Every process under Linux is dynamically allocated a struct task_struct
structure. The maximum number of processes which can be created on Linux
is limited only by the amount of physical memory present, and is
equal to (see kernel/fork.c:fork_init()):
/*
* The default maximum number of threads is set to a safe
* value: the thread structures can take up at most half
* of memory.
*/
max_threads = mempages / (THREAD_SIZE/PAGE_SIZE) / 2;
which, on IA32 architecture, basically means num_physpages/4. As an example,
on a 512M machine, you can create 32k threads. This is a considerable
improvement over the 4k-epsilon limit for older (2.2 and earlier) kernels.
Moreover, this can be changed at runtime using the KERN_MAX_THREADS sysctl(2),
or simply using procfs interface to kernel tunables:
# cat /proc/sys/kernel/threads-max
32764
# echo 100000 > /proc/sys/kernel/threads-max
# cat /proc/sys/kernel/threads-max
100000
# gdb -q vmlinux /proc/kcore
Core was generated by `BOOT_IMAGE=240ac18 ro root=306 video=matrox:vesa:0x118'.
#0 0x0 in ?? ()
(gdb) p max_threads
$1 = 100000
The set of processes on the Linux system is represented as a collection of
struct task_struct structures which are linked in two ways:
p->next_task and p->prev_task
pointers.The hashtable is called pidhash[] and is defined in
include/linux/sched.h:
/* PID hashing. (shouldnt this be dynamic?) */
#define PIDHASH_SZ (4096 >> 2)
extern struct task_struct *pidhash[PIDHASH_SZ];
#define pid_hashfn(x) ((((x) >> 8) ^ (x)) & (PIDHASH_SZ - 1))
The tasks are hashed by their pid value and the above hashing function is
supposed to distribute the elements uniformly in their domain
(0 to PID_MAX-1). The hashtable is used to quickly find a task by given pid,
using find_task_pid() inline from include/linux/sched.h:
static inline struct task_struct *find_task_by_pid(int pid)
{
struct task_struct *p, **htable = &pidhash[pid_hashfn(pid)];
for(p = *htable; p && p->pid != pid; p = p->pidhash_next)
;
return p;
}
The tasks on each hashlist (i.e. hashed to the same value) are linked
by p->pidhash_next/pidhash_pprev which are used by hash_pid() and
unhash_pid() to insert and remove a given process into the hashtable.
These are done under protection of the read-write spinlock called tasklist_lock
taken for WRITE.
The circular doubly-linked list that uses p->next_task/prev_task is
maintained so that one could go through all tasks on the system easily.
This is achieved by the for_each_task() macro from include/linux/sched.h:
#define for_each_task(p) \
for (p = &init_task ; (p = p->next_task) != &init_task ; )
Users of for_each_task() should take tasklist_lock for READ.
Note that for_each_task() is using init_task to mark the beginning (and end)
of the list - this is safe because the idle task (pid 0) never exits.
The modifiers of the process hashtable or/and the process table links,
notably fork(), exit() and ptrace(), must take tasklist_lock for WRITE. What is
more interesting is that the writers must also disable interrupts on the
local CPU. The reason for this is not trivial: the send_sigio() function walks the
task list and thus takes tasklist_lock for READ, and it is called from
kill_fasync() in interrupt context. This is why writers must disable
interrupts while readers don't need to.
Now that we understand how the task_struct structures are linked together,
let us examine the members of task_struct. They loosely correspond to the
members of UNIX 'struct proc' and 'struct user' combined together.
The other versions of UNIX separated the task state information into one part which should be kept memory-resident at all times (called 'proc structure' which includes process state, scheduling information etc.) and another part which is only needed when the process is running (called 'u area' which includes file descriptor table, disk quota information etc.). The only reason for such ugly design was that memory was a very scarce resource. Modern operating systems (well, only Linux at the moment but others, e.g. FreeBSD seem to improve in this direction towards Linux) do not need such separation and therefore maintain process state in a kernel memory-resident data structure at all times.
The task_struct structure is declared in include/linux/sched.h and is
currently 1680 bytes in size.
The state field is declared as:
volatile long state; /* -1 unrunnable, 0 runnable, >0 stopped */
#define TASK_RUNNING 0
#define TASK_INTERRUPTIBLE 1
#define TASK_UNINTERRUPTIBLE 2
#define TASK_ZOMBIE 4
#define TASK_STOPPED 8
#define TASK_EXCLUSIVE 32
Why is TASK_EXCLUSIVE defined as 32 and not 16? Because 16 was used up by
TASK_SWAPPING and I forgot to shift TASK_EXCLUSIVE up when I removed
all references to TASK_SWAPPING (sometime in 2.3.x).
The volatile in p->state declaration means it can be modified
asynchronously (from interrupt handler):
TASK_RUNNING and placing it on the runqueue is not atomic. You need to hold
the runqueue_lock read-write spinlock for read in order to look at the
runqueue. If you do so, you will then see that every task on the runqueue is in
TASK_RUNNING state. However, the converse is not true for the reason explained
above. Similarly, drivers can mark themselves (or rather the process context they
run in) as TASK_INTERRUPTIBLE (or TASK_UNINTERRUPTIBLE) and then call schedule(),
which will then remove it from the runqueue (unless there is a pending signal, in which
case it is left on the runqueue). TASK_INTERRUPTIBLE, except it cannot
be woken up.wait()-ed for) by the parent (natural or by adoption).TASK_INTERRUPTIBLE or TASK_UNINTERRUPTIBLE.
