Released Wednesday, February 16, 2011
Part A due on Friday, March 04, 2011, 9:00 PM
Part B due on Friday, March 25, 2011, 9:00 PM
Part C due on Friday, April 01, 2011, 9:00 PM
In this lab you will implement preemptive multitasking among multiple simultaneously active user-mode environments. In part A you will implement round-robin scheduling and the basic environment management system calls (calls that create and destroy environments, and allocate/map memory).
In part B, you will implement a Unix-like fork(),
which allows a user-mode environment to create copies of
itself.
Finally, in part C you will add support for inter-process communication (IPC), allowing different user-mode environments to communicate and synchronize with each other explicitly. You will also add support for hardware clock interrupts and preemption.
Use Git to commit your Lab 3 source, fetch the latest version of the course repository, and then create a local branch called lab4 based on our lab4 branch, origin/lab4:
tig% cd ~/CS372H/lab tig% git commit -am 'my solution to lab3' Created commit 734fab7: my solution to lab3 4 files changed, 42 insertions(+), 9 deletions(-) tig% git pull Already up-to-date. tig% git checkout -b lab4 origin/lab4 Branch lab4 set up to track remote branch refs/remotes/origin/lab4. Switched to a new branch "lab4" tig% git merge lab3 Merge made by recursive. kern/env.c | 42 +++++++++++++++++++ 1 files changed, 42 insertions(+), 0 deletions(-) tig%
This lab is divided into three parts, A, B, and C. You should make turnin-partA your lab before the Part A deadline, at which point your code must pass all of the part A tests. By the part B deadline, your code must pass all of the part B tests, etc.
As before, do all of the regular exercises described in the lab and at least one challenge problem. (Some challenge problems are more challenging than others, of course!) Additionally, write up brief answers to the questions posed in the lab and a short (e.g., one or two paragraph) description of what you did to solve your chosen challenge problem. If you implement more than one challenge problem, you only need to describe one of them in the write-up, though of course you are welcome to do more. Furthermore, challenge problems are due at the same time as the last part of a lab, so you do not need to have one completed and written up until you are completely done with the lab. Place the write-up in a file called answers.txt (plain text) in the top level of your lab directory before handing in your work.
Please include a header that contains your name, UTCS username, and lab number and make sure the file is named correctly. If you do not, your answer may not be graded. Please include in your answers.txt the write-up for all of the parts of the lab that you have completed so far (so the lab2b write-up should include the lab2a write-up), but do not include the write-up of previous labs. We may print these out and we would rather not waste paper.
When you are ready to hand in your lab code and write-up, create a file called slack.txt noting how many slack hours you have used for this and each of the previous assignments -- each lab part counts as a different assignment -- as well as how many you have used in total. (This is to help us agree on the number that you have used.) You must include this file even if you didn't use any slack hours. Then run make turnin-parti where i is the part you want to turnin in the lab directory. This will first do a make clean to clean out any .o files and executables, and then create a tar file called lab4i-handin.tar.gz with the entire contents of your lab directory and submit it via the CS turnin utility. If you submit multiple times, we will take the latest submission and count slack hours accordingly.
As before, we will be grading your solutions with a grading program. You can run make grade in the lab directory to test your kernel with the grading program. You may change any of the kernel source and header files you need to in order to complete the lab, but needless to say you must not change or otherwise subvert the grading code.
In the first part of this lab, you will implement some new JOS kernel system calls to allow user-level environments to create additional new environments. You will also implement cooperative round-robin scheduling, allowing the kernel to switch from one environment to another when the current environment voluntarily relinquishes the CPU (or exits). Later in part C you will implement preemptive scheduling, which allows the kernel to re-take control of the CPU from an environment after a certain time has passed even if the environment does not cooperate.
Your first task in this lab is to change the JOS kernel
so that it does not always just run the environment in envs[0],
but instead can alternate between multiple environments
in "round-robin" fashion.
Round-robin scheduling in JOS works as follows:
envs[0],
will from now on always be a special idle environment,
which always runs the program user/idle.c.
The purpose of this program is simply to "waste time"
whenever the processor has nothing better to do -
it just perpetually attempts to give up the CPU
to another environment.
Read the code and comments in user/idle.c
for other useful details.
