new_node makes the assumption that when it does last_dir on a path, a
successive advance would not yield a lock on a vmnt, because last_dir
already locked the vmnt. This is true except when last_dir resolves
to a directory on the parent vmnt of the file that was the result of
advance. For example,
# cd /
# echo foo > home
where home is on a different (sub) partition than / is (default
install). last_dir would resolve to / and advance would resolve to
/home.
With this change, last_dir resolves to the root node on the /home
partition, making the assumption valid again.
. 'anonymous' cache blocks (retrieved with NO_DEV as dev
parameter) were used to implement read()s from holes in
inodes that should return zeroes
. this is an awkward special case in the cache code though
and there's a more direct way to implement the same functionality:
instead of copying from a new, anonymous, zero block, to
the user target buffer, simply sys_safememset the user target
buffer directly. as this was the only use of this feature,
this is all that's needed to simplify the cache code a little.
- CHOOSETRAP define makes impossible to use some common words
like send, receive and notify in any other context, for
instance as members or structures
- any reasonable compiler inlines the static inline functions so
no extra function call overhead is introduced by this change
- this gets us back to the situation before the SYSCALL/SYSENTER
change. It is not perfect, but it used to work and still does.
The tested targets are the followgin ones:
* tools
* distribution
* sets
* release
The remaining NetBSD targets have not been disabled nor tested
*at all*. Try them at your own risk, they may reboot the earth.
For all compliant Makefiles, objects and generated files are put in
MAKEOBJDIR, which means you can now keep objects between two branch
switching. Same for DESTDIR, please refer to build.sh options.
Regarding new or modifications of Makefiles a few things:
* Read share/mk/bsd.README
* If you add a subdirectory, add a Makefile in it, and have it called
by the parent through the SUBDIR variable.
* Do not add arbitrary inclusion which crosses to another branch of
the hierarchy; If you can't do without it, put a comment on why.
If possible, do not use inclusion at all.
* Use as much as possible the infrastructure, it is here to make
life easier, do not fight it.
Sets and package are now used to track files.
We have one set called "minix", composed of one package called "minix-sys"
The VFS/FS protocol does not require the file server to supply a
special device node number in response to a REQ_CREATE request, as
this call creates only regular files. Therefore, VFS should not
erroneously save this piece of information from the REQ_CREATE reply
either.
Upon reboot VFS semi-exits all processes and unmounts the file system.
However, upon unmount, exiting FUSE file systems might need service from
the file system (due to libc). As the FUSE process is halfway the exit
procedure, it doesn't have a valid root directory and working directory.
Trying to do system calls then triggers a sanity check in VFS.
This fix first exits normal processes which should then allow for
unmounting FUSE file systems. Then VFS exits all processes including
File Servers and unmounts the rest of the file system.
There is a deadlock vulnerability when there are no worker threads
available and all of them blocked on a worker thread that's waiting for a
reply from a driver or a reply from an FS that needs to make a back call. In
these cases the deadlock resolver thread should kick in, but didn't in all
cases. Moreover, POSIX calls from File Servers weren't handled properly
anymore, which also could lead to deadlocks.
. also make other out-of-memory conditions less fatal
. add a test case for a user program using all the memory
it can
. remove some diagnostic prints for situations that are normal
when running out of memory so running the test isn't noisy
Add primary cache management feature to libminixfs as mfs and ext2
currently do separately, remove cache code from mfs and ext2, and make
them use the libminixfs interface. This makes all fields of the buf
struct private to libminixfs and FS clients aren't supposed to access
them at all. Only the opaque 'void *data' field (the FS block contents,
used to be called bp) is to be accessed by the FS client.
The main purpose is to implement the interface to the 2ndary vm cache
just once, get rid of some code duplication, and add a little
abstraction to reduce the code inertia of the whole caching business.
Some minor sanity checking and prohibition done by mfs in this code
as removed from the generic primary cache code as a result:
- checking all inodes are not in use when allocating/resizing
the cache
- checking readonly filesystems aren't written to
- checking the superblock isn't written to on mounted filesystems
The minixfslib code relies on fs_blockstats() in the client filesystem to
return some FS usage information.