This means that when
this task is sleeping on a wait queue with many other tasks, it will be
woken up alone instead of causing "thundering herd" problem by waking up all
the waiters.Task flags contain information about the process states which are not mutually exclusive:
unsigned long flags; /* per process flags, defined below */
/*
* Per process flags
*/
#define PF_ALIGNWARN 0x00000001 /* Print alignment warning msgs */
/* Not implemented yet, only for 486*/
#define PF_STARTING 0x00000002 /* being created */
#define PF_EXITING 0x00000004 /* getting shut down */
#define PF_FORKNOEXEC 0x00000040 /* forked but didn't exec */
#define PF_SUPERPRIV 0x00000100 /* used super-user privileges */
#define PF_DUMPCORE 0x00000200 /* dumped core */
#define PF_SIGNALED 0x00000400 /* killed by a signal */
#define PF_MEMALLOC 0x00000800 /* Allocating memory */
#define PF_VFORK 0x00001000 /* Wake up parent in mm_release */
#define PF_USEDFPU 0x00100000 /* task used FPU this quantum (SMP) */
The fields p->has_cpu, p->processor, p->counter, p->priority, p->policy and
p->rt_priority are related to the scheduler and will be looked at later.
The fields p->mm and p->active_mm point respectively to the process' address space
described by mm_struct structure and to the active address space if the
process doesn't have a real one (e.g. kernel threads). This helps minimise
TLB flushes on switching address spaces when the task is scheduled out.
So, if we are scheduling-in the kernel thread (which has no p->mm) then its
next->active_mm will be set to the prev->active_mm of the task that was
scheduled-out, which will be the same as prev->mm if prev->mm != NULL.
The address space can be shared between threads if CLONE_VM flag is passed
to the clone(2) system call or by means of vfork(2) system call.
The fields p->exec_domain and p->personality relate to the personality of
the task, i.e. to the way certain system calls behave in order to emulate the
"personality" of foreign flavours of UNIX.
The field p->fs contains filesystem information, which under Linux means
three pieces of information:
This structure also includes a reference count because it can be shared
between cloned tasks when CLONE_FS flag is passed to the clone(2) system
call.
The field p->files contains the file descriptor table. This too can be
shared between tasks, provided CLONE_FILES is specified with clone(2) system
call.
The field p->sig contains signal handlers and can be shared between cloned
tasks by means of CLONE_SIGHAND.
Different books on operating systems define a "process" in different ways, starting from "instance of a program in execution" and ending with "that which is produced by clone(2) or fork(2) system calls". Under Linux, there are three kinds of processes:
The idle thread is created at compile time for the first CPU; it is then
"manually" created for each CPU by means of arch-specific
fork_by_hand() in arch/i386/kernel/smpboot.c, which unrolls the fork(2) system
call by hand (on some archs). Idle tasks share one init_task structure but
have a private TSS structure, in the per-CPU array init_tss. Idle tasks all have
pid = 0 and no other task can share pid, i.e. use CLONE_PID flag to clone(2).
Kernel threads are created using kernel_thread() function which invokes
the clone(2) system call in kernel mode. Kernel threads usually have no user
address space, i.e. p->mm = NULL, because they explicitly do exit_mm(), e.g.
via daemonize() function. Kernel threads can always access kernel address
space directly. They are allocated pid numbers in the low range. Running at
processor's ring 0 (on x86, that is) implies that the kernel threads enjoy all I/O privileges
and cannot be pre-empted by the scheduler.
User tasks are created by means of clone(2) or fork(2) system calls, both of which internally invoke kernel/fork.c:do_fork().
Let us understand what happens when a user process makes a fork(2) system
call. Although fork(2) is architecture-dependent due to the
different ways of passing user stack and registers, the actual underlying
function do_fork() that does the job is portable and is located at
kernel/fork.c.
The following steps are done:
retval is set to -ENOMEM, as this is the value which errno
should be set to if fork(2) fails to allocate a new task structure.
CLONE_PID is set in clone_flags then return an error (-EPERM), unless
the caller is the idle thread (during boot only). So, normal user
threads cannot pass CLONE_PID to clone(2) and expect it to succeed.
For fork(2), this is irrelevant as clone_flags is set to SIFCHLD - this
is only relevant when do_fork() is invoked from sys_clone() which
passes the clone_flags from the value requested from userspace.
current->vfork_sem is initialised (it is later cleared in the child).
This is used by sys_vfork() (vfork(2) system call, corresponds to
clone_flags = CLONE_VFORK|CLONE_VM|SIGCHLD) to make the parent sleep
until the child does mm_release(), for example as a result of exec()ing
another program or exit(2)-ing.
alloc_task_struct() macro. On x86 it is just a gfp at GFP_KERNEL
priority. This is the first reason why fork(2) system call may sleep.
If this allocation fails, we return -ENOMEM.
*p = *current. Perhaps this
should be replaced by a memset? Later on, the fields that should not
be inherited by the child are set to the correct values.
RLIMIT_NPROC soft limit - if so, fail with -EAGAIN, if
not, increment the count of processes by given uid p->user->count.
-EAGAIN.
p->did_exec = 0)
p->swappable = 0)
p->state = TASK_UNINTERRUPTIBLE (TODO: why is this done?