We have modified kern/init.c for you
to create this special idle environment in envs[0]
before creating the first "real" environment in envs[1].sched_yield() in the new kern/sched.c
is responsible for selecting a new environment to run.
It searches sequentially through the envs[] array
in circular fashion,
starting just after the previously running environment
(or at the beginning of the array
if there was no previously running environment),
picks the first environment it finds
with a status of ENV_RUNNABLE
(see inc/env.h),
and calls env_run() to jump into that environment.
However, sched_yield() is aware
that envs[0] is the special idle environment,
and never picks it unless
there are no other runnable environments.sys_yield(),
which user environments can call
to invoke the kernel's sched_yield() function
and thereby voluntarily give up the CPU to a different environment.
As you can see in user/idle.c,
the idle environment does this routinely.Exercise 1.
Implement round-robin scheduling in sched_yield()
as described above. Don't forget to modify
syscall() to dispatch sys_yield().
Modify kern/init.c to create three (or more!) environments that all run the program user/yield.c. You should see the environments switch back and forth between each other five times before terminating, like this:
... Hello, I am environment 00001001. Hello, I am environment 00001002. Hello, I am environment 00001003. Back in environment 00001001, iteration 0. Back in environment 00001002, iteration 0. Back in environment 00001003, iteration 0. Back in environment 00001001, iteration 1. Back in environment 00001002, iteration 1. Back in environment 00001003, iteration 1. ...
After the yield programs exit, the idle environment should run and invoke the JOS kernel debugger. If any of this does not happen, then fix your code before proceeding.
Question
env_run() you should have
called lcr3(). Before and after the call to
lcr3(), your code makes references (at least it should)
to the variable e, the argument to env_run.
Upon loading the %cr3 register, the addressing context
used by the MMU is instantly changed. But a virtual
address (namely e) has meaning relative to a given
address context--the address context specifies the physical address to
which the virtual address maps. Why can the pointer e be
dereferenced both before and after the addressing switch?
Challenge! Add a less trivial scheduling policy to the kernel, such as a fixed-priority scheduler that allows each environment to be assigned a priority and ensures that higher-priority environments are always chosen in preference to lower-priority environments. If you're feeling really adventurous, try implementing a Unix-style adjustable-priority scheduler or even a lottery or stride scheduler. (Look up "lottery scheduling" and "stride scheduling" in Google.)
Write a test program or two
that verifies that your scheduling algorithm is working correctly
(i.e., the right environments get run in the right order).
It may be easier to write these test programs
once you have implemented fork() and IPC
in parts B and C of this lab.
Challenge!
The JOS kernel currently does not allow applications
to use the x86 processor's x87 floating-point unit (FPU),
MMX instructions, or Streaming SIMD Extensions (SSE).
Extend the Env structure
to provide a save area for the processor's floating point state,
and extend the context switching code
to save and restore this state properly
when switching from one environment to another.
The FXSAVE and FXRSTOR instructions may be useful,
but note that these are not in the old i386 user's manual
because they were introduced in more recent processors.
Write a user-level test program
that does something cool with floating-point.
Although your kernel is now capable of running and switching between multiple user-level environments, it is still limited to running environments that the kernel initially set up. You will now implement the necessary JOS system calls to allow user environments to create and start other new user environments.
Unix provides the fork() system call
as its process creation primitive.
Unix fork() copies
the entire address space of calling process (the parent)
to create a new process (the child).
The only differences between the two observable from user space
are their process IDs and parent process IDs
(as returned by getpid and getppid).
In the parent,
fork() returns the child's process ID,
while in the child, fork() returns 0.
By default, each process gets its own private address space, and
neither process's modifications to memory are visible to the other.
You will provide a different, more primitive
set of JOS system calls
for creating new user-mode environments.
With these system calls you will be able to implement
a Unix-like fork() entirely in user space,
in addition to other styles of environment creation.
The new system calls you will write for JOS are as follows:
sys_exofork:sys_exofork call.
In the parent, sys_exofork
will return the envid_t of the newly created
environment
(or a negative error code if the environment allocation failed).
In the child, however, it will return 0.
