Rusty Russell

rusty@rustcorp.com.au

2003 Rusty Russell This documentation is free software; you can redistribute it and/or modify it under the terms of the GNU General Public License as published by the Free Software Foundation; either version 2 of the License, or (at your option) any later version. This program is distributed in the hope that it will be useful, but WITHOUT ANY WARRANTY; without even the implied warranty of MERCHANTABILITY or FITNESS FOR A PARTICULAR PURPOSE. See the GNU General Public License for more details. You should have received a copy of the GNU General Public License along with this program; if not, write to the Free Software Foundation, Inc., 59 Temple Place, Suite 330, Boston, MA 02111-1307 USA For more details see the file COPYING in the source distribution of Linux. Welcome, to Rusty's Remarkably Unreliable Guide to Kernel Locking issues. This document describes the locking systems in the Linux Kernel in 2.6. With the wide availability of HyperThreading, and preemption in the Linux Kernel, everyone hacking on the kernel needs to know the fundamentals of concurrency and locking for SMP. (Skip this if you know what a Race Condition is). In a normal program, you can increment a counter like so: very_important_count++; This is what they would expect to happen: Instance 1 Instance 2 read very_important_count (5) add 1 (6) write very_important_count (6) read very_important_count (6) add 1 (7) write very_important_count (7) This is what might happen: Instance 1 Instance 2 read very_important_count (5) read very_important_count (5) add 1 (6) add 1 (6) write very_important_count (6) write very_important_count (6) This overlap, where the result depends on the relative timing of multiple tasks, is called a race condition. The piece of code containing the concurrency issue is called a critical region. And especially since Linux starting running on SMP machines, they became one of the major issues in kernel design and implementation. Preemption can have the same effect, even if there is only one CPU: by preempting one task during the critical region, we have exactly the same race condition. In this case the thread which preempts might run the critical region itself. The solution is to recognize when these simultaneous accesses occur, and use locks to make sure that only one instance can enter the critical region at any time. There are many friendly primitives in the Linux kernel to help you do this. And then there are the unfriendly primitives, but I'll pretend they don't exist. If I could give you one piece of advice: never sleep with anyone crazier than yourself. But if I had to give you advice on locking: keep it simple. Be reluctant to introduce new locks. Strangely enough, this last one is the exact reverse of my advice when you have slept with someone crazier than yourself. And you should think about getting a big dog. There are two main types of kernel locks. The fundamental type is the spinlock (include/asm/spinlock.h), which is a very simple single-holder lock: if you can't get the spinlock, you keep trying (spinning) until you can. Spinlocks are very small and fast, and can be used anywhere. The second type is a mutex (include/linux/mutex.h): it is like a spinlock, but you may block holding a mutex. If you can't lock a mutex, your task will suspend itself, and be woken up when the mutex is released. This means the CPU can do something else while you are waiting. There are many cases when you simply can't sleep (see ), and so have to use a spinlock instead. Neither type of lock is recursive: see . For kernels compiled without CONFIG_SMP, and without CONFIG_PREEMPT spinlocks do not exist at all. This is an excellent design decision: when no-one else can run at the same time, there is no reason to have a lock. If the kernel is compiled without CONFIG_SMP, but CONFIG_PREEMPT is set, then spinlocks simply disable preemption, which is sufficient to prevent any races. For most purposes, we can think of preemption as equivalent to SMP, and not worry about it separately. You should always test your locking code with CONFIG_SMP and CONFIG_PREEMPT enabled, even if you don't have an SMP test box, because it will still catch some kinds of locking bugs. Mutexes still exist, because they are required for synchronization between user contexts, as we will see below. If you have a data structure which is only ever accessed from user context, then you can use a simple mutex (include/linux/mutex.h) to protect it. This is the most trivial case: you initialize the mutex. Then you can call mutex_lock_interruptible() to grab the mutex, and mutex_unlock() to release it. There is also a mutex_lock(), which should be avoided, because it will not return if a signal is received. Example: net/netfilter/nf_sockopt.c allows registration of new setsockopt() and getsockopt() calls, with nf_register_sockopt(). Registration and de-registration are only done on module load and unload (and boot time, where there is no concurrency), and the list of registrations is only consulted for an unknown setsockopt() or getsockopt() system call. The nf_sockopt_mutex is perfect to protect this, especially since the setsockopt and getsockopt calls may well sleep. If a softirq shares data with user context, you have two problems. Firstly, the current user context can be interrupted by a softirq, and secondly, the critical region could be entered from another CPU. This is where spin_lock_bh() (include/linux/spinlock.h) is used. It disables softirqs on that CPU, then grabs the lock. spin_unlock_bh() does the reverse. (The '_bh' suffix is a historical reference to "Bottom Halves", the old name for software interrupts. It should really be called spin_lock_softirq()' in a perfect world). Note that you can also use spin_lock_irq() or spin_lock_irqsave() here, which stop hardware interrupts as well: see . This works perfectly for UP as well: the spin lock vanishes, and this macro simply becomes local_bh_disable() (include/linux/interrupt.h), which protects you from the softirq being run. This is exactly the same as above, because tasklets are actually run from a softirq. This, too, is exactly the same as above, because timers are actually run from a softirq. From a locking point of view, tasklets and timers are identical. Sometimes a tasklet or timer might want to share data with another tasklet or timer. Since a tasklet is never run on two CPUs at once, you don't need to worry about your tasklet being reentrant (running twice at once), even on SMP. If another tasklet/timer wants to share data with your tasklet or timer , you will both need to use spin_lock() and spin_unlock() calls. spin_lock_bh() is unnecessary here, as you are already in a tasklet, and none will be run on the same CPU. Often a softirq might want to share data with itself or a tasklet/timer. The same softirq can run on the other CPUs: you can use a per-CPU array (see ) for better performance. If you're going so far as to use a softirq, you probably care about scalable performance enough to justify the extra complexity. You'll need to use spin_lock() and spin_unlock() for shared data. You'll need to use spin_lock() and spin_unlock() for shared data, whether it be a timer, tasklet, different softirq or the same or another softirq: any of them could be running on a