I think it's not needed - get rid of it, Linus confirms it is not
needed)
p->flags are set according to the value of clone_flags;
for plain fork(2), this will be p->flags = PF_FORKNOEXEC.
p->pid is set using the fast algorithm in
kernel/fork.c:get_pid() (TODO: lastpid_lock spinlock can be made
redundant since get_pid() is always called under big kernel lock
from do_fork(), also remove flags argument of get_pid(), patch sent
to Alan on 20/06/2000 - followup later).
do_fork() initialises the rest of child's
task structure. At the very end, the child's task structure is
hashed into the pidhash hashtable and the child is woken up (TODO:
wake_up_process(p) sets p->state = TASK_RUNNING and adds the process
to the runq, therefore we probably didn't need to set p->state to
TASK_RUNNING earlier on in do_fork()). The interesting part is
setting p->exit_signal to clone_flags & CSIGNAL, which for fork(2)
means just SIGCHLD and setting p->pdeath_signal to 0. The
pdeath_signal is used when a process 'forgets' the original parent
(by dying) and can be set/get by means of PR_GET/SET_PDEATHSIG
commands of prctl(2) system call (You might argue that the way the
value of pdeath_signal is returned via userspace pointer argument
in prctl(2) is a bit silly - mea culpa, after Andries Brouwer
updated the manpage it was too late to fix ;)Thus tasks are created. There are several ways for tasks to terminate:
func == 1 (this is Linux-specific, for
compatibility with old distributions that still had the 'update'
line in /etc/inittab - nowadays the work of update is done by
kernel thread kupdate).Functions implementing system calls under Linux are prefixed with sys_,
but they are usually concerned only with argument checking or arch-specific
ways to pass some information and the actual work is done by do_ functions.
So it is with sys_exit() which calls do_exit() to do the work. Although,
other parts of the kernel sometimes invoke sys_exit() while they should really
call do_exit().
The function do_exit() is found in kernel/exit.c. The points to note about
do_exit():
schedule() at the end, which never returns.
TASK_ZOMBIE.
current->pdeath_signal, if not 0.
current->exit_signal, which is usually
equal to SIGCHLD.
The job of a scheduler is to arbitrate access to the current CPU between
multiple processes. The scheduler is implemented in the 'main kernel file'
kernel/sched.c. The corresponding header file include/linux/sched.h is
included (either explicitly or indirectly) by virtually every kernel source
file.
The fields of task structure relevant to scheduler include:
p->need_resched: this field is set if schedule() should be invoked at
the 'next opportunity'.
p->counter: number of clock ticks left to run in this scheduling
slice, decremented by a timer. When this field becomes lower than or equal to zero, it is reset
to 0 and p->need_resched is set. This is also sometimes called 'dynamic
priority' of a process because it can change by itself.
p->priority: the process' static priority, only changed through well-known
system calls like nice(2), POSIX.1b sched_setparam(2) or 4.4BSD/SVR4
setpriority(2).
p->rt_priority: realtime priority
p->policy: the scheduling policy, specifies which scheduling class the
task belongs to. Tasks can change their scheduling class using the
sched_setscheduler(2) system call. The valid values are SCHED_OTHER
(traditional UNIX process), SCHED_FIFO (POSIX.1b FIFO realtime
process) and SCHED_RR (POSIX round-robin realtime process). One can
also OR SCHED_YIELD to any of these values to signify that the process
decided to yield the CPU, for example by calling sched_yield(2) system
call. A FIFO realtime process will run until either a) it blocks on I/O,
b) it explicitly yields the CPU or c) it is preempted by another realtime
process with a higher p->rt_priority value. SCHED_RR is the same as
SCHED_FIFO, except that when its timeslice expires it goes back to
the end of the runqueue.The scheduler's algorithm is simple, despite the great apparent complexity
of the schedule() function. The function is complex because it implements
three scheduling algorithms in one and also because of the subtle
SMP-specifics.
The apparently 'useless' gotos in schedule() are there for a purpose - to
generate the best optimised (for i386) code. Also, note that scheduler
(like most of the kernel) was completely rewritten for 2.4, therefore the
discussion below does not apply to 2.2 or earlier kernels.
Let us look at the function in detail:
current->active_mm == NULL then something is wrong. Current
process, even a kernel thread (current->mm == NULL) must have a valid
p->active_mm at all times.
tq_scheduler task queue, process it
now. Task queues provide a kernel mechanism to schedule execution of
functions at a later time. We shall look at it in details elsewhere.
prev and this_cpu to current task and
current CPU respectively.
schedule() was invoked from interrupt handler (due to a bug)
and panic if so.
struct schedule_data *sched_data to point
to per-CPU (cacheline-aligned to prevent cacheline ping-pong)
scheduling data area, which contains the TSC value of last_schedule and the
pointer to last scheduled task structure (TODO: sched_data is used on
SMP only but why does init_idle() initialises it on UP as well?).
runqueue_lock spinlock is taken. Note that we use spin_lock_irq()
because in schedule() we guarantee that interrupts are enabled. Therefore,
when we unlock runqueue_lock, we can just re-enable them instead of
saving/restoring eflags (spin_lock_irqsave/restore variant).
TASK_RUNNING state, it is left
alone; if it is in TASK_INTERRUPTIBLE state and a signal is pending,
it is moved into TASK_RUNNING state. In all other cases, it is deleted
from the runqueue.
next (best candidate to be scheduled) is set to the idle task of
this cpu. However, the goodness of this candidate is set to a very
low value (-1000), in hope that there is someone better than that.
prev (current) task is in TASK_RUNNING state, then the
current goodness is set to its goodness and it is marked as a better
candidate to be scheduled than the idle task.
goodness(), which
treats realtime processes by making their goodness very high
(1000 + p->rt_priority), this being greater than 1000 guarantees that
no SCHED_OTHER process can win; so they only contend with other
realtime processes that may have a greater p->rt_priority. The
goodness function returns 0 if the process' time slice (p->counter)
is over. For non-realtime processes, the initial value of goodness is
set to p->counter - this way, the process is less likely to get CPU if
it already had it for a while, i.e. interactive processes are favoured
more than CPU bound number crunchers. The arch-specific constant
PROC_CHANGE_PENALTY attempts to implement "cpu affinity" (i.e. give
advantage to a process on the same CPU). It also gives a slight
advantage to processes with mm pointing to current active_mm or to
processes with no (user) address space, i.e. kernel threads.