(Since the child starts out marked as not runnable,
sys_exofork will not actually return in the child
until the parent has explicitly allowed this
by marking the child runnable using....)sys_env_set_status:ENV_RUNNABLE or ENV_NOT_RUNNABLE.
This system call is typically used
to mark a new environment ready to run,
once its address space and register state
has been fully initialized.sys_page_alloc:sys_page_map:sys_page_unmap:
For all of the system calls above that accept environment IDs,
the JOS kernel supports the convention
that a value of 0 means "the current environment."
This convention is implemented by envid2env()
in kern/env.c.
We have provided a very primitive implementation
of a Unix-like fork()
in the test program user/dumbfork.c.
This test program uses the above system calls
to create and run a child environment
with a copy of its own address space.
The two environments
then switch back and forth using sys_yield
as in the previous exercise.
The parent exits after 10 iterations,
whereas the child exits after 20.
Exercise 2.
Implement the system calls described above
in kern/syscall.c.
You will need to use various functions
in kern/pmap.c and kern/env.c,
particularly envid2env().
For now, whenever you call envid2env(),
pass 1 in the checkperm parameter.
Be sure you check for any invalid system call arguments,
returning -E_INVAL in that case.
Test your JOS kernel with user/dumbfork
and make sure it works before proceeding.
Challenge!
Add the additional system calls necessary
to read all of the vital state of an existing environment
as well as set it up.
Then implement a user mode program that forks off a child environment,
runs it for a while (e.g., a few iterations of sys_yield()),
then takes a complete snapshot or checkpoint
of the child environment,
runs the child for a while longer,
and finally restores the child environment to the state it was in
at the checkpoint
and continues it from there.
Thus, you are effectively "replaying"
the execution of the child environment from an intermediate state.
Make the child environment perform some interaction with the user
using sys_cgetc() or readline()
so that the user can view and mutate its internal state,
and verify that with your checkpoint/restart
you can give the child environment a case of selective amnesia,
making it "forget" everything that happened beyond a certain point.
This completes Part A of the lab; check it using make grade and hand it in using make turnin-partA as usual.
As mentioned earlier,
Unix provides the fork() system call
as its primary process creation primitive.
The fork() system call
copies the address space of the calling process (the parent)
to create a new process (the child).
dumbfork() copies all data from the
parent's pages into new pages allocated for the child.
The copying of the parent's address space into the child is
the most expensive part of the fork() operation.
However, a call to fork()
is frequently followed almost immediately
by a call to exec() in the child process,
which replaces the child's memory with a new program.
This is what the the shell typically does, for example.
In this case,
the time spent copying the parent's address space is largely wasted,
because the child process will use
very little of its memory before calling exec().
For this reason,
later versions of Unix took advantage
of virtual memory hardware
to allow the parent and child to share
the memory mapped into their respective address spaces
until one of the processes actually modifies it.
This technique is known as copy-on-write.
To do this,
on fork() the kernel would
copy the address space mappings
from the parent to the child
instead of the contents of the mapped pages,
and at the same time mark the now-shared pages read-only.
When one of the two processes tries to write to one of these shared pages,
the process takes a page fault.
At this point, the Unix kernel realizes that the page
was really a "virtual" or "copy-on-write" copy,
and so it makes a new, private, writable copy of the page for the
faulting process.
In this way, the contents of individual pages aren't actually copied
until they are actually written to.
This optimization makes a fork() followed by
an exec() in the child much cheaper:
the child will probably only need to copy one page
(the current page of its stack)
before it calls exec().
In the next piece of this lab, you will implement a "proper"
Unix-like fork() with copy-on-write,
as a user space library routine.
Implementing fork() and copy-on-write support in user space
has the benefit that the kernel remains much simpler
and thus more likely to be correct.
It also lets individual user-mode programs
define their own semantics for fork().
A program that wants a slightly different implementation
(for example, the expensive always-copy version like dumbfork(),
or one in which the parent and child actually share memory afterward)
can easily provide its own.
A user-level copy-on-write fork() needs to know about
page faults on write-protected pages, so that's what you'll
implement first.
Copy-on-write is only one of many possible uses
for user-level page fault handling.
