I picked it up again. The scheduler is forked from ULE, but the
algorithm to detect an interactive process is almost completely
different with ULE, it comes from Linux paper "Understanding the
Linux 2.6.8.1 CPU Scheduler", although I still use same word
"score" as a priority boost in ULE scheduler.
Briefly, the scheduler has following characteristic:
1. Timesharing process's nice value is seriously respected,
timeslice and interaction detecting algorithm are based
on nice value.
2. per-cpu scheduling queue and load balancing.
3. O(1) scheduling.
4. Some cpu affinity code in wakeup path.
5. Support POSIX SCHED_FIFO and SCHED_RR.
Unlike scheduler 4BSD and ULE which using fuzzy RQ_PPQ, the scheduler
uses 256 priority queues. Unlike ULE which using pull and push, the
scheduelr uses pull method, the main reason is to let relative idle
cpu do the work, but current the whole scheduler is protected by the
big sched_lock, so the benefit is not visible, it really can be worse
than nothing because all other cpu are locked out when we are doing
balancing work, which the 4BSD scheduelr does not have this problem.
The scheduler does not support hyperthreading very well, in fact,
the scheduler does not make the difference between physical CPU and
logical CPU, this should be improved in feature. The scheduler has
priority inversion problem on MP machine, it is not good for
realtime scheduling, it can cause realtime process starving.
As a result, it seems the MySQL super-smack runs better on my
Pentium-D machine when using libthr, despite on UP or SMP kernel.
the arm to compile without all the extras that don't appear, at least
not in the flavors of ARM I deal with. This helps us save about 100k.
If I've botched the available devices on a platform, please let me
know and I'll correct ASAP.
that it just warns the user with a printf when it misaligns a piece
of memory that was requested through a busdma tag.
Some drivers (such as mpt, and probably others) were asking for alignments
that could not be satisfied, but as far as driver operation was concerned,
that did not matter. In the theory that other drivers will fall into
this same category, we agreed that panicing or making the allocation
fail will cause more hardship than is necessary. The printf should
be sufficient motivation to get the driver glitch fixed.
Add a quick hack to ensure that bus_dmamem_alloc properly aligns
small allocations with large alignment requirements.
Add a panic to detect cases where we've still failed to properly align.
conformance with the mbuf and uio load routines. ENOMEM can only happen
with BUS_DMA_NOWAIT is passed in, thus the deferals are disabled. I don't
like doing this, but fixing this fixes assumptions in other important drivers,
which is a net benefit for now.
o Properly use rman(9) to manage resources. This eliminates the
need to puc-specific hacks to rman. It also allows devinfo(8)
to be used to find out the specific assignment of resources to
serial/parallel ports.
o Compress the PCI device "database" by optimizing for the common
case and to use a procedural interface to handle the exceptions.
The procedural interface also generalizes the need to setup the
hardware (program chipsets, program clock frequencies).
o Eliminate the need for PUC_FASTINTR. Serdev devices are fast by
default and non-serdev devices are handled by the bus.
o Use the serdev I/F to collect interrupt status and to handle
interrupts across ports in priority order.
o Sync the PCI device configuration to include devices found in
NetBSD and not yet merged to FreeBSD.
o Add support for Quatech 2, 4 and 8 port UARTs.
o Add support for a couple dozen Timedia serial cards as found
in Linux.
entry (PTE) have the same meaning. The exception to this rule is the
eighth bit (0x080). It is the PS bit in a PDE and the PAT bit in a
PTE. This change avoids the possibility that pmap_enter() confuses a
PAT bit with a PS bit, avoiding a panic().
Eliminate a diagnostic printf() from the i386 pmap_enter() that serves
no current purpose, i.e., I've seen no bug reports in the last two
years that are helped by this printf().
Reviewed by: jhb
caches are dangerous" to "a shared L1 data cache is dangerous". This
is a compromise between paranoia and performance: Unlike the L1 cache,
nobody has publicly demonstrated a cryptographic side channel which
exploits the L2 cache -- this is harder due to the larger size, lower
bandwidth, and greater associativity -- and prohibiting shared L2
caches turns Intel Core Duo processors into Intel Core Solo processors.
As before, the 'machdep.hyperthreading_allowed' sysctl will allow even
the L1 data cache to be shared.
Discussed with: jhb, scottl
Security: See FreeBSD-SA-05:09.htt for background material.
via the debug.minidump sysctl and tunable.
