tigran@veritas.com(quintela@fi.udc.es),
Francis Galiegue (fg@mandrakesoft.com),
Hakjun Mun (juniorm@orgio.net),
Matt Kraai (kraai@alumni.carnegiemellon.edu),
Nicholas Dronen (ndronen@frii.com),
Samuel S Chessman (chessman@tux.org),
Nadeem Hasan (nhasan@nadmm.com),
Michael Svetlik (m.svetlik@ssi-schaefer-peem.com)
for various corrections and suggestions.
The Linux Page Cache chapter was written by: Christoph Hellwig (hch@caldera.de).
The IPC Mechanisms chapter was written by: Russell Weight (weightr@us.ibm.com) and Mingming Cao (mcao@us.ibm.com)
This section explains the steps taken during compilation of the Linux kernel and the output produced at each stage. The build process depends on the architecture so I would like to emphasize that we only consider building a Linux/x86 kernel.
When the user types 'make zImage' or 'make bzImage' the resulting bootable
kernel image is stored as
arch/i386/boot/zImage or
arch/i386/boot/bzImage respectively.
Here is how the image is built:
vmlinux which is a
statically linked, non-stripped ELF 32-bit LSB 80386 executable file.
System.map is produced by nm vmlinux, irrelevant or uninteresting
symbols are grepped out.
arch/i386/boot.
bootsect.S is preprocessed either with or without
-D__BIG_KERNEL__, depending on whether the target is
bzImage or zImage, into bbootsect.s or bootsect.s respectively.
bbootsect.s is assembled and then converted into 'raw binary' form
called bbootsect (or bootsect.s assembled and raw-converted into
bootsect for zImage).
setup.S (setup.S includes video.S) is preprocessed into
bsetup.s for bzImage or setup.s for zImage. In the same way as the
bootsector code, the difference is marked by -D__BIG_KERNEL__ present
for bzImage. The result is then converted into 'raw binary' form
called bsetup.
arch/i386/boot/compressed and convert
/usr/src/linux/vmlinux to $tmppiggy (tmp filename) in raw binary
format, removing .note and .comment ELF sections.
piggy.o.
head.S and misc.c (still in
arch/i386/boot/compressed directory) into ELF objects head.o and
misc.o.
head.o, misc.o and piggy.o into bvmlinux (or vmlinux for
zImage, don't mistake this for /usr/src/linux/vmlinux!). Note the
difference between -Ttext 0x1000 used for vmlinux and -Ttext 0x100000
for bvmlinux, i.e. for bzImage compression loader is high-loaded.
bvmlinux to 'raw binary' bvmlinux.out removing .note and
.comment ELF sections.
arch/i386/boot directory and, using the program tools/build,
cat together bbootsect, bsetup and compressed/bvmlinux.out into bzImage
(delete extra 'b' above for zImage). This writes important variables
like setup_sects and root_dev at the end of the bootsector.0x4000 bytes >= 512 + setup_sects * 512 + room for stack while running bootsector/setup
We will see later where this limitation comes from.
The upper limit on the bzImage size produced at this step is about 2.5M for booting with LILO and 0xFFFF paragraphs (0xFFFF0 = 1048560 bytes) for booting raw image, e.g. from floppy disk or CD-ROM (El-Torito emulation mode).
Note that while tools/build does validate the size of boot sector, kernel image
and lower bound of setup size, it does not check the *upper* bound of said
setup size. Therefore it is easy to build a broken kernel by just adding some
large ".space" at the end of setup.S.
The boot process details are architecture-specific, so we shall focus our attention on the IBM PC/IA32 architecture. Due to old design and backward compatibility, the PC firmware boots the operating system in an old-fashioned manner. This process can be separated into the following six logical stages:
The bootsector used to boot Linux kernel could be either:
arch/i386/boot/bootsect.S),We consider here the Linux bootsector in detail. The first few lines initialise the convenience macros to be used for segment values:
29 SETUPSECS = 4 /* default nr of setup-sectors */
30 BOOTSEG = 0x07C0 /* original address of boot-sector */
31 INITSEG = DEF_INITSEG /* we move boot here - out of the way */
32 SETUPSEG = DEF_SETUPSEG /* setup starts here */
33 SYSSEG = DEF_SYSSEG /* system loaded at 0x10000 (65536) */
34 SYSSIZE = DEF_SYSSIZE /* system size: # of 16-byte clicks */
(the numbers on the left are the line numbers of bootsect.S file)
The values of DEF_INITSEG, DEF_SETUPSEG, DEF_SYSSEG and DEF_SYSSIZE are taken
from include/asm/boot.h:
/* Don't touch these, unless you really know what you're doing. */
#define DEF_INITSEG 0x9000
#define DEF_SYSSEG 0x1000
#define DEF_SETUPSEG 0x9020
#define DEF_SYSSIZE 0x7F00
Now, let us consider the actual code of bootsect.S:
54 movw $BOOTSEG, %ax
55 movw %ax, %ds
56 movw $INITSEG, %ax
57 movw %ax, %es
58 movw $256, %cx
59 subw %si, %si
60 subw %di, %di
61 cld
62 rep
63 movsw
64 ljmp $INITSEG, $go
65 # bde - changed 0xff00 to 0x4000 to use debugger at 0x6400 up (bde). We
66 # wouldn't have to worry about this if we checked the top of memory. Also
67 # my BIOS can be configured to put the wini drive tables in high memory
68 # instead of in the vector table. The old stack might have clobbered the
69 # drive table.
70 go: movw $0x4000-12, %di # 0x4000 is an arbitrary value >=
71 # length of bootsect + length of
72 # setup + room for stack;
73 # 12 is disk parm size.
74 movw %ax, %ds # ax and es already contain INITSEG
75 movw %ax, %ss
76 movw %di, %sp # put stack at INITSEG:0x4000-12.
Lines 54-63 move the bootsector code from address 0x7C00 to 0x90000. This is achieved by:
The reason this code does not use rep movsd is intentional (hint - .code16).
Line 64 jumps to label go: in the newly made copy of the
bootsector, i.e. in segment 0x9000. This and the following three
instructions (lines 64-76) prepare the stack at $INITSEG:0x4000-0xC, i.e.
%ss = $INITSEG (0x9000) and %sp = 0x3FF4 (0x4000-0xC). This is where the
limit on setup size comes from that we mentioned earlier (see Building the
Linux Kernel Image).
Lines 77-103 patch the disk parameter table for the first disk to allow multi-sector reads:
77 # Many BIOS's default disk parameter tables will not recognise
78 # multi-sector reads beyond the maximum sector number specified
79 # in the default diskette parameter tables - this may mean 7
80 # sectors in some cases.
81 #
82 # Since single sector reads are slow and out of the question,
83 # we must take care of this by creating new parameter tables
84 # (for the first disk) in RAM. We will set the maximum sector
85 # count to 36 - the most we will encounter on an ED 2.88.
86 #
87 # High doesn't hurt. Low does.
88 #
89 # Segments are as follows: ds = es = ss = cs - INITSEG, fs = 0,
90 # and gs is unused.
91 movw %cx, %fs # set fs to 0
92 movw $0x78, %bx # fs:bx is parameter table address
93 pushw %ds
94 ldsw %fs:(%bx), %si # ds:si is source
95 movb $6, %cl # copy 12 bytes
96 pushw %di # di = 0x4000-12.
97 rep # don't need cld -> done on line 66
98 movsw
99 popw %di
100 popw %ds
101 movb $36, 0x4(%di) # patch sector count
102 movw %di, %fs:(%bx)
103 movw %es, %fs:2(%bx)
The floppy disk controller is reset using BIOS service int 0x13 function 0 (reset FDC) and setup sectors are loaded immediately after the bootsector, i.e. at physical address 0x90200 ($INITSEG:0x200), again using BIOS service int 0x13, function 2 (read sector(s)). This happens during lines 107-124:
107 load_setup:
108 xorb %ah, %ah # reset FDC
109 xorb %dl, %dl
110 int $0x13
111 xorw %dx, %dx # drive 0, head 0
112 movb $0x02, %cl # sector 2, track 0
113 movw $0x0200, %bx # address = 512, in INITSEG
114 movb $0x02, %ah # service 2, "read sector(s)"
115 movb setup_sects, %al # (assume all on head 0, track 0)
116 int $0x13 # read it
117 jnc ok_load_setup # ok - continue
118 pushw %ax # dump error code
119 call print_nl
120 movw %sp, %bp
121 call print_hex
122 popw %ax
123 jmp load_setup
124 ok_load_setup:
If loading setup_sects sectors of setup code succeeded we jump to label
ok_load_setup:.
Then we proceed to load the compressed kernel image at physical
address 0x10000. This
is done to preserve the firmware data areas in low memory (0-64K).
After the kernel is loaded, we jump to $SETUPSEG:0 (arch/i386/boot/setup.S).
Once the data is no longer needed (e.g. no more calls to BIOS) it is
overwritten by moving the entire (compressed) kernel image from 0x10000 to
0x1000 (physical addresses, of course).
This is done by setup.S which sets things up for protected mode and jumps
to 0x1000 which is the head of the compressed kernel, i.e.
arch/386/boot/compressed/{head.S,misc.c}.
This sets up stack and calls decompress_kernel() which uncompresses the
kernel to address 0x100000 and jumps to it.
Note that old bootloaders (old versions of LILO) could only load the first 4 sectors of setup, which is why there is code in setup to load the rest of itself if needed. Also, the code in setup has to take care of various combinations of loader type/version vs zImage/bzImage and is therefore highly complex.
Let us examine the kludge in the bootsector code that allows to load a big
kernel, known also as "bzImage".
The setup sectors are loaded as usual at 0x90200, but the kernel is loaded
64K chunk at a time using a special helper routine that calls BIOS to move
data from low to high memory. This helper routine is referred to by
bootsect_kludge in bootsect.S and is defined as bootsect_helper in setup.S.
The bootsect_kludge label in setup.S contains the value of setup segment
and the offset of bootsect_helper code in it so that bootsector can use the lcall
instruction to jump to it (inter-segment jump).
The reason why it is in setup.S is simply because there is no more space left
in bootsect.S (which is strictly not true - there are approximately 4 spare bytes
and at least 1 spare byte in bootsect.S but that is not enough, obviously).
This routine uses BIOS service int 0x15 (ax=0x8700) to move to high memory
and resets %es to always point to 0x10000. This ensures that the code in bootsect.S
doesn't run out of low memory when copying data from disk.
There are several advantages in using a specialised bootloader (LILO) over a bare bones Linux bootsector:
The last thing LILO does is to jump to setup.S and things proceed as normal.
By "high-level initialisation" we consider anything which is not directly
related to bootstrap, even though parts of the code to perform this are
written in asm, namely arch/i386/kernel/head.S which is the head of the
uncompressed kernel. The following steps are performed:
start_kernel(), all others call
arch/i386/kernel/smpboot.c:initialize_secondary() if ready=1,
which just reloads esp/eip and doesn't return.The init/main.c:start_kernel() is written in C and does the following:
kmem_cache_init(), initialise most of slab allocator.mem_init() which calculates max_mapnr, totalram_pages and
high_memory and prints out the "Memory: ..." line.kmem_cache_sizes_init(), finish slab allocator initialisation.fork_init(), create uid_cache, initialise max_threads based on
the amount of memory available and configure RLIMIT_NPROC for
init_task to be max_threads/2.init() which execs
execute_command if supplied via "init=" boot parameter, or tries to
exec /sbin/init, /etc/init, /bin/init, /bin/sh in this order; if
all these fail, panic with "suggestion" to use "init=" parameter.Important thing to note here that the init() kernel thread calls
do_basic_setup() which in turn calls do_initcalls() which goes through the
list of functions registered by means of __initcall or module_init() macros
and invokes them. These functions either do not depend on each other
or their dependencies have been manually fixed by the link order in the
Makefiles. This means that, depending on
the position of directories in the trees and the structure of the Makefiles,
the order in which initialisation functions are invoked can change. Sometimes, this
is important because you can imagine two subsystems A and B with B depending
on some initialisation done by A. If A is compiled statically and B is a
module then B's entry point is guaranteed to be invoked after A prepared
all the necessary environment. If A is a module, then B is also necessarily
a module so there are no problems. But what if both A and B are statically
linked into the kernel? The order in which they are invoked depends on the relative
entry point offsets in the .initcall.init ELF section of the kernel image.
Rogier Wolff proposed to introduce a hierarchical "priority" infrastructure
whereby modules could let the linker know in what (relative) order they
should be linked, but so far there are no patches available that implement
this in a sufficiently elegant manner to be acceptable into the kernel.
Therefore, make sure your link order is correct. If, in the example above,
A and B work fine when compiled statically once, they will always work,
provided they are listed sequentially in the same Makefile. If they don't
work, change the order in which their object files are listed.
Another thing worth noting is Linux's ability to execute an "alternative
init program" by means of passing "init=" boot commandline. This is useful
for recovering from accidentally overwritten /sbin/init or debugging the
initialisation (rc) scripts and /etc/inittab by hand, executing them
one at a time.
