->saved_state
+
+ Other types of wakeups would normally unconditionally set the task state
+ to RUNNING, but that does not work here because the task must remain
+ blocked until the lock becomes available. Therefore, when a non-lock
+ wakeup attempts to awaken a task blocked waiting for a spinlock, it
+ instead sets the saved state to RUNNING. Then, when the lock
+ acquisition completes, the lock wakeup sets the task state to the saved
+ state, in this case setting it to RUNNING::
+
+ task->state = TASK_INTERRUPTIBLE
+ lock()
+ block()
+ task->saved_state = task->state
+ task->state = TASK_UNINTERRUPTIBLE
+ schedule()
+ non lock wakeup
+ task->saved_state = TASK_RUNNING
+
+ lock wakeup
+ task->state = task->saved_state
+
+ This ensures that the real wakeup cannot be lost.
+
+
+rwlock_t
+========
+
+rwlock_t is a multiple readers and single writer lock mechanism.
+
+Non-PREEMPT_RT kernels implement rwlock_t as a spinning lock and the
+suffix rules of spinlock_t apply accordingly. The implementation is fair,
+thus preventing writer starvation.
+
+rwlock_t and PREEMPT_RT
+-----------------------
+
+PREEMPT_RT kernels map rwlock_t to a separate rt_mutex-based
+implementation, thus changing semantics:
+
+ - All the spinlock_t changes also apply to rwlock_t.
+
+ - Because an rwlock_t writer cannot grant its priority to multiple
+ readers, a preempted low-priority reader will continue holding its lock,
+ thus starving even high-priority writers. In contrast, because readers
+ can grant their priority to a writer, a preempted low-priority writer
+ will have its priority boosted until it releases the lock, thus
+ preventing that writer from starving readers.
+
+
+PREEMPT_RT caveats
+==================
+
+local_lock on RT
+----------------
+
+The mapping of local_lock to spinlock_t on PREEMPT_RT kernels has a few
+implications. For example, on a non-PREEMPT_RT kernel the following code
+sequence works as expected::
+
+ local_lock_irq(&local_lock);
+ raw_spin_lock(&lock);
+
+and is fully equivalent to::
+
+ raw_spin_lock_irq(&lock);
+
+On a PREEMPT_RT kernel this code sequence breaks because local_lock_irq()
+is mapped to a per-CPU spinlock_t which neither disables interrupts nor
+preemption. The following code sequence works perfectly correct on both
+PREEMPT_RT and non-PREEMPT_RT kernels::
+
+ local_lock_irq(&local_lock);
+ spin_lock(&lock);
+
+Another caveat with local locks is that each local_lock has a specific
+protection scope. So the following substitution is wrong::
+
+ func1()
+ {
+ local_irq_save(flags); -> local_lock_irqsave(&local_lock_1, flags);
+ func3();
+ local_irq_restore(flags); -> local_unlock_irqrestore(&local_lock_1, flags);
+ }
+
+ func2()
+ {
+ local_irq_save(flags); -> local_lock_irqsave(&local_lock_2, flags);
+ func3();
+ local_irq_restore(flags); -> local_unlock_irqrestore(&local_lock_2, flags);
+ }
+
+ func3()
+ {
+ lockdep_assert_irqs_disabled();
+ access_protected_data();
+ }
+
+On a non-PREEMPT_RT kernel this works correctly, but on a PREEMPT_RT kernel
+local_lock_1 and local_lock_2 are distinct and cannot serialize the callers
+of func3(). Also the lockdep assert will trigger on a PREEMPT_RT kernel
+because local_lock_irqsave() does not disable interrupts due to the
+PREEMPT_RT-specific semantics of spinlock_t. The correct substitution is::
+
+ func1()
+ {
+ local_irq_save(flags); -> local_lock_irqsave(&local_lock, flags);
+ func3();
+ local_irq_restore(flags); -> local_unlock_irqrestore(&local_lock, flags);
+ }
+
+ func2()
+ {
+ local_irq_save(flags); -> local_lock_irqsave(&local_lock, flags);
+ func3();
+ local_irq_restore(flags); -> local_unlock_irqrestore(&local_lock, flags);
+ }
+
+ func3()
+ {
+ lockdep_assert_held(&local_lock);
+ access_protected_data();
+ }
+
+
+spinlock_t and rwlock_t
+-----------------------
+
+The changes in spinlock_t and rwlock_t semantics on PREEMPT_RT kernels
+have a few implications. For example, on a non-PREEMPT_RT kernel the
+following code sequence works as expected::
+
+ local_irq_disable();
+ spin_lock(&lock);
+
+and is fully equivalent to::
+
+ spin_lock_irq(&lock);
+
+Same applies to rwlock_t and the _irqsave() suffix variants.
+
+On PREEMPT_RT kernel this code sequence breaks because RT-mutex requires a
+fully preemptible context. Instead, use spin_lock_irq() or
+spin_lock_irqsave() and their unlock counterparts. In cases where the
+interrupt disabling and locking must remain separate, PREEMPT_RT offers a
+local_lock mechanism. Acquiring the local_lock pins the task to a CPU,
+allowing things like per-CPU interrupt disabled locks to be acquired.
+However, this approach should be used only where absolutely necessary.
+
+A typical scenario is protection of per-CPU variables in thread context::
+
+ struct foo *p = get_cpu_ptr(&var1);
+
+ spin_lock(&p->lock);
+ p->count += this_cpu_read(var2);
+
+This is correct code on a non-PREEMPT_RT kernel, but on a PREEMPT_RT kernel
+this breaks. The PREEMPT_RT-specific change of spinlock_t semantics does
+not allow to acquire p->lock because get_cpu_ptr() implicitly disables
+preemption. The following substitution works on both kernels::
+
+ struct foo *p;
+
+ migrate_disable();
+ p = this_cpu_ptr(&var1);
+ spin_lock(&p->lock);
+ p->count += this_cpu_read(var2);
+
+migrate_disable() ensures that the task is pinned on the current CPU which
+in turn guarantees that the per-CPU access to var1 and var2 are staying on
+the same CPU while the task remains preemptible.
+
+The migrate_disable() substitution is not valid for the following
+scenario::
+
+ func()
+ {
+ struct foo *p;
+
+ migrate_disable();
+ p = this_cpu_ptr(&var1);
+ p->val = func2();
+
+This breaks because migrate_disable() does not protect against reentrancy from
+a preempting task. A correct substitution for this case is::
+
+ func()
+ {
+ struct foo *p;
+
+ local_lock(&foo_lock);
+ p = this_cpu_ptr(&var1);
+ p->val = func2();
+
+On a non-PREEMPT_RT kernel this protects against reentrancy by disabling
+preemption. On a PREEMPT_RT kernel this is achieved by acquiring the
+underlying per-CPU spinlock.