different CPU. Hardware interrupts usually communicate with a tasklet or softirq. Frequently this involves putting work in a queue, which the softirq will take out. If a hardware irq handler shares data with a softirq, you have two concerns. Firstly, the softirq processing can be interrupted by a hardware interrupt, and secondly, the critical region could be entered by a hardware interrupt on another CPU. This is where spin_lock_irq() is used. It is defined to disable interrupts on that cpu, then grab the lock. spin_unlock_irq() does the reverse. The irq handler does not to use spin_lock_irq(), because the softirq cannot run while the irq handler is running: it can use spin_lock(), which is slightly faster. The only exception would be if a different hardware irq handler uses the same lock: spin_lock_irq() will stop that from interrupting us. This works perfectly for UP as well: the spin lock vanishes, and this macro simply becomes local_irq_disable() (include/asm/smp.h), which protects you from the softirq/tasklet/BH being run. spin_lock_irqsave() (include/linux/spinlock.h) is a variant which saves whether interrupts were on or off in a flags word, which is passed to spin_unlock_irqrestore(). This means that the same code can be used inside an hard irq handler (where interrupts are already off) and in softirqs (where the irq disabling is required). Note that softirqs (and hence tasklets and timers) are run on return from hardware interrupts, so spin_lock_irq() also stops these. In that sense, spin_lock_irqsave() is the most general and powerful locking function. It is rare to have to share data between two IRQ handlers, but if you do, spin_lock_irqsave() should be used: it is architecture-specific whether all interrupts are disabled inside irq handlers themselves. Pete Zaitcev gives the following summary: If you are in a process context (any syscall) and want to lock other process out, use a mutex. You can take a mutex and sleep (copy_from_user*( or kmalloc(x,GFP_KERNEL)). Otherwise (== data can be touched in an interrupt), use spin_lock_irqsave() and spin_unlock_irqrestore(). Avoid holding spinlock for more than 5 lines of code and across any function call (except accessors like readb). The following table lists the minimum locking requirements between various contexts. In some cases, the same context can only be running on one CPU at a time, so no locking is required for that context (eg. a particular thread can only run on one CPU at a time, but if it needs shares data with another thread, locking is required). Remember the advice above: you can always use spin_lock_irqsave(), which is a superset of all other spinlock primitives. IRQ Handler A IRQ Handler B Softirq A Softirq B Tasklet A Tasklet B Timer A Timer B User Context A User Context B IRQ Handler A None IRQ Handler B SLIS None Softirq A SLI SLI SL Softirq B SLI SLI SL SL Tasklet A SLI SLI SL SL None Tasklet B SLI SLI SL SL SL None Timer A SLI SLI SL SL SL SL None Timer B SLI SLI SL SL SL SL SL None User Context A SLI SLI SLBH SLBH SLBH SLBH SLBH SLBH None User Context B SLI SLI SLBH SLBH SLBH SLBH SLBH SLBH MLI None SLIS spin_lock_irqsave SLI spin_lock_irq SL spin_lock SLBH spin_lock_bh MLI mutex_lock_interruptible There are functions that try to acquire a lock only once and immediately return a value telling about success or failure to acquire the lock. They can be used if you need no access to the data protected with the lock when some other thread is holding the lock. You should acquire the lock later if you then need access to the data protected with the lock. spin_trylock() does not spin but returns non-zero if it acquires the spinlock on the first try or 0 if not. This function can be used in all contexts like spin_lock: you must have disabled the contexts that might interrupt you and acquire the spin lock. mutex_trylock() does not suspend your task but returns non-zero if it could lock the mutex on the first try or 0 if not. This function cannot be safely used in hardware or software interrupt contexts despite not sleeping. Let's step through a simple example: a cache of number to name mappings. The cache keeps a count of how often each of the objects is used, and when it gets full, throws out the least used one. For our first example, we assume that all operations are in user context (ie. from system calls), so we can sleep. This means we can use a mutex to protect the cache and all the objects within it. Here's the code: #include <linux/list.h> #include <linux/slab.h> #include <linux/string.h> #include <linux/mutex.h> #include <asm/errno.h> struct object { struct list_head list; int id; char name[32]; int popularity; }; /* Protects the cache, cache_num, and the objects within it */ static DEFINE_MUTEX(cache_lock); static LIST_HEAD(cache); static unsigned int cache_num = 0; #define MAX_CACHE_SIZE 10 /* Must be holding cache_lock */ static struct object *__cache_find(int id) { struct object *i; list_for_each_entry(i, &cache, list) if (i->id == id) { i->popularity++; return i; } return NULL; } /* Must be holding cache_lock */ static void __cache_delete(struct object *obj) { BUG_ON(!obj); list_del(&obj->list); kfree(obj); cache_num--; } /* Must be holding cache_lock */ static void __cache_add(struct object *obj) { list_add(&obj->list, &cache); if (++cache_num > MAX_CACHE_SIZE) { struct object *i, *outcast = NULL; list_for_each_entry(i, &cache, list) { if (!outcast || i->popularity < outcast->popularity) outcast = i; } __cache_delete(outcast); } } int cache_add(int id, const char *name) { struct object *obj; if ((obj = kmalloc(sizeof(*obj), GFP_KERNEL)) == NULL) return -ENOMEM; strlcpy(obj->name, name, sizeof(obj->name)); obj->id = id; obj->popularity = 0; mutex_lock(&cache_lock); __cache_add(obj); mutex_unlock(&cache_lock); return 0; } void cache_delete(int id) { mutex_lock(&cache_lock); __cache_delete(__cache_find(id)); mutex_unlock(&cache_lock); } int cache_find(int id, char *name) { struct object *obj; int ret = -ENOENT; mutex_lock(&cache_lock); obj = __cache_find(id); if (obj) { ret = 0; strcpy(name, obj->name); } mutex_unlock(&cache_lock); return ret; } Note that we always make sure we have the cache_lock when we add, delete, or look up the cache: both the cache infrastructure itself and the contents of the objects are protected by the lock. In this case it's easy, since we copy the data for the user, and never let them access the objects directly. There is a slight (and common) optimization here: in cache_add we set up the fields of the object before grabbing the lock. This is safe, as no-one else can access it until we put it in cache. Now consider the case where cache_find can be called from interrupt context: either a hardware interrupt or a softirq. An example would be a timer which deletes object from the cache. The change is shown below, in standard patch format: the - are lines which are taken away, and the + are lines which are added. --- cache.c.usercontext 2003-12-09 13:58:54.000000000 +1100 +++ cache.c.interrupt 2003-12-09 14:07:49.000000000 +1100 @@ -12,7 +12,7 @@ int popularity; }; -static DEFINE_MUTEX(cache_lock); +static DEFINE_SPINLOCK(cache_lock); static LIST_HEAD(cache); static unsigned int cache_num = 0; #define MAX_CACHE_SIZE 10 @@ -55,6 +55,7 @@ int cache_add(int id, const char *name) { struct object *obj; + unsigned