It's common to set up an address space so that page faults indicate when some action needs to take place. For example, most Unix kernels initially map only a single page in a new process's stack region, and allocate and map additional stack pages later "on demand" as the process's stack consumption increases and causes page faults on stack addresses that are not yet mapped. A typical Unix kernel must keep track of what action to take when a page fault occurs in each region of a process's space. For example, a fault in the stack region will typically allocate and map new page of physical memory. A fault in the program's BSS region will typically allocate a new page, fill it with zeroes, and map it. In systems with demand-paged executables, a fault in the text region will read the corresponding page of the binary off of disk and then map it.
This is a lot of information for the kernel to keep track of. Instead of taking the traditional Unix approach, you will decide what to do about each page fault in user space, where bugs are less damaging. This design has the added benefit of allowing programs great flexibility in defining their memory regions; you'll use user-level page fault handling later for mapping and accessing files on a disk-based file system.
In order to handle its own page faults,
a user environment will need to register
a page fault handler entrypoint with the JOS kernel.
The user environment registers its page fault entrypoint
via the new sys_env_set_pgfault_upcall system call.
We have added a new member to the Env structure,
env_pgfault_upcall,
to record this information.
Exercise 3.
Implement the sys_env_set_pgfault_upcall system call.
Be sure to enable permission checking
when looking up the environment ID of the target environment,
since this is a "dangerous" system call.
During normal execution,
a user environment in JOS
will run on the normal user stack:
its ESP register starts out pointing at USTACKTOP,
and the stack data it pushes resides on the page
between USTACKTOP-PGSIZE and USTACKTOP-1 inclusive.
When a page fault occurs in user mode,
however,
the kernel will restart the user environment
running a designated user-level page fault handler
on a different stack,
namely the user exception stack.
In essence, we will make the JOS kernel
implement automatic "stack switching"
on behalf of the user environment,
in much the same way that the x86 processor
already implements stack switching on behalf of JOS
when transferring from user mode to kernel mode!
The JOS user exception stack is also one page in size,
and its top is defined to be at virtual address UXSTACKTOP,
so the valid bytes of the user exception stack
are from UXSTACKTOP-PGSIZE through UXSTACKTOP-1 inclusive.
While running on this exception stack,
the user-level page fault handler
can use JOS's regular system calls to map new pages or adjust mappings
so as to fix whatever problem originally caused the page fault.
Then the user-level page fault handler returns,
via an assembly language stub,
to the faulting code on the original stack.
Each user environment that wants to support user-level page fault handling
will need to allocate memory for its own exception stack,
using the sys_page_alloc() system call introduced in part A.
You will now need to change the page fault handling code in kern/trap.c to handle page faults from user mode as follows. We will call the state of the user environment at the time of the fault the trap-time state.
If there is no page fault handler registered,
the JOS kernel destroys the user environment with a message as before.
Otherwise,
the kernel sets up a trap frame on the exception stack that looks like
a struct UTrapframe from inc/trap.h:
<-- UXSTACKTOP
trap-time esp
trap-time eflags
trap-time eip
trap-time eax start of struct PushRegs
trap-time ecx
trap-time edx
trap-time ebx
trap-time esp
trap-time ebp
trap-time esi
trap-time edi end of struct PushRegs
tf_err (error code)
fault_va <-- %esp when handler is run
The kernel then arranges for the user environment to resume execution with the page fault handler running on the exception stack with this stack frame; you must figure out how to make this happen. The fault_va is the virtual address that caused the page fault.
If the user environment is already running on the user exception stack
when an exception occurs,
then the page fault handler itself has faulted.
In this case,
you should start the new stack frame just under the current
tf->tf_esp rather than at UXSTACKTOP.
You should first push an empty 32-bit word, then a struct UTrapframe.
To test whether tf->tf_esp is already on the user
exception stack, check whether it is in the range
between UXSTACKTOP-PGSIZE and UXSTACKTOP-1, inclusive.
Exercise 4.
Implement the code in page_fault_handler in
kern/trap.c
required to dispatch page faults to the user-mode handler.
Be sure to take appropriate precautions
when writing into the exception stack.
(What happens if the user environment runs out of space
on the exception stack?)
Next, you need to implement the assembly routine that will
take care of calling the C page fault handler and resume
execution at the original faulting instruction.
This assembly routine is the handler that will be registered
with the kernel using sys_env_set_pgfault_upcall().