recalculate:
{
struct task_struct *p;
spin_unlock_irq(&runqueue_lock);
read_lock(&tasklist_lock);
for_each_task(p)
p->counter = (p->counter >> 1) + p->priority;
read_unlock(&tasklist_lock);
spin_lock_irq(&runqueue_lock);
}
runqueue_lock before we recalculate. The
reason is that we go through entire set of processes; this can take
a long time, during which the schedule() could be called on another CPU and
select a process with goodness good enough for that CPU, whilst we on
this CPU were forced to recalculate. Ok, admittedly this is somewhat
inconsistent because while we (on this CPU) are selecting a process with
the best goodness, schedule() running on another CPU could be
recalculating dynamic priorities.
next points to the task to
be scheduled, so we initialise next->has_cpu to 1 and next->processor
to this_cpu. The runqueue_lock can now be unlocked.
next == prev) then we can
simply reacquire the global kernel lock and return, i.e. skip all the
hardware-level (registers, stack etc.) and VM-related (switch page
directory, recalculate active_mm etc.) stuff.
switch_to() is architecture specific. On i386, it is
concerned with a) FPU handling, b) LDT handling, c) reloading segment
registers, d) TSS handling and e) reloading debug registers.
Before we go on to examine implementation of wait queues, we must
acquaint ourselves with the Linux standard doubly-linked list implementation.
Wait queues (as well as everything else in Linux) make heavy use
of them and they are called in jargon "list.h implementation" because the
most relevant file is include/linux/list.h.
The fundamental data structure here is struct list_head:
struct list_head {
struct list_head *next, *prev;
};
#define LIST_HEAD_INIT(name) { &(name), &(name) }
#define LIST_HEAD(name) \
struct list_head name = LIST_HEAD_INIT(name)
#define INIT_LIST_HEAD(ptr) do { \
(ptr)->next = (ptr); (ptr)->prev = (ptr); \
} while (0)
#define list_entry(ptr, type, member) \
((type *)((char *)(ptr)-(unsigned long)(&((type *)0)->member)))
#define list_for_each(pos, head) \
for (pos = (head)->next; pos != (head); pos = pos->next)
The first three macros are for initialising an empty list by pointing both
next and prev pointers to itself. It is obvious from C syntactical
restrictions which ones should be used where - for example, LIST_HEAD_INIT()
can be used for structure's element initialisation in declaration, the second
can be used for static variable initialising declarations and the third can
be used inside a function.
The macro list_entry() gives access to individual list element, for example
(from fs/file_table.c:fs_may_remount_ro()):
struct super_block {
...
struct list_head s_files;
...
} *sb = &some_super_block;
struct file {
...
struct list_head f_list;
...
} *file;
struct list_head *p;
for (p = sb->s_files.next; p != &sb->s_files; p = p->next) {
struct file *file = list_entry(p, struct file, f_list);
do something to 'file'
}
A good example of the use of list_for_each() macro is in the scheduler where
we walk the runqueue looking for the process with highest goodness:
static LIST_HEAD(runqueue_head);
struct list_head *tmp;
struct task_struct *p;
list_for_each(tmp, &runqueue_head) {
p = list_entry(tmp, struct task_struct, run_list);
if (can_schedule(p)) {
int weight = goodness(p, this_cpu, prev->active_mm);
if (weight > c)
c = weight, next = p;
}
}
Here, p->run_list is declared as struct list_head run_list inside
task_struct structure and serves as anchor to the list. Removing an element
from the list and adding (to head or tail of the list) is done by
list_del()/list_add()/list_add_tail() macros. The examples below are adding
and removing a task from runqueue:
static inline void del_from_runqueue(struct task_struct * p)
{
nr_running--;
list_del(&p->run_list);
p->run_list.next = NULL;
}
static inline void add_to_runqueue(struct task_struct * p)
{
list_add(&p->run_list, &runqueue_head);
nr_running++;
}
static inline void move_last_runqueue(struct task_struct * p)
{
list_del(&p->run_list);
list_add_tail(&p->run_list, &runqueue_head);
}
static inline void move_first_runqueue(struct task_struct * p)
{
list_del(&p->run_list);
list_add(&p->run_list, &runqueue_head);
}
When a process requests the kernel to do something which is currently impossible but that may become possible later, the process is put to sleep and is woken up when the request is more likely to be satisfied. One of the kernel mechanisms used for this is called a 'wait queue'.
Linux implementation allows wake-on semantics using TASK_EXCLUSIVE flag.
With waitqueues, you can either use a well-known queue and then simply
sleep_on/sleep_on_timeout/interruptible_sleep_on/interruptible_sleep_on_timeout,
or you can define your own waitqueue and use add/remove_wait_queue to add and
remove yourself from it and wake_up/wake_up_interruptible to wake up
when needed.
An example of the first usage of waitqueues is interaction between the page
allocator (in mm/page_alloc.c:__alloc_pages()) and the kswapd kernel daemon (in
mm/vmscan.c:kswap()), by means of wait queue kswapd_wait, declared in
mm/vmscan.c; the kswapd daemon sleeps on this queue, and it is woken up
whenever the page allocator needs to free up some pages.
An example of autonomous waitqueue usage is interaction between
user process requesting data via read(2) system call and kernel running in
the interrupt context to supply the data. An interrupt handler might look
like (simplified drivers/char/rtc_interrupt()):
static DECLARE_WAIT_QUEUE_HEAD(rtc_wait);
void rtc_interrupt(int irq, void *dev_id, struct pt_regs *regs)
{
spin_lock(&rtc_lock);
rtc_irq_data = CMOS_READ(RTC_INTR_FLAGS);
spin_unlock(&rtc_lock);
wake_up_interruptible(&rtc_wait);
}
So, the interrupt handler obtains the data by reading from some
device-specific I/O port (CMOS_READ() macro turns into a couple outb/inb) and
then wakes up whoever is sleeping on the rtc_wait wait queue.