long flags; if ((obj = kmalloc(sizeof(*obj), GFP_KERNEL)) == NULL) return -ENOMEM; @@ -63,30 +64,33 @@ obj->id = id; obj->popularity = 0; - mutex_lock(&cache_lock); + spin_lock_irqsave(&cache_lock, flags); __cache_add(obj); - mutex_unlock(&cache_lock); + spin_unlock_irqrestore(&cache_lock, flags); return 0; } void cache_delete(int id) { - mutex_lock(&cache_lock); + unsigned long flags; + + spin_lock_irqsave(&cache_lock, flags); __cache_delete(__cache_find(id)); - mutex_unlock(&cache_lock); + spin_unlock_irqrestore(&cache_lock, flags); } int cache_find(int id, char *name) { struct object *obj; int ret = -ENOENT; + unsigned long flags; - mutex_lock(&cache_lock); + spin_lock_irqsave(&cache_lock, flags); obj = __cache_find(id); if (obj) { ret = 0; strcpy(name, obj->name); } - mutex_unlock(&cache_lock); + spin_unlock_irqrestore(&cache_lock, flags); return ret; } Note that the spin_lock_irqsave will turn off interrupts if they are on, otherwise does nothing (if we are already in an interrupt handler), hence these functions are safe to call from any context. Unfortunately, cache_add calls kmalloc with the GFP_KERNEL flag, which is only legal in user context. I have assumed that cache_add is still only called in user context, otherwise this should become a parameter to cache_add. If our objects contained more information, it might not be sufficient to copy the information in and out: other parts of the code might want to keep pointers to these objects, for example, rather than looking up the id every time. This produces two problems. The first problem is that we use the cache_lock to protect objects: we'd need to make this non-static so the rest of the code can use it. This makes locking trickier, as it is no longer all in one place. The second problem is the lifetime problem: if another structure keeps a pointer to an object, it presumably expects that pointer to remain valid. Unfortunately, this is only guaranteed while you hold the lock, otherwise someone might call cache_delete and even worse, add another object, re-using the same address. As there is only one lock, you can't hold it forever: no-one else would get any work done. The solution to this problem is to use a reference count: everyone who has a pointer to the object increases it when they first get the object, and drops the reference count when they're finished with it. Whoever drops it to zero knows it is unused, and can actually delete it. Here is the code: --- cache.c.interrupt 2003-12-09 14:25:43.000000000 +1100 +++ cache.c.refcnt 2003-12-09 14:33:05.000000000 +1100 @@ -7,6 +7,7 @@ struct object { struct list_head list; + unsigned int refcnt; int id; char name[32]; int popularity; @@ -17,6 +18,35 @@ static unsigned int cache_num = 0; #define MAX_CACHE_SIZE 10 +static void __object_put(struct object *obj) +{ + if (--obj->refcnt == 0) + kfree(obj); +} + +static void __object_get(struct object *obj) +{ + obj->refcnt++; +} + +void object_put(struct object *obj) +{ + unsigned long flags; + + spin_lock_irqsave(&cache_lock, flags); + __object_put(obj); + spin_unlock_irqrestore(&cache_lock, flags); +} + +void object_get(struct object *obj) +{ + unsigned long flags; + + spin_lock_irqsave(&cache_lock, flags); + __object_get(obj); + spin_unlock_irqrestore(&cache_lock, flags); +} + /* Must be holding cache_lock */ static struct object *__cache_find(int id) { @@ -35,6 +65,7 @@ { BUG_ON(!obj); list_del(&obj->list); + __object_put(obj); cache_num--; } @@ -63,6 +94,7 @@ strlcpy(obj->name, name, sizeof(obj->name)); obj->id = id; obj->popularity = 0; + obj->refcnt = 1; /* The cache holds a reference */ spin_lock_irqsave(&cache_lock, flags); __cache_add(obj); @@ -79,18 +111,15 @@ spin_unlock_irqrestore(&cache_lock, flags); } -int cache_find(int id, char *name) +struct object *cache_find(int id) { struct object *obj; - int ret = -ENOENT; unsigned long flags; spin_lock_irqsave(&cache_lock, flags); obj = __cache_find(id); - if (obj) { - ret = 0; - strcpy(name, obj->name); - } + if (obj) + __object_get(obj); spin_unlock_irqrestore(&cache_lock, flags); - return ret; + return obj; } We encapsulate the reference counting in the standard 'get' and 'put' functions. Now we can return the object itself from cache_find which has the advantage that the user can now sleep holding the object (eg. to copy_to_user to name to userspace). The other point to note is that I said a reference should be held for every pointer to the object: thus the reference count is 1 when first inserted into the cache. In some versions the framework does not hold a reference count, but they are more complicated. In practice, atomic_t would usually be used for refcnt. There are a number of atomic operations defined in include/asm/atomic.h: these are guaranteed to be seen atomically from all CPUs in the system, so no lock is required. In this case, it is simpler than using spinlocks, although for anything non-trivial using spinlocks is clearer. The atomic_inc and atomic_dec_and_test are used instead of the standard increment and decrement operators, and the lock is no longer used to protect the reference count itself. --- cache.c.refcnt 2003-12-09 15:00:35.000000000 +1100 +++ cache.c.refcnt-atomic 2003-12-11 15:49:42.000000000 +1100 @@ -7,7 +7,7 @@ struct object { struct list_head list; - unsigned int refcnt; + atomic_t refcnt; int id; char name[32]; int popularity; @@ -18,33 +18,15 @@ static unsigned int cache_num = 0; #define MAX_CACHE_SIZE 10 -static void __object_put(struct object *obj) -{ - if (--obj->refcnt == 0) - kfree(obj); -} - -static void __object_get(struct object *obj) -{ - obj->refcnt++; -} - void object_put(struct object *obj) { - unsigned long flags; - - spin_lock_irqsave(&cache_lock, flags); - __object_put(obj); - spin_unlock_irqrestore(&cache_lock, flags); + if (atomic_dec_and_test(&obj->refcnt)) + kfree(obj); } void object_get(struct object *obj) { - unsigned long flags; - - spin_lock_irqsave(&cache_lock, flags); - __object_get(obj); - spin_unlock_irqrestore(&cache_lock, flags); + atomic_inc(&obj->refcnt); } /* Must be holding cache_lock */ @@ -65,7 +47,7 @@ { BUG_ON(!obj); list_del(&obj->list); - __object_put(obj); + object_put(obj); cache_num--; } @@ -94,7 +76,7 @@ strlcpy(obj->name, name, sizeof(obj->name)); obj->id = id; obj->popularity = 0; - obj->refcnt = 1; /* The cache holds a reference */ + atomic_set(&obj->refcnt, 1); /* The cache holds a reference */ spin_lock_irqsave(&cache_lock, flags); __cache_add(obj); @@ -119,7 +101,7 @@ spin_lock_irqsave(&cache_lock, flags); obj = __cache_find(id); if (obj) - __object_get(obj); + object_get(obj); spin_unlock_irqrestore(&cache_lock, flags); return obj; } In these examples, we assumed that the objects (except the reference counts) never changed once they are created. If we wanted to allow the name to change, there are three possibilities: You can make cache_lock non-static, and tell people to grab that lock before changing the name in any object. You can provide a cache_obj_rename which grabs this lock and changes the name for the caller, and tell everyone to use that function. You can make the cache_lock protect only the cache itself, and use another lock to protect the name. Theoretically, you can make the locks as fine-grained as one lock for every field, for every object. In practice, the most common variants are: One lock