+
+
+raw_spinlock_t on RT
+--------------------
+
+Acquiring a raw_spinlock_t disables preemption and possibly also
+interrupts, so the critical section must avoid acquiring a regular
+spinlock_t or rwlock_t, for example, the critical section must avoid
+allocating memory. Thus, on a non-PREEMPT_RT kernel the following code
+works perfectly::
+
+ raw_spin_lock(&lock);
+ p = kmalloc(sizeof(*p), GFP_ATOMIC);
+
+But this code fails on PREEMPT_RT kernels because the memory allocator is
+fully preemptible and therefore cannot be invoked from truly atomic
+contexts. However, it is perfectly fine to invoke the memory allocator
+while holding normal non-raw spinlocks because they do not disable
+preemption on PREEMPT_RT kernels::
+
+ spin_lock(&lock);
+ p = kmalloc(sizeof(*p), GFP_ATOMIC);
+
+
+bit spinlocks
+-------------
+
+PREEMPT_RT cannot substitute bit spinlocks because a single bit is too
+small to accommodate an RT-mutex. Therefore, the semantics of bit
+spinlocks are preserved on PREEMPT_RT kernels, so that the raw_spinlock_t
+caveats also apply to bit spinlocks.
+
+Some bit spinlocks are replaced with regular spinlock_t for PREEMPT_RT
+using conditional (#ifdef'ed) code changes at the usage site. In contrast,
+usage-site changes are not needed for the spinlock_t substitution.
+Instead, conditionals in header files and the core locking implemementation
+enable the compiler to do the substitution transparently.
+
+
+Lock type nesting rules
+=======================
+
+The most basic rules are:
+
+ - Lock types of the same lock category (sleeping, CPU local, spinning)
+ can nest arbitrarily as long as they respect the general lock ordering
+ rules to prevent deadlocks.
+
+ - Sleeping lock types cannot nest inside CPU local and spinning lock types.
+
+ - CPU local and spinning lock types can nest inside sleeping lock types.
+
+ - Spinning lock types can nest inside all lock types
+
+These constraints apply both in PREEMPT_RT and otherwise.
+
+The fact that PREEMPT_RT changes the lock category of spinlock_t and
+rwlock_t from spinning to sleeping and substitutes local_lock with a
+per-CPU spinlock_t means that they cannot be acquired while holding a raw
+spinlock. This results in the following nesting ordering:
+
+ 1) Sleeping locks
+ 2) spinlock_t, rwlock_t, local_lock
+ 3) raw_spinlock_t and bit spinlocks
+
+Lockdep will complain if these constraints are violated, both in
+PREEMPT_RT and otherwise.
diff --git a/Documentation/locking/mutex-design.rst b/Documentation/locking/mutex-design.rst
new file mode 100644
index 000000000..78540cd7f
--- /dev/null
+++ b/Documentation/locking/mutex-design.rst
@@ -0,0 +1,152 @@
+=======================
+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] https://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 an atomic variable
+(->owner) to keep track of the lock state during its lifetime. Field owner
+actually contains `struct task_struct *` to the current lock owner and it is
+therefore NULL if not currently owned. Since task_struct pointers are aligned
+to at least L1_CACHE_BYTES, low bits (3) are used to store extra state (e.g.,
+if waiter list is non-empty). In its most basic form it also includes a
+wait-queue and a spinlock that serializes access to it. Furthermore,
+CONFIG_MUTEX_SPIN_ON_OWNER=y systems use a spinner MCS lock (->osq), 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 cmpxchg()ing the owner with
+ the current task. This only works in the uncontended case (cmpxchg() checks
+ against 0UL, so all 3 state bits above have to be 0). If the lock is
+ contended it goes to the next possible path.
+
+(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 among the largest
+locks in the kernel. E.g: on x86-64 it is 32 bytes, where 'struct semaphore'
+is 24 bytes and rw_semaphore is 40 bytes. Larger structure sizes mean more CPU
+cache and memory footprint.
+
+When to use mutexes
+-------------------
+
+Unless the strict semantics of mutexes are unsuitable and/or the critical
+region prevents the lock from being shared, always prefer them to any other
+locking primitive.
diff --git a/Documentation/locking/percpu-rw-semaphore.rst b/Documentation/locking/percpu-rw-semaphore.rst
new file mode 100644
index 000000000..247de6410
--- /dev/null
+++ b/Documentation/locking/percpu-rw-semaphore.rst
@@ -0,0 +1,28 @@
+====================
+Percpu rw semaphores
+====================
+
+Percpu rw semaphores is a new read-write semaphore design that is
+optimized for locking for reading.
+
+The problem with traditional read-write semaphores is that when multiple
+cores take the lock for reading, the cache line containing the semaphore
+is bouncing between L1 caches of the cores, causing performance
+degradation.
+
+Locking for reading is very fast, it uses RCU and it avoids any atomic
+instruction in the lock and unlock path. On the other hand, locking for
+writing is very expensive, it calls synchronize_rcu() that can take
+hundreds of milliseconds.
+
+The lock is declared with "struct percpu_rw_semaphore" type.
+The lock is initialized percpu_init_rwsem, it returns 0 on success and
+-ENOMEM on allocation failure.
+The lock must be freed with percpu_free_rwsem to avoid memory leak.
+
+The lock is locked for read with percpu_down_read, percpu_up_read and
+for write with percpu_down_write, percpu_up_write.
+
+The idea of using RCU for optimized rw-lock was introduced by
+Eric Dumazet <eric.dumazet@gmail.com>.
+The code was written by Mikulas Patocka <mpatocka@redhat.com>
diff --git a/Documentation/locking/pi-futex.rst b/Documentation/locking/pi-futex.rst
new file mode 100644
index 000000000..c33ba2bef
--- /dev/null
+++ b/Documentation/locking/pi-futex.rst
@@ -0,0 +1,122 @@
+======================
+Lightweight PI-futexes
+======================
+
+We are calling them lightweight for 3 reasons:
+
+ - in the user-space fastpath a PI-enabled futex involves no kernel work
+ (or any other PI complexity) at all. No registration, no extra kernel
+ calls - just pure fast atomic ops in userspace.
+
+ - even in the slowpath, the system call and scheduling pattern is very
+ similar to normal futexes.
+
+ - the in-kernel PI implementation is streamlined around the mutex
+ abstraction, with strict rules that keep the implementation
+ relatively simple: only a single owner may own a lock (i.e. no
+ read-write lock support), only the owner may unlock a lock, no
+ recursive locking, etc.
+
+Priority Inheritance - why?
+---------------------------
+
+The short reply: user-space PI helps achieving/improving determinism for
+user-space applications. In the best-case, it can help achieve
+determinism and well-bound latencies. Even in the worst-case, PI will
+improve the statistical distribution of locking related application
+delays.