Exercise 5.
Implement the _pgfault_upcall routine
in lib/pfentry.S.
The interesting part is returning to the original point in
the user code that caused the page fault.
You'll return directly there, without going back through
the kernel.
The hard part is simultaneously switching stacks and
re-loading the EIP.
Finally, you need to implement the C user library side of the user-level page fault handling mechanism.
Exercise 6.
Finish set_pgfault_handler()
in lib/pgfault.c.
Run user/faultread. You should see:
[00000000] new env 00001000 [00000000] new env 00001001 [00001001] user fault va 00000000 ip 0080003a TRAP frame ... [00001001] free env 00001001
Run user/faultdie. You should see:
[00000000] new env 00001000 [00000000] new env 00001001 i faulted at va deadbeef, err 6 [00001001] exiting gracefully [00001001] free env 00001001
Run user/faultalloc. You should see:
[00000000] new env 00001000 [00000000] new env 00001001 fault deadbeef this string was faulted in at deadbeef fault cafebffe fault cafec000 this string was faulted in at cafebffe [00001001] exiting gracefully [00001001] free env 00001001
If you see only the first "this string" line, it means you are not handling recursive page faults properly.
Run user/faultallocbad. You should see:
[00000000] new env 00001000 [00000000] new env 00001001 [00001001] user_mem_check assertion failure for va deadbeef [00001001] free env 00001001
Make sure you understand why user/faultalloc and user/faultallocbad behave differently.
Challenge! Extend your kernel so that not only page faults, but all types of processor exceptions that code running in user space can generate, can be redirected to a user-mode exception handler. Write user-mode test programs to test user-mode handling of various exceptions such as divide-by-zero, general protection fault, and illegal opcode.
You now have the kernel facilities
to implement copy-on-write fork()
entirely in user space.
We have provided a skeleton for your fork()
in lib/fork.c.
Like dumbfork(),
fork() creates a new environment,
then scans through the parent environment's entire address space
and sets up corresponding page mappings in the child.
The key difference is that,
while dumbfork() copied pages,
fork() will initially only copy page mappings.
fork() will
copy each page only when one of the environments tries to write it.
The basic control flow for fork() is as follows:
pgfault()
as the C-level page fault handler,
using the set_pgfault_handler() function
you implemented above.sys_exofork() to create
a child environment.duppage, which should
map the page copy-on-write into the address
space of the child and then remap the page copy-on-write
in its own address space. duppage sets both PTEs so that the
page is not writeable, and to contain PTE_COW in the
"avail" field to distinguish copy-on-write pages from genuine
read-only pages.
The exception stack is not remapped this way, however. Instead the parent needs to allocate a fresh page in the child for the exception stack. Since the page fault handler will be doing the actual copying and the page fault handler runs on the exception stack, the exception stack cannot be made copy-on-write: who would copy it?
fork() also needs to handle pages that are
present, but not writable or copy-on-write.
Each time one of the environments writes a copy-on-write page that it hasn't yet written, it will take a page fault. Here's the control flow for the user page fault handler:
_pgfault_upcall,
which calls fork()'s pgfault() handler.pgfault() checks that the fault is a write
(check for FEC_WR in the error code) and that the
PTE for the page is marked PTE_COW.
If not, panic.pgfault() allocates a new page mapped
at a temporary location and copies
the contents of the faulting page contents into it.
Then the fault handler maps the new page at the
appropriate address with read/write permissions,
in place of the old read-only mapping.Exercise 7.
Implement fork, duppage and
pgfault in lib/fork.c.
Test your code with the forktree program. It should produce the following messages, with interspersed 'new env', 'free env', and 'exiting gracefully' messages. The messages may not appear in this order, and the environment IDs may be different.
1001: I am '' 1802: I am '0' 2801: I am '00' 3802: I am '000' 2003: I am '1' 5001: I am '11' 4802: I am '10' 6801: I am '100' 5803: I am '110' 3004: I am '01' 8001: I am '011' 7803: I am '010' 4005: I am '001' 6006: I am '111' 7007: I am '101'
Challenge!
Implement a shared-memory fork()
called sfork(). This version should have the parent
and child share all their memory pages
(so writes in one environment appear in the other)
except for pages in the stack area,
which should be treated in the usual copy-on-write manner.