which protects the infrastructure (the cache list in this example) and all the objects. This is what we have done so far. One lock which protects the infrastructure (including the list pointers inside the objects), and one lock inside the object which protects the rest of that object. Multiple locks to protect the infrastructure (eg. one lock per hash chain), possibly with a separate per-object lock. Here is the "lock-per-object" implementation: --- cache.c.refcnt-atomic 2003-12-11 15:50:54.000000000 +1100 +++ cache.c.perobjectlock 2003-12-11 17:15:03.000000000 +1100 @@ -6,11 +6,17 @@ struct object { + /* These two protected by cache_lock. */ struct list_head list; + int popularity; + atomic_t refcnt; + + /* Doesn't change once created. */ int id; + + spinlock_t lock; /* Protects the name */ char name[32]; - int popularity; }; static DEFINE_SPINLOCK(cache_lock); @@ -77,6 +84,7 @@ obj->id = id; obj->popularity = 0; atomic_set(&obj->refcnt, 1); /* The cache holds a reference */ + spin_lock_init(&obj->lock); spin_lock_irqsave(&cache_lock, flags); __cache_add(obj); Note that I decide that the popularity count should be protected by the cache_lock rather than the per-object lock: this is because it (like the struct list_head inside the object) is logically part of the infrastructure. This way, I don't need to grab the lock of every object in __cache_add when seeking the least popular. I also decided that the id member is unchangeable, so I don't need to grab each object lock in __cache_find() to examine the id: the object lock is only used by a caller who wants to read or write the name field. Note also that I added a comment describing what data was protected by which locks. This is extremely important, as it describes the runtime behavior of the code, and can be hard to gain from just reading. And as Alan Cox says, Lock data, not code. There is a coding bug where a piece of code tries to grab a spinlock twice: it will spin forever, waiting for the lock to be released (spinlocks, rwlocks and mutexes are not recursive in Linux). This is trivial to diagnose: not a stay-up-five-nights-talk-to-fluffy-code-bunnies kind of problem. For a slightly more complex case, imagine you have a region shared by a softirq and user context. If you use a spin_lock() call to protect it, it is possible that the user context will be interrupted by the softirq while it holds the lock, and the softirq will then spin forever trying to get the same lock. Both of these are called deadlock, and as shown above, it can occur even with a single CPU (although not on UP compiles, since spinlocks vanish on kernel compiles with CONFIG_SMP=n. You'll still get data corruption in the second example). This complete lockup is easy to diagnose: on SMP boxes the watchdog timer or compiling with DEBUG_SPINLOCK set (include/linux/spinlock.h) will show this up immediately when it happens. A more complex problem is the so-called 'deadly embrace', involving two or more locks. Say you have a hash table: each entry in the table is a spinlock, and a chain of hashed objects. Inside a softirq handler, you sometimes want to alter an object from one place in the hash to another: you grab the spinlock of the old hash chain and the spinlock of the new hash chain, and delete the object from the old one, and insert it in the new one. There are two problems here. First, if your code ever tries to move the object to the same chain, it will deadlock with itself as it tries to lock it twice. Secondly, if the same softirq on another CPU is trying to move another object in the reverse direction, the following could happen: CPU 1 CPU 2 Grab lock A -> OK Grab lock B -> OK Grab lock B -> spin Grab lock A -> spin The two CPUs will spin forever, waiting for the other to give up their lock. It will look, smell, and feel like a crash. Textbooks will tell you that if you always lock in the same order, you will never get this kind of deadlock. Practice will tell you that this approach doesn't scale: when I create a new lock, I don't understand enough of the kernel to figure out where in the 5000 lock hierarchy it will fit. The best locks are encapsulated: they never get exposed in headers, and are never held around calls to non-trivial functions outside the same file. You can read through this code and see that it will never deadlock, because it never tries to grab another lock while it has that one. People using your code don't even need to know you are using a lock. A classic problem here is when you provide callbacks or hooks: if you call these with the lock held, you risk simple deadlock, or a deadly embrace (who knows what the callback will do?). Remember, the other programmers are out to get you, so don't do this. Deadlocks are problematic, but not as bad as data corruption. Code which grabs a read lock, searches a list, fails to find what it wants, drops the read lock, grabs a write lock and inserts the object has a race condition. If you don't see why, please stay the fuck away from my code. Timers can produce their own special problems with races. Consider a collection of objects (list, hash, etc) where each object has a timer which is due to destroy it. If you want to destroy the entire collection (say on module removal), you might do the following: /* THIS CODE BAD BAD BAD BAD: IF IT WAS ANY WORSE IT WOULD USE HUNGARIAN NOTATION */ spin_lock_bh(&list_lock); while (list) { struct foo *next = list->next; del_timer(&list->timer); kfree(list); list = next; } spin_unlock_bh(&list_lock); Sooner or later, this will crash on SMP, because a timer can have just gone off before the spin_lock_bh(), and it will only get the lock after we spin_unlock_bh(), and then try to free the element (which has already been freed!). This can be avoided by checking the result of del_timer(): if it returns 1, the timer has been deleted. If 0, it means (in this case) that it is currently running, so we can do: retry: spin_lock_bh(&list_lock); while (list) { struct foo *next = list->next; if (!del_timer(&list->timer)) { /* Give timer a chance to delete this */ spin_unlock_bh(&list_lock); goto retry; } kfree(list); list = next; } spin_unlock_bh(&list_lock); Another common problem is deleting timers which restart themselves (by calling add_timer() at the end of their timer function). Because this is a fairly common case which is prone to races, you should use del_timer_sync() (include/linux/timer.h) to handle this case. It returns the number of times the timer had to be deleted before we finally stopped it from adding itself back in. There are three main things to worry about when considering speed of some code which does locking. First is concurrency: how many things are going to be waiting while someone else is holding a lock. Second is the time taken to actually acquire and release an uncontended lock. Third is using fewer, or smarter locks. I'm assuming that the lock is used fairly often: otherwise, you wouldn't be concerned about efficiency. Concurrency depends on how long the lock is usually held: you should hold the lock for as long as needed, but no longer. In the cache example, we always create the object without the lock held, and then grab the lock only when we are ready to insert it in the list. Acquisition times depend on how much damage the lock operations do to the pipeline (pipeline