Now, the read(2) system call could be implemented as:
ssize_t rtc_read(struct file file, char *buf, size_t count, loff_t *ppos)
{
DECLARE_WAITQUEUE(wait, current);
unsigned long data;
ssize_t retval;
add_wait_queue(&rtc_wait, &wait);
current->state = TASK_INTERRUPTIBLE;
do {
spin_lock_irq(&rtc_lock);
data = rtc_irq_data;
rtc_irq_data = 0;
spin_unlock_irq(&rtc_lock);
if (data != 0)
break;
if (file->f_flags & O_NONBLOCK) {
retval = -EAGAIN;
goto out;
}
if (signal_pending(current)) {
retval = -ERESTARTSYS;
goto out;
}
schedule();
} while(1);
retval = put_user(data, (unsigned long *)buf);
if (!retval)
retval = sizeof(unsigned long);
out:
current->state = TASK_RUNNING;
remove_wait_queue(&rtc_wait, &wait);
return retval;
}
What happens in rtc_read() is this:
rtc_wait waitqueue.
TASK_INTERRUPTIBLE which means it will not
be rescheduled after the next time it sleeps.
TASK_RUNNING, remove
ourselves from the wait queue and return
EAGAIN (which is the same as EWOULDBLOCK)
TASK_INTERRUPTIBLE
then the scheduler could schedule us sooner than when the data is
available, thus causing unneeded processing.It is also worth pointing out that, using wait queues, it is rather easy to implement the poll(2) system call:
static unsigned int rtc_poll(struct file *file, poll_table *wait)
{
unsigned long l;
poll_wait(file, &rtc_wait, wait);
spin_lock_irq(&rtc_lock);
l = rtc_irq_data;
spin_unlock_irq(&rtc_lock);
if (l != 0)
return POLLIN | POLLRDNORM;
return 0;
}
All the work is done by the device-independent function poll_wait() which does
the necessary waitqueue manipulations; all we need to do is point it to the
waitqueue which is woken up by our device-specific interrupt handler.
Now let us turn our attention to kernel timers. Kernel timers are used to
dispatch execution of a particular function (called 'timer handler') at a
specified time in the future. The main data structure is struct timer_list
declared in include/linux/timer.h:
struct timer_list {
struct list_head list;
unsigned long expires;
unsigned long data;
void (*function)(unsigned long);
volatile int running;
};
The list field is for linking into the internal list, protected by the
timerlist_lock spinlock. The expires field is the value of jiffies when
the function handler should be invoked with data passed as a parameter.
The running field is used on SMP to test if the timer handler is currently
running on another CPU.
The functions add_timer() and del_timer() add and remove a given timer to the
list. When a timer expires, it is removed automatically. Before a timer is
used, it MUST be initialised by means of init_timer() function. And before it
is added, the fields function and expires must be set.
Sometimes it is reasonable to split the amount of work to be performed inside an interrupt handler into immediate work (e.g. acknowledging the interrupt, updating the stats etc.) and work which can be postponed until later, when interrupts are enabled (e.g. to do some postprocessing on data, wake up processes waiting for this data, etc).
Bottom halves are the oldest mechanism for deferred execution of kernel tasks and have been available since Linux 1.x. In Linux 2.0, a new mechanism was added, called 'task queues', which will be the subject of next section.
Bottom halves are serialised by the global_bh_lock spinlock, i.e.
there can only be one bottom half running on any CPU at a time. However,
when attempting to execute the handler, if global_bh_lock is not available,
the bottom half is marked (i.e. scheduled) for execution - so processing can
continue, as opposed to a busy loop on global_bh_lock.
There can only be 32 bottom halves registered in total. The functions required to manipulate bottom halves are as follows (all exported to modules):
void init_bh(int nr, void (*routine)(void)): installs a bottom half
handler pointed to by routine argument into slot nr. The slot
ought to be enumerated in include/linux/interrupt.h in the form
XXXX_BH, e.g. TIMER_BH or TQUEUE_BH. Typically, a subsystem's
initialisation routine (init_module() for modules) installs the
required bottom half using this function.
void remove_bh(int nr): does the opposite of init_bh(), i.e.
de-installs bottom half installed at slot nr. There is no error
checking performed there, so, for example remove_bh(32) will
panic/oops the system. Typically, a subsystem's cleanup routine
(cleanup_module() for modules) uses this function to free up the slot
that can later be reused by some other subsystem. (TODO: wouldn't it
be nice to have /proc/bottom_halves list all registered bottom
halves on the system? That means global_bh_lock must be made
read/write, obviously)
void mark_bh(int nr): marks bottom half in slot nr for execution. Typically,
an interrupt handler will mark its bottom half (hence the name!) for
execution at a "safer time".
Bottom halves are globally locked tasklets, so the question "when are bottom
half handlers executed?" is really "when are tasklets executed?". And the
answer is, in two places: a) on each schedule() and b) on each
interrupt/syscall return path in entry.S (TODO: therefore, the schedule()
case is really boring - it like adding yet another very very slow interrupt,
why not get rid of handle_softirq label from schedule() altogether?).
Task queues can be though of as a dynamic extension to old bottom halves. In fact, in the source code they are sometimes referred to as "new" bottom halves. More specifically, the old bottom halves discussed in previous section have these limitations:
So, with task queues, arbitrary number of functions can be chained and
processed one after another at a later time. One creates a new task queue
using the DECLARE_TASK_QUEUE() macro and queues a task onto it using
the queue_task() function. The task queue then can be processed using
run_task_queue(). Instead of creating your own task queue (and
having to consume it manually) you can use one of Linux' predefined
task queues which are consumed at well-known points:
tq_timer tasks also run in interrupt context and thus cannot block.
tq_timer). Since the scheduler executed
in the context of the process being re-scheduled, the tq_scheduler
tasks can do anything they like, i.e. block, use process context data
(but why would they want to), etc.