Modify user/forktree.c
to use sfork() instead of regular fork().
Also, once you have finished implementing IPC in part C,
use your sfork() to run user/pingpongs.
You will have to find a new way to provide the functionality
of the global env pointer.
Challenge!
Your implementation of fork
makes a huge number of system calls. On the x86, switching into
the kernel using interrupts has non-trivial cost. Augment the
system call interface
so that it is possible to send a batch of system calls at once.
Then change fork to use this interface.
How much faster is your new fork?
You can answer this (roughly) by using analytical
arguments to estimate how much of an improvement batching
system calls will make to the performance of your
fork: How expensive is an int 0x30
instruction? How many times do you execute int 0x30
in your fork? Is accessing the TSS stack
switch also expensive? And so on...
Alternatively, you can boot your kernel on real hardware
and really benchmark your code. See the RDTSC
(read time-stamp counter) instruction, defined in the IA32
manual, which counts the number of clock cycles that have
elapsed since the last processor reset. QEMU doesn't emulate
this instruction faithfully (it can either count the number of
virtual instructions executed or use the host TSC, neither of
which reflects the number of cycles a real CPU would
require).
This ends part B. As usual, you can grade your submission with make grade and hand it in with make turnin-partB.
In the final part of lab 4 you will modify the kernel to preempt uncooperative environments and to allow environments to pass messages to each other explicitly.
Run the user/spin test program. This test program forks off a child environment, which simply spins forever in a tight loop once it receives control of the CPU. Neither the parent environment nor the kernel ever regains the CPU. This is obviously not an ideal situation in terms of protecting the system from bugs or malicious code in user-mode environments, because any user-mode environment can bring the whole system to a halt simply by getting into an infinite loop and never giving back the CPU. In order to allow the kernel to preempt a running environment, forcefully retaking control of the CPU from it, we must extend the JOS kernel to support external hardware interrupts from the clock hardware.
External interrupts (i.e., device interrupts) are referred to as IRQs.
There are 16 possible IRQs, numbered 0 through 15.
The mapping from IRQ number to IDT entry is not fixed.
Pic_init in picirq.c maps IRQs 0-15
to IDT entries IRQ_OFFSET through IRQ_OFFSET+15.
In kern/picirq.h,
IRQ_OFFSET is defined to be decimal 32.
Thus the IDT entries 32-47 correspond to the IRQs 0-15.
For example, the clock interrupt is IRQ 0.
Thus, IDT[IRQ_OFFSET+0] (i.e., IDT[32]) contains the address of
the clock's interrupt handler routine in the kernel.
This IRQ_OFFSET is chosen so that the device interrupts
do not overlap with the processor exceptions,
which could obviously cause confusion.
(In fact, in the early days of PCs running MS-DOS,
the IRQ_OFFSET effectively was zero,
which indeed caused massive confusion between handling hardware interrupts
and handling processor exceptions!)
In JOS,
external device interrupts are always disabled
when in the kernel (and enabled when in user space).
External interrupts are controlled by the FL_IF flag bit
of the %eflags register
(see inc/mmu.h).
When this bit is set, external interrupts are enabled.
While the bit can be modified in several ways,
because of our simplification, we will handle it solely
through the process of saving and restoring the %eflags register
as we enter and leave user mode.
You will have to ensure that the FL_IF flag is set in
user environments when they run so that when an interrupt arrives, it
gets passed through to the processor and handled by your interrupt code.
Otherwise, interrupts are masked,
or ignored until interrupts are re-enabled.
We masked interrupts with the very first instruction of the bootloader,
and so far we have never gotten around to re-enabling them.
Exercise 8.
Modify kern/trapentry.S and kern/trap.c to
initialize the appropriate entries in the IDT and provide
handlers for IRQs 0 through 15. Then modify the code
in env_alloc() in kern/env.c to ensure
that user environments are always run with interrupts enabled.
The processor never pushes an error code or checks the Descriptor Privilege Level (DPL) of the IDT entry when invoking a hardware interrupt handler. You might want to re-read section 9.2 of the 80386 Reference Manual, or section 5.8 of the IA-32 Intel Architecture Software Developer's Manual, Volume 3, at this time.