stalls) and how likely it is that this CPU was the last one to grab the lock (ie. is the lock cache-hot for this CPU): on a machine with more CPUs, this likelihood drops fast. Consider a 700MHz Intel Pentium III: an instruction takes about 0.7ns, an atomic increment takes about 58ns, a lock which is cache-hot on this CPU takes 160ns, and a cacheline transfer from another CPU takes an additional 170 to 360ns. (These figures from Paul McKenney's Linux Journal RCU article). These two aims conflict: holding a lock for a short time might be done by splitting locks into parts (such as in our final per-object-lock example), but this increases the number of lock acquisitions, and the results are often slower than having a single lock. This is another reason to advocate locking simplicity. The third concern is addressed below: there are some methods to reduce the amount of locking which needs to be done. Both spinlocks and mutexes have read/write variants: rwlock_t and struct rw_semaphore. These divide users into two classes: the readers and the writers. If you are only reading the data, you can get a read lock, but to write to the data you need the write lock. Many people can hold a read lock, but a writer must be sole holder. If your code divides neatly along reader/writer lines (as our cache code does), and the lock is held by readers for significant lengths of time, using these locks can help. They are slightly slower than the normal locks though, so in practice rwlock_t is not usually worthwhile. There is a special method of read/write locking called Read Copy Update. Using RCU, the readers can avoid taking a lock altogether: as we expect our cache to be read more often than updated (otherwise the cache is a waste of time), it is a candidate for this optimization. How do we get rid of read locks? Getting rid of read locks means that writers may be changing the list underneath the readers. That is actually quite simple: we can read a linked list while an element is being added if the writer adds the element very carefully. For example, adding new to a single linked list called list: new->next = list->next; wmb(); list->next = new; The wmb() is a write memory barrier. It ensures that the first operation (setting the new element's next pointer) is complete and will be seen by all CPUs, before the second operation is (putting the new element into the list). This is important, since modern compilers and modern CPUs can both reorder instructions unless told otherwise: we want a reader to either not see the new element at all, or see the new element with the next pointer correctly pointing at the rest of the list. Fortunately, there is a function to do this for standard struct list_head lists: list_add_rcu() (include/linux/list.h). Removing an element from the list is even simpler: we replace the pointer to the old element with a pointer to its successor, and readers will either see it, or skip over it. list->next = old->next; There is list_del_rcu() (include/linux/list.h) which does this (the normal version poisons the old object, which we don't want). The reader must also be careful: some CPUs can look through the next pointer to start reading the contents of the next element early, but don't realize that the pre-fetched contents is wrong when the next pointer changes underneath them. Once again, there is a list_for_each_entry_rcu() (include/linux/list.h) to help you. Of course, writers can just use list_for_each_entry(), since there cannot be two simultaneous writers. Our final dilemma is this: when can we actually destroy the removed element? Remember, a reader might be stepping through this element in the list right now: if we free this element and the next pointer changes, the reader will jump off into garbage and crash. We need to wait until we know that all the readers who were traversing the list when we deleted the element are finished. We use call_rcu() to register a callback which will actually destroy the object once all pre-existing readers are finished. Alternatively, synchronize_rcu() may be used to block until all pre-existing are finished. But how does Read Copy Update know when the readers are finished? The method is this: firstly, the readers always traverse the list inside rcu_read_lock()/rcu_read_unlock() pairs: these simply disable preemption so the reader won't go to sleep while reading the list. RCU then waits until every other CPU has slept at least once: since readers cannot sleep, we know that any readers which were traversing the list during the deletion are finished, and the callback is triggered. The real Read Copy Update code is a little more optimized than this, but this is the fundamental idea. --- cache.c.perobjectlock 2003-12-11 17:15:03.000000000 +1100 +++ cache.c.rcupdate 2003-12-11 17:55:14.000000000 +1100 @@ -1,15 +1,18 @@ #include <linux/list.h> #include <linux/slab.h> #include <linux/string.h> +#include <linux/rcupdate.h> #include <linux/mutex.h> #include <asm/errno.h> struct object { - /* These two protected by cache_lock. */ + /* This is protected by RCU */ struct list_head list; int popularity; + struct rcu_head rcu; + atomic_t refcnt; /* Doesn't change once created. */ @@ -40,7 +43,7 @@ { struct object *i; - list_for_each_entry(i, &cache, list) { + list_for_each_entry_rcu(i, &cache, list) { if (i->id == id) { i->popularity++; return i; @@ -49,19 +52,25 @@ return NULL; } +/* Final discard done once we know no readers are looking. */ +static void cache_delete_rcu(void *arg) +{ + object_put(arg); +} + /* Must be holding cache_lock */ static void __cache_delete(struct object *obj) { BUG_ON(!obj); - list_del(&obj->list); - object_put(obj); + list_del_rcu(&obj->list); cache_num--; + call_rcu(&obj->rcu, cache_delete_rcu); } /* Must be holding cache_lock */ static void __cache_add(struct object *obj) { - list_add(&obj->list, &cache); + list_add_rcu(&obj->list, &cache); if (++cache_num > MAX_CACHE_SIZE) { struct object *i, *outcast = NULL; list_for_each_entry(i, &cache, list) { @@ -104,12 +114,11 @@ struct object *cache_find(int id) { struct object *obj; - unsigned long flags; - spin_lock_irqsave(&cache_lock, flags); + rcu_read_lock(); obj = __cache_find(id); if (obj) object_get(obj); - spin_unlock_irqrestore(&cache_lock, flags); + rcu_read_unlock(); return obj; } Note that the reader will alter the popularity member in __cache_find(), and now it doesn't hold a lock. One solution would be to make it an atomic_t, but for this usage, we don't really care about races: an approximate result is good enough, so I didn't change it. The result is that cache_find() requires no synchronization with any other functions, so is almost as fast on SMP as it would be on UP. There is a further optimization possible here: remember our original cache code, where there were no reference counts and the caller simply held the lock whenever using the object? This is still possible: if you hold the lock, no one can delete the object, so you don't need to get and put the reference count. Now, because the 'read lock' in RCU is simply disabling preemption, a caller which always has preemption disabled between calling cache_find() and object_put() does not need to actually get and put the reference count: we could expose __cache_find() by making it non-static, and such callers could simply call that. The