+
+The longer reply
+----------------
+
+Firstly, sharing locks between multiple tasks is a common programming
+technique that often cannot be replaced with lockless algorithms. As we
+can see it in the kernel [which is a quite complex program in itself],
+lockless structures are rather the exception than the norm - the current
+ratio of lockless vs. locky code for shared data structures is somewhere
+between 1:10 and 1:100. Lockless is hard, and the complexity of lockless
+algorithms often endangers to ability to do robust reviews of said code.
+I.e. critical RT apps often choose lock structures to protect critical
+data structures, instead of lockless algorithms. Furthermore, there are
+cases (like shared hardware, or other resource limits) where lockless
+access is mathematically impossible.
+
+Media players (such as Jack) are an example of reasonable application
+design with multiple tasks (with multiple priority levels) sharing
+short-held locks: for example, a highprio audio playback thread is
+combined with medium-prio construct-audio-data threads and low-prio
+display-colory-stuff threads. Add video and decoding to the mix and
+we've got even more priority levels.
+
+So once we accept that synchronization objects (locks) are an
+unavoidable fact of life, and once we accept that multi-task userspace
+apps have a very fair expectation of being able to use locks, we've got
+to think about how to offer the option of a deterministic locking
+implementation to user-space.
+
+Most of the technical counter-arguments against doing priority
+inheritance only apply to kernel-space locks. But user-space locks are
+different, there we cannot disable interrupts or make the task
+non-preemptible in a critical section, so the 'use spinlocks' argument
+does not apply (user-space spinlocks have the same priority inversion
+problems as other user-space locking constructs). Fact is, pretty much
+the only technique that currently enables good determinism for userspace
+locks (such as futex-based pthread mutexes) is priority inheritance:
+
+Currently (without PI), if a high-prio and a low-prio task shares a lock
+[this is a quite common scenario for most non-trivial RT applications],
+even if all critical sections are coded carefully to be deterministic
+(i.e. all critical sections are short in duration and only execute a
+limited number of instructions), the kernel cannot guarantee any
+deterministic execution of the high-prio task: any medium-priority task
+could preempt the low-prio task while it holds the shared lock and
+executes the critical section, and could delay it indefinitely.
+
+Implementation
+--------------
+
+As mentioned before, the userspace fastpath of PI-enabled pthread
+mutexes involves no kernel work at all - they behave quite similarly to
+normal futex-based locks: a 0 value means unlocked, and a value==TID
+means locked. (This is the same method as used by list-based robust
+futexes.) Userspace uses atomic ops to lock/unlock these mutexes without
+entering the kernel.
+
+To handle the slowpath, we have added two new futex ops:
+
+ - FUTEX_LOCK_PI
+ - FUTEX_UNLOCK_PI
+
+If the lock-acquire fastpath fails, [i.e. an atomic transition from 0 to
+TID fails], then FUTEX_LOCK_PI is called. The kernel does all the
+remaining work: if there is no futex-queue attached to the futex address
+yet then the code looks up the task that owns the futex [it has put its
+own TID into the futex value], and attaches a 'PI state' structure to
+the futex-queue. The pi_state includes an rt-mutex, which is a PI-aware,
+kernel-based synchronization object. The 'other' task is made the owner
+of the rt-mutex, and the FUTEX_WAITERS bit is atomically set in the
+futex value. Then this task tries to lock the rt-mutex, on which it
+blocks. Once it returns, it has the mutex acquired, and it sets the
+futex value to its own TID and returns. Userspace has no other work to
+perform - it now owns the lock, and futex value contains
+FUTEX_WAITERS|TID.
+
+If the unlock side fastpath succeeds, [i.e. userspace manages to do a
+TID -> 0 atomic transition of the futex value], then no kernel work is
+triggered.
+
+If the unlock fastpath fails (because the FUTEX_WAITERS bit is set),
+then FUTEX_UNLOCK_PI is called, and the kernel unlocks the futex on the
+behalf of userspace - and it also unlocks the attached
+pi_state->rt_mutex and thus wakes up any potential waiters.
+
+Note that under this approach, contrary to previous PI-futex approaches,
+there is no prior 'registration' of a PI-futex. [which is not quite
+possible anyway, due to existing ABI properties of pthread mutexes.]
+
+Also, under this scheme, 'robustness' and 'PI' are two orthogonal
+properties of futexes, and all four combinations are possible: futex,
+robust-futex, PI-futex, robust+PI-futex.
+
+More details about priority inheritance can be found in
+Documentation/locking/rt-mutex.rst.
diff --git a/Documentation/locking/preempt-locking.rst b/Documentation/locking/preempt-locking.rst
new file mode 100644
index 000000000..dce336134
--- /dev/null
+++ b/Documentation/locking/preempt-locking.rst
@@ -0,0 +1,144 @@
+===========================================================================
+Proper Locking Under a Preemptible Kernel: Keeping Kernel Code Preempt-Safe
+===========================================================================
+
+:Author: Robert Love <rml@tech9.net>
+
+
+Introduction
+============
+
+
+A preemptible kernel creates new locking issues. The issues are the same as
+those under SMP: concurrency and reentrancy. Thankfully, the Linux preemptible
+kernel model leverages existing SMP locking mechanisms. Thus, the kernel
+requires explicit additional locking for very few additional situations.
+
+This document is for all kernel hackers. Developing code in the kernel
+requires protecting these situations.
+
+
+RULE #1: Per-CPU data structures need explicit protection
+^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^
+
+
+Two similar problems arise. An example code snippet::
+
+ struct this_needs_locking tux[NR_CPUS];
+ tux[smp_processor_id()] = some_value;
+ /* task is preempted here... */
+ something = tux[smp_processor_id()];
+
+First, since the data is per-CPU, it may not have explicit SMP locking, but
+require it otherwise. Second, when a preempted task is finally rescheduled,
+the previous value of smp_processor_id may not equal the current. You must
+protect these situations by disabling preemption around them.
+
+You can also use put_cpu() and get_cpu(), which will disable preemption.
+
+
+RULE #2: CPU state must be protected.
+^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^
+
+
+Under preemption, the state of the CPU must be protected. This is arch-
+dependent, but includes CPU structures and state not preserved over a context
+switch. For example, on x86, entering and exiting FPU mode is now a critical
+section that must occur while preemption is disabled. Think what would happen
+if the kernel is executing a floating-point instruction and is then preempted.
+Remember, the kernel does not save FPU state except for user tasks. Therefore,
+upon preemption, the FPU registers will be sold to the lowest bidder. Thus,
+preemption must be disabled around such regions.
+
+Note, some FPU functions are already explicitly preempt safe. For example,
+kernel_fpu_begin and kernel_fpu_end will disable and enable preemption.