Traditional dumps store all physical memory. This was once a good thing
when machines had a maximum of 64M of ram and 1GB of kvm. These days,
machines often have many gigabytes of ram and a smaller amount of kvm.
libkvm+kgdb don't have a way to access physical ram that is not mapped
into kvm at the time of the crash dump, so the extra ram being dumped
is mostly wasted.
Minidumps invert the process. Instead of dumping physical memory in
in order to guarantee that all of kvm's backing is dumped, minidumps
instead dump only memory that is actively mapped into kvm.
amd64 has a direct map region that things like UMA use. Obviously we
cannot dump all of the direct map region because that is effectively
an old style all-physical-memory dump. Instead, introduce a bitmap
and two helper routines (dump_add_page(pa) and dump_drop_page(pa)) that
allow certain critical direct map pages to be included in the dump.
uma_machdep.c's allocator is the intended consumer.
Dumps are a custom format. At the very beginning of the file is a header,
then a copy of the message buffer, then the bitmap of pages present in
the dump, then the final level of the kvm page table trees (2MB mappings
are expanded into a 4K page mappings), then the sparse physical pages
according to the bitmap. libkvm can now conveniently access the kvm
page table entries.
Booting my test 8GB machine, forcing it into ddb and forcing a dump
leads to a 48MB minidump. While this is a best case, I expect minidumps
to be in the 100MB-500MB range. Obviously, never larger than physical
memory of course.
minidumps are on by default. It would want be necessary to turn them off
if it was necessary to debug corrupt kernel page table management as that
would mess up minidumps as well.
Both minidumps and regular dumps are supported on the same machine.
to reduce the pv_entry_count counter. This was found by Tor Egge. In the
same email, Tor also pointed out the pv_stats problem in the previous
commit, but I'd forgotten about it until I went looking for this email
about this allocation problem.
create managed mappings within the clean submap. To prevent regressions,
add assertions blocking the creation of managed mappings within the clean
submap.
Reviewed by: tegge
so that we only have to do an ioapic_write() instead of an ioapic_read()
followed by an ioapic_write() every time we mask and unmask level triggered
interrupts. This cuts the execution time for these operations roughly in
half.
Profiled by: Paolo Pisati <p.pisati@oltrelinux.com>
MFC after: 1 week
PCB in which the context of stopped CPUs is stored. To access this
PCB from KDB, we introduce a new define, called KDB_STOPPEDPCB. The
definition, when present, lives in <machine/kdb.h> and abstracts
where MD code saves the context. Define KDB_STOPPEDPCB on i386,
amd64, alpha and sparc64 in accordance to previous code.
with large mmap files mapped into many processes, this saves hundreds of
megabytes of ram.
pv entries were individually allocated and had two tailq entries and two
pointers (or addresses). Each pv entry was linked to a vm_page_t and
a process's address space (pmap). It had the virtual address and a
pointer to the pmap.
This change replaces the individual allocation with a per-process
allocation system. A page ("pv chunk") is allocated and this provides
168 pv entries for that process. We can now eliminate one of the 16 byte
tailq entries because we can simply iterate through the pv chunks to find
all the pv entries for a process. We can eliminate one of the 8 byte
pointers because the location of the pv entry implies the containing
pv chunk, which has the pointer. After overheads from the pv chunk
bitmap and tailq linkage, this works out that each pv entry has an
effective size of 24.38 bytes.
Future work still required, and other problems:
* when running low on pv entries or system ram, we may need to defrag
the chunk pages and free any spares. The stats (vm.pmap.*) show that
this doesn't seem to be that much of a problem, but it can be done if
needed.
* running low on pv entries is now a much bigger problem. The old
get_pv_entry() routine just needed to reclaim one other pv entry.
Now, since they are per-process, we can only use pv entries that are
assigned to our current process, or by stealing an entire page worth
from another process. Under normal circumstances, the pmap_collect()
code should be able to dislodge some pv entries from the current
process. But if needed, it can still reclaim entire pv chunk pages
from other processes.
* This should port to i386 really easily, except there it would reduce
pv entries from 24 bytes to about 12 bytes.
(I have integrated Alan's recent changes.)
a pv entry if the number of entries is below the high water mark for pv
entries.
Use pmap_try_insert_pv_entry() in pmap_copy() instead of
pmap_insert_entry(). This avoids possible recursion on a pmap lock in
get_pv_entry().
Eliminate the explicit low-memory checks in pmap_copy(). The check that
the number of pv entries was below the high water mark was largely
ineffective because it was located in the outer loop rather than the
inner loop where pv entries were allocated. Instead of checking, we
attempt the allocation and handle the failure.
Reviewed by: tegge
Reported by: kris
MFC after: 5 days