benefit here is that the reference count is not written to: the object is not altered in any way, which is much faster on SMP machines due to caching. Another technique for avoiding locking which is used fairly widely is to duplicate information for each CPU. For example, if you wanted to keep a count of a common condition, you could use a spin lock and a single counter. Nice and simple. If that was too slow (it's usually not, but if you've got a really big machine to test on and can show that it is), you could instead use a counter for each CPU, then none of them need an exclusive lock. See DEFINE_PER_CPU(), get_cpu_var() and put_cpu_var() (include/linux/percpu.h). Of particular use for simple per-cpu counters is the local_t type, and the cpu_local_inc() and related functions, which are more efficient than simple code on some architectures (include/asm/local.h). Note that there is no simple, reliable way of getting an exact value of such a counter, without introducing more locks. This is not a problem for some uses. If data is always accessed from within the same IRQ handler, you don't need a lock at all: the kernel already guarantees that the irq handler will not run simultaneously on multiple CPUs. Manfred Spraul points out that you can still do this, even if the data is very occasionally accessed in user context or softirqs/tasklets. The irq handler doesn't use a lock, and all other accesses are done as so: spin_lock(&lock); disable_irq(irq); ... enable_irq(irq); spin_unlock(&lock); The disable_irq() prevents the irq handler from running (and waits for it to finish if it's currently running on other CPUs). The spinlock prevents any other accesses happening at the same time. Naturally, this is slower than just a spin_lock_irq() call, so it only makes sense if this type of access happens extremely rarely. Many functions in the kernel sleep (ie. call schedule()) directly or indirectly: you can never call them while holding a spinlock, or with preemption disabled. This also means you need to be in user context: calling them from an interrupt is illegal. The most common ones are listed below, but you usually have to read the code to find out if other calls are safe. If everyone else who calls it can sleep, you probably need to be able to sleep, too. In particular, registration and deregistration functions usually expect to be called from user context, and can sleep. Accesses to userspace: copy_from_user() copy_to_user() get_user() put_user() kmalloc(GFP_KERNEL) mutex_lock_interruptible() and mutex_lock() There is a mutex_trylock() which does not sleep. Still, it must not be used inside interrupt context since its implementation is not safe for that. mutex_unlock() will also never sleep. It cannot be used in interrupt context either since a mutex must be released by the same task that acquired it. Some functions are safe to call from any context, or holding almost any lock. printk() kfree() add_timer() and del_timer() Kernel Hackers Manual July 2017 mutex_init 9 4.1.27 mutex_init initialize the mutex mutex_init mutex mutex the mutex to be initialized Initialize the mutex to unlocked state. It is not allowed to initialize an already locked mutex. Kernel Hackers Manual July 2017 mutex_is_locked 9 4.1.27 mutex_is_locked is the mutex locked int mutex_is_locked struct mutex * lock lock the mutex to be queried Returns 1 if the mutex is locked, 0 if unlocked. Kernel Hackers Manual July 2017 mutex_lock 9 4.1.27 mutex_lock acquire the mutex void __sched mutex_lock struct mutex * lock lock the mutex to be acquired Lock the mutex exclusively for this task. If the mutex is not available right now, it will sleep until it can get it. The mutex must later on be released by the same task that acquired it. Recursive locking is not allowed. The task may not exit without first unlocking the mutex. Also, kernel memory where the mutex resides must not be freed with the mutex still locked. The mutex must first be initialized (or statically defined) before it can be locked. memset-ing the mutex to 0 is not allowed. ( The CONFIG_DEBUG_MUTEXES .config option turns on debugging checks that will enforce the restrictions and will also do deadlock debugging. ) This function is similar to (but not equivalent to) down. Kernel Hackers Manual July 2017 mutex_unlock 9 4.1.27 mutex_unlock release the mutex void __sched mutex_unlock struct mutex * lock lock the mutex to be released Unlock a mutex that has been locked by this task previously. This function must not be used in interrupt context. Unlocking of a not locked mutex is not allowed. This function is similar to (but not equivalent to) up. Kernel Hackers Manual July 2017 ww_mutex_unlock 9 4.1.27 ww_mutex_unlock release the w/w mutex void __sched ww_mutex_unlock struct ww_mutex * lock lock the mutex to be released Unlock a mutex that has been locked by this task previously with any of the ww_mutex_lock* functions (with or without an acquire context). It is forbidden to release the locks after releasing the acquire context. This function must not be used in interrupt context. Unlocking of a unlocked mutex is not allowed. Kernel Hackers Manual July 2017 mutex_lock_interruptible 9 4.1.27 mutex_lock_interruptible acquire the mutex, interruptible int __sched mutex_lock_interruptible struct mutex * lock lock the mutex to be acquired Lock the mutex like mutex_lock, and return 0 if the mutex has been acquired or sleep until the mutex becomes available. If a signal arrives while waiting for the lock then this function returns -EINTR. This function is similar to (but not equivalent to) down_interruptible. Kernel Hackers Manual July 2017 mutex_trylock 9 4.1.27 mutex_trylock try to acquire the mutex, without waiting int __sched mutex_trylock struct mutex * lock lock the mutex to be acquired Try to acquire the mutex atomically. Returns 1 if the mutex has been acquired successfully, and 0 on contention. this function follows the spin_trylock convention, so it is negated from the down_trylock return values! Be careful about this when converting semaphore users to mutexes. This function must not be used in interrupt context. The mutex must be released by the same task that acquired it. Kernel Hackers Manual July 2017 atomic_dec_and_mutex_lock 9 4.1.27 atomic_dec_and_mutex_lock return holding mutex if we dec to 0 int atomic_dec_and_mutex_lock atomic_t * cnt struct mutex * lock cnt the atomic which we are to dec lock the mutex to return holding if we dec to 0 return true and hold lock if we dec to 0, return false otherwise Kernel Hackers Manual July 2017 struct futex_q 9 4.1.27 struct futex_q The hashed futex queue entry, one per waiting task struct futex_q { struct plist_node list; struct task_struct * task; spinlock_t * lock_ptr; union futex_key key; struct futex_pi_state * pi_state; struct rt_mutex_waiter * rt_waiter; union futex_key * requeue_pi_key; u32 bitset; }; list priority-sorted list of tasks waiting on this futex task the task waiting on the futex lock_ptr the hash bucket lock key the key the futex is hashed on pi_state optional priority inheritance state rt_waiter rt_waiter storage for use with requeue_pi requeue_pi_key the requeue_pi target futex key bitset bitset for the optional bitmasked wakeup We use this hashed waitqueue, instead of a normal wait_queue_t, so we can wake only the relevant