+
+
+RULE #3: Lock acquire and release must be performed by same task
+^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^
+
+
+A lock acquired in one task must be released by the same task. This
+means you can't do oddball things like acquire a lock and go off to
+play while another task releases it. If you want to do something
+like this, acquire and release the task in the same code path and
+have the caller wait on an event by the other task.
+
+
+Solution
+========
+
+
+Data protection under preemption is achieved by disabling preemption for the
+duration of the critical region.
+
+::
+
+ preempt_enable() decrement the preempt counter
+ preempt_disable() increment the preempt counter
+ preempt_enable_no_resched() decrement, but do not immediately preempt
+ preempt_check_resched() if needed, reschedule
+ preempt_count() return the preempt counter
+
+The functions are nestable. In other words, you can call preempt_disable
+n-times in a code path, and preemption will not be reenabled until the n-th
+call to preempt_enable. The preempt statements define to nothing if
+preemption is not enabled.
+
+Note that you do not need to explicitly prevent preemption if you are holding
+any locks or interrupts are disabled, since preemption is implicitly disabled
+in those cases.
+
+But keep in mind that 'irqs disabled' is a fundamentally unsafe way of
+disabling preemption - any cond_resched() or cond_resched_lock() might trigger
+a reschedule if the preempt count is 0. A simple printk() might trigger a
+reschedule. So use this implicit preemption-disabling property only if you
+know that the affected codepath does not do any of this. Best policy is to use
+this only for small, atomic code that you wrote and which calls no complex
+functions.
+
+Example::
+
+ cpucache_t *cc; /* this is per-CPU */
+ preempt_disable();
+ cc = cc_data(searchp);
+ if (cc && cc->avail) {
+ __free_block(searchp, cc_entry(cc), cc->avail);
+ cc->avail = 0;
+ }
+ preempt_enable();
+ return 0;
+
+Notice how the preemption statements must encompass every reference of the
+critical variables. Another example::
+
+ int buf[NR_CPUS];
+ set_cpu_val(buf);
+ if (buf[smp_processor_id()] == -1) printf(KERN_INFO "wee!\n");
+ spin_lock(&buf_lock);
+ /* ... */
+
+This code is not preempt-safe, but see how easily we can fix it by simply
+moving the spin_lock up two lines.
+
+
+Preventing preemption using interrupt disabling
+===============================================
+
+
+It is possible to prevent a preemption event using local_irq_disable and
+local_irq_save. Note, when doing so, you must be very careful to not cause
+an event that would set need_resched and result in a preemption check. When
+in doubt, rely on locking or explicit preemption disabling.
+
+Note in 2.5 interrupt disabling is now only per-CPU (e.g. local).
+
+An additional concern is proper usage of local_irq_disable and local_irq_save.
+These may be used to protect from preemption, however, on exit, if preemption
+may be enabled, a test to see if preemption is required should be done. If
+these are called from the spin_lock and read/write lock macros, the right thing
+is done. They may also be called within a spin-lock protected region, however,
+if they are ever called outside of this context, a test for preemption should
+be made. Do note that calls from interrupt context or bottom half/ tasklets
+are also protected by preemption locks and so may use the versions which do
+not check preemption.
diff --git a/Documentation/locking/robust-futex-ABI.rst b/Documentation/locking/robust-futex-ABI.rst
new file mode 100644
index 000000000..f24904f1c
--- /dev/null
+++ b/Documentation/locking/robust-futex-ABI.rst
@@ -0,0 +1,184 @@
+====================
+The robust futex ABI
+====================
+
+:Author: Started by Paul Jackson <pj@sgi.com>
+
+
+Robust_futexes provide a mechanism that is used in addition to normal
+futexes, for kernel assist of cleanup of held locks on task exit.
+
+The interesting data as to what futexes a thread is holding is kept on a
+linked list in user space, where it can be updated efficiently as locks
+are taken and dropped, without kernel intervention. The only additional
+kernel intervention required for robust_futexes above and beyond what is
+required for futexes is:
+
+ 1) a one time call, per thread, to tell the kernel where its list of
+ held robust_futexes begins, and
+ 2) internal kernel code at exit, to handle any listed locks held
+ by the exiting thread.
+
+The existing normal futexes already provide a "Fast Userspace Locking"
+mechanism, which handles uncontested locking without needing a system
+call, and handles contested locking by maintaining a list of waiting
+threads in the kernel. Options on the sys_futex(2) system call support
+waiting on a particular futex, and waking up the next waiter on a
+particular futex.
+
+For robust_futexes to work, the user code (typically in a library such
+as glibc linked with the application) has to manage and place the
+necessary list elements exactly as the kernel expects them. If it fails
+to do so, then improperly listed locks will not be cleaned up on exit,
+probably causing deadlock or other such failure of the other threads
+waiting on the same locks.
+
+A thread that anticipates possibly using robust_futexes should first
+issue the system call::
+
+ asmlinkage long
+ sys_set_robust_list(struct robust_list_head __user *head, size_t len);
+
+The pointer 'head' points to a structure in the threads address space
+consisting of three words. Each word is 32 bits on 32 bit arch's, or 64
+bits on 64 bit arch's, and local byte order. Each thread should have
+its own thread private 'head'.
+
+If a thread is running in 32 bit compatibility mode on a 64 native arch
+kernel, then it can actually have two such structures - one using 32 bit
+words for 32 bit compatibility mode, and one using 64 bit words for 64
+bit native mode. The kernel, if it is a 64 bit kernel supporting 32 bit
+compatibility mode, will attempt to process both lists on each task
+exit, if the corresponding sys_set_robust_list() call has been made to
+setup that list.
+
+ The first word in the memory structure at 'head' contains a
+ pointer to a single linked list of 'lock entries', one per lock,
+ as described below. If the list is empty, the pointer will point
+ to itself, 'head'. The last 'lock entry' points back to the 'head'.
+
+ The second word, called 'offset', specifies the offset from the
+ address of the associated 'lock entry', plus or minus, of what will
+ be called the 'lock word', from that 'lock entry'. The 'lock word'
+ is always a 32 bit word, unlike the other words above. The 'lock
+ word' holds 2 flag bits in the upper 2 bits, and the thread id (TID)
+ of the thread holding the lock in the bottom 30 bits. See further
+ below for a description of the flag bits.
+
+ The third word, called 'list_op_pending', contains transient copy of
+ the address of the 'lock entry', during list insertion and removal,
+ and is needed to correctly resolve races should a thread exit while
+ in the middle of a locking or unlocking operation.
+
+Each 'lock entry' on the single linked list starting at 'head' consists
+of just a single word, pointing to the next 'lock entry', or back to
+'head' if there are no more entries. In addition, nearby to each 'lock
+entry', at an offset from the 'lock entry' specified by the 'offset'
+word, is one 'lock word'.