IMMEDIATE_BH, so
drivers can queue_task(task, &tq_immediate) and then
mark_bh(IMMEDIATE_BH) to be consumed in interrupt context.
Unless a driver uses its own task queues, it does not need to call
run_tasks_queues() to process the queue, except under circumstances explained
below.
The reason tq_timer/tq_scheduler task queues are consumed not only in the
usual places but elsewhere (closing tty device is but one example) becomes
clear if one remembers that the driver can schedule tasks on the queue, and these tasks
only make sense while a particular instance of the device is still valid
- which usually means until the application closes it. So, the driver may
need to call run_task_queue() to flush the tasks it (and anyone else) has
put on the queue, because allowing them to run at a later time may make no
sense - i.e. the relevant data structures may have been freed/reused by a
different instance. This is the reason you see run_task_queue() on tq_timer
and tq_scheduler in places other than timer interrupt and schedule()
respectively.
Not yet, will be in future revision.
Not yet, will be in future revision.
There are two mechanisms under Linux for implementing system calls:
Native Linux programs use int 0x80 whilst binaries from foreign flavours of UNIX (Solaris, UnixWare 7 etc.) use the lcall7 mechanism. The name 'lcall7' is historically misleading because it also covers lcall27 (e.g. Solaris/x86), but the handler function is called lcall7_func.
When the system boots, the function arch/i386/kernel/traps.c:trap_init() is
called which sets up the IDT so that vector 0x80 (of type 15, dpl 3) points to
the address of system_call entry from arch/i386/kernel/entry.S.
When a userspace application makes a system call, the arguments are passed via registers
and the application executes 'int 0x80' instruction. This causes a trap into
kernel mode and processor jumps to system_call entry point in entry.S.
What this does is:
NR_syscalls (currently 256),
fail with ENOSYS error.
tsk->ptrace & PF_TRACESYS), do special
processing. This is to support programs like strace (analogue of
SVR4 truss(1)) or debuggers.
sys_call_table+4*(syscall_number from %eax). This table is
initialised in the same file (arch/i386/kernel/entry.S) to point to
individual system call handlers which under Linux are (usually)
prefixed with sys_, e.g. sys_open, sys_exit, etc. These C system
call handlers will find their arguments on the stack where SAVE_ALL
stored them.
schedule() is needed (tsk->need_resched != 0), checking if there
are signals pending and if so handling them.Linux supports up to 6 arguments for system calls. They are passed in
%ebx, %ecx, %edx, %esi, %edi (and %ebp used temporarily, see _syscall6() in
asm-i386/unistd.h). The system call number is passed via %eax.
There are two types of atomic operations: bitmaps and atomic_t. Bitmaps are
very convenient for maintaining a concept of "allocated" or "free" units
from some large collection where each unit is identified by some number, for
example free inodes or free blocks. They are also widely used for simple
locking, for example to provide exclusive access to open a device. An example
of this can be found in arch/i386/kernel/microcode.c:
/*
* Bits in microcode_status. (31 bits of room for future expansion)
*/
#define MICROCODE_IS_OPEN 0 /* set if device is in use */
static unsigned long microcode_status;
There is no need to initialise microcode_status to 0 as BSS is zero-cleared
under Linux explicitly.
/*
* We enforce only one user at a time here with open/close.
*/
static int microcode_open(struct inode *inode, struct file *file)
{
if (!capable(CAP_SYS_RAWIO))
return -EPERM;
/* one at a time, please */
if (test_and_set_bit(MICROCODE_IS_OPEN, µcode_status))
return -EBUSY;
MOD_INC_USE_COUNT;
return 0;
}
The operations on bitmaps are:
nr
in the bitmap pointed to by addr.
nr in the bitmap pointed to by addr.
nr (if set clear, if clear set) in the bitmap pointed to by addr.
nr and return the old bit value.
nr and return the old bit value.
nr and return the old bit value.These operations use the LOCK_PREFIX macro, which on SMP kernels evaluates to
bus lock instruction prefix and to nothing on UP. This guarantees atomicity
of access in SMP environment.
Sometimes bit manipulations are not convenient, but instead we need to perform
arithmetic operations - add, subtract, increment decrement. The typical cases
are reference counts (e.g. for inodes). This facility is provided by the
atomic_t data type and the following operations:
atomic_t variable v.
atomic_t variable
v to integer i.
i to the value of atomic variable pointed to by v.
i from the value of atomic variable pointed to by v.
i from the value of atomic variable pointed to by
v; return 1 if the new value is 0, return 0 otherwise.
i to v and return 1 if the result is negative. Return
0 if the result is greater than or equal to 0. This operation is used
for implementing semaphores.
Since the early days of Linux support (early 90s, this century), developers were faced with the classical problem of accessing shared data between different types of context (user process vs interrupt) and different instances of the same context from multiple cpus.
SMP support was added to Linux 1.3.42 on 15 Nov 1995 (the original patch was made to 1.3.37 in October the same year).
If the critical region of code may be executed by either process context
and interrupt context, then the way to protect it using cli/sti instructions
on UP is:
unsigned long flags;
save_flags(flags);
cli();
/* critical code */
restore_flags(flags);
While this is ok on UP, it obviously is of no use on SMP because the same
code sequence may be executed simultaneously on another cpu, and while cli()
provides protection against races with interrupt context on each CPU individually, it
provides no protection at all against races between contexts running on different
CPUs. This is where spinlocks are useful for.