ones (hashed queues may be shared). A futex_q has a woken state, just like tasks have TASK_RUNNING. It is considered woken when plist_node_empty(q->list) || q->lock_ptr == 0. The order of wakeup is always to make the first condition true, then the second. PI futexes are typically woken before they are removed from the hash list via the rt_mutex code. See unqueue_me_pi. Kernel Hackers Manual July 2017 get_futex_key 9 4.1.27 get_futex_key Get parameters which are the keys for a futex int get_futex_key u32 __user * uaddr int fshared union futex_key * key int rw uaddr virtual address of the futex fshared 0 for a PROCESS_PRIVATE futex, 1 for PROCESS_SHARED key address where result is stored. rw mapping needs to be read/write (values: VERIFY_READ, VERIFY_WRITE) a negative error code or 0 The key words are stored in *key on success. For shared mappings, it's (page->index, file_inode(vma->vm_file), offset_within_page). For private mappings, it's (uaddr, current->mm). We can usually work out the index without swapping in the page. lock_page might sleep, the caller should not hold a spinlock. Kernel Hackers Manual July 2017 fault_in_user_writeable 9 4.1.27 fault_in_user_writeable Fault in user address and verify RW access int fault_in_user_writeable u32 __user * uaddr uaddr pointer to faulting user space address Slow path to fixup the fault we just took in the atomic write access to uaddr. We have no generic implementation of a non-destructive write to the user address. We know that we faulted in the atomic pagefault disabled section so we can as well avoid the #PF overhead by calling get_user_pages right away. Kernel Hackers Manual July 2017 futex_top_waiter 9 4.1.27 futex_top_waiter Return the highest priority waiter on a futex struct futex_q * futex_top_waiter struct futex_hash_bucket * hb union futex_key * key hb the hash bucket the futex_q's reside in key the futex key (to distinguish it from other futex futex_q's) Must be called with the hb lock held. Kernel Hackers Manual July 2017 futex_lock_pi_atomic 9 4.1.27 futex_lock_pi_atomic Atomic work required to acquire a pi aware futex int futex_lock_pi_atomic u32 __user * uaddr struct futex_hash_bucket * hb union futex_key * key struct futex_pi_state ** ps struct task_struct * task int set_waiters uaddr the pi futex user address hb the pi futex hash bucket key the futex key associated with uaddr and hb ps the pi_state pointer where we store the result of the lookup task the task to perform the atomic lock work for. This will be current except in the case of requeue pi. set_waiters force setting the FUTEX_WAITERS bit (1) or not (0) 0 - ready to wait; 1 - acquired the lock; <0 - error The hb->lock and futex_key refs shall be held by the caller. Kernel Hackers Manual July 2017 __unqueue_futex 9 4.1.27 __unqueue_futex Remove the futex_q from its futex_hash_bucket void __unqueue_futex struct futex_q * q q The futex_q to unqueue The q->lock_ptr must not be NULL and must be held by the caller. Kernel Hackers Manual July 2017 requeue_futex 9 4.1.27 requeue_futex Requeue a futex_q from one hb to another void requeue_futex struct futex_q * q struct futex_hash_bucket * hb1 struct futex_hash_bucket * hb2 union futex_key * key2 q the futex_q to requeue hb1 the source hash_bucket hb2 the target hash_bucket key2 the new key for the requeued futex_q Kernel Hackers Manual July 2017 requeue_pi_wake_futex 9 4.1.27 requeue_pi_wake_futex Wake a task that acquired the lock during requeue void requeue_pi_wake_futex struct futex_q * q union futex_key * key struct futex_hash_bucket * hb q the futex_q key the key of the requeue target futex hb the hash_bucket of the requeue target futex During futex_requeue, with requeue_pi=1, it is possible to acquire the target futex if it is uncontended or via a lock steal. Set the futex_q key to the requeue target futex so the waiter can detect the wakeup on the right futex, but remove it from the hb and NULL the rt_waiter so it can detect atomic lock acquisition. Set the q->lock_ptr to the requeue target hb->lock to protect access to the pi_state to fixup the owner later. Must be called with both q->lock_ptr and hb->lock held. Kernel Hackers Manual July 2017 futex_proxy_trylock_atomic 9 4.1.27 futex_proxy_trylock_atomic Attempt an atomic lock for the top waiter int futex_proxy_trylock_atomic u32 __user * pifutex struct futex_hash_bucket * hb1 struct futex_hash_bucket * hb2 union futex_key * key1 union futex_key * key2 struct futex_pi_state ** ps int set_waiters pifutex the user address of the to futex hb1 the from futex hash bucket, must be locked by the caller hb2 the to futex hash bucket, must be locked by the caller key1 the from futex key key2 the to futex key ps address to store the pi_state pointer set_waiters force setting the FUTEX_WAITERS bit (1) or not (0) Try and get the lock on behalf of the top waiter if we can do it atomically. Wake the top waiter if we succeed. If the caller specified set_waiters, then direct futex_lock_pi_atomic to force setting the FUTEX_WAITERS bit. hb1 and hb2 must be held by the caller. 0 - failed to acquire the lock atomically; >0 - acquired the lock, return value is vpid of the top_waiter <0 - error Kernel Hackers Manual July 2017 futex_requeue 9 4.1.27 futex_requeue Requeue waiters from uaddr1 to uaddr2 int futex_requeue u32 __user * uaddr1 unsigned int flags u32 __user * uaddr2 int nr_wake int nr_requeue u32 * cmpval int requeue_pi uaddr1 source futex user address flags futex flags (FLAGS_SHARED, etc.) uaddr2 target futex user address nr_wake number of waiters to wake (must be 1 for requeue_pi) nr_requeue number of waiters to requeue (0-INT_MAX) cmpval uaddr1 expected value (or NULL) requeue_pi if we are attempting to requeue from a non-pi futex to a pi futex (pi to pi requeue is not supported) Requeue waiters on uaddr1 to uaddr2. In the requeue_pi case, try to acquire uaddr2 atomically on behalf of the top waiter. >=0 - on success, the number of tasks requeued or woken; <0 - on error Kernel Hackers Manual July 2017 queue_me 9 4.1.27 queue_me Enqueue the futex_q on the futex_hash_bucket void queue_me struct futex_q * q struct futex_hash_bucket * hb q The futex_q to enqueue hb The destination hash bucket The hb->lock must be held by the caller, and is released here. A call to queue_me is typically paired with exactly one call to unqueue_me. The exceptions involve the PI related operations, which may use unqueue_me_pi or nothing if the unqueue is done as part of the wake process and the unqueue state is implicit in the state of woken task (see futex_wait_requeue_pi for an example). Kernel Hackers Manual July 2017 unqueue_me 9 4.1.27 unqueue_me Remove the futex_q from its futex_hash_bucket int unqueue_me struct futex_q * q q The futex_q to unqueue The q->lock_ptr must not be held by the caller. A call to unqueue_me must be paired with exactly one earlier call to queue_me. 1 - if the futex_q was still queued (and we removed unqueued it); 0 - if the futex_q was already removed by the waking thread Kernel Hackers Manual July 2017 fixup_owner 9 4.1.27 fixup_owner Post lock pi_state and corner case management int fixup_owner u32 __user * uaddr struct futex_q * q int locked uaddr user address of the futex q futex_q (contains pi_state and access to the rt_mutex) locked if the attempt to take the rt_mutex succeeded (1) or not (0) After attempting to lock an rt_mutex, this function is called to cleanup the pi_state owner as well as handle race conditions that may allow us to acquire the lock. Must be called with the hb lock held. 1 - success, lock taken; 0 - success, lock not taken; <0 - on error (-EFAULT) Kernel Hackers Manual July 2017 futex_wait_queue_me 9 4.1.27 futex_wait_queue_me queue_me and wait for wakeup, timeout, or signal void futex_wait_queue_me struct futex_hash_bucket * hb struct futex_q * q struct hrtimer_sleeper * timeout hb the futex hash bucket, must be locked by the caller q the futex_q to queue up on timeout the prepared hrtimer_sleeper, or null for no timeout Kernel Hackers Manual July 2017 futex_wait_setup 9 4.1.27 futex_wait_setup Prepare to wait on a futex int futex_wait_setup u32 __user * uaddr u32 val unsigned int flags struct futex_q * q struct futex_hash_bucket ** hb uaddr the futex userspace address val the expected value flags futex flags (FLAGS_SHARED, etc.) q the associated futex_q hb storage for hash_bucket pointer to be returned to caller Setup the futex_q and locate the hash_bucket. Get the futex value and compare it with the expected value. Handle atomic faults internally. Return with the hb lock held and a q.key reference on success, and unlocked with no q.key reference on failure. 0 - uaddr contains val and hb has been locked; <1 - -EFAULT or -EWOULDBLOCK (uaddr does not contain val) and hb is unlocked Kernel Hackers Manual July 2017 handle_early_requeue_pi_wakeup 9 4.1.27 handle_early_requeue_pi_wakeup Detect early wakeup on the initial futex int handle_early_requeue_pi_wakeup struct futex_hash_bucket * hb struct futex_q * q union futex_key * key2 struct hrtimer_sleeper * timeout hb the hash_bucket futex_q was original enqueued on q the futex_q woken while waiting to be requeued key2 the futex_key of the requeue target futex timeout the timeout associated with the wait (NULL if none) Detect if the task was woken on the initial futex as opposed to the requeue target futex. If so, determine if it was a timeout or a signal that caused the wakeup and return the appropriate error code to the caller. Must be called with the hb lock held. 0 = no early wakeup detected; <0 = -ETIMEDOUT or -ERESTARTNOINTR Kernel Hackers Manual July 2017 futex_wait_requeue_pi 9 4.1.27 futex_wait_requeue_pi Wait on uaddr and take uaddr2 int futex_wait_requeue_pi u32 __user * uaddr unsigned int flags u32 val ktime_t * abs_time u32 bitset u32 __user * uaddr2 uaddr the futex we initially wait on (non-pi) flags futex flags (FLAGS_SHARED, FLAGS_CLOCKRT, etc.), they must be the same type, no requeueing from private to shared, etc. val the expected value of uaddr abs_time absolute timeout bitset 32 bit wakeup bitset set by userspace, defaults to all uaddr2 the pi futex we will take prior to returning to user-space The caller will wait on uaddr and will be requeued by futex_requeue to uaddr2 which must be PI aware and unique from uaddr. Normal wakeup will wake on uaddr2 and complete the acquisition of the rt_mutex prior to returning to userspace. This ensures the rt_mutex maintains an owner when it has waiters; without one, the pi logic would not know which task to boost/deboost, if there was a need to. We call schedule in futex_wait_queue_me when we enqueue and return there via the following-- 1) wakeup on uaddr2 after an atomic lock acquisition by futex_requeue 2) wakeup on uaddr2 after a requeue 3) signal 4) timeout If 3, cleanup and return -ERESTARTNOINTR. If 2, we may then block on trying to take the rt_mutex and return via: 5) successful lock 6) signal 7) timeout 8) other lock acquisition failure If 6, return -EWOULDBLOCK (restarting the syscall would do the same). If 4 or 7, we cleanup and return with -ETIMEDOUT. 0 - On success; <0 - On error Kernel Hackers Manual July 2017 sys_set_robust_list 9 4.1.27 sys_set_robust_list Set the robust-futex list head of a task long sys_set_robust_list struct robust_list_head __user * head size_t len head pointer to the list-head len length of the list-head, as userspace expects Kernel Hackers Manual July 2017 sys_get_robust_list 9 4.1.27 sys_get_robust_list Get the robust-futex list head of a task long sys_get_robust_list int pid struct robust_list_head __user *__user * head_ptr size_t __user * len_ptr pid pid of the process [zero for current task] head_ptr pointer to a list-head pointer, the kernel fills it in len_ptr pointer to a length field, the kernel fills in the header size Documentation/locking/spinlocks.txt: Linus Torvalds' spinlocking tutorial in the kernel sources. Unix Systems for Modern Architectures: Symmetric Multiprocessing and Caching for Kernel Programmers: Curt Schimmel's very good introduction to kernel level locking (not written for Linux, but nearly everything applies). The book is expensive, but really worth every penny to understand SMP locking. [ISBN: 0201633388] Thanks to Telsa Gwynne for DocBooking, neatening and adding style. Thanks to Martin Pool, Philipp Rumpf, Stephen Rothwell, Paul Mackerras, Ruedi Aschwanden, Alan Cox, Manfred Spraul, Tim Waugh, Pete Zaitcev, James Morris, Robert Love, Paul McKenney, John Ashby for proofreading, correcting, flaming, commenting. Thanks to the cabal for having no influence on this document. preemption Prior to 2.5, or when CONFIG_PREEMPT is unset, processes in user context inside the kernel would not preempt each other (ie. you had that CPU until you gave it up, except for interrupts). With the addition of CONFIG_PREEMPT in 2.5.4, this changed: when in user context, higher priority tasks can "cut in": spinlocks were changed to disable preemption, even on UP. bh Bottom Half: for historical reasons, functions with '_bh' in them often now refer to any software interrupt, e.g. spin_lock_bh() blocks any software interrupt on the current CPU. Bottom halves are deprecated, and will eventually be replaced by tasklets. Only one bottom half will be running at any time. Hardware Interrupt / Hardware IRQ Hardware interrupt request. in_irq() returns true in a hardware interrupt handler. Interrupt Context Not user context: processing a hardware irq or software irq. Indicated by the in_interrupt() macro returning true. SMP Symmetric Multi-Processor: kernels compiled for multiple-CPU machines. (CONFIG_SMP=y). Software Interrupt / softirq Software interrupt handler. in_irq() returns false; in_softirq() returns true. Tasklets and softirqs both fall into the category of 'software interrupts'. Strictly speaking a softirq is one of up to 32 enumerated software interrupts which can run on multiple CPUs at once. Sometimes used to refer to tasklets as well (ie. all software interrupts). tasklet A dynamically-registrable software interrupt, which is guaranteed to only run on one CPU at a time. timer A dynamically-registrable software interrupt, which is run at (or close to) a given time. When running, it is just like a tasklet (in fact, they are called from the TIMER_SOFTIRQ). UP Uni-Processor: Non-SMP. (CONFIG_SMP=n). User Context The kernel executing on behalf of a particular process (ie. a system call or trap) or kernel thread. You can tell which process with the current macro.) Not to be confused with userspace. Can be interrupted by software or hardware interrupts. Userspace A process executing its own code outside the kernel.