+
+The 'lock word' is always 32 bits, and is intended to be the same 32 bit
+lock variable used by the futex mechanism, in conjunction with
+robust_futexes. The kernel will only be able to wakeup the next thread
+waiting for a lock on a threads exit if that next thread used the futex
+mechanism to register the address of that 'lock word' with the kernel.
+
+For each futex lock currently held by a thread, if it wants this
+robust_futex support for exit cleanup of that lock, it should have one
+'lock entry' on this list, with its associated 'lock word' at the
+specified 'offset'. Should a thread die while holding any such locks,
+the kernel will walk this list, mark any such locks with a bit
+indicating their holder died, and wakeup the next thread waiting for
+that lock using the futex mechanism.
+
+When a thread has invoked the above system call to indicate it
+anticipates using robust_futexes, the kernel stores the passed in 'head'
+pointer for that task. The task may retrieve that value later on by
+using the system call::
+
+ asmlinkage long
+ sys_get_robust_list(int pid, struct robust_list_head __user **head_ptr,
+ size_t __user *len_ptr);
+
+It is anticipated that threads will use robust_futexes embedded in
+larger, user level locking structures, one per lock. The kernel
+robust_futex mechanism doesn't care what else is in that structure, so
+long as the 'offset' to the 'lock word' is the same for all
+robust_futexes used by that thread. The thread should link those locks
+it currently holds using the 'lock entry' pointers. It may also have
+other links between the locks, such as the reverse side of a double
+linked list, but that doesn't matter to the kernel.
+
+By keeping its locks linked this way, on a list starting with a 'head'
+pointer known to the kernel, the kernel can provide to a thread the
+essential service available for robust_futexes, which is to help clean
+up locks held at the time of (a perhaps unexpectedly) exit.
+
+Actual locking and unlocking, during normal operations, is handled
+entirely by user level code in the contending threads, and by the
+existing futex mechanism to wait for, and wakeup, locks. The kernels
+only essential involvement in robust_futexes is to remember where the
+list 'head' is, and to walk the list on thread exit, handling locks
+still held by the departing thread, as described below.
+
+There may exist thousands of futex lock structures in a threads shared
+memory, on various data structures, at a given point in time. Only those
+lock structures for locks currently held by that thread should be on
+that thread's robust_futex linked lock list a given time.
+
+A given futex lock structure in a user shared memory region may be held
+at different times by any of the threads with access to that region. The
+thread currently holding such a lock, if any, is marked with the threads
+TID in the lower 30 bits of the 'lock word'.
+
+When adding or removing a lock from its list of held locks, in order for
+the kernel to correctly handle lock cleanup regardless of when the task
+exits (perhaps it gets an unexpected signal 9 in the middle of
+manipulating this list), the user code must observe the following
+protocol on 'lock entry' insertion and removal:
+
+On insertion:
+
+ 1) set the 'list_op_pending' word to the address of the 'lock entry'
+ to be inserted,
+ 2) acquire the futex lock,
+ 3) add the lock entry, with its thread id (TID) in the bottom 30 bits
+ of the 'lock word', to the linked list starting at 'head', and
+ 4) clear the 'list_op_pending' word.
+
+On removal:
+
+ 1) set the 'list_op_pending' word to the address of the 'lock entry'
+ to be removed,
+ 2) remove the lock entry for this lock from the 'head' list,
+ 3) release the futex lock, and
+ 4) clear the 'lock_op_pending' word.
+
+On exit, the kernel will consider the address stored in
+'list_op_pending' and the address of each 'lock word' found by walking
+the list starting at 'head'. For each such address, if the bottom 30
+bits of the 'lock word' at offset 'offset' from that address equals the
+exiting threads TID, then the kernel will do two things:
+
+ 1) if bit 31 (0x80000000) is set in that word, then attempt a futex
+ wakeup on that address, which will waken the next thread that has
+ used to the futex mechanism to wait on that address, and
+ 2) atomically set bit 30 (0x40000000) in the 'lock word'.
+
+In the above, bit 31 was set by futex waiters on that lock to indicate
+they were waiting, and bit 30 is set by the kernel to indicate that the
+lock owner died holding the lock.
+
+The kernel exit code will silently stop scanning the list further if at
+any point:
+
+ 1) the 'head' pointer or an subsequent linked list pointer
+ is not a valid address of a user space word
+ 2) the calculated location of the 'lock word' (address plus
+ 'offset') is not the valid address of a 32 bit user space
+ word
+ 3) if the list contains more than 1 million (subject to
+ future kernel configuration changes) elements.
+
+When the kernel sees a list entry whose 'lock word' doesn't have the
+current threads TID in the lower 30 bits, it does nothing with that
+entry, and goes on to the next entry.
diff --git a/Documentation/locking/robust-futexes.rst b/Documentation/locking/robust-futexes.rst
new file mode 100644
index 000000000..6361fb01c
--- /dev/null
+++ b/Documentation/locking/robust-futexes.rst
@@ -0,0 +1,221 @@
+========================================
+A description of what robust futexes are
+========================================
+
+:Started by: Ingo Molnar <mingo@redhat.com>
+
+Background
+----------
+
+what are robust futexes? To answer that, we first need to understand
+what futexes are: normal futexes are special types of locks that in the
+noncontended case can be acquired/released from userspace without having
+to enter the kernel.
+
+A futex is in essence a user-space address, e.g. a 32-bit lock variable
+field. If userspace notices contention (the lock is already owned and
+someone else wants to grab it too) then the lock is marked with a value
+that says "there's a waiter pending", and the sys_futex(FUTEX_WAIT)
+syscall is used to wait for the other guy to release it. The kernel
+creates a 'futex queue' internally, so that it can later on match up the
+waiter with the waker - without them having to know about each other.
+When the owner thread releases the futex, it notices (via the variable
+value) that there were waiter(s) pending, and does the
+sys_futex(FUTEX_WAKE) syscall to wake them up. Once all waiters have
+taken and released the lock, the futex is again back to 'uncontended'
+state, and there's no in-kernel state associated with it. The kernel
+completely forgets that there ever was a futex at that address. This
+method makes futexes very lightweight and scalable.
+
+"Robustness" is about dealing with crashes while holding a lock: if a
+process exits prematurely while holding a pthread_mutex_t lock that is
+also shared with some other process (e.g. yum segfaults while holding a
+pthread_mutex_t, or yum is kill -9-ed), then waiters for that lock need
+to be notified that the last owner of the lock exited in some irregular
+way.
+
+To solve such types of problems, "robust mutex" userspace APIs were
+created: pthread_mutex_lock() returns an error value if the owner exits
+prematurely - and the new owner can decide whether the data protected by
+the lock can be recovered safely.
+
+There is a big conceptual problem with futex based mutexes though: it is
+the kernel that destroys the owner task (e.g. due to a SEGFAULT), but
+the kernel cannot help with the cleanup: if there is no 'futex queue'
+(and in most cases there is none, futexes being fast lightweight locks)
+then the kernel has no information to clean up after the held lock!