On SMP, the BP goes through the normal sequence of bootsector, setup etc
until it reaches the start_kernel(), and then on to smp_init() and
especially src/i386/kernel/smpboot.c:smp_boot_cpus(). The smp_boot_cpus()
goes in a loop for each apicid (until NR_CPUS) and calls do_boot_cpu() on
it. What do_boot_cpu() does is create (i.e. fork_by_hand) an idle task for
the target cpu and write in well-known locations defined by the Intel MP
spec (0x467/0x469) the EIP of trampoline code found in trampoline.S. Then
it generates STARTUP IPI to the target cpu which makes this AP execute the
code in trampoline.S.
The boot CPU creates a copy of trampoline code for each CPU in low memory. The AP code writes a magic number in its own code which is verified by the BP to make sure that AP is executing the trampoline code. The requirement that trampoline code must be in low memory is enforced by the Intel MP specification.
The trampoline code simply sets %bx register to 1, enters protected mode
and jumps to startup_32 which is the main entry to arch/i386/kernel/head.S.
Now, the AP starts executing head.S and discovering that it is not a BP,
it skips the code that clears BSS and then enters initialize_secondary()
which just enters the idle task for this CPU - recall that init_tasks[cpu]
was already initialised by BP executing do_boot_cpu(cpu).
Note that init_task can be shared but each idle thread must have its own
TSS. This is why init_tss[NR_CPUS] is an array.
When the operating system initialises itself, most of the code and data structures are never needed again. Most operating systems (BSD, FreeBSD etc.) cannot dispose of this unneeded information, thus wasting precious physical kernel memory. The excuse they use (see McKusick's 4.4BSD book) is that "the relevant code is spread around various subsystems and so it is not feasible to free it". Linux, of course, cannot use such excuses because under Linux "if something is possible in principle, then it is already implemented or somebody is working on it".
So, as I said earlier, Linux kernel can only be compiled as an ELF binary, and now we find out the reason (or one of the reasons) for that. The reason related to throwing away initialisation code/data is that Linux provides two macros to be used:
__init - for initialisation code__initdata - for dataThese evaluate to gcc attribute specificators (also known as "gcc magic")
as defined in include/linux/init.h:
#ifndef MODULE
#define __init __attribute__ ((__section__ (".text.init")))
#define __initdata __attribute__ ((__section__ (".data.init")))
#else
#define __init
#define __initdata
#endif
What this means is that if the code is compiled statically into the kernel
(i.e. MODULE is not defined) then it is placed in the special ELF section
.text.init, which is declared in the linker map in arch/i386/vmlinux.lds.
Otherwise (i.e. if it is a module) the macros evaluate to nothing.
What happens during boot is that the "init" kernel thread (function
init/main.c:init()) calls the arch-specific function free_initmem() which
frees all the pages between addresses __init_begin and __init_end.
On a typical system (my workstation), this results in freeing about 260K of memory.
The functions registered via module_init() are placed in .initcall.init
which is also freed in the static case. The current trend in Linux, when
designing a subsystem (not necessarily a module), is to provide
init/exit entry points from the early stages of design so that in the
future, the subsystem in question can be modularised if needed. Example of
this is pipefs, see fs/pipe.c. Even if a given subsystem will never become a
module, e.g. bdflush (see fs/buffer.c), it is still nice and tidy to use
the module_init() macro against its initialisation function, provided it does
not matter when exactly is the function called.
There are two more macros which work in a similar manner, called __exit and
__exitdata, but they are more directly connected to the module support and
therefore will be explained in a later section.
Let us recall what happens to the commandline passed to kernel during boot:
arch/i386/kernel/head.S copies the first 2k of it out to the zeropage.
arch/i386/kernel/setup.c:parse_mem_cmdline() (called by
setup_arch(), itself called by start_kernel()) copies 256 bytes from zeropage
into saved_command_line which is displayed by /proc/cmdline. This
same routine processes the "mem=" option if present and makes appropriate
adjustments to VM parameters.
parse_options() (called by start_kernel())
which processes some "in-kernel" parameters (currently "init=" and
environment/arguments for init) and passes each word to checksetup().
checksetup() goes through the code in ELF section .setup.init and
invokes each function, passing it the word if it matches. Note that
using the return value of 0 from the function registered via __setup(),
it is possible to pass the same "variable=value" to more than one
function with "value" invalid to one and valid to another.
Jeff Garzik commented: "hackers who do that get spanked :)"
Why? Because this is clearly ld-order specific, i.e. kernel linked
in one order will have functionA invoked before functionB and another
will have it in reversed order, with the result depending on the order.
So, how do we write code that processes boot commandline? We use the __setup()
macro defined in include/linux/init.h:
/*
* Used for kernel command line parameter setup
*/
struct kernel_param {
const char *str;
int (*setup_func)(char *);
};
extern struct kernel_param __setup_start, __setup_end;
#ifndef MODULE
#define __setup(str, fn) \
static char __setup_str_##fn[] __initdata = str; \
static struct kernel_param __setup_##fn __initsetup = \
{ __setup_str_##fn, fn }
#else
#define __setup(str,func) /* nothing */
endif
So, you would typically use it in your code like this
(taken from code of real driver, BusLogic HBA drivers/scsi/BusLogic.c):
static int __init
BusLogic_Setup(char *str)
{
int ints[3];
(void)get_options(str, ARRAY_SIZE(ints), ints);
if (ints[0] != 0) {
BusLogic_Error("BusLogic: Obsolete Command Line Entry "
"Format Ignored\n", NULL);
return 0;
}
if (str == NULL || *str == '\0')
return 0;
return BusLogic_ParseDriverOptions(str);
}
__setup("BusLogic=", BusLogic_Setup);
Note that __setup() does nothing for modules, so the code that wishes to
process boot commandline and can be either a module or statically linked
must invoke its parsing function manually in the module initialisation
routine. This also means that it is possible to write code that
processes parameters when compiled as a module but not when it is static or
vice versa.
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 memcpy? 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.
In order to support multiple filesystems, Linux contains a special kernel interface level called VFS (Virtual Filesystem Switch). This is similar to the vnode/vfs interface found in SVR4 derivatives (originally it came from BSD and Sun original implementations).
Linux inode cache is implemented in a single file, fs/inode.c, which consists
of 977 lines of code. It is interesting to note that not many changes have been
made to it for the last 5-7 years: one can still recognise some of the code
comparing the latest version with, say, 1.3.42.
The structure of Linux inode cache is as follows:
inode_hashtable, where each inode is hashed by the
value of the superblock pointer and 32bit inode number. Inodes without a
superblock (inode->i_sb == NULL) are added to a doubly linked list
headed by anon_hash_chain instead. Examples of anonymous inodes
are sockets created by net/socket.c:sock_alloc(), by calling
fs/inode.c:get_empty_inode().
inode_in_use), which contains valid inodes
with i_count>0 and i_nlink>0. Inodes newly allocated by
get_empty_inode() and get_new_inode() are added to the inode_in_use list.
inode_unused), which contains valid inodes
with i_count = 0.
sb->s_dirty) which contains valid
inodes with i_count>0, i_nlink>0 and i_state & I_DIRTY.
When inode is marked
dirty, it is added to the sb->s_dirty list if it is also hashed.
Maintaining a per-superblock dirty list of inodes allows to quickly
sync inodes.
inode_cachep. As inode
objects are allocated and freed, they are taken from and returned to
this SLAB cache.The type lists are anchored from inode->i_list, the hashtable from
inode->i_hash. Each inode can be on a hashtable and one and only one type
(in_use, unused or dirty) list.
All these lists are protected by a single spinlock: inode_lock.
The inode cache subsystem is initialised when inode_init() function is called from
init/main.c:start_kernel(). The function is marked as __init, which means
its code is thrown away later on. It is passed a single argument - the
number of physical pages on the system. This is so that the inode cache can
configure itself depending on how much memory is available, i.e. create
a larger hashtable if there is enough memory.
The only stats information about inode cache is the number of unused inodes,
stored in inodes_stat.nr_unused and accessible to user programs via files
/proc/sys/fs/inode-nr and /proc/sys/fs/inode-state.
We can examine one of the lists from gdb running on a live kernel thus:
(gdb) printf "%d\n", (unsigned long)(&((struct inode *)0)->i_list)
8
(gdb) p inode_unused
$34 = 0xdfa992a8
(gdb) p (struct list_head)inode_unused
$35 = {next = 0xdfa992a8, prev = 0xdfcdd5a8}
(gdb) p ((struct list_head)inode_unused).prev
$36 = (struct list_head *) 0xdfcdd5a8
(gdb) p (((struct list_head)inode_unused).prev)->prev
$37 = (struct list_head *) 0xdfb5a2e8
(gdb) set $i = (struct inode *)0xdfb5a2e0
(gdb) p $i->i_ino
$38 = 0x3bec7
(gdb) p $i->i_count
$39 = {counter = 0x0}
Note that we deducted 8 from the address 0xdfb5a2e8 to obtain the address of
the struct inode (0xdfb5a2e0) according to the definition of list_entry()
macro from include/linux/list.h.
To understand how inode cache works, let us trace a lifetime of an inode of a regular file on ext2 filesystem as it is opened and closed:
fd = open("file", O_RDONLY);
close(fd);
The open(2) system call is implemented in fs/open.c:sys_open function and
the real work is done by fs/open.c:filp_open() function, which is split into
two parts:
open_namei(): fills in the nameidata structure containing the dentry
and vfsmount structures.
dentry_open(): given a dentry and vfsmount, this function allocates a new
struct file and links them together; it also invokes the filesystem
specific f_op->open() method which was set in inode->i_fop when inode
was read in open_namei() (which provided inode via dentry->d_inode).The open_namei() function interacts with dentry cache via path_walk(), which
in turn calls real_lookup(), which invokes the filesystem specific inode_operations->lookup() method.
The role of this method is to find the entry in the parent
directory with the matching name and then do iget(sb, ino) to get the
corresponding inode - which brings us to the inode cache. When the inode is
read in, the dentry is instantiated by means of d_add(dentry, inode). While
we are at it, note that for UNIX-style filesystems which have the concept of
on-disk inode number, it is the lookup method's job to map its endianness
to current CPU format, e.g. if the inode number in raw (fs-specific) dir
entry is in little-endian 32 bit format one could do:
unsigned long ino = le32_to_cpu(de->inode);
inode = iget(sb, ino);
d_add(dentry, inode);
So, when we open a file we hit iget(sb, ino) which is really
iget4(sb, ino, NULL, NULL), which does:
inode_lock. If inode is found,
its reference count (i_count) is incremented; if it
was 0 prior to incrementation and the inode is not dirty, it is removed from whatever
type list (inode->i_list) it is currently on (it has to be
inode_unused list, of course) and inserted into
inode_in_use type list; finally, inodes_stat.nr_unused is decremented.
iget4() is guaranteed to return an unlocked inode.
get_new_inode(), passing it the pointer
to the place in the hashtable where it should be inserted to.
get_new_inode() allocates a new inode from the inode_cachep SLAB
cache but this operation can block (GFP_KERNEL allocation), so it
must drop the inode_lock spinlock which guards the hashtable. Since it
has dropped the spinlock, it must retry searching the inode in the
hashtable afterwards; if it is found this time, it returns (after incrementing
the reference by __iget) the one found in the hashtable and destroys
the newly allocated one. If it is still not found in the hashtable,
then the new inode we have just allocated is the one to be used;
therefore it is initialised to the required values and the fs-specific
sb->s_op->read_inode() method is invoked to populate the rest of the
inode. This brings us from inode cache back to the filesystem code -
remember that we came to the inode cache when filesystem-specific
lookup() method invoked iget(). While the s_op->read_inode() method
is reading the inode from disk, the inode is locked (i_state = I_LOCK);
it is unlocked after the read_inode() method returns and all the waiters for it are
woken up.Now, let's see what happens when we close this file descriptor. The close(2)
system call is implemented in fs/open.c:sys_close() function, which calls
do_close(fd, 1) which rips (replaces with NULL) the descriptor of the
process' file descriptor table and invokes the filp_close() function which does
most of the work. The interesting things happen in fput(), which checks if
this was the last reference to the file, and if so calls
fs/file_table.c:_fput() which calls __fput() which is where interaction with
dcache (and therefore with inode cache - remember dcache is a Master of inode
cache!) happens. The fs/dcache.c:dput() does dentry_iput() which brings us
back to inode cache via iput(inode) so let us understand
fs/inode.c:iput(inode):
sb->s_op->put_inode() method, it is invoked
immediately with no spinlocks held (so it can block).
inode_lock spinlock is taken and i_count is decremented. If this was
NOT the last reference to this inode then we simply check if
there are too many references to it and so i_count can wrap around
the 32 bits allocated to it and if so we print a warning and return.
Note that we call printk() while holding the inode_lock spinlock -
this is fine because printk() can never block, therefore it may be called in
absolutely any context (even from interrupt handlers!).
The work performed by iput() on the last inode reference is rather complex
so we separate it into a list of its own:
i_nlink == 0 (e.g. the file was unlinked while we held it open)
then the inode is removed from hashtable and from its type list; if
there are any data pages held in page cache for this inode, they are
removed by means of truncate_all_inode_pages(&inode->i_data). Then
the filesystem-specific s_op->delete_inode() method is invoked,
which typically deletes the on-disk copy of the inode. If there is no
s_op->delete_inode() method registered by the filesystem (e.g. ramfs)
then we call clear_inode(inode), which invokes s_op->clear_inode() if
registered and if inode corresponds to a block device, this device's
reference count is dropped by bdput(inode->i_bdev).
i_nlink != 0 then we check if there are other inodes in the same
hash bucket and if there is none, then if inode is not dirty we delete
it from its type list and add it to inode_unused list, incrementing
inodes_stat.nr_unused. If there are inodes in the same hashbucket
then we delete it from the type list and add to inode_unused list.