There are three types of spinlocks: vanilla (basic), read-write and
big-reader spinlocks. Read-write spinlocks should be used when there is a
natural tendency of 'many readers and few writers'. Example of this is
access to the list of registered filesystems (see fs/super.c). The list is
guarded by the file_systems_lock read-write spinlock because one needs exclusive
access only when registering/unregistering a filesystem, but any process can
read the file /proc/filesystems or use the sysfs(2) system call to force a
read-only scan of the file_systems list. This makes it sensible to use
read-write spinlocks. With read-write spinlocks, one can have multiple
readers at a time but only one writer and there can be no readers while
there is
a writer. Btw, it would be nice if new readers would not get a lock while
there
is a writer trying to get a lock, i.e. if Linux could correctly deal with
the issue of potential writer starvation by multiple readers.
This would mean that readers must be blocked while there is a writer
attempting to get the lock. This is not
currently the case and it is not obvious whether this should be fixed - the
argument to the contrary is - readers usually take the lock for a very short
time so should they really be starved while the writer takes the lock for
potentially longer periods?
Big-reader spinlocks are a form of read-write spinlocks heavily optimised for very light read access, with a penalty for writes. There is a limited number of big-reader spinlocks - currently only two exist, of which one is used only on sparc64 (global irq) and the other is used for networking. In all other cases where the access pattern does not fit into any of these two scenarios, one should use basic spinlocks. You cannot block while holding any kind of spinlock.
Spinlocks come in three flavours: plain, _irq() and _bh().
spin_lock()/spin_unlock(): if you know the interrupts are always
disabled or if you do not race with interrupt context (e.g. from
within interrupt handler), then you can use this one. It does not
touch interrupt state on the current CPU.
spin_lock_irq()/spin_unlock_irq(): if you know that interrupts are
always enabled then you can use this version, which simply disables
(on lock) and re-enables (on unlock) interrupts on the current CPU.
For example, rtc_read() uses
spin_lock_irq(&rtc_lock) (interrupts are always enabled inside
read()) whilst rtc_interrupt() uses
spin_lock(&rtc_lock) (interrupts are always disabled inside
interrupt handler). Note that rtc_read() uses spin_lock_irq() and not
the more generic spin_lock_irqsave() because on entry to any system
call interrupts are always enabled.
spin_lock_irqsave()/spin_unlock_irqrestore(): the strongest form,
to be used when the interrupt state is not known, but only if
interrupts matter at all, i.e. there is no point in using it if
our interrupt handlers don't execute any critical code.The reason you cannot use plain spin_lock() if you race against interrupt handlers is because if you take it and then
an interrupt comes in on the same CPU, it will busy wait for the lock forever:
the lock holder, having been interrupted, will not continue until the
interrupt handler returns.
The most common usage of a spinlock is to access a data structure shared between user process context and interrupt handlers:
spinlock_t my_lock = SPIN_LOCK_UNLOCKED;
my_ioctl()
{
spin_lock_irq(&my_lock);
/* critical section */
spin_unlock_irq(&my_lock);
}
my_irq_handler()
{
spin_lock(&lock);
/* critical section */
spin_unlock(&lock);
}
There are a couple of things to note about this example:
ioctl() (arguments and return values omitted for clarity), must
use spin_lock_irq() because it knows that interrupts are always
enabled while executing the device ioctl() method.
my_irq_handler() (again
arguments omitted for clarity) can use plain spin_lock() form because
interrupts are disabled inside an interrupt handler.
Sometimes, while accessing a shared data structure, one must perform operations that can block, for example copy data to userspace. The locking primitive available for such scenarios under Linux is called a semaphore. There are two types of semaphores: basic and read-write semaphores. Depending on the initial value of the semaphore, they can be used for either mutual exclusion (initial value of 1) or to provide more sophisticated type of access.
Read-write semaphores differ from basic semaphores in the same way as read-write spinlocks differ from basic spinlocks: one can have multiple readers at a time but only one writer and there can be no readers while there are writers - i.e. the writer blocks all readers and new readers block while a writer is waiting.
Also, basic semaphores can be interruptible - just use the operations
down/up_interruptible() instead of the plain down()/up() and check the
value returned from down_interruptible(): it will be non zero if the operation was
interrupted.
Using semaphores for mutual exclusion is ideal in situations where a critical code section may call by reference unknown functions registered by other subsystems/modules, i.e. the caller cannot know apriori whether the function blocks or not.
A simple example of semaphore usage is in kernel/sys.c, implementation of
gethostname(2)/sethostname(2) system calls.
asmlinkage long sys_sethostname(char *name, int len)
{
int errno;
if (!capable(CAP_SYS_ADMIN))
return -EPERM;
if (len < 0 || len > __NEW_UTS_LEN)
return -EINVAL;
down_write(&uts_sem);
errno = -EFAULT;
if (!copy_from_user(system_utsname.nodename, name, len)) {
system_utsname.nodename[len] = 0;
errno = 0;
}
up_write(&uts_sem);
return errno;
}
asmlinkage long sys_gethostname(char *name, int len)
{
int i, errno;
if (len < 0)
return -EINVAL;
down_read(&uts_sem);
i = 1 + strlen(system_utsname.nodename);
if (i > len)
i = len;
errno = 0;
if (copy_to_user(name, system_utsname.nodename, i))
errno = -EFAULT;
up_read(&uts_sem);
return errno;
}
The points to note about this example are:
copy_from_user()/copy_to_user(). Therefore they could not use any form
of spinlock here.
Although Linux implementation of semaphores and read-write semaphores is very sophisticated, there are possible scenarios one can think of which are not yet implemented, for example there is no concept of interruptible read-write semaphores. This is obviously because there are no real-world situations which require these exotic flavours of the primitives.
Linux is a monolithic operating system and despite all the modern hype about some "advantages" offered by operating systems based on micro-kernel design, the truth remains (quoting Linus Torvalds himself):
... message passing as the fundamental operation of the OS is just an
exercise in computer science masturbation. It may feel good, but you
don't actually get anything DONE.