+Userspace has no chance to clean up after the lock either - userspace is
+the one that crashes, so it has no opportunity to clean up. Catch-22.
+
+In practice, when e.g. yum is kill -9-ed (or segfaults), a system reboot
+is needed to release that futex based lock. This is one of the leading
+bugreports against yum.
+
+To solve this problem, the traditional approach was to extend the vma
+(virtual memory area descriptor) concept to have a notion of 'pending
+robust futexes attached to this area'. This approach requires 3 new
+syscall variants to sys_futex(): FUTEX_REGISTER, FUTEX_DEREGISTER and
+FUTEX_RECOVER. At do_exit() time, all vmas are searched to see whether
+they have a robust_head set. This approach has two fundamental problems
+left:
+
+ - it has quite complex locking and race scenarios. The vma-based
+ approach had been pending for years, but they are still not completely
+ reliable.
+
+ - they have to scan _every_ vma at sys_exit() time, per thread!
+
+The second disadvantage is a real killer: pthread_exit() takes around 1
+microsecond on Linux, but with thousands (or tens of thousands) of vmas
+every pthread_exit() takes a millisecond or more, also totally
+destroying the CPU's L1 and L2 caches!
+
+This is very much noticeable even for normal process sys_exit_group()
+calls: the kernel has to do the vma scanning unconditionally! (this is
+because the kernel has no knowledge about how many robust futexes there
+are to be cleaned up, because a robust futex might have been registered
+in another task, and the futex variable might have been simply mmap()-ed
+into this process's address space).
+
+This huge overhead forced the creation of CONFIG_FUTEX_ROBUST so that
+normal kernels can turn it off, but worse than that: the overhead makes
+robust futexes impractical for any type of generic Linux distribution.
+
+So something had to be done.
+
+New approach to robust futexes
+------------------------------
+
+At the heart of this new approach there is a per-thread private list of
+robust locks that userspace is holding (maintained by glibc) - which
+userspace list is registered with the kernel via a new syscall [this
+registration happens at most once per thread lifetime]. At do_exit()
+time, the kernel checks this user-space list: are there any robust futex
+locks to be cleaned up?
+
+In the common case, at do_exit() time, there is no list registered, so
+the cost of robust futexes is just a simple current->robust_list != NULL
+comparison. If the thread has registered a list, then normally the list
+is empty. If the thread/process crashed or terminated in some incorrect
+way then the list might be non-empty: in this case the kernel carefully
+walks the list [not trusting it], and marks all locks that are owned by
+this thread with the FUTEX_OWNER_DIED bit, and wakes up one waiter (if
+any).
+
+The list is guaranteed to be private and per-thread at do_exit() time,
+so it can be accessed by the kernel in a lockless way.
+
+There is one race possible though: since adding to and removing from the
+list is done after the futex is acquired by glibc, there is a few
+instructions window for the thread (or process) to die there, leaving
+the futex hung. To protect against this possibility, userspace (glibc)
+also maintains a simple per-thread 'list_op_pending' field, to allow the
+kernel to clean up if the thread dies after acquiring the lock, but just
+before it could have added itself to the list. Glibc sets this
+list_op_pending field before it tries to acquire the futex, and clears
+it after the list-add (or list-remove) has finished.
+
+That's all that is needed - all the rest of robust-futex cleanup is done
+in userspace [just like with the previous patches].
+
+Ulrich Drepper has implemented the necessary glibc support for this new
+mechanism, which fully enables robust mutexes.
+
+Key differences of this userspace-list based approach, compared to the
+vma based method:
+
+ - it's much, much faster: at thread exit time, there's no need to loop
+ over every vma (!), which the VM-based method has to do. Only a very
+ simple 'is the list empty' op is done.
+
+ - no VM changes are needed - 'struct address_space' is left alone.
+
+ - no registration of individual locks is needed: robust mutexes don't
+ need any extra per-lock syscalls. Robust mutexes thus become a very
+ lightweight primitive - so they don't force the application designer
+ to do a hard choice between performance and robustness - robust
+ mutexes are just as fast.
+
+ - no per-lock kernel allocation happens.
+
+ - no resource limits are needed.
+
+ - no kernel-space recovery call (FUTEX_RECOVER) is needed.
+
+ - the implementation and the locking is "obvious", and there are no
+ interactions with the VM.
+
+Performance
+-----------
+
+I have benchmarked the time needed for the kernel to process a list of 1
+million (!) held locks, using the new method [on a 2GHz CPU]:
+
+ - with FUTEX_WAIT set [contended mutex]: 130 msecs
+ - without FUTEX_WAIT set [uncontended mutex]: 30 msecs
+
+I have also measured an approach where glibc does the lock notification
+[which it currently does for !pshared robust mutexes], and that took 256
+msecs - clearly slower, due to the 1 million FUTEX_WAKE syscalls
+userspace had to do.
+
+(1 million held locks are unheard of - we expect at most a handful of
+locks to be held at a time. Nevertheless it's nice to know that this
+approach scales nicely.)
+
+Implementation details
+----------------------
+
+The patch adds two new syscalls: one to register the userspace list, and
+one to query the registered list pointer::
+
+ asmlinkage long
+ sys_set_robust_list(struct robust_list_head __user *head,
+ size_t len);
+
+ asmlinkage long
+ sys_get_robust_list(int pid, struct robust_list_head __user **head_ptr,
+ size_t __user *len_ptr);
+
+List registration is very fast: the pointer is simply stored in
+current->robust_list. [Note that in the future, if robust futexes become
+widespread, we could extend sys_clone() to register a robust-list head
+for new threads, without the need of another syscall.]
+
+So there is virtually zero overhead for tasks not using robust futexes,
+and even for robust futex users, there is only one extra syscall per
+thread lifetime, and the cleanup operation, if it happens, is fast and
+straightforward. The kernel doesn't have any internal distinction between
+robust and normal futexes.
+
+If a futex is found to be held at exit time, the kernel sets the
+following bit of the futex word::
+
+ #define FUTEX_OWNER_DIED 0x40000000
+
+and wakes up the next futex waiter (if any). User-space does the rest of
+the cleanup.
+
+Otherwise, robust futexes are acquired by glibc by putting the TID into
+the futex field atomically. Waiters set the FUTEX_WAITERS bit::
+
+ #define FUTEX_WAITERS 0x80000000
+
+and the remaining bits are for the TID.
+
+Testing, architecture support
+-----------------------------
+
+I've tested the new syscalls on x86 and x86_64, and have made sure the
+parsing of the userspace list is robust [ ;-) ] even if the list is
+deliberately corrupted.
+
+i386 and x86_64 syscalls are wired up at the moment, and Ulrich has
+tested the new glibc code (on x86_64 and i386), and it works for his
+robust-mutex testcases.