After doing this exercise, if you run your kernel with any test program that runs for a non-trivial length of time (e.g., spin), you should see the kernel print trap frames for hardware interrupts. While interrupts are now enabled in the processor, JOS isn't yet handling them, so you should see it misattribute each interrupt to the currently running user environment and destroy it. Eventually it should run out of environments to destroy and drop into the monitor. If you see an infinite cascade of hardware interrupts, you're probably running with interrupts enabled in the kernel and you should fix this before proceeding.
In the user/spin program, after the child environment was first run, it just spun in a loop, and the kernel never got control back. We need to program the hardware to generate clock interrupts periodically, which will force control back to the kernel where we can switch control to a different user environment.
The calls to pic_init and kclock_init
(from i386_init in init.c),
which we have written for you,
set up the clock and the interrupt controller to generate interrupts.
You now need to write the code to handle these interrupts.
Exercise 9.
Modify the kernel's trap_dispatch() function
so that it calls sched_yield()
to find and run a different environment
whenever a clock interrupt takes place.
You should now be able to get the user/spin test to work:
the parent environment should fork off the child,
sys_yield() to it a couple times
but in each case regain control of the CPU after one time slice,
and finally kill the child environment and terminate gracefully.
This is a great time to do some regression testing. Make sure that you haven't broken any earlier part of that lab that used to work (e.g. forktree) by enabling interrupts. Run make grade to see for sure. You should now get a total score of 60/70 points on this lab.
(Technically in JOS this is "inter-environment communication" or "IEC", but everyone else calls it IPC, so we'll use the standard term.)
We've been focusing on the isolation aspects of the operating system, the ways it provides the illusion that each program has a machine all to itself. Another important service of an operating system is to allow programs to communicate with each other when they want to. It can be quite powerful to let programs interact with other programs. The Unix pipe model is the canonical example.
There are many models for interprocess communication. Even today there are still debates about which models are best. We won't get into that debate. Instead, we'll implement a simple IPC mechanism and then try it out.
You will implement a few additional JOS kernel system calls
that collectively provide a simple interprocess communication mechanism.
You will implement two
system calls, sys_ipc_recv and
sys_ipc_try_send.
Then you will implement two library wrappers
ipc_recv and ipc_send.
The "messages" that user environments can send to each other using JOS's IPC mechanism consist of two components: a single 32-bit value, and optionally a single page mapping. Allowing environments to pass page mappings in messages provides an efficient way to transfer more data than will fit into a single 32-bit integer, and also allows environments to set up shared memory arrangements easily.
To receive a message, an environment calls
sys_ipc_recv.
This system call de-schedules the current
environment and does not run it again until a message has
been received.
When an environment is waiting to receive a message,
any other environment can send it a message -
not just a particular environment,
and not just environments that have a parent/child arrangement
with the receiving environment.
In other words, the permission checking that you implemented in Part A
will not apply to IPC,
because the IPC system calls are carefully designed so as to be "safe":
an environment cannot cause another environment to malfunction
simply by sending it messages
(unless the target environment is also buggy).
To try to send a value, an environment calls
sys_ipc_try_send with both the receiver's
environment id and the value to be sent. If the named
environment is actually receiving (it has called
sys_ipc_recv and not gotten a value yet),
then the send delivers the message and returns 0. Otherwise
the send returns -E_IPC_NOT_RECV to indicate
that the target environment is not currently expecting
to receive a value. (Other errors are possible too; see the comments in
the source files.)
A library function ipc_recv in user space will take care
of calling sys_ipc_recv and then looking up
the information about the received values in the current
environment's struct Env.
Similarly, a library function ipc_send will
take care of repeatedly calling sys_ipc_try_send
until the send succeeds.
When an environment calls sys_ipc_recv
with a dstva parameter in the user address space,
the environment is stating that it is willing to receive a page mapping.
If the sender sends a page,
then that page should be mapped at dstva
in the receiver's address space.
If the receiver already had a page mapped at dstva,
then that previous page is unmapped.
When an environment calls sys_ipc_try_send
with a srcva parameter in the user address space,
it means the sender wants to send the page
currently mapped at srcva to the receiver,
with permissions perm.
After a successful IPC,
the sender keeps its original mapping
for