If this was an anonymous inode (NetApp .snapshot) then we delete it
from the type list and clear/destroy it completely.The Linux kernel provides a mechanism for new filesystems to be written with minimum effort. The historical reasons for this are:
Let us consider the steps required to implement a filesystem under Linux.
The code to implement a filesystem can be either a dynamically loadable
module or statically linked into the kernel, and the way it is done under
Linux is very transparent. All that is needed is to fill in a
struct file_system_type structure and register it with the VFS using
the register_filesystem() function as in the following example from
fs/bfs/inode.c:
#include <linux/module.h>
#include <linux/init.h>
static struct super_block *bfs_read_super(struct super_block *, void *, int);
static DECLARE_FSTYPE_DEV(bfs_fs_type, "bfs", bfs_read_super);
static int __init init_bfs_fs(void)
{
return register_filesystem(&bfs_fs_type);
}
static void __exit exit_bfs_fs(void)
{
unregister_filesystem(&bfs_fs_type);
}
module_init(init_bfs_fs)
module_exit(exit_bfs_fs)
The module_init()/module_exit() macros ensure that, when BFS is compiled as a
module, the functions init_bfs_fs() and exit_bfs_fs() turn into init_module()
and cleanup_module() respectively; if BFS is statically linked into the kernel,
the exit_bfs_fs() code vanishes as it is unnecessary.
The struct file_system_type is declared in include/linux/fs.h:
struct file_system_type {
const char *name;
int fs_flags;
struct super_block *(*read_super) (struct super_block *, void *, int);
struct module *owner;
struct vfsmount *kern_mnt; /* For kernel mount, if it's FS_SINGLE fs */
struct file_system_type * next;
};
The fields thereof are explained thus:
/proc/filesystems file
and is used as a key to find a filesystem by its name; this same name is
used for the filesystem type in mount(2), and it should be unique: there
can (obviously) be only one filesystem with a given name. For modules,
name points to module's address spaces and not copied: this means cat
/proc/filesystems can oops if the module was unloaded but filesystem is
still registered.
FS_REQUIRES_DEV
for filesystems that can only be mounted on a block device, FS_SINGLE
for filesystems that can have only one superblock, FS_NOMOUNT for
filesystems that cannot be mounted from userspace by means of mount(2)
system call: they can however be mounted internally using kern_mount()
interface, e.g. pipefs.
FS_SINGLE case where it will Oops in
get_sb_single(), trying to dereference a NULL pointer in
fs_type->kern_mnt->mnt_sb with (fs_type->kern_mnt = NULL).
THIS_MODULE
does the right thing automatically.
FS_SINGLE filesystems only. This is set by
kern_mount() (TODO: kern_mount() should refuse to mount filesystems
if FS_SINGLE is not set).
file_systems
(see fs/super.c). The list is protected by the file_systems_lock
read-write spinlock and functions register/unregister_filesystem()
modify it by linking and unlinking the entry from the list.The job of the read_super() function is to fill in the fields of the superblock,
allocate root inode and initialise any fs-private information associated with
this mounted instance of the filesystem. So, typically the read_super() would
do:
sb->s_dev argument,
using buffer cache bread() function. If it anticipates to read a few
more subsequent metadata blocks immediately then it makes sense to
use breada() to schedule reading extra blocks asynchronously.
sb->s_op to point to struct super_block_operations
structure. This structure contains filesystem-specific functions
implementing operations like "read inode", "delete inode", etc.
d_alloc_root().
sb->s_dirt to 1
and mark the buffer containing superblock dirty (TODO: why do we
do this? I did it in BFS because MINIX did it...)Under Linux there are several levels of indirection between user file
descriptor and the kernel inode structure. When a process makes open(2)
system call, the kernel returns a small non-negative integer which can be
used for subsequent I/O operations on this file. This integer is an index
into an array of pointers to struct file. Each file structure points to
a dentry via file->f_dentry. And each dentry points to an inode via
dentry->d_inode.
Each task contains a field tsk->files which is a pointer to
struct files_struct defined in include/linux/sched.h:
/*
* Open file table structure
*/
struct files_struct {
atomic_t count;
rwlock_t file_lock;
int max_fds;
int max_fdset;
int next_fd;
struct file ** fd; /* current fd array */
fd_set *close_on_exec;
fd_set *open_fds;
fd_set close_on_exec_init;
fd_set open_fds_init;
struct file * fd_array[NR_OPEN_DEFAULT];
};
The file->count is a reference count, incremented by get_file() (usually
called by fget()) and decremented by fput() and by put_filp(). The difference
between fput() and put_filp() is that fput() does more work usually needed
for regular files, such as releasing flock locks, releasing dentry, etc, while
put_filp() is only manipulating file table structures, i.e. decrements the
count, removes the file from the anon_list and adds it to the free_list,
under protection of files_lock spinlock.
The tsk->files can be shared between parent and child if the child thread
was created using clone() system call with CLONE_FILES set in the clone flags
argument. This can be seen in kernel/fork.c:copy_files() (called by
do_fork()) which only increments the file->count if CLONE_FILES is set
instead of the usual copying file descriptor table in time-honoured
tradition of classical UNIX fork(2).
When a file is opened, the file structure allocated for it is installed into
current->files->fd[fd] slot and a fd bit is set in the bitmap
current->files->open_fds . All this is done under the write protection of
current->files->file_lock read-write spinlock. When the descriptor is
closed a fd bit is cleared in current->files->open_fds and
current->files->next_fd is set equal to fd as a hint for finding the
first unused descriptor next time this process wants to open a file.
The file structure is declared in include/linux/fs.h:
struct fown_struct {
int pid; /* pid or -pgrp where SIGIO should be sent */
uid_t uid, euid; /* uid/euid of process setting the owner */
int signum; /* posix.1b rt signal to be delivered on IO */
};
struct file {
struct list_head f_list;
struct dentry *f_dentry;
struct vfsmount *f_vfsmnt;
struct file_operations *f_op;
atomic_t f_count;
unsigned int f_flags;
mode_t f_mode;
loff_t f_pos;
unsigned long f_reada, f_ramax, f_raend, f_ralen, f_rawin;
struct fown_struct f_owner;
unsigned int f_uid, f_gid;
int f_error;
unsigned long f_version;
/* needed for tty driver, and maybe others */
void *private_data;
};
Let us look at the various fields of struct file:
sb->s_files list of all open files on this filesystem,
if the corresponding inode is not anonymous, then dentry_open() (called
by filp_open()) links the file into this list;
b) fs/file_table.c:free_list, containing unused file structures;
c) fs/file_table.c:anon_list, when a new file structure is created by
get_empty_filp() it is placed on this list. All these lists are
protected by the files_lock spinlock.
open_namei() (or
rather path_walk()
which it calls) but the actual file->f_dentry field is set by
dentry_open() to contain the dentry thus found.
vfsmount structure of the filesystem
containing the file. This is set by dentry_open() but is found as part
of nameidata lookup by open_namei() (or rather path_init() which it
calls).
file_operations which contains various
methods that can be invoked on the file. This is copied from
inode->i_fop which is placed there by filesystem-specific
s_op->read_inode() method during nameidata lookup. We will look at
file_operations methods in detail later on in this section.
get_file/put_filp/fput.
O_XXX flags from open(2) system call copied there
(with slight modifications by filp_open()) by dentry_open() and after
clearing O_CREAT, O_EXCL, O_NOCTTY, O_TRUNC - there is no point in
storing these flags permanently since they cannot be modified by
F_SETFL (or queried by F_GETFL) fcntl(2) calls.
dentry_open(). The point of the conversion is to store read and
write access in separate bits so one could do easy checks like
(f_mode & FMODE_WRITE) and (f_mode & FMODE_READ).
long long, i.e. a 64bit value.
SIGIO mechanism (see fs/fcntl.c:kill_fasync()).
get_empty_filp(). If the file is a socket, used by ipv4 netfilter.
fs/nfs/file.c and checked in mm/filemap.c:generic_file_write().
event) whenever f_pos changes.
file->f_dentry->d_inode->i_rdev.
Now let us look at file_operations structure which contains the methods that
can be invoked on files. Let us recall that it is copied from inode->i_fop
where it is set by s_op->read_inode() method. It is declared in
include/linux/fs.h:
struct file_operations {
struct module *owner;
loff_t (*llseek) (struct file *, loff_t, int);
ssize_t (*read) (struct file *, char *, size_t, loff_t *);
ssize_t (*write) (struct file *, const char *, size_t, loff_t *);
int (*readdir) (struct file *, void *, filldir_t);
unsigned int (*poll) (struct file *, struct poll_table_struct *);
int (*ioctl) (struct inode *, struct file *, unsigned int, unsigned long);
int (*mmap) (struct file *, struct vm_area_struct *);
int (*open) (struct inode *, struct file *);
int (*flush) (struct file *);
int (*release) (struct inode *, struct file *);
int (*fsync) (struct file *, struct dentry *, int datasync);
int (*fasync) (int, struct file *, int);
int (*lock) (struct file *, int, struct file_lock *);
ssize_t (*readv) (struct file *, const struct iovec *, unsigned long, loff_t *);
ssize_t (*writev) (struct file *, const struct iovec *, unsigned long, loff_t *);
};
THIS_MODULE, filesystems can
happily ignore it because their module counts are controlled at
mount/umount time whilst the drivers need to control it at open/release
time.
fs/read_write.c:default_llseek() is used, which does the
right thing (TODO: force all those who set it to NULL currently to use
default_llseek - that way we save an if() in llseek())
read(2) system call. Filesystems can use
mm/filemap.c:generic_file_read() for regular files and
fs/read_write.c:generic_read_dir() (which simply returns -EISDIR)
for directories here.
mm/filemap.c:generic_file_write() for regular files and ignore it for
directories here.
FIBMAP, FIGETBSZ, FIONREAD
are implemented by higher levels so they never read f_op->ioctl()
method.
dentry_open(). Filesystems
rarely use this, e.g. coda tries to cache the file locally at open
time.
release() method below). The only filesystem that
uses this is NFS client to flush all dirty pages. Note that this can
return an error which will be passed back to userspace which made the
close(2) system call.
file->f_count reaches 0. Although defined as returning int, the return
value is ignored by VFS (see fs/file_table.c:__fput()).
file = fget(fd)) and down/up
inode->i_sem semaphore. Ext2 filesystem currently ignores the last
argument and does exactly the same for fsync(2) and fdatasync(2).
file->f_flags & FASYNC
changes.
posix_lock_file()), if it
succeeds but the standard POSIX lock code fails then it will never be
unlocked on fs-dependent level..
Under Linux, information about mounted filesystems is kept in two separate
structures - super_block and vfsmount. The reason for this is that Linux
allows to mount the same filesystem (block device) under multiple mount
points, which means that the same super_block can correspond to multiple
vfsmount structures.
Let us look at struct super_block first, declared in include/linux/fs.h:
struct super_block {
struct list_head s_list; /* Keep this first */
kdev_t s_dev;
unsigned long s_blocksize;
unsigned char s_blocksize_bits;
unsigned char s_lock;
unsigned char s_dirt;
struct file_system_type *s_type;
struct super_operations *s_op;
struct dquot_operations *dq_op;
unsigned long s_flags;
unsigned long s_magic;
struct dentry *s_root;
wait_queue_head_t s_wait;
struct list_head s_dirty; /* dirty inodes */
struct list_head s_files;
struct block_device *s_bdev;
struct list_head s_mounts; /* vfsmount(s) of this one */
struct quota_mount_options s_dquot; /* Diskquota specific options */
union {
struct minix_sb_info minix_sb;
struct ext2_sb_info ext2_sb;
..... all filesystems that need sb-private info ...
void *generic_sbp;
} u;
/*
* The next field is for VFS *only*. No filesystems have any business
* even looking at it. You had been warned.
*/
struct semaphore s_vfs_rename_sem; /* Kludge */
/* The next field is used by knfsd when converting a (inode number based)
* file handle into a dentry. As it builds a path in the dcache tree from
* the bottom up, there may for a time be a subpath of dentrys which is not
* connected to the main tree. This semaphore ensure that there is only ever
* one such free path per filesystem. Note that unconnected files (or other
* non-directories) are allowed, but not unconnected diretories.
*/
struct semaphore s_nfsd_free_path_sem;
};
The various fields in the super_block structure are:
FS_REQUIRES_DEV filesystems, this is the i_dev of the
block device. For others (called anonymous filesystems) this is an
integer MKDEV(UNNAMED_MAJOR, i) where i is the first unset bit in
unnamed_dev_in_use array, between 1 and 255 inclusive. See
fs/super.c:get_unnamed_dev()/put_unnamed_dev(). It has been suggested
many times that anonymous filesystems should not use s_dev field.
lock_super()/unlock_super().
struct file_system_type of the
corresponding filesystem. Filesystem's read_super() method doesn't need
to set it as VFS fs/super.c:read_super() sets it for you if
fs-specific read_super() succeeds and resets to NULL if it fails.
super_operations structure which contains
fs-specific methods to read/write inodes etc. It is the job of
filesystem's read_super() method to initialise s_op correctly.
read_super() to read the root inode from the disk and pass it to
d_alloc_root() to allocate the dentry and instantiate it. Some
filesystems spell "root" other than "/" and so use more generic
d_alloc() function to bind the dentry to a name, e.g. pipefs mounts
itself on "pipe:" as its own root instead of "/".
inode->i_state & I_DIRTY) then it is on superblock-specific
dirty list linked via inode->i_list.
fs/file_table.c:fs_may_remount_ro() which goes through sb->s_files list
and denies remounting if there are files opened for write
(file->f_mode & FMODE_WRITE) or files with pending
unlink (inode->i_nlink == 0).