Therefore, Linux is and will always be based on a monolithic design, which means that all subsystems run in the same privileged mode and share the same address space; communication between them is achieved by the usual C function call means.
However, although separating kernel functionality into separate "processes"
as is done in micro-kernels is definitely a bad idea, separating it into
dynamically loadable on demand kernel modules is desirable in some
circumstances (e.g. on machines with low memory or for installation kernels
which could otherwise contain ISA auto-probing device drivers that are
mutually exclusive). The decision whether to include support for loadable
modules is made at compile time and is determined by the CONFIG_MODULES
option. Support for module autoloading via request_module() mechanism is
a separate compilation option (CONFIG_KMOD).
The following functionality can be implemented as loadable modules under Linux:
/proc and in devfs (e.g. /dev/cpu/microcode
vs /dev/misc/microcode).
There a few things that cannot be implemented as modules under Linux (probably because it makes no sense for them to be modularised):
Linux provides several system calls to assist in loading modules:
caddr_t create_module(const char *name, size_t size): allocates
size bytes using vmalloc() and maps a module structure at the
beginning thereof. This new module is then linked into the list headed
by module_list. Only a process with CAP_SYS_MODULE can invoke this
system call, others will get EPERM returned.
long init_module(const char *name, struct module *image): loads the
relocated module image and causes the module's initialisation routine
to be invoked. Only a process with CAP_SYS_MODULE can invoke this
system call, others will get EPERM returned.
long delete_module(const char *name): attempts to unload the module.
If name == NULL, attempt is made to unload all unused modules.
long query_module(const char *name, int which, void *buf, size_t bufsize, size_t *ret): returns information about a module
(or about all modules).The command interface available to users consists of:
Apart from being able to load a module manually using either insmod or modprobe,
it is also possible to have the module inserted automatically by the kernel
when a particular functionality is required. The kernel interface for this
is the function called request_module(name) which is exported to modules,
so that modules can load other modules as well. The request_module(name)
internally creates a kernel thread which execs the userspace command
modprobe -s -k module_name, using the standard exec_usermodehelper() kernel
interface (which is also exported to modules). The function returns 0 on
success, however it is usually not worth checking the return code from
request_module(). Instead, the programming idiom is:
if (check_some_feature() == NULL)
request_module(module);
if (check_some_feature() == NULL)
return -ENODEV;
For example, this is done by fs/block_dev.c:get_blkfops() to load a module
block-major-N when attempt is made to open a block device with major N.
Obviously, there is no such module called block-major-N (Linux developers
only chose sensible names for their modules) but it is mapped to a proper
module name using the file /etc/modules.conf. However, for most well-known
major numbers (and other kinds of modules) the modprobe/insmod commands
know which real module to load without needing an explicit alias statement
in /etc/modules.conf.
A good example of loading a module is inside the mount(2) system call. The
mount(2) system call accepts the filesystem type as a string which
fs/super.c:do_mount() then passes on to fs/super.c:get_fs_type():
static struct file_system_type *get_fs_type(const char *name)
{
struct file_system_type *fs;
read_lock(&file_systems_lock);
fs = *(find_filesystem(name));
if (fs && !try_inc_mod_count(fs->owner))
fs = NULL;
read_unlock(&file_systems_lock);
if (!fs && (request_module(name) == 0)) {
read_lock(&file_systems_lock);
fs = *(find_filesystem(name));
if (fs && !try_inc_mod_count(fs->owner))
fs = NULL;
read_unlock(&file_systems_lock);
}
return fs;
}
A few things to note in this function:
file_systems_lock taken for read (as we are not modifying the list
of registered filesystems).
try_inc_mod_count() returned 0 then
we consider it a failure - i.e. if the module is there but is being
deleted, it is as good as if it were not there at all.
file_systems_lock because what we are about to do next
(request_module()) is a blocking operation, and therefore we can't
hold a spinlock over it. Actually, in this specific case, we would
have to drop file_systems_lock anyway, even if request_module() were
guaranteed to be non-blocking and the module loading were executed
in the same context atomically. The reason for this is that the module's
initialisation function will try to call register_filesystem(), which will
take the same file_systems_lock read-write spinlock for write.
file_systems_lock spinlock and try to locate the newly registered
filesystem in the list. Note that this is slightly wrong because
it is in principle possible for a bug in modprobe command to cause
it to coredump after it successfully loaded the requested module, in
which case request_module() will fail even though the new filesystem will be
registered, and yet get_fs_type() won't find it.
When a module is loaded into the kernel, it can refer to any symbols that
are exported as public by the kernel using EXPORT_SYMBOL() macro or by
other currently loaded modules. If the module uses symbols from another
module, it is marked as depending on that module during dependency
recalculation, achieved by running depmod -a command on boot (e.g. after
installing a new kernel).
Usually, one must match the set of modules with the version of the kernel
interfaces they use, which under Linux simply means the "kernel version" as
there is no special kernel interface versioning mechanism in general.
However, there is a limited functionality called "module versioning" or
CONFIG_MODVERSIONS which allows to avoid recompiling modules when switching
to a new kernel. What happens here is that the kernel symbol table is treated
differently for internal access and for access from modules. The elements of
public (i.e. exported) part of the symbol table are built by 32bit
checksumming the C declaration. So, in order to resolve a symbol used by a
module during loading, the loader must match the full representation of the
symbol that includes the checksum; it will refuse to load the module if these
symbols differ. This
only happens when both the kernel and the module are compiled with module
versioning enabled. If either one of them uses the original symbol names,
the loader simply tries to match the kernel version declared by the module
and the one exported by the kernel and refuses to load if they differ.