Introduce explicit abstractions for different mapping types,
handling the instantiation, forking, pagefaults and freeing of
anonymous memory, direct physical mappings, shared memory and
physically contiguous anonymous memory as separate types, making
region.c more generic.
Also some other genericification like merging the 3 munmap cases
into one.
COW and SMAP safemap code is still implicit in region.c.
The check_bsf() macro uses assert(mutex_trylock(&bsf_lock)) and
assumes bsf_lock is locked afterwards. This breaks when compiling
with NOASSERTS="yes". Also: macro to function transition.
. add cpufeature detection of both
. use it for both ipc and kernelcall traps, using a register
for call number
. SYSENTER/SYSCALL does not save any context, therefore userland
has to save it
. to accomodate multiple kernel entry/exit types, the entry
type is recorded in the process struct. hitherto all types
were interrupt (soft int, exception, hard int); now SYSENTER/SYSCALL
is new, with the difference that context is not fully restored
from proc struct when running the process again. this can't be
done as some information is missing.
. complication: cases in which the kernel has to fully change
process context (i.e. sigreturn). in that case the exit type
is changed from SYSENTER/SYSEXIT to soft-int (i.e. iret) and
context is fully restored from the proc struct. this does mean
the PC and SP must change, as the sysenter/sysexit userland code
will otherwise try to restore its own context. this is true in the
sigreturn case.
. override all usage by setting libc_ipc=1
. whenever this function is called, pm will expect
the process to be cleaned up
. so don't abort the process entirely on error
. fixes a later 'forking on top of in-use child' vfs panic
fixes an assert() firing when starting X. thanks to the report by pikpik.
. NO_MEM was 0, which is actually an existing piece
of physical memory. it can't be allocated because it's reserved
for bios data (by the kernel), but it can be mapped in (e.g.
by X), causing sanity check disaster.
. NONCONTIGUOUS is also obsolete as all allocations are single-page
now, i.e. NONCONTIGUOUS is really the default and only mode.
complete munmap implementation; single-page references made
a general munmap() implementation possible to write cleanly.
. memory: let the MIOCRAMSIZE ioctl set the imgrd device
size (but only to 0)
. let the ramdisk command set sizes to 0
. use this command to set /dev/imgrd to 0 after mounting /usr
in /etc/rc, so the boot time ramdisk is freed (about 4MB
currently)
. only reference single pages in process data structures
to simplify page faults, copy-on-write, etc.
. this breaks the secondary cache for objects that are
not one-page-sized; restored in a next commit
By decoupling synchronous drivers from VFS, we are a big step closer to
supporting driver crashes under all circumstances. That is, VFS can't
become stuck on IPC with a synchronous driver (e.g., INET) and can
recover from crashing block drivers during open/close/ioctl or during
communication with an FS.
In order to maintain serialized communication with a synchronous driver,
the communication is wrapped by a mutex on a per driver basis (not major
numbers as there can be multiple majors with identical endpoints). Majors
that share a driver endpoint point to a single mutex object.
In order to support crashes from block drivers, the file reopen tactic
had to be changed; first reopen files associated with the crashed
driver, then send the new driver endpoint to FSes. This solves a
deadlock between the FS and the block driver;
- VFS would send REQ_NEW_DRIVER to an FS, but he FS only receives it
after retrying the current request to the newly started driver.
- The block driver would refuse the retried request until all files
had been reopened.
- VFS would reopen files only after getting a reply from the initial
REQ_NEW_DRIVER.
When a character special driver crashes, all associated files have to
be marked invalid and closed (or reopened if flagged as such). However,
they can only be closed if a thread holds exclusive access to it. To
obtain exclusive access, the worker thread (which handles the new driver
endpoint event from DS) schedules a new job to garbage collect invalid
files. This way, we can signal the worker thread that was talking to the
crashed driver and will release exclusive access to a file associated
with the crashed driver and prevent the garbage collecting worker thread
from dead locking on that file.