FS_REQUIRES_DEV, this points to the block_device
structure describing the device the filesystem is mounted on.
vfsmount structures, one for each
mounted instance of this superblock.
The superblock operations are described in the super_operations structure
declared in include/linux/fs.h:
struct super_operations {
void (*read_inode) (struct inode *);
void (*write_inode) (struct inode *, int);
void (*put_inode) (struct inode *);
void (*delete_inode) (struct inode *);
void (*put_super) (struct super_block *);
void (*write_super) (struct super_block *);
int (*statfs) (struct super_block *, struct statfs *);
int (*remount_fs) (struct super_block *, int *, char *);
void (*clear_inode) (struct inode *);
void (*umount_begin) (struct super_block *);
};
fs/inode.c:get_new_inode() from iget4() (and therefore
iget()). If a filesystem wants to use iget() then read_inode() must be
implemented - otherwise get_new_inode() will panic.
While inode is being read it is locked (inode->i_state = I_LOCK). When
the function returns, all waiters on inode->i_wait are woken up. The job
of the filesystem's read_inode() method is to locate the disk block which
contains the inode to be read and use buffer cache bread() function to
read it in and initialise the various fields of inode structure, for
example the inode->i_op and inode->i_fop so that VFS level knows what
operations can be performed on the inode or corresponding file.
Filesystems that don't implement read_inode() are ramfs and
pipefs. For example, ramfs has its own inode-generating function
ramfs_get_inode() with all the inode operations calling it as needed.
read_inode() in that it needs to locate the relevant block on
disk and interact with buffer cache by calling
mark_buffer_dirty(bh). This method is called on dirty inodes
(those marked dirty with mark_inode_dirty()) when the inode needs
to be sync'd either individually or as part of syncing the
entire filesystem.
inode->i_count and
inode->i_nlink reach 0. Filesystem deletes the on-disk copy of the
inode and calls clear_inode() on VFS inode to "terminate it with
extreme prejudice".
brelse() the block containing the superblock and kfree() any
bitmaps allocated for free blocks, inodes, etc.
sb-private area) and
mark_buffer_dirty(bh) . It should also clear sb->s_dirt flag.
struct statfs passed as argument is a kernel
pointer, not a user pointer so we don't need to do any I/O to
userspace. If not implemented then statfs(2) will fail with ENOSYS.
clear_inode(). Filesystems
that attach private data to inode structure (via generic_ip field) must
free it here.
So, let us look at what happens when we mount a on-disk (FS_REQUIRES_DEV)
filesystem. The implementation of the mount(2) system call is in
fs/super.c:sys_mount() which is the just a wrapper that copies the options,
filesystem type and device name for the do_mount() function which does the
real work:
do_mount()
calling get_fs_type() and once by get_sb_dev() calling get_filesystem()
if read_super() was successful. The first increment is to prevent
module unloading while we are inside read_super() method and the second
increment is to indicate that the module is in use by this mounted
instance. Obviously, do_mount() decrements the count before returning, so
overall the count only grows by 1 after each mount.
fs_type->fs_flags & FS_REQUIRES_DEV is true, the
superblock is initialised by a call to get_sb_bdev() which obtains
the reference to the block device and interacts with the filesystem's
read_super() method to fill in the superblock. If all goes well, the
super_block structure is initialised and we have an extra reference
to the filesystem's module and a reference to the underlying block
device.
vfsmount structure is allocated and linked to sb->s_mounts list
and to the global vfsmntlist list. The vfsmount field mnt_instances
allows to find all instances mounted on the same superblock as this
one. The mnt_list field allows to find all instances for all
superblocks system-wide. The mnt_sb field
points to this superblock and mnt_root has a new reference to the
sb->s_root dentry.As a simple example of Linux filesystem that does not require a block device
for mounting, let us consider pipefs from fs/pipe.c. The filesystem's preamble
is rather straightforward and requires little explanation:
static DECLARE_FSTYPE(pipe_fs_type, "pipefs", pipefs_read_super,
FS_NOMOUNT|FS_SINGLE);
static int __init init_pipe_fs(void)
{
int err = register_filesystem(&pipe_fs_type);
if (!err) {
pipe_mnt = kern_mount(&pipe_fs_type);
err = PTR_ERR(pipe_mnt);
if (!IS_ERR(pipe_mnt))
err = 0;
}
return err;
}
static void __exit exit_pipe_fs(void)
{
unregister_filesystem(&pipe_fs_type);
kern_umount(pipe_mnt);
}
module_init(init_pipe_fs)
module_exit(exit_pipe_fs)
The filesystem is of type FS_NOMOUNT|FS_SINGLE, which means it cannot be
mounted from userspace and can only have one superblock system-wide. The
FS_SINGLE file also means that it must be mounted via kern_mount() after
it is successfully registered via register_filesystem(), which is exactly
what happens in init_pipe_fs(). The only bug in this function is that if
kern_mount() fails (e.g. because kmalloc() failed in add_vfsmnt()) then the
filesystem is left as registered but module initialisation fails. This will
cause cat /proc/filesystems to Oops. (have just sent a patch to Linus
mentioning that although this is not a real bug today as pipefs can't be
compiled as a module, it should be written with the view that in the future
it may become modularised).
The result of register_filesystem() is that pipe_fs_type is linked into
the file_systems list so one can read /proc/filesystems and find "pipefs"
entry in there with "nodev" flag indicating that FS_REQUIRES_DEV was not set.
The /proc/filesystems file should really be enhanced to support all the new
FS_ flags (and I made a patch to do so) but it cannot be done because it will
break all the user applications that use it. Despite Linux kernel interfaces
changing every minute (only for the better) when it comes to the userspace
compatibility, Linux is a very conservative operating system which allows
many applications to be used for a long time without being recompiled.
The result of kern_mount() is that:
unnamed_dev_in_use bitmap; if there are no more bits then kern_mount()
fails with EMFILE.
get_empty_super().
The get_empty_super() function walks the list of superblocks headed
by super_block and looks for empty entry, i.e. s->s_dev == 0. If no
such empty superblock is found then a new one is allocated using
kmalloc() at GFP_USER priority. The maximum system-wide number of
superblocks is checked in get_empty_super() so if it starts failing,
one can adjust the tunable /proc/sys/fs/super-max.
pipe_fs_type->read_super() method, i.e.
pipefs_read_super(), is invoked which allocates root inode and root
dentry sb->s_root, and sets sb->s_op to be &pipefs_ops.
kern_mount() calls add_vfsmnt(NULL, sb->s_root, "none") which
allocates a new vfsmount structure and links it into vfsmntlist and
sb->s_mounts.
pipe_fs_type->kern_mnt is set to this new vfsmount structure and
it is returned. The reason why the return value of kern_mount() is a
vfsmount structure is because even FS_SINGLE filesystems can be mounted
multiple times and so their mnt->mnt_sb will point to the same thing
which would be silly to return from multiple calls to kern_mount().Now that the filesystem is registered and inkernel-mounted we can use it.
The entry point into the pipefs filesystem is the pipe(2) system call,
implemented in arch-dependent function sys_pipe() but the real work is done
by a portable fs/pipe.c:do_pipe() function. Let us look at do_pipe() then.
The interaction with pipefs happens when do_pipe() calls get_pipe_inode()
to allocate a new pipefs inode. For this inode, inode->i_sb is set to
pipefs' superblock pipe_mnt->mnt_sb, the file operations i_fop is set to
rdwr_pipe_fops and the number of readers and writers (held in inode->i_pipe)
is set to 1. The reason why there is a separate inode field i_pipe instead
of keeping it in the fs-private union is that pipes and FIFOs share the same
code and FIFOs can exist on other filesystems which use the other access
paths within the same union which is very bad C and can work only by pure
luck. So, yes, 2.2.x kernels work only by pure luck and will stop working
as soon as you slightly rearrange the fields in the inode.
Each pipe(2) system call increments a reference count on the pipe_mnt
mount instance.
Under Linux, pipes are not symmetric (bidirection or STREAM pipes), i.e.
two sides of the file have different file->f_op operations - the
read_pipe_fops and write_pipe_fops respectively. The write on read side
returns EBADF and so does read on write side.
As a simple example of ondisk Linux filesystem, let us consider BFS. The
preamble of the BFS module is in fs/bfs/inode.c:
static DECLARE_FSTYPE_DEV(bfs_fs_type, "bfs", bfs_read_super);
static int __init init_bfs_fs(void)
{
return register_filesystem(&bfs_fs_type);
}
static void __exit exit_bfs_fs(void)
{
unregister_filesystem(&bfs_fs_type);
}
module_init(init_bfs_fs)
module_exit(exit_bfs_fs)
A special fstype declaration macro DECLARE_FSTYPE_DEV() is used which
sets the fs_type->flags to FS_REQUIRES_DEV to signify that BFS requires a
real block device to be mounted on.
The module's initialisation function registers the filesystem with VFS and the cleanup function (only present when BFS is configured to be a module) unregisters it.
With the filesystem registered, we can proceed to mount it, which would
invoke out fs_type->read_super() method which is implemented in
fs/bfs/inode.c:bfs_read_super(). It does the following:
set_blocksize(s->s_dev, BFS_BSIZE): since we are about to interact
with the block device layer via the buffer cache, we must initialise a few
things, namely set the block size and also inform VFS via fields
s->s_blocksize and s->s_blocksize_bits.
bh = bread(dev, 0, BFS_BSIZE): we read block 0 of the device
passed via s->s_dev. This block is the filesystem's superblock.
BFS_MAGIC number and, if valid, stored
in the sb-private field s->su_sbh (which is really s->u.bfs_sb.si_sbh).
kmalloc(GFP_KERNEL) and clear all
bits to 0 except the first two which we set to 1 to indicate that we
should never allocate inodes 0 and 1. Inode 2 is root and the
corresponding bit will be set to 1 a few lines later anyway - the
filesystem should have a valid root inode at mounting time!
s->s_op, which means that we can from this point
invoke inode cache via iget() which results in s_op->read_inode() to
be invoked. This finds the block that contains the specified (by
inode->i_ino and inode->i_dev) inode and reads it in. If we fail to
get root inode then we free the inode bitmap and release superblock
buffer back to buffer cache and return NULL. If root inode was read OK,
then we allocate a dentry with name / (as becometh root) and
instantiate it with this inode.
iput() - we don't hold a reference
to it longer than needed.
s->s_dirt flag (TODO: why do I do this?
Originally, I did it because minix_read_super() did but neither minix
nor BFS seem to modify superblock in the read_super()).
fs/super.c:read_super().After the read_super() function returns successfully, VFS obtains the
reference to the filesystem module via call to get_filesystem(fs_type) in
fs/super.c:get_sb_bdev() and a reference to the block device.
Now, let us examine what happens when we do I/O on the filesystem. We already
examined how inodes are read when iget() is called and how they are released
on iput(). Reading inodes sets up, among other things, inode->i_op and
inode->i_fop; opening a file will propagate inode->i_fop into file->f_op.
Let us examine the code path of the link(2) system call. The implementation
of the system call is in fs/namei.c:sys_link():
getname()
function which does the error checking.
path_init()/path_walk()
interaction with dcache. The result is stored in old_nd and nd
structures.
old_nd.mnt != nd.mnt then "cross-device link" EXDEV is returned -
one cannot link between filesystems, in Linux this translates into -
one cannot link between mounted instances of a filesystem (or, in
particular between filesystems).
nd by lookup_create() .
vfs_link() function is called which checks if we can
create a new entry in the directory and invokes the dir->i_op->link()
method which brings us back to filesystem-specific
fs/bfs/dir.c:bfs_link() function.
bfs_link(), we check if we are trying to link a directory and
if so, refuse with EPERM error. This is the same behaviour as standard (ext2).
bfs_add_entry() which goes through all
entries looking for unused slot (de->ino == 0) and, when found, writes
out the name/inode pair into the corresponding block and marks it
dirty (at non-superblock priority).
inode->i_nlink, update
inode->i_ctime and mark this inode dirty as well as instantiating the
new dentry with the inode.Other related inode operations like unlink()/rename() etc work in a similar
way, so not much is gained by examining them all in details.
Linux supports loading user application binaries from disk. More interestingly, the binaries can be stored in different formats and the operating system's response to programs via system calls can deviate from norm (norm being the Linux behaviour) as required, in order to emulate formats found in other flavours of UNIX (COFF, etc) and also to emulate system calls behaviour of other flavours (Solaris, UnixWare, etc). This is what execution domains and binary formats are for.
Each Linux task has a personality stored in its task_struct (p->personality).
