diff options
Diffstat (limited to 'Documentation/locking')
| -rw-r--r-- | Documentation/locking/futex-requeue-pi.rst | 132 | ||||
| -rw-r--r-- | Documentation/locking/hwspinlock.rst | 485 | ||||
| -rw-r--r-- | Documentation/locking/index.rst | 33 | ||||
| -rw-r--r-- | Documentation/locking/lockdep-design.rst | 663 | ||||
| -rw-r--r-- | Documentation/locking/lockstat.rst | 204 | ||||
| -rw-r--r-- | Documentation/locking/locktorture.rst | 169 | ||||
| -rw-r--r-- | Documentation/locking/locktypes.rst | 537 | ||||
| -rw-r--r-- | Documentation/locking/mutex-design.rst | 152 | ||||
| -rw-r--r-- | Documentation/locking/percpu-rw-semaphore.rst | 28 | ||||
| -rw-r--r-- | Documentation/locking/pi-futex.rst | 122 | ||||
| -rw-r--r-- | Documentation/locking/preempt-locking.rst | 144 | ||||
| -rw-r--r-- | Documentation/locking/robust-futex-ABI.rst | 184 | ||||
| -rw-r--r-- | Documentation/locking/robust-futexes.rst | 221 | ||||
| -rw-r--r-- | Documentation/locking/rt-mutex-design.rst | 574 | ||||
| -rw-r--r-- | Documentation/locking/rt-mutex.rst | 77 | ||||
| -rw-r--r-- | Documentation/locking/seqlock.rst | 239 | ||||
| -rw-r--r-- | Documentation/locking/spinlocks.rst | 165 | ||||
| -rw-r--r-- | Documentation/locking/ww-mutex-design.rst | 393 |
18 files changed, 4522 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..14ab5787b --- /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-desgin.txt] 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..9f3cfca9f --- /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. + + +---+---+---+---+ + | | E | r | R | + +---+---+---+---+ + | E | 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..4fd7b70fc --- /dev/null +++ b/Documentation/locking/locktypes.rst @@ -0,0 +1,537 @@ +.. 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 regardless of the +kernel configuration including PREEMPT_RT enabled kernels. + +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 disabled. 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..64405e5da --- /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 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..54d9c17bb --- /dev/null +++ b/Documentation/locking/ww-mutex-design.rst @@ -0,0 +1,393 @@ +====================================== +Wound/Wait Deadlock-Proof Mutex Design +====================================== + +Please read mutex-design.txt first, as it applies to wait/wound mutexes too. + +Motivation for WW-Mutexes +------------------------- + +GPU's do operations that commonly involve many buffers. Those buffers +can be shared across contexts/processes, exist in different memory +domains (for example VRAM vs system memory), and so on. And with +PRIME / dmabuf, they can even be shared across devices. So there are +a handful of situations where the driver needs to wait for buffers to +become ready. If you think about this in terms of waiting on a buffer +mutex for it to become available, this presents a problem because +there is no way to guarantee that buffers appear in a execbuf/batch in +the same order in all contexts. That is directly under control of +userspace, and a result of the sequence of GL calls that an application +makes. Which results in the potential for deadlock. The problem gets +more complex when you consider that the kernel may need to migrate the +buffer(s) into VRAM before the GPU operates on the buffer(s), which +may in turn require evicting some other buffers (and you don't want to +evict other buffers which are already queued up to the GPU), but for a +simplified understanding of the problem you can ignore this. + +The algorithm that the TTM graphics subsystem came up with for dealing with +this problem is quite simple. For each group of buffers (execbuf) that need +to be locked, the caller would be assigned a unique reservation id/ticket, +from a global counter. In case of deadlock while locking all the buffers +associated with a execbuf, the one with the lowest reservation ticket (i.e. +the oldest task) wins, and the one with the higher reservation id (i.e. the +younger task) unlocks all of the buffers that it has already locked, and then +tries again. + +In the RDBMS literature, 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 the a task +trying to acquire locks doesn't grab a new reservation id, but keeps the one it +acquired when starting the lock acquisition. This ticket is stored in the +acquire context. Furthermore the acquire context keeps track of debugging state +to catch w/w mutex interface abuse. 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. |