+
+All other architectures should build just fine too - but they won't have
+the new syscalls yet.
+
+Architectures need to implement the new futex_atomic_cmpxchg_inatomic()
+inline function before writing up the syscalls.
diff --git a/Documentation/locking/rt-mutex-design.rst b/Documentation/locking/rt-mutex-design.rst
new file mode 100644
index 000000000..59c2a64ef
--- /dev/null
+++ b/Documentation/locking/rt-mutex-design.rst
@@ -0,0 +1,574 @@
+==============================
+RT-mutex implementation design
+==============================
+
+Copyright (c) 2006 Steven Rostedt
+
+Licensed under the GNU Free Documentation License, Version 1.2
+
+
+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/locking/rt-mutex.rst. 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.rst 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/Documentation/locking/rt-mutex.rst b/Documentation/locking/rt-mutex.rst
new file mode 100644
index 000000000..3b5097a38
--- /dev/null
+++ b/Documentation/locking/rt-mutex.rst
@@ -0,0 +1,77 @@
+==================================
+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/locking/pi-futex.rst 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 Notes
+ ============ ======= ================================================
+ 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 [1]_
+ taskpointer 0 lock is held (fast release possible)
+ taskpointer 1 lock is held and has waiters [2]_
+ ============ ======= ================================================
+
+The fast atomic compare exchange based acquire and release is only
+possible when bit 0 of lock->owner is 0.
+
+.. [1] 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.
+
+.. [2] 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/Documentation/locking/seqlock.rst b/Documentation/locking/seqlock.rst
new file mode 100644
index 000000000..bfda1a5fe
--- /dev/null
+++ b/Documentation/locking/seqlock.rst
@@ -0,0 +1,239 @@
+======================================
+Sequence counters and sequential locks
+======================================
+
+Introduction
+============
+
+Sequence counters are a reader-writer consistency mechanism with
+lockless readers (read-only retry loops), and no writer starvation. They
+are used for data that's rarely written to (e.g. system time), where the
+reader wants a consistent set of information and is willing to retry if
+that information changes.
+
+A data set is consistent when the sequence count at the beginning of the
+read side critical section is even and the same sequence count value is
+read again at the end of the critical section. The data in the set must
+be copied out inside the read side critical section. If the sequence
+count has changed between the start and the end of the critical section,
+the reader must retry.
+
+Writers increment the sequence count at the start and the end of their
+critical section. After starting the critical section the sequence count
+is odd and indicates to the readers that an update is in progress. At
+the end of the write side critical section the sequence count becomes
+even again which lets readers make progress.
+
+A sequence counter write side critical section must never be preempted
+or interrupted by read side sections. Otherwise the reader will spin for
+the entire scheduler tick due to the odd sequence count value and the
+interrupted writer. If that reader belongs to a real-time scheduling
+class, it can spin forever and the kernel will livelock.
+
+This mechanism cannot be used if the protected data contains pointers,
+as the writer can invalidate a pointer that the reader is following.
+
+
+.. _seqcount_t:
+
+Sequence counters (``seqcount_t``)
+==================================
+
+This is the raw counting mechanism, which does not protect against
+multiple writers. Write side critical sections must thus be serialized
+by an external lock.
+
+If the write serialization primitive is not implicitly disabling
+preemption, preemption must be explicitly disabled before entering the
+write side section. If the read section can be invoked from hardirq or
+softirq contexts, interrupts or bottom halves must also be respectively
+disabled before entering the write section.
+
+If it's desired to automatically handle the sequence counter
+requirements of writer serialization and non-preemptibility, use
+:ref:`seqlock_t` instead.
+
+Initialization::
+
+ /* dynamic */
+ seqcount_t foo_seqcount;
+ seqcount_init(&foo_seqcount);
+
+ /* static */
+ static seqcount_t foo_seqcount = SEQCNT_ZERO(foo_seqcount);
+
+ /* C99 struct init */
+ struct {
+ .seq = SEQCNT_ZERO(foo.seq),
+ } foo;
+
+Write path::
+
+ /* Serialized context with disabled preemption */
+
+ write_seqcount_begin(&foo_seqcount);
+
+ /* ... [[write-side critical section]] ... */
+
+ write_seqcount_end(&foo_seqcount);
+
+Read path::
+
+ do {
+ seq = read_seqcount_begin(&foo_seqcount);
+
+ /* ... [[read-side critical section]] ... */
+
+ } while (read_seqcount_retry(&foo_seqcount, seq));
+
+
+.. _seqcount_locktype_t:
+
+Sequence counters with associated locks (``seqcount_LOCKNAME_t``)
+-----------------------------------------------------------------
+
+As discussed at :ref:`seqcount_t`, sequence count write side critical
+sections must be serialized and non-preemptible. This variant of
+sequence counters associate the lock used for writer serialization at
+initialization time, which enables lockdep to validate that the write
+side critical sections are properly serialized.
+
+This lock association is a NOOP if lockdep is disabled and has neither
+storage nor runtime overhead. If lockdep is enabled, the lock pointer is
+stored in struct seqcount and lockdep's "lock is held" assertions are
+injected at the beginning of the write side critical section to validate
+that it is properly protected.
+
+For lock types which do not implicitly disable preemption, preemption
+protection is enforced in the write side function.
+
+The following sequence counters with associated locks are defined:
+
+ - ``seqcount_spinlock_t``
+ - ``seqcount_raw_spinlock_t``
+ - ``seqcount_rwlock_t``
+ - ``seqcount_mutex_t``
+ - ``seqcount_ww_mutex_t``
+
+The sequence counter read and write APIs can take either a plain
+seqcount_t or any of the seqcount_LOCKNAME_t variants above.
+
+Initialization (replace "LOCKNAME" with one of the supported locks)::
+
+ /* dynamic */
+ seqcount_LOCKNAME_t foo_seqcount;
+ seqcount_LOCKNAME_init(&foo_seqcount, &lock);
+
+ /* static */
+ static seqcount_LOCKNAME_t foo_seqcount =
+ SEQCNT_LOCKNAME_ZERO(foo_seqcount, &lock);
+
+ /* C99 struct init */
+ struct {
+ .seq = SEQCNT_LOCKNAME_ZERO(foo.seq, &lock),
+ } foo;
+
+Write path: same as in :ref:`seqcount_t`, while running from a context
+with the associated write serialization lock acquired.
+
+Read path: same as in :ref:`seqcount_t`.
+
+
+.. _seqcount_latch_t:
+
+Latch sequence counters (``seqcount_latch_t``)
+----------------------------------------------
+
+Latch sequence counters are a multiversion concurrency control mechanism
+where the embedded seqcount_t counter even/odd value is used to switch
+between two copies of protected data. This allows the sequence counter
+read path to safely interrupt its own write side critical section.
