. 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
. 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/)
these two functions will be used to support all exec() functionality
going into a single library shared by RS and VFS and exec() knowledge
leaving VM.
. third-party mmap: allow certain processes (VFS, RS) to
do mmap() on behalf of another process
. PROCCTL: used to free and clear a process' address space
This patch provides basic protection against damage resulting from
differently compiled servers blindly copying tables to one another.
In every getsysinfo() call, the caller is provided with the expected
size of the requested data structure. The callee fails the call if
the expected size does not match the data structure's actual size.
This patch fixes most of current reasons to generate compiler warnings.
The changes consist of:
- adding missing casts
- hiding or unhiding function declarations
- including headers where missing
- add __UNCONST when assigning a const char * to a char *
- adding missing return statements
- changing some types from unsigned to signed, as the code seems to want
signed ints
- converting old-style function definitions to current style (i.e.,
void func(param1, param2) short param1, param2; {...} to
void func (short param1, short param2) {...})
- making the compiler silent about signed vs unsigned comparisons. We
have too many of those in the new libc to fix.
A number of bugs in the test set were fixed. These bugs were never
triggered with our old libc. Consequently, these tests are now forced to
link with the new libc or they will generate errors (in particular tests 43
and 55).
Most changes in NetBSD libc are limited to moving aroudn "#ifndef __minix"
or stuff related to Minix-specific things (code in sys-minix or gen/minix).
. it's a good extra interface to have but doesn't
meet standardised functionality
. applications (in pkgsrc) find it and expect
full functionality the minix mmap doesn't offter
. on the whole probably better to hide these functions
(mmap and friends) until they are grown up; the base system
can use the new minix_* names
Now users can choose between libsys, libsys + libminc and
libsys + libc. E.g. PUFFS/FUSE servers need libsys + libc while
old servers can use libsys + libminc.
3 sets of libraries are built now:
. ack: all libraries that ack can compile (/usr/lib/i386/)
. clang+elf: all libraries with minix headers (/usr/lib/)
. clang+elf: all libraries with netbsd headers (/usr/netbsd/)
Once everything can be compiled with netbsd libraries and headers, the
/usr/netbsd hierarchy will be obsolete and its libraries compiled with
netbsd headers will be installed in /usr/lib, and its headers
in /usr/include. (i.e. minix libc and current minix headers set
will be gone.)
To use the NetBSD libc system (libraries + headers) before
it is the default libc, see:
http://wiki.minix3.org/en/DevelopersGuide/UsingNetBSDCode
This wiki page also documents the maintenance of the patch
files of minix-specific changes to imported NetBSD code.
Changes in this commit:
. libsys: Add NBSD compilation and create a safe NBSD-based libc.
. Port rest of libraries (except libddekit) to new header system.
. Enable compilation of libddekit with new headers.
. Enable kernel compilation with new headers.
. Enable drivers compilation with new headers.
. Port legacy commands to new headers and libc.
. Port servers to new headers.
. Add <sys/sigcontext.h> in compat library.
. Remove dependency file in tree.
. Enable compilation of common/lib/libc/atomic in libsys
. Do not generate RCSID strings in libc.
. Temporarily disable zoneinfo as they are incompatible with NetBSD format
. obj-nbsd for .gitignore
. Procfs: use only integer arithmetic. (Antoine Leca)
. Increase ramdisk size to create NBSD-based images.
. Remove INCSYMLINKS handling hack.
. Add nbsd_include/sys/exec_elf.h
. Enable ELF compilation with NBSD libc.
. Add 'make nbsdsrc' in tools to download reference NetBSD sources.
. Automate minix-port.patch creation.
. Avoid using fstavfs() as it is *extremely* slow and unneeded.
. Set err() as PRIVATE to avoid name clash with libc.
. [NBSD] servers/vm: remove compilation warnings.
. u32 is not a long in NBSD headers.
. UPDATING info on netbsd hierarchy
. commands fixes for netbsd libc
sys_umap now supports only:
- looking up the physical address of a virtual address in the address space
of the caller;
- looking up the physical address of a grant for which the caller is the
grantee.
This is enough for nearly all umap users. The new sys_umap_remote supports
lookups in arbitrary address spaces and grants for arbitrary grantees.
Before safecopies, the IO_ENDPT and DL_ENDPT message fields were needed
to know which actual process to copy data from/to, as that process may
not always be the caller. Now that we have full safecopy support, these
fields have become useless for that purpose: the owner of the grant is
*always* the caller. Allowing the caller to supply another endpoint is
in fact dangerous, because the callee may then end up using a grant
from a third party. One could call this a variant of the confused
deputy problem.
From now on, safecopy calls should always use the caller's endpoint as
grant owner. This fully obsoletes the DL_ENDPT field in the
inet/ethernet protocol. IO_ENDPT has other uses besides identifying the
grant owner though. This patch renames IO_ENDPT to USER_ENDPT, not only
because that is a more fitting name (it should never be used for I/O
after all), but also in order to intentionally break any old system
source code outside the base system. If this patch breaks your code,
fixing it is fairly simple:
- DL_ENDPT should be replaced with m_source;
- IO_ENDPT should be replaced with m_source when used for safecopies;
- IO_ENDPT should be replaced with USER_ENDPT for any other use, e.g.
when setting REP_ENDPT, matching requests in CANCEL calls, getting
DEV_SELECT flags, and retrieving of the real user process's endpoint
in DEV_OPEN.
The changes in this patch are binary backward compatible.
asynchronous message resulted in an error.
The model here is that:
- Iff a sender wishes to be notified, the sender MUST check for errors
BEFORE sending another asynchronous message.
The reason is that in order to remember the error code, we can't clean up
the message table and hence we risk running out of table space. This is
less of a problem when the sender enables notifications only for errors.