The currently existing (either in the official kernel or as addon patch)
personalities include support for FreeBSD, Solaris, UnixWare, OpenServer and
many other popular operating systems.
The value of current->personality is split into two parts:
STICKY_TIMEOUTS, WHOLE_SECONDS, etc.By changing the personality, we can change
the way the operating system treats certain system calls, for example
adding a STICKY_TIMEOUT to current->personality makes select(2) system call
preserve the value of last argument (timeout) instead of storing the
unslept time. Some buggy programs rely on buggy operating systems (non-Linux)
and so Linux provides a way to emulate bugs in cases where the source code
is not available and so bugs cannot be fixed.
Execution domain is a contiguous range of personalities implemented by a single module. Usually a single execution domain implements a single personality but sometimes it is possible to implement "close" personalities in a single module without too many conditionals.
Execution domains are implemented in kernel/exec_domain.c and were completely
rewritten for 2.4 kernel, compared with 2.2.x. The list of execution domains
currently supported by the kernel, along with the range of personalities
they support, is available by reading the /proc/execdomains file. Execution
domains, except the PER_LINUX one, can be implemented as dynamically
loadable modules.
The user interface is via personality(2) system call, which sets the current
process' personality or returns the value of current->personality if the
argument is set to impossible personality 0xffffffff. Obviously, the
behaviour of this system call itself does not depend on personality..
The kernel interface to execution domains registration consists of two functions:
int register_exec_domain(struct exec_domain *): registers the
execution domain by linking it into single-linked list exec_domains
under the write protection of the read-write spinlock exec_domains_lock.
Returns 0 on success, non-zero on failure.
int unregister_exec_domain(struct exec_domain *): unregisters the
execution domain by unlinking it from the exec_domains list, again using
exec_domains_lock spinlock in write mode. Returns 0 on success.The reason why exec_domains_lock is a read-write is that only registration
and unregistration requests modify the list, whilst doing
cat /proc/filesystems calls fs/exec_domain.c:get_exec_domain_list(), which
needs only read access to the list. Registering a new execution domain
defines a "lcall7 handler" and a signal number conversion map. Actually,
ABI patch extends this concept of exec domain to include extra information
(like socket options, socket types, address family and errno maps).
The binary formats are implemented in a similar manner, i.e. a single-linked
list formats is defined in fs/exec.c and is protected by a read-write lock
binfmt_lock. As with exec_domains_lock, the binfmt_lock is taken read on
most occasions except for registration/unregistration of binary formats.
Registering a new binary format enhances the execve(2) system call with new
load_binary()/load_shlib() functions as well as ability to core_dump() . The
load_shlib() method is used only by the old uselib(2) system call while
the load_binary() method is called by the search_binary_handler() from
do_execve() which implements execve(2) system call.
The personality of the process is determined at binary format loading by
the corresponding format's load_binary() method using some heuristics.
For example to determine UnixWare7 binaries one first marks the binary
using the elfmark(1) utility, which sets the ELF header's e_flags to the magic
value 0x314B4455 which is detected at ELF loading time and
current->personality is set to PER_UW7. If this heuristic fails, then a more
generic one, such as treat ELF interpreter paths like /usr/lib/ld.so.1 or
/usr/lib/libc.so.1 to
indicate a SVR4 binary, is used and personality is set to PER_SVR4. One
could write a little utility program that uses Linux's ptrace(2) capabilities
to single-step the code and force a running program into any personality.
Once personality (and therefore current->exec_domain) is known, the system
calls are handled as follows. Let us assume that a process makes a system
call by means of lcall7 gate instruction. This transfers control to
ENTRY(lcall7) of arch/i386/kernel/entry.S because it was prepared in
arch/i386/kernel/traps.c:trap_init(). After appropriate stack layout
conversion, entry.S:lcall7 obtains the pointer to exec_domain from current
and then an offset of lcall7 handler within the exec_domain (which is
hardcoded as 4 in asm code so you can't shift the handler field around in
C declaration of struct exec_domain) and jumps to it. So, in C, it would
look like this:
static void UW7_lcall7(int segment, struct pt_regs * regs)
{
abi_dispatch(regs, &uw7_funcs[regs->eax & 0xff], 1);
}
where abi_dispatch() is a wrapper around the table of function pointers that
implement this personality's system calls uw7_funcs.
In this chapter we describe the Linux 2.4 pagecache. The pagecache is - as the name suggests - a cache of physical pages. In the UNIX world the concept of a pagecache became popular with the introduction of SVR4 UNIX, where it replaced the buffercache for data IO operations.
While the SVR4 pagecache is only used for filesystem data cache and thus uses
the struct vnode and an offset into the file as hash parameters, the Linux page
cache is designed to be more generic, and therefore uses a struct address_space
(explained below) as first parameter. Because the Linux pagecache is tightly
coupled to the notation of address spaces, you will need at least a basic
understanding of adress_spaces to understand the way the pagecache works.
An address_space is some kind of software MMU that maps all pages of one object
(e.g. inode) to an other concurrency (typically physical disk blocks).
The struct address_space is defined in include/linux/fs.h as:
struct address_space {
struct list_head clean_pages;
struct list_head dirty_pages;
struct list_head locked_pages;
unsigned long nrpages;
struct address_space_operations *a_ops;
struct inode *host;
struct vm_area_struct *i_mmap;
struct vm_area_struct *i_mmap_shared;
spinlock_t i_shared_lock;
};
To understand the way address_spaces works, we only need to look at a few of this fields:
clean_pages, dirty_pages and locked_pages are double linked lists
of all clean, dirty and locked pages that belong to this address_space, nrpages
is the total number of pages in this address_space. a_ops defines the methods of
this object and host is an pointer to the inode this address_space belongs to -
it may also be NULL, e.g. in the case of the swapper address_space
(mm/swap_state.c,).
The usage of clean_pages, dirty_pages, locked_pages and
nrpages is obvious, so we will take a tighter look at the
address_space_operations structure, defined in the same header:
struct address_space_operations {
int (*writepage)(struct page *);
int (*readpage)(struct file *, struct page *);
int (*sync_page)(struct page *);
int (*prepare_write)(struct file *, struct page *, unsigned, unsigned);
int (*commit_write)(struct file *, struct page *, unsigned, unsigned);
int (*bmap)(struct address_space *, long);
};
For a basic view at the principle of address_spaces (and the pagecache) we need
to take a look at ->writepage and ->readpage, but in practice we need
to take a look at ->prepare_write and ->commit_write, too.
You can probably guess what the address_space_operations methods do by virtue of their names alone; nevertheless, they do require some explanation. Their use in the course of filesystem data I/O, by far the most common path through the pagecache, provides a good way of understanding them. Unlike most other UNIX-like operating systems, Linux has generic file operations (a subset of the SYSVish vnode operations) for data IO through the pagecache. This means that the data will not directly interact with the file- system on read/write/mmap, but will be read/written from/to the pagecache whenever possible. The pagecache has to get data from the actual low-level filesystem in case the user wants to read from a page not yet in memory, or write data to disk in case memory gets low.
In the read path the generic methods will first try to find a page that matches the wanted inode/index tuple.
hash = page_hash(inode->i_mapping, index);
Then we test whether the page actually exists.
hash = page_hash(inode->i_mapping, index);
page = __find_page_nolock(inode->i_mapping, index, *hash);
When it does not exist, we allocate a new free page, and add it to the page- cache hash.
page = page_cache_alloc();
__add_to_page_cache(page, mapping, index, hash);
After the page is hashed we use the ->readpage address_space operation to
actually fill the page with data. (file is an open instance of inode).
error = mapping->a_ops->readpage(file, page);
Finally we can copy the data to userspace.
For writing to the filesystem two pathes exist: one for writable mappings (mmap) and one for the write(2) family of syscalls. The mmap case is very simple, so it will be discussed first. When a user modifies mappings, the VM subsystem marks the page dirty.
SetPageDirty(page);
The bdflush kernel thread that is trying to free pages, either as background
activity or because memory gets low will try to call ->writepage on the pages
that are explicitly marked dirty. The ->writepage method does now have to
write the pages content back to disk and free the page.
The second write path is _much_ more complicated. For each page the user
writes to, we are basically doing the following:
(for the full code see mm/filemap.c:generic_file_write()).
page = __grab_cache_page(mapping, index, &cached_page);
mapping->a_ops->prepare_write(file, page, offset, offset+bytes);
copy_from_user(kaddr+offset, buf, bytes);
mapping->a_ops->commit_write(file, page, offset, offset+bytes);
So first we try to find the hashed page or allocate a new one, then we call the
->prepare_write address_space method, copy the user buffer to kernel memory and
finally call the ->commit_write method. As you probably have seen
->prepare_write and ->commit_write are fundamentally different from ->readpage
and ->writepage, because they are not only called when physical IO is actually
wanted but everytime the user modifies the file.
There are two (or more?) ways to handle this, the first one uses the Linux
buffercache to delay the physical IO, by filling a page->buffers pointer with
buffer_heads, that will be used in try_to_free_buffers (fs/buffers.c) to
request IO once memory gets low, and is used very widespread in the current
kernel. The other way just sets the page dirty and relies on ->writepage to do
all the work. Due to the lack of a validitity bitmap in struct page this does
not work with filesystem that have a smaller granuality then PAGE_SIZE.
This chapter describes the semaphore, shared memory, and message queue IPC mechanisms as implemented in the Linux 2.4 kernel. It is organized into four sections. The first three sections cover the interfaces and support functions for semaphores, message queues, and shared memory respectively. The last section describes a set of common functions and data structures that are shared by all three mechanisms.
The functions described in this section implement the user level semaphore mechanisms. Note that this implementation relies on the use of kernel splinlocks and kernel semaphores. To avoid confusion, the term "kernel semaphore" will be used in reference to kernel semaphores. All other uses of the word "sempahore" will be in reference to the user level semaphores.
The entire call to sys_semget() is protected by the global sem_ids.sem kernel semaphore.
In the case where a new set of semaphores must be created, the newary() function is called to create and initialize a new semaphore set. The ID of the new set is returned to the caller.
In the case where a key value is provided for an existing semaphore set, ipc_findkey() is invoked to look up the corresponding semaphore descriptor array index. The parameters and permissions of the caller are verified before returning the semaphore set ID.
For the IPC_INFO, SEM_INFO, and SEM_STAT commands, semctl_nolock() is called to perform the necessary functions.
For the GETALL, GETVAL, GETPID, GETNCNT, GETZCNT, IPC_STAT, SETVAL,and SETALL commands, semctl_main() is called to perform the necessary functions.
For the IPC_RMID and IPC_SET command, semctl_down() is called to perform the necessary functions. Throughout both of these operations, the global sem_ids.sem kernel semaphore is held.
After validating the call parameters, the semaphore operations data is copied from user space to a temporary buffer. If a small temporary buffer is sufficient, then a stack buffer is used. Otherwise, a larger buffer is allocated. After copying in the semaphore operations data, the global semaphores spinlock is locked, and the user-specified semaphore set ID is validated. Access permissions for the semaphore set are also validated.
All of the user-specified semaphore operations are parsed.
During this process, a count is maintained of all the operations that
have the SEM_UNDO flag set. A decrease flag is set if any of the
operations subtract from a semaphore value, and an alter flag is set
if any of the semaphore values are modified (i.e. increased or
decreased). The number of each
semaphore to be modified is validated.
If SEM_UNDO was asserted for any of the semaphore operations, then the undo list for the current task is searched for an undo structure associated with this semaphore set. During this search, if the semaphore set ID of any of the undo structures is found to be -1, then freeundos() is called to free the undo structure and remove it from the list. If no undo structure is found for this semaphore set then alloc_undo() is called to allocate and initialize one.
The
try_atomic_semop()
function is called with the do_undo
parameter equal to 0 in order to execute the sequence of
operations. The return value indicates that either the
operations passed, failed, or were not executed because
they need to block. Each of these cases are further described below:
The try_atomic_semop() function returns zero to indicate that all operations in the sequence succeeded. In this case, update_queue() is called to traverse the queue of pending semaphore operations for the semaphore set and awaken any sleeping tasks that no longer need to block. This completes the execution of the sys_semop() system call for this case.
If try_atomic_semop() returns a negative value, then a failure condition was encountered. In this case, none of the operations have been executed. This occurs when either a semaphore operation would cause an invalid semaphore value, or an operation marked IPC_NOWAIT is unable to complete. The error condition is then returned to the caller of sys_semop().
Before sys_semop() returns, a call is made to update_queue() to traverse the queue of pending semaphore operations for the semaphore set and awaken any sleeping tasks that no longer need to block.
The try_atomic_semop() function returns 1 to indicate that the sequence of semaphore operations was not executed because one of the semaphores would block. For this case, a new sem_queue element is initialized containing these semaphore operations. If any of these operations would alter the state of the semaphore, then the new queue element is added at the tail of the queue. Otherwise, the new queue element is added at the head of the queue.
The semsleeping element of the current
task is set to indicate that the task is sleeping on this
sem_queue element.
The current task is marked as TASK_INTERRUPTIBLE, and the
sleeper element of the
sem_queue
is set to identify this task as the sleeper. The
global semaphore spinlock is then unlocked, and schedule() is called
to put the current task to sleep.
When awakened, the task re-locks the global semaphore spinlock, determines why it was awakened, and how it should respond. The following cases are handled:
status element of the
sem_queue structure
is set to 1, then the task was awakened in order to retry the
semaphore operations. Another call to
try_atomic_semop() is
made to execute the sequence of semaphore operations. If
try_atomic_sweep() returns 1, then the task must block again
as described above. Otherwise, 0 is returned for success,
or an appropriate error code is returned in case of failure.
Before sys_semop() returns, current->semsleeping is cleared,
and the
sem_queue
is removed from the queue. If any of the specified semaphore
operations were altering operations (increase or decrease),
then
update_queue() is
called to traverse the queue of pending semaphore operations
for the semaphore set and awaken any sleeping tasks that no
longer need to block.
status element of the
sem_queue structure is
NOT set to 1, and the
sem_queue element has
not been dequeued, then the task was awakened by an interrupt.
In this case, the system call fails with EINTR. Before
returning, current->semsleeping is cleared, and the
sem_queue is removed
from the queue. Also,
update_queue() is called
if any of the operations were altering operations.
status element of the
sem_queue structure is
NOT set to 1, and the
sem_queue element
has been dequeued,
then the semaphore operations have already been executed by
update_queue(). The
queue status, which could be 0 for success
or a negated error code for failure, becomes the return value of
the system call.
The following structures are used specifically for semaphore support:
/* One sem_array data structure for each set of semaphores in the system. */
struct sem_array {
struct kern_ipc_perm sem_perm; /* permissions .. see ipc.h */
time_t sem_otime; /* last semop time */
time_t sem_ctime; /* last change time */
struct sem *sem_base; /* ptr to first semaphore in array */
struct sem_queue *sem_pending; /* pending operations to be processed */
struct sem_queue **sem_pending_last; /* last pending operation */
struct sem_undo *undo; /* undo requests on this array * /
unsigned long sem_nsems; /* no. of semaphores in array */
};
/* One semaphore structure for each semaphore in the system. */
struct sem {
int semval; /* current value */
int sempid; /* pid of last operation */
};
struct seminfo {
int semmap;
int semmni;
int semmns;
int semmnu;
int semmsl;
int semopm;
int semume;
int semusz;
int semvmx;
int semaem;
};
struct semid64_ds {
struct ipc64_perm sem_perm; /* permissions .. see
ipc.h */
__kernel_time_t sem_otime; /* last semop time */
unsigned long __unused1;
__kernel_time_t sem_ctime; /* last change time */
unsigned long __unused2;
unsigned long sem_nsems; /* no. of semaphores in
array */
unsigned long __unused3;
unsigned long __unused4;
};
/* One queue for each sleeping process in the system. */
struct sem_queue {
struct sem_queue * next; /* next entry in the queue */
struct sem_queue ** prev; /* previous entry in the queue, *(q->pr
ev) == q */
struct task_struct* sleeper; /* this process */
struct sem_undo * undo; /* undo structure */
int pid; /* process id of requesting process */
int status; /* completion status of operation */
struct sem_array * sma; /* semaphore array for operations */
int id; /* internal sem id */
struct sembuf * sops; /* array of pending operations */
int nsops; /* number of operations */
int alter; /* operation will alter semaphore */
};
/* semop system calls takes an array of these. */
struct sembuf {
unsigned short sem_num; /* semaphore index in array */
short sem_op; /* semaphore operation */
short sem_flg; /* operation flags */
};
/* Each task has a list of undo requests. They are executed automatically
* when the process exits.
*/
struct sem_undo {
struct sem_undo * proc_next; /* next entry on this process */
struct sem_undo * id_next; /* next entry on this semaphore set */
int semid; /* semaphore set identifier */
short * semadj; /* array of adjustments, one per
semaphore */
};
The following functions are used specifically in support of semaphores:
newary() relies on the
ipc_alloc()
function to allocate the memory
required for the new semaphore set. It allocates enough memory
for the semaphore set descriptor and for each of the semaphores
in the set. The allocated memory is cleared, and the address of the
first element of the semaphore set descriptor is passed to
ipc_addid().
ipc_addid() reserves an array entry
for the new semaphore set descriptor and initializes the
(
struct kern_ipc_perm) data for the set.
The global used_sems variable is updated by the number of
semaphores in the new set and the initialization of the
(
struct kern_ipc_perm)
data for the new set is completed. Other
initialization for this set performed are listed below:
sem_base element for the set is initialized
to the address immediately following the
(
struct sem_array)
portion of the newly allocated data. This corresponds to
the location of the first semaphore in the set.
sem_pending queue is initialized as empty.All of the operations following the call to ipc_addid() are performed while holding the global semaphores spinlock. After unlocking the global semaphores spinlock, newary() calls ipc_buildid() (via sem_buildid()). This function uses the index of the semaphore set descriptor to create a unique ID, that is then returned to the caller of newary().
freeary() is called by semctl_down() to perform the functions listed below. It is called with the global semaphores spinlock locked and it returns with the spinlock unlocked
semctl_down() provides the IPC_RMID and IPC_SET operations of the semctl() system call. The semaphore set ID and the access permissions are verified prior to either of these operations, and in either case, the global semaphore spinlock is held throughout the operation.
The IPC_RMID operation calls freeary() to remove the semaphore set.
The IPC_SET operation updates the uid, gid,
mode, and ctime elements of the semaphore set.
semctl_nolock() is called by sys_semctl() to perform the IPC_INFO, SEM_INFO and SEM_STAT functions.
IPC_INFO and SEM_INFO cause a temporary
seminfo
buffer to be initialized and loaded with unchanging semaphore
statistical data. Then, while holding the global sem_ids.sem
kernel semaphore, the semusz and semaem elements of
the
seminfo structure are
updated according to the given command (IPC_INFO or SEM_INFO).
The return value of the system call is set to the maximum
semaphore set ID.
SEM_STAT causes a temporary
semid64_ds
buffer to be initialized. The global
semaphore spinlock is then held while copying the sem_otime,
sem_ctime, and sem_nsems values into the buffer. This data is
then copied to user space.
semctl_main() is called by sys_semctl() to perform many of the supported functions, as described in the subsections below. Prior to performing any of the following operations, semctl_main() locks the global semaphore spinlock and validates the semaphore set ID and the permissions. The spinlock is released before returning.
The GETALL operation loads the current semaphore values into a temporary kernel buffer and copies them out to user space. The small stack buffer is used if the semaphore set is small. Otherwise, the spinlock is temporarily dropped in order to allocate a larger buffer. The spinlock is held while copying the semaphore values in to the temporary buffer.
The SETALL operation copies semaphore values from user space into a temporary buffer, and then into the semaphore set. The spinlock is dropped while copying the values from user space into the temporary buffer, and while verifying reasonable values. If the semaphore set is small, then a stack buffer is used, otherwise a larger buffer is allocated. The spinlock is regained and held while the following operations are performed on the semaphore set:
sem_ctime value for the semaphore set is set.
In the IPC_STAT operation, the sem_otime,
sem_ctime, and sem_nsems value are copied into
a stack buffer. The data is then copied to user space after
dropping the spinlock.
For GETVAL in the non-error case, the return value for the system call is set to the value of the specified semaphore.
For GETPID in the non-error case, the return value for the system call is
set to the pid associated with the last operation on the
semaphore.
For GETNCNT in the non-error case, the return value for the system call is set to the number of processes waiting on the semaphore being less than zero. This number is calculated by the count_semncnt() function.
For GETZCNT in the non-error case, the return value for the system call is set to the number of processes waiting on the semaphore being set to zero. This number is calculated by the count_semzcnt() function.
After validating the new semaphore value, the following functions are performed:
sem_ctime value for the semaphore set is updated.count_semncnt() counts the number of tasks waiting on the value of a semaphore to be less than zero.
count_semzcnt() counts the number of tasks waiting on the value of a semaphore to be zero.
update_queue() traverses the queue of pending semops for
a semaphore set and calls
try_atomic_semop()
to determine which sequences of semaphore operations
would succeed. If the status of the queue element
indicates that blocked tasks have already
been awakened, then the queue element is skipped over. For other
elements of the queue, the q-alter flag
is passed as the undo parameter to
try_atomic_semop(),
indicating that any
altering operations should be undone before returning.
If the sequence of operations would block, then update_queue() returns without making any changes.
A sequence of operations can fail if one of the semaphore operations would cause an invalid semaphore value, or an operation marked IPC_NOWAIT is unable to complete. In such a case, the task that is blocked on the sequence of semaphore operations is awakened, and the queue status is set with an appropriate error code. The queue element is also dequeued.
If the sequence of operations is non-altering, then
they would have passed a zero value as the undo parameter to
try_atomic_semop().
If these operations succeeded, then they
are considered complete and are removed from the queue.
The blocked task is awakened, and the queue element
status is set to indicate success.
If the sequence of operations would alter the semaphore values, but can succeed, then sleeping tasks that no longer need to be blocked are awakened. The queue status is set to 1 to indicate that the blocked task has been awakened. The operations have not been performed, so the queue element is not removed from the queue. The semaphore operations would be executed by the awakened task.
try_atomic_semop() is called by sys_semop() and update_queue() to determine if a sequence of semaphore operations will all succeed. It determines this by attempting to perform each of the operations.
If a blocking operation is encountered, then the process is aborted and all operations are reversed. -EAGAIN is returned if IPC_NOWAIT is set. Otherwise 1 is returned to indicate that the sequence of semaphore operations is blocked.
If a semaphore value is adjusted beyond system limits, then then all operations are reversed, and -ERANGE is returned.
If all operations in the sequence succeed, and the do_undo
parameter is non-zero, then all operations are reversed, and 0
is returned. If the do_undo parameter is zero, then all operations
succeeded and remain in force, and the sem_otime, field of the
semaphore set is updated.
sem_revalidate() is called when the global semaphores spinlock has been temporarily dropped and needs to be locked again. It is called by semctl_main() and alloc_undo(). It validates the semaphore ID and permissions and on success, returns with the global semaphores spinlock locked.
freeundos() traverses the process undo list in search of the desired undo structure. If found, the undo structure is removed from the list and freed. A pointer to the next undo structure on the process list is returned.
alloc_undo() expects to be called with the global semaphores spinlock locked. In the case of an error, it returns with it unlocked.
The global semaphores spinlock is unlocked, and kmalloc() is called to allocate sufficient memory for both the sem_undo structure, and also an array of one adjustment value for each semaphore in the set. On success, the global spinlock is regained with a call to sem_revalidate().
The new semundo structure is then initialized, and the address of this structure is placed at the address provided by the caller. The new undo structure is then placed at the head of undo list for the current task.
sem_exit() is called by do_exit(), and is responsible for executing all of the undo adjustments for the exiting task.
If the current process was blocked on a semaphore, then it is removed from the sem_queue list while holding the global semaphores spinlock.
The undo list for the current task is then traversed, and the following operations are performed while holding and releasing the the global semaphores spinlock around the processing of each element of the list. The following operations are performed for each of the undo elements:
sem_otime parameter of the semaphore set is updated.When the processing of the list is complete, the current->semundo value is cleared.
The entire call to sys_msgget() is protected by the global message queue semaphore ( msg_ids.sem).
In the case where a new message queue must be created, the newque() function is called to create and initialize a new message queue, and the new queue ID is returned to the caller.
If a key value is provided for an existing message queue, then ipc_findkey() is called to look up the corresponding index in the global message queue descriptor array (msg_ids.entries). The parameters and permissions of the caller are verified before returning the message queue ID. The look up operation and verification are performed while the global message queue spinlock(msg_ids.ary) is held.
The parameters passed to sys_msgctl() are: a message
queue ID (msqid), the operation
(cmd), and a pointer to a user space buffer of type
msgid_ds
(buf). Six operations are
provided in this function: IPC_INFO, MSG_INFO,IPC_STAT,
MSG_STAT, IPC_SET and IPC_RMID. The message queue
ID and the operation parameters are validated; then, the operation(cmd)
is performed as follows:
The global message queue information is copied to user space.
A temporary buffer of type struct msqid64_ds is initialized and the global message queue spinlock is locked. After verifying the access permissions of the calling process, the message queue information associated with the message queue ID is loaded into the temporary buffer, the global message queue spinlock is unlocked, and the contents of the temporary buffer are copied out to user space by copy_msqid_to_user().
The user data is copied in via copy_msqid_to_user(). The global message queue semaphore and spinlock are obtained and released at the end. After the message queue ID and the current process access permissions are validated, the message queue information is updated with the user provided data. Later, expunge_all() and ss_wakeup() are called to wake up all processes sleeping on the receiver and sender waiting queues of the message queue. This is because some receivers may now be excluded by stricter access permissions and some senders may now be able to send the message due to an increased queue size.
The global message queue semaphore is obtained and the global message queue spinlock is locked. After validating the message queue ID and the current task access permissions, freeque() is called to free the resources related to the message queue ID. The global message queue semaphore and spinlock are released.
sys_msgsnd() receives as parameters a message queue ID
(msqid), a pointer to a buffer of type
struct msg_msg
(msgp), the size of the message to be sent
(msgsz), and a flag indicating wait vs.
not wait (msgflg). There are two task waiting
queues and one message waiting queue associated with the message
queue ID. If there is a task in the receiver waiting queue
that is waiting for this message, then the message is
delivered directly to the receiver, and the receiver is
awakened. Otherwise, if there is enough space available in
the message waiting queue, the message is saved in this
queue. As a last resort, the sending task enqueues itself
on the sender waiting queue. A more in-depth discussion of the
operations performed by sys_msgsnd() follows:
msg of type
struct msg_msg.
The message type and message size fields
of msg are also initialized.msgflg the global message
queue spinlock is unlocked, the memory
resources for the message are freed, and EAGAIN
is returned.msg into the message waiting
queue(msq->q_messages). Updates the
q_cbytes and
the q_qnum fields of the message
queue descriptor, as well as the global variables
msg_bytes and
msg_hdrs, which indicate the total
number of bytes used for messages and the total number
of messages system wide.q_lspid
and the q_stime fields
of the message queue descriptor and releases the global
message queue spinlock.The sys_msgrcv() function receives as parameters
a message queue ID
(msqid), a pointer to a buffer of type
msg_msg
(msgp), the desired
message size(msgsz), the message type
(msgtyp), and the flags
(msgflg). It searches the message waiting queue
associated with the message queue ID, finds the first
message in the queue which matches the request type, and
copies it into the given user buffer. If no such message
is found in the message waiting queue, the requesting task
is enqueued into the receiver waiting queue until the
desired message is available. A more in-depth discussion of the
operations performed by sys_msgrcv() follows:
msgtyp. sys_msgrcv() then locks
the global message
queue spinlock and obtains the message queue descriptor
associated with the message queue ID. If no such
message queue exists, it returns EINVAL.msgtyp is searched.msgflg
indicates no error allowed, unlocks the global
message queue spinlock and returns E2BIG.msgflg
is checked. If IPC_NOWAIT is set, then the global message
queue spinlock is unlocked and ENOMSG is returned. Otherwise,
the receiver is enqueued on the receiver waiting queue as
follows:
msr is allocated and is
added to the head of waiting queue.r_tsk field of msr
is set to current task.r_msgtype and
r_mode fields are
initialized with the desired message type and
mode respectively.msgflg indicates
MSG_NOERROR, then the r_maxsize field of
msr is set to be the
value of msgsz otherwise
it is set to be INT_MAX.r_msg field
is initialized to indicate that
no message has been received yet.r_msg field of
msr is checked. This field is used to
store the pipelined message or in the case of an error,
to store the error status.
If the r_msg field is filled
with the desired message, then go to the
last step Otherwise,
the global message queue spinlock is locked again.r_msg field is
re-checked to see if the message was received while
waiting for the spinlock. If the message has been
received, the
last step
occurs.r_msg field remains
unchanged, then the task was
awakened in order to retry. In this case,
msr is dequeued. If there is a
signal pending for the task, then the global message
queue spinlock is unlocked and EINTR is returned.
Otherwise, the function needs to go
back and retry.r_msg field shows
that an error occurred
while sleeping, the global message queue spinlock
is unlocked and the error is returned.msp is valid, message type is loaded
into the mtype field of
msp,and
store_msg()
is invoked to copy the message contents to
the mtext field of
msp. Finally the memory for the message is
freed by function
free_msg().Data structures for message queues are defined in msg.c.
/* one msq_queue structure for each present queue on the system */
struct msg_queue {
struct kern_ipc_perm q_perm;
time_t q_stime; /* last msgsnd time */
time_t q_rtime; /* last msgrcv time */
time_t q_ctime; /* last change time */
unsigned long q_cbytes; /* current number of bytes on queue */
unsigned long q_qnum; /* number of messages in queue */
unsigned long q_qbytes; /* max number of bytes on queue */
pid_t q_lspid; /* pid of last msgsnd */
pid_t q_lrpid; /* last receive pid */
struct list_head q_messages;
struct list_head q_receivers;
struct list_head q_senders;
};
/* one msg_msg structure for each message */
struct msg_msg {
struct list_head m_list;
long m_type;
int m_ts; /* message text size */
struct msg_msgseg* next;
/* the actual message follows immediately */
};
/* message segment for each message */
struct msg_msgseg {
struct msg_msgseg* next;
/* the next part of the message follows immediately */
};
/* one msg_sender for each sleeping sender */
struct msg_sender {
struct list_head list;
struct task_struct* tsk;
};
/* one msg_receiver structure for each sleeping receiver */
struct msg_receiver {
struct list_head r_list;
struct task_struct* r_tsk;
int r_mode;
long r_msgtype;
long r_maxsize;
struct msg_msg* volatile r_msg;
};
struct msqid64_ds {
struct ipc64_perm msg_perm;
__kernel_time_t msg_stime; /* last msgsnd time */
unsigned long __unused1;
__kernel_time_t msg_rtime; /* last msgrcv time */
unsigned long __unused2;
__kernel_time_t msg_ctime; /* last change time */
unsigned long __unused3;
unsigned long msg_cbytes; /* current number of bytes on queue */
unsigned long msg_qnum; /* number of messages in queue */
unsigned long msg_qbytes; /* max number of bytes on queue */
__kernel_pid_t msg_lspid; /* pid of last msgsnd */
__kernel_pid_t msg_lrpid; /* last receive pid */
unsigned long __unused4;
unsigned long __unused5;
};
struct msqid_ds {
struct ipc_perm msg_perm;
struct msg *msg_first; /* first message on queue,unused */
struct msg *msg_last; /* last message in queue,unused */
__kernel_time_t msg_stime; /* last msgsnd time */
__kernel_time_t msg_rtime; /* last msgrcv time */
__kernel_time_t msg_ctime; /* last change time */
unsigned long msg_lcbytes; /* Reuse junk fields for 32 bit */
unsigned long msg_lqbytes; /* ditto */
unsigned short msg_cbytes; /* current number of bytes on queue */
unsigned short msg_qnum; /* number of messages in queue */
unsigned short msg_qbytes; /* max number of bytes on queue */
__kernel_ipc_pid_t msg_lspid; /* pid of last msgsnd */
__kernel_ipc_pid_t msg_lrpid; /* last receive pid */
};
struct msq_setbuf {
unsigned long qbytes;
uid_t uid;
gid_t gid;
mode_t mode;
};
newque() allocates the memory for a new message queue descriptor ( struct msg_queue) and then calls ipc_addid(), which reserves a message queue array entry for the new message queue descriptor. The message queue descriptor is initialized as follows:
q_stime and q_rtime fields of the message
queue descriptor are initialized as 0. The q_ctime
field is set to be CURRENT_TIME.q_qbytes) is set to be MSGMNB,
and the number of bytes currently used by the queue
(q_cbytes) is initialized as 0.q_messages),
the receiver waiting queue (q_receivers),
and the sender waiting queue (q_senders)
are each initialized as empty.All the operations following the call to ipc_addid() are performed while holding the global message queue spinlock. After unlocking the spinlock, newque() calls msg_buildid(), which maps directly to ipc_buildid(). ipc_buildid() uses the index of the message queue descriptor to create a unique message queue ID that is then returned to the caller of newque().
When a message queue is going to be removed, the freeque() function is called. This function assumes that the global message queue spinlock is already locked by the calling function. It frees all kernel resources associated with that message queue. First, it calls ipc_rmid() (via msg_rmid()) to remove the message queue descriptor from the array of global message queue descriptors. Then it calls expunge_all to wake up all receivers and ss_wakeup() to wake up all senders sleeping on this message queue. Later the global message queue spinlock is released. All messages stored in this message queue are freed and the memory for the message queue descriptor is freed.
ss_wakeup() wakes up all the tasks waiting in the given message sender waiting queue. If this function is called by freeque(), then all senders in the queue are dequeued.
ss_add() receives as parameters a message queue descriptor
and a message sender data structure. It fills the
tsk field of the message sender data
structure with the current process, changes the status of
current process to TASK_INTERRUPTIBLE,
then inserts the message sender data structure at the head of
the sender waiting queue of the given message queue.
If the given message sender data structure
(mss) is still in the associated sender
waiting queue, then ss_del() removes
mss from the queue.
expunge_all() receives as parameters a message queue
descriptor(msq) and an integer value
(res) indicating the reason for waking up the
receivers. For each sleeping receiver associated with
msq, the r_msg
field is set to the indicated
wakeup reason (res), and the associated receiving
task is awakened. This function is called when a message queue is
removed or a message control operation has been performed.
When a process sends a message, the sys_msgsnd() function first invokes the load_msg() function to load the message from user space to kernel space. The message is represented in kernel memory as a linked list of data blocks. Associated with the first data block is a msg_msg structure that describes the overall message. The datablock associated with the msg_msg structure is limited to a size of DATA_MSG_LEN. The data block and the structure are allocated in one contiguous memory block that can be as large as one page in memory. If the full message will not fit into this first data block, then additional data blocks are allocated and are organized into a linked list. These additional data blocks are limited to a size of DATA_SEG_LEN, and each include an associated msg_msgseg) structure. The msg_msgseg structure and the associated data block are allocated in one contiguous memory block that can be as large as one page in memory. This function returns the address of the new msg_msg structure on success.
The store_msg() function is called by sys_msgrcv() to reassemble a received message into the user space buffer provided by the caller. The data described by the msg_msg structure and any msg_msgseg structures are sequentially copied to the user space buffer.
The free_msg() function releases the memory for a message data structure msg_msg, and the message segments.
convert_mode() is called by
sys_msgrcv().
It receives as parameters the address of the specified message
type (msgtyp) and a flag (msgflg).
It returns the search mode to the caller based on the value of
msgtyp and msgflg. If
msgtyp is null, then SEARCH_ANY is returned.
If msgtyp is less than 0, then msgtyp is
set to it's absolute value and SEARCH_LESSEQUAL is returned.
If MSG_EXCEPT is specified in msgflg, then SEARCH_NOTEQUAL is returned.
Otherwise SEARCH_EQUAL is returned.
The testmsg() function checks whether a message meets the criteria specified by the receiver. It returns 1 if one of the following conditions is true:
pipelined_send() allows a process to directly send a message
to a waiting receiver rather than deposit the message in the
associated message waiting queue. The
testmsg() function is
invoked to find the first receiver which is waiting for the
given message. If found, the waiting receiver is removed from
the receiver waiting queue, and the associated receiving task is
awakened. The message is stored in the r_msg
field of the receiver, and 1 is returned. In the case where no
receiver is waiting for the message, 0 is returned.
In the process of searching for a receiver, potential
receivers may be found which have requested a size that is too small
for the given message. Such receivers are removed from the queue,
and are awakened with an error status of E2BIG, which is stored in the
r_msg field. The search then continues until
either a valid receiver is found, or the queue is exhausted.
copy_msqid_to_user() copies the contents of a kernel buffer to the user buffer. It receives as parameters a user buffer, a kernel buffer of type msqid64_ds, and a version flag indicating the new IPC version vs. the old IPC version. If the version flag equals IPC_64, then copy_to_user() is invoked to copy from the kernel buffer to the user buffer directly. Otherwise a temporary buffer of type struct msqid_ds is initialized, and the kernel data is translated to this temporary buffer. Later copy_to_user() is called to copy the contents of the temporary buffer to the user buffer.
The function copy_msqid_from_user() receives as parameters
a kernel message buffer of type struct msq_setbuf, a user buffer
and a version flag indicating the new IPC version vs. the old IPC
version. In the case of the new IPC version, copy_from_user()
is called to copy the contents of the user buffer
to a temporary buffer of type
msqid64_ds.
Then, the qbytes,uid, gid,
and mode fields of the kernel buffer are
filled with the values of the
corresponding fields from the temporary buffer. In the case of the
old IPC version, a temporary buffer of type struct
msqid_ds is used instead.
The entire call to sys_shmget() is protected by the global shared memory semaphore.
In the case where a new shared memory segment must be created, the newseg() function is called to create and initialize a new shared memory segment. The ID of the new segment is returned to the caller.
In the case where a key value is provided for an existing shared memory segment, the corresponding index in the shared memory descriptors array is looked up, and the parameters and permissions of the caller are verified before returning the shared memory segment ID. The look up operation and verification are performed while the global shared memory spinlock is held.
A temporary shminfo64 buffer is loaded with system-wide shared memory parameters and is copied out to user space for access by the calling application.
The global shared memory semaphore and the global shared
memory spinlock are held while gathering system-wide statistical
information for shared memory. The
shm_get_stat() function is called
to calculate both the number of shared memory pages that are
resident in memory and the number of shared memory pages that are
swapped out. Other statistics include the total number of shared
memory pages and the number of shared memory segments in use.
The counts of swap_attempts and swap_successes
are hard-coded to zero. These statistics are stored in a temporary
shm_info buffer and copied out
to user space for the calling application.
For SHM_STAT and IPC_STATA, a temporary buffer of type struct shmid64_ds is initialized, and the global shared memory spinlock is locked.
For the SHM_STAT case, the shared memory segment ID parameter is expected to be a straight index (i.e. 0 to n where n is the number of shared memory IDs in the system). After validating the index, ipc_buildid() is called (via shm_buildid()) to convert the index into a shared memory ID. In the passing case of SHM_STAT, the shared memory ID will be the return value. Note that this is an undocumented feature, but is maintained for the ipcs(8) program.
For the IPC_STAT case, the shared memory segment ID parameter is expected to be an ID that was generated by a call to shmget(). The ID is validated before proceeding. In the passing case of IPC_STAT, 0 will be the return value.