Also, when a character special driver crashes, RS will unmap the driver
and remap it upon restart. During unmapping, associated files are marked
invalid instead of waiting for an endpoint up event from DS, as that
event might come later than new read/write/select requests and thus
cause confusion in the freshly started driver.
When locking a filp, the usage counters are no longer checked. The usage
counter can legally go down to zero during filp invalidation while there
are locks pending.
DS events are handled by a separate worker thread instead of the main
thread as reopening files could lead to another crash and a stuck thread.
An additional worker thread is then necessary to unlock it.
Finally, with everything asynchronous a race condition in do_select
surfaced. A select entry was only marked in use after succesfully sending
initial select requests to drivers and having to wait. When multiple
select() calls were handled there was opportunity that these entries
were overwritten. This had as effect that some select results were
ignored (and select() remained blocking instead if returning) or do_select
tried to access filps that were not present (because thrown away by
secondary select()). This bug manifested itself with sendrecs, but was
very hard to reproduce. However, it became awfully easy to trigger with
asynsends only.
Instead of using a loop to find a matching ipc (inter process
communication) system call type, the offset in the call table can be
simply calculated in constant time.
Also, when the interprocess communication server receives an ipc
system call from a process, ipc should tell VM to watch the process
only once. This patch fixes that also.
(Patch and commit message slightly edited by committer.)
. ld.so is linked at 0 but it can relocate itself; we
wish to load ld.so higher though to trap NULL dereferences.
if we know we have to execute ld.so, vfs tells libexec to put it
higher.
. done by RS to reduce/remove dependency on VM for recovery
. RS has the default stack size of 64MB since the nosegments
change, using a huge amount of unused memory to pre-allocate
. ignore these requests until actually required (i.e. being able
to survive VM crashes)
Thanks to pikpik for investigating why RS was so huge.
When VFS runs out of vnodes after closing a vnode in opcl, common_open
will try to unlock a vnode through unlock_filp that has already been
unlocked in clone_opcl. By first obtaining and locking a new vnode this
situation is prevented; if there are no free vnodes, common_open will
unlock a still locked vnode.
.enable all compile time warnings and make them errors
.refactor functions with unused parameters
.fix null pointer dereference before checking for null
.proper variable initialization
.use safe string copy functions
.fix massive memory corruption bug in fs_getdents
. map all objects named usermapped_*.o with globally visible
pages; usermapped_glo_*.o with the VM 'global' bit on, i.e.
permanently in tlb (very scarce resource!)
. added kinfo, machine, kmessages and loadinfo for a start
. modified log, tty to make use of the shared messages struct
. some strncpy/strcpy to strlcpy conversions
. new <minix/param.h> to avoid including other minix headers
that have colliding definitions with library and commands code,
causing parse warnings
. removed some dead code / assignments
This commit removes all traces of Minix segments (the text/data/stack
memory map abstraction in the kernel) and significance of Intel segments
(hardware segments like CS, DS that add offsets to all addressing before
page table translation). This ultimately simplifies the memory layout
and addressing and makes the same layout possible on non-Intel
architectures.
There are only two types of addresses in the world now: virtual
and physical; even the kernel and processes have the same virtual
address space. Kernel and user processes can be distinguished at a
glance as processes won't use 0xF0000000 and above.
No static pre-allocated memory sizes exist any more.
Changes to booting:
. The pre_init.c leaves the kernel and modules exactly as
they were left by the bootloader in physical memory
. The kernel starts running using physical addressing,
loaded at a fixed location given in its linker script by the
bootloader. All code and data in this phase are linked to
this fixed low location.
. It makes a bootstrap pagetable to map itself to a
fixed high location (also in linker script) and jumps to
the high address. All code and data then use this high addressing.
. All code/data symbols linked at the low addresses is prefixed by
an objcopy step with __k_unpaged_*, so that that code cannot
reference highly-linked symbols (which aren't valid yet) or vice
versa (symbols that aren't valid any more).
. The two addressing modes are separated in the linker script by
collecting the unpaged_*.o objects and linking them with low
addresses, and linking the rest high. Some objects are linked
twice, once low and once high.
. The bootstrap phase passes a lot of information (e.g. free memory
list, physical location of the modules, etc.) using the kinfo
struct.
. After this bootstrap the low-linked part is freed.
. The kernel maps in VM into the bootstrap page table so that VM can
begin executing. Its first job is to make page tables for all other
boot processes. So VM runs before RS, and RS gets a fully dynamic,
VM-managed address space. VM gets its privilege info from RS as usual
but that happens after RS starts running.
. Both the kernel loading VM and VM organizing boot processes happen
using the libexec logic. This removes the last reason for VM to
still know much about exec() and vm/exec.c is gone.
Further Implementation:
. All segments are based at 0 and have a 4 GB limit.
. The kernel is mapped in at the top of the virtual address
space so as not to constrain the user processes.
. Processes do not use segments from the LDT at all; there are
no segments in the LDT any more, so no LLDT is needed.
. The Minix segments T/D/S are gone and so none of the
user-space or in-kernel copy functions use them. The copy
functions use a process endpoint of NONE to realize it's
a physical address, virtual otherwise.
. The umap call only makes sense to translate a virtual address
to a physical address now.
. Segments-related calls like newmap and alloc_segments are gone.
. All segments-related translation in VM is gone (vir2map etc).
. Initialization in VM is simpler as no moving around is necessary.
. VM and all other boot processes can be linked wherever they wish
and will be mapped in at the right location by the kernel and VM
respectively.
Other changes:
. The multiboot code is less special: it does not use mb_print
for its diagnostics any more but uses printf() as normal, saving
the output into the diagnostics buffer, only printing to the
screen using the direct print functions if a panic() occurs.
. The multiboot code uses the flexible 'free memory map list'
style to receive the list of free memory if available.
. The kernel determines the memory layout of the processes to
a degree: it tells VM where the kernel starts and ends and
where the kernel wants the top of the process to be. VM then
uses this entire range, i.e. the stack is right at the top,
and mmap()ped bits of memory are placed below that downwards,
and the break grows upwards.
Other Consequences:
. Every process gets its own page table as address spaces
can't be separated any more by segments.
. As all segments are 0-based, there is no distinction between
virtual and linear addresses, nor between userspace and
kernel addresses.
. Less work is done when context switching, leading to a net
performance increase. (8% faster on my machine for 'make servers'.)
. The layout and configuration of the GDT makes sysenter and syscall
possible.
. sys_vircopy always uses D for both src and dst
. sys_physcopy uses PHYS_SEG if and only if corresponding
endpoint is NONE, so we can derive the mode (PHYS_SEG or D)
from the endpoint arg in the kernel, dropping the seg args
. fields in msg still filled in for backwards compatability,
using same NONE-logic in the library
. all invocations were S or D, so can safely be dropped
to prepare for the segmentless world
. still assign D to the SCP_SEG field in the message
to make previous kernels usable
. new mode for sys_memset: include process so memset can be
done in physical or virtual address space.
. add a mode to mmap() that lets a process allocate uninitialized
memory.
. this allows an exec()er (RS, VFS, etc.) to request uninitialized
memory from VM and selectively clear the ranges that don't come
from a file, leaving no uninitialized memory left for the process
to see.
. use callbacks for clearing the process, clearing memory in the
process, and copying into the process; so that the libexec code
can be used from rs, vfs, and in the future, kernel (to load vm)
and vm (to load boot-time processes)
. make exec() callers (i.e. vfs and rs) determine the
memory layout by explicitly reserving regions using
mmap() calls on behalf of the exec()ing process,
i.e. handling all of the exec logic, thereby eliminating
all special exec() knowledge from VM.
. the new procedure is: clear the exec()ing process
first, then call third-party mmap()s to reserve memory, then
copy the executable file section contents in, all using callbacks
tailored to the caller's way of starting an executable
. i.e. no more explicit EXEC_NEWMEM-style calls in PM or VM
as with rigid 2-section arguments
. this naturally allows generalizing exec() by simply loading
all ELF sections
. drop/merge of lots of duplicate exec() code into libexec
. not copying the code sections to vfs and into the executable
again is a measurable performance improvement (about 3.3% faster
for 'make' in src/servers/)