+
+Use seqcount_latch_t when the write side sections cannot be protected
+from interruption by readers. This is typically the case when the read
+side can be invoked from NMI handlers.
+
+Check `raw_write_seqcount_latch()` for more information.
+
+
+.. _seqlock_t:
+
+Sequential locks (``seqlock_t``)
+================================
+
+This contains the :ref:`seqcount_t` mechanism earlier discussed, plus an
+embedded spinlock for writer serialization and non-preemptibility.
+
+If the read side section can be invoked from hardirq or softirq context,
+use the write side function variants which disable interrupts or bottom
+halves respectively.
+
+Initialization::
+
+ /* dynamic */
+ seqlock_t foo_seqlock;
+ seqlock_init(&foo_seqlock);
+
+ /* static */
+ static DEFINE_SEQLOCK(foo_seqlock);
+
+ /* C99 struct init */
+ struct {
+ .seql = __SEQLOCK_UNLOCKED(foo.seql)
+ } foo;
+
+Write path::
+
+ write_seqlock(&foo_seqlock);
+
+ /* ... [[write-side critical section]] ... */
+
+ write_sequnlock(&foo_seqlock);
+
+Read path, three categories:
+
+1. Normal Sequence readers which never block a writer but they must
+ retry if a writer is in progress by detecting change in the sequence
+ number. Writers do not wait for a sequence reader::
+
+ do {
+ seq = read_seqbegin(&foo_seqlock);
+
+ /* ... [[read-side critical section]] ... */
+
+ } while (read_seqretry(&foo_seqlock, seq));
+
+2. Locking readers which will wait if a writer or another locking reader
+ is in progress. A locking reader in progress will also block a writer
+ from entering its critical section. This read lock is
+ exclusive. Unlike rwlock_t, only one locking reader can acquire it::
+
+ read_seqlock_excl(&foo_seqlock);
+
+ /* ... [[read-side critical section]] ... */
+
+ read_sequnlock_excl(&foo_seqlock);
+
+3. Conditional lockless reader (as in 1), or locking reader (as in 2),
+ according to a passed marker. This is used to avoid lockless readers
+ starvation (too much retry loops) in case of a sharp spike in write
+ activity. First, a lockless read is tried (even marker passed). If
+ that trial fails (odd sequence counter is returned, which is used as
+ the next iteration marker), the lockless read is transformed to a
+ full locking read and no retry loop is necessary::
+
+ /* marker; even initialization */
+ int seq = 0;
+ do {
+ read_seqbegin_or_lock(&foo_seqlock, &seq);
+
+ /* ... [[read-side critical section]] ... */
+
+ } while (need_seqretry(&foo_seqlock, seq));
+ done_seqretry(&foo_seqlock, seq);
+
+
+API documentation
+=================
+
+.. kernel-doc:: include/linux/seqlock.h
diff --git a/Documentation/locking/spinlocks.rst b/Documentation/locking/spinlocks.rst
new file mode 100644
index 000000000..bec96f7a9
--- /dev/null
+++ b/Documentation/locking/spinlocks.rst
@@ -0,0 +1,165 @@
+===============
+Locking lessons
+===============
+
+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) ACQUIRE operations.
+
+ (6) RELEASE 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.rst).
+
+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.rst 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).
+
+ Linus
+
+----
+
+Reference information:
+======================
+
+For dynamic initialization, use spin_lock_init() or rwlock_init() as
+appropriate::
+
+ spinlock_t xxx_lock;
+ rwlock_t xxx_rw_lock;
+
+ static int __init xxx_init(void)
+ {
+ spin_lock_init(&xxx_lock);
+ rwlock_init(&xxx_rw_lock);
+ ...
+ }
+
+ module_init(xxx_init);
+
+For static initialization, use DEFINE_SPINLOCK() / DEFINE_RWLOCK() or
+__SPIN_LOCK_UNLOCKED() / __RW_LOCK_UNLOCKED() as appropriate.
diff --git a/Documentation/locking/ww-mutex-design.rst b/Documentation/locking/ww-mutex-design.rst
new file mode 100644
index 000000000..6a8f8beb9
--- /dev/null
+++ b/Documentation/locking/ww-mutex-design.rst
@@ -0,0 +1,393 @@
+======================================
+Wound/Wait Deadlock-Proof Mutex Design
+======================================
+
+Please read mutex-design.rst 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, a reservation ticket is associated with a transaction.
+and the deadlock handling approach is called Wait-Die. The name is based on
+the actions of a locking thread when it encounters an already locked mutex.
+If the transaction holding the lock is younger, the locking transaction waits.
+If the transaction holding the lock is older, the locking transaction backs off
+and dies. Hence Wait-Die.
+There is also another algorithm called Wound-Wait:
+If the transaction holding the lock is younger, the locking transaction
+wounds the transaction holding the lock, requesting it to die.
+If the transaction holding the lock is older, it waits for the other
+transaction. Hence Wound-Wait.
+The two algorithms are both fair in that a transaction will eventually succeed.
+However, the Wound-Wait algorithm is typically stated to generate fewer backoffs
+compared to Wait-Die, but is, on the other hand, associated with more work than
+Wait-Die when recovering from a backoff. Wound-Wait is also a preemptive
+algorithm in that transactions are wounded by other transactions, and that
+requires a reliable way to pick up the wounded condition and preempt the
+running transaction. Note that this is not the same as process preemption. A
+Wound-Wait transaction is considered preempted when it dies (returning
+-EDEADLK) following a wound.
+
+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 that 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. An acquire context is representing a
+transaction.
+
+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. The lock
+class also specifies what algorithm to use, Wound-Wait or Wait-Die.
+
+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 task that just
+ killed its transaction 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
+-----
+
+The algorithm (Wait-Die vs Wound-Wait) is chosen by using either
+DEFINE_WW_CLASS() (Wound-Wait) or DEFINE_WD_CLASS() (Wait-Die)
+As a rough rule of thumb, use Wound-Wait iff you
+expect the number of simultaneous competing transactions to be typically small,
+and you want to reduce the number of rollbacks.
+
+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 die 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) For Wait-Die, 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.
+
+ The Wound-Wait preemption is implemented with a lazy-preemption scheme:
+ The wounded status of the transaction is checked only when there is
+ contention for a new lock and hence a true chance of deadlock. In that
+ situation, if the transaction is wounded, it backs off, clears the
+ wounded status and retries. A great benefit of implementing preemption in
+ this way is that the wounded transaction can identify a contending lock to
+ wait for before restarting the transaction. Just blindly restarting the
+ transaction would likely make the transaction end up in a situation where
+ it would have to back off again.
+
+ In general, not much contention is expected. The locks are typically used to
+ serialize access to resources for devices, and optimization focus should
+ therefore be directed towards the uncontended cases.
+
+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.