For both SHM_STAT and IPC_STAT, the access permissions of the caller are verified. The desired statistics are loaded into the temporary buffer and then copied out to the calling application.
After validating access permissions, the global shared memory spinlock is locked, and the shared memory segment ID is validated. For both SHM_LOCK and SHM_UNLOCK, shmem_lock() is called to perform the function. The parameters for shmem_lock() identify the function to be performed.
During IPC_RMID the global shared memory semaphore and the global shared memory spinlock are held throughout this function. The Shared Memory ID is validated, and then if there are no current attachments, shm_destroy() is called to destroy the shared memory segment. Otherwise, the SHM_DEST flag is set to mark it for destruction, and the IPC_PRIVATE flag is set to prevent other processes from being able to reference the shared memory ID.
After validating the shared memory segment ID and the user
access permissions, the uid, gid, and mode flags of the
shared memory segment are updated with the user data. The
shm_ctime field is also updated. These changes are made
while holding the global shared memory semaphore and the
global share memory spinlock.
sys_shmat() takes as parameters, a shared memory segment ID,
an address at which the shared memory segment should be
attached(shmaddr), and flags which will be described below.
If shmaddr is non-zero, and the SHM_RND flag is
specified, then shmaddr is rounded down to a multiple of
SHMLBA. If shmaddr is not a multiple of SHMLBA and SHM_RND
is not specified, then EINVAL is returned.
The access permissions of the caller are validated and
the shm_nattch field for the shared memory segment is
incremented. Note that this increment guarantees that the
attachment count is non-zero and prevents the shared memory
segment from being destroyed during the process of attaching
to the segment. These operations are performed while holding the
global shared memory spinlock.
The do_mmap() function is called to create a virtual memory
mapping to the shared memory segment pages. This is done while
holding the mmap_sem semaphore of the current task. The
MAP_SHARED flag is passed to do_mmap(). If an address was
provided by the caller, then the MAP_FIXED flag is also passed
to do_mmap(). Otherwise, do_mmap() will select the virtual
address at which to map the shared memory segment.
NOTE
shm_inc() will be invoked within the do_mmap()
function call via the shm_file_operations structure. This
function is called to set the PID, to set the current time, and
to increment the number of attachments to this shared memory
segment.
After the call to do_mmap(), the global shared memory semaphore and the global shared memory spinlock are both obtained. The attachment count is then decremented. The the net change to the attachment count is 1 for a call to shmat() because of the call to shm_inc(). If, after decrementing the attachment count, the resulting count is found to be zero, and if the segment is marked for destruction (SHM_DEST), then shm_destroy() is called to release the shared memory segment resources.
Finally, the virtual address at which the shared memory is mapped is returned to the caller at the user specified address. If an error code had been returned by do_mmap(), then this failure code is passed on as the return value for the system call.
The global shared memory semaphore is held while performing
sys_shmdt(). The mm_struct of the current
process is searched for the vm_area_struct associated with
the shared memory address. When it is found, do_munmap() is
called to undo the virtual address mapping for the shared memory segment.
Note also that do_munmap() performs a call-back to shm_close(), which performs the shared-memory book keeping functions, and releases the shared memory segment resources if there are no other attachments.
sys_shmdt() unconditionally returns 0.
struct shminfo64 {
unsigned long shmmax;
unsigned long shmmin;
unsigned long shmmni;
unsigned long shmseg;
unsigned long shmall;
unsigned long __unused1;
unsigned long __unused2;
unsigned long __unused3;
unsigned long __unused4;
};
struct shm_info {
int used_ids;
unsigned long shm_tot; /* total allocated shm */
unsigned long shm_rss; /* total resident shm */
unsigned long shm_swp; /* total swapped shm */
unsigned long swap_attempts;
unsigned long swap_successes;
};
struct shmid_kernel /* private to the kernel */
{
struct kern_ipc_perm shm_perm;
struct file * shm_file;
int id;
unsigned long shm_nattch;
unsigned long shm_segsz;
time_t shm_atim;
time_t shm_dtim;
time_t shm_ctim;
pid_t shm_cprid;
pid_t shm_lprid;
};
struct shmid64_ds {
struct ipc64_perm shm_perm; /* operation perms */
size_t shm_segsz; /* size of segment (bytes) */
__kernel_time_t shm_atime; /* last attach time */
unsigned long __unused1;
__kernel_time_t shm_dtime; /* last detach time */
unsigned long __unused2;
__kernel_time_t shm_ctime; /* last change time */
unsigned long __unused3;
__kernel_pid_t shm_cpid; /* pid of creator */
__kernel_pid_t shm_lpid; /* pid of last operator */
unsigned long shm_nattch; /* no. of current attaches */
unsigned long __unused4;
unsigned long __unused5;
};
struct shmem_inode_info {
spinlock_t lock;
unsigned long max_index;
swp_entry_t i_direct[SHMEM_NR_DIRECT]; /* for the first blocks */
swp_entry_t **i_indirect; /* doubly indirect blocks */
unsigned long swapped;
int locked; /* into memory */
struct list_head list;
};
The newseg() function is called when a new shared memory
segment needs to be created. It acts on three parameters for
the new segment the key, the flag, and the size. After
validating that the size of the shared memory segment to be
created is between SHMMIN and SHMMAX and that the total number
of shared memory segments does not exceed SHMALL, it allocates
a new shared memory segment descriptor.
The
shmem_file_setup()
function is invoked later to create an unlinked file of type
tmpfs. The returned file pointer is saved in the shm_file field
of the associated shared memory segment descriptor. The files
size is set to be the same as the size of the segment. The
new shared memory segment descriptor is initialized and inserted
into the global IPC shared memory descriptors array. The shared
memory segment ID is created by shm_buildid()
(via
ipc_buildid()).
This segment ID is saved in the id field of the shared memory
segment descriptor, as well as in the i_ino field of the associated
inode. In addition, the address of the shared memory operations
defined in structure shm_file_operation is stored in the associated
file. The value of the global variable shm_tot, which indicates
the total number of shared memory segments system wide, is also
increased to reflect this change. On success, the segment ID is
returned to the caller application.
shm_get_stat() cycles through all of the shared memory structures, and calculates the total number of memory pages in use by shared memory and the total number of shared memory pages that are swapped out. There is a file structure and an inode structure for each shared memory segment. Since the required data is obtained via the inode, the spinlock for each inode structure that is accessed is locked and unlocked in sequence.
shmem_lock() receives as parameters a pointer to the shared memory segment descriptor and a flag indicating lock vs. unlock.The locking state of the shared memory segment is stored in an associated inode. This state is compared with the desired locking state; shmem_lock() simply returns if they match.
While holding the semaphore of the associated inode, the locking state of the inode is set. The following list of items occur for each page in the shared memory segment:
During shm_destroy() the total number of shared memory pages
is adjusted to account for the removal of the shared memory segment.
ipc_rmid() is called
(via shm_rmid()) to remove the Shared Memory ID.
shmem_lock is
called to unlock the shared memory pages, effectively decrementing
the reference counts to zero for each page. fput() is called to
decrement the usage counter f_count for the associated file object,
and if necessary, to release the file object resources. kfree() is
called to free the shared memory segment descriptor.
shm_inc() sets the PID, sets the current time, and increments the number of attachments for the given shared memory segment. These operations are performed while holding the global shared memory spinlock.
shm_close() updates the shm_lprid and the shm_dtim fields
and decrements the number of attached shared memory segments. If
there are no other attachments to the shared memory segment,
then
shm_destroy() is called to
release the shared memory segment resources. These operations are
all performed while holding both the global shared memory semaphore
and the global shared memory spinlock.
The function shmem_file_setup() sets up an unlinked file living
in the tmpfs file system with the given name and size. If there
are enough systen memory resource for this file, it creates a new
dentry under the mount root of tmpfs, and allocates a new file
descriptor and a new inode object of tmpfs type. Then it associates
the new dentry object with the new inode object by calling
d_instantiate() and saves the address of the dentry object in the
file descriptor. The i_size field of the inode object is set to
be the file size and the i_nlink field is set to be 0 in order to
mark the inode unlinked. Also, shmem_file_setup() stores the
address of the shmem_file_operations structure in the f_op field,
and initializes f_mode and f_vfsmnt fields of the file descriptor
properly. The function shmem_truncate() is called to complete the
initialization of the inode object. On success, shmem_file_setup()
returns the new file descriptor.
The semaphores, messages, and shared memory mechanisms of Linux are built on a set of common primitives. These primitives are described in the sections below.
If the memory allocation is greater than PAGE_SIZE, then vmalloc() is used to allocate memory. Otherwise, kmalloc() is called with GFP_KERNEL to allocate the memory.
When a new semaphore set, message queue, or shared memory segment is added, ipc_addid() first calls grow_ary() to insure that the size of the corresponding descriptor array is sufficiently large for the system maximum. The array of descriptors is searched for the first unused element. If an unused element is found, the count of descriptors which are in use is incremented. The kern_ipc_perm structure for the new resource descriptor is then initialized, and the array index for the new descriptor is returned. When ipc_addid() succeeds, it returns with the global spinlock for the given IPC type locked.
ipc_rmid() removes the IPC descriptor from the global descriptor array of the IPC type, updates the count of IDs which are in use, and adjusts the maximum ID in the corresponding descriptor array if necessary. A pointer to the IPC descriptor associated with given IPC ID is returned.
ipc_buildid() creates a unique ID to be associated with each descriptor within a given IPC type. This ID is created at the time a new IPC element is added (e.g. a new shared memory segment or a new semaphore set). The IPC ID converts easily into the corresponding descriptor array index. Each IPC type maintains a sequence number which is incremented each time a descriptor is added. An ID is created by multiplying the sequence number with SEQ_MULTIPLIER and adding the product to the descriptor array index. The sequence number used in creating a particular IPC ID is then stored in the corresponding descriptor. The existence of the sequence number makes it possible to detect the use of a stale IPC ID.
ipc_checkid() divides the given IPC ID by the SEQ_MULTIPLIER and compares the quotient with the seq value saved corresponding descriptor. If they are equal, then the IPC ID is considered to be valid and 1 is returned. Otherwise, 0 is returned.
grow_ary() handles the possibility that the maximum (tunable) number of IDs for a given IPC type can be dynamically changed. It enforces the current maximum limit so that it is no greater than the permanent system limit (IPCMNI) and adjusts it down if necessary. It also insures that the existing descriptor array is large enough. If the existing array size is sufficiently large, then the current maximum limit is returned. Otherwise, a new larger array is allocated, the old array is copied into the new array, and the old array is freed. The corresponding global spinlock is held when updating the descriptor array for the given IPC type.
ipc_findkey() searches through the descriptor array of the specified ipc_ids object, and searches for the specified key. Once found, the index of the corresponding descriptor is returned. If the key is not found, then -1 is returned.
ipcperms() checks the user, group, and other permissions for access to the IPC resources. It returns 0 if permission is granted and -1 otherwise.
ipc_lock() takes an IPC ID as one of its parameters. It locks the global spinlock for the given IPC type, and returns a pointer to the descriptor corresponding to the specified IPC ID.
ipc_unlock() releases the global spinlock for the indicated IPC type.
ipc_lockall() locks the global spinlock for the given IPC mechanism (i.e. shared memory, semaphores, and messaging).
ipc_unlockall() unlocks the global spinlock for the given IPC mechanism (i.e. shared memory, semaphores, and messaging).
ipc_get() takes a pointer to a particular IPC type (i.e. shared memory, semaphores, or message queues) and a descriptor ID, and returns a pointer to the corresponding IPC descriptor. Note that although the descriptors for each IPC type are of different data types, the common kern_ipc_perm structure type is embedded as the first entity in every case. The ipc_get() function returns this common data type. The expected model is that ipc_get() is called through a wrapper function (e.g. shm_get()) which casts the data type to the correct descriptor data type.
ipc_parse_version() removes the IPC_64 flag from the command if it is present and returns either IPC_64 or IPC_OLD.
The semaphores, messages, and shared memory mechanisms all make use of the following common structures:
Each of the IPC descriptors has a data object of this type as the first element. This makes it possible to access any descriptor from any of the generic IPC functions using a pointer of this data type.
/* used by in-kernel data structures */
struct kern_ipc_perm {
key_t key;
uid_t uid;
gid_t gid;
uid_t cuid;
gid_t cgid;
mode_t mode;
unsigned long seq;
};
The ipc_ids structure describes the common data for semaphores,
message queues, and shared memory. There are three global instances of
this data structure-- semid_ds,
msgid_ds and shmid_ds-- for
semaphores, messages and shared memory respectively. In each
instance, the sem semaphore is used to
protect access to the structure.
The entries field points to an IPC
descriptor array, and the
ary spinlock protects access to this array. The
seq field is a global sequence number which will
be incremented when a new IPC resource is created.
struct ipc_ids {
int size;
int in_use;
int max_id;
unsigned short seq;
unsigned short seq_max;
struct semaphore sem;
spinlock_t ary;
struct ipc_id* entries;
};
An array of struct ipc_id exists in each instance of the ipc_ids structure. The array is dynamically allocated and may be replaced with larger array by grow_ary() as required. The array is sometimes referred to as the descriptor array, since the kern_ipc_perm data type is used as the common descriptor data type by the IPC generic functions.
struct ipc_id {
struct kern_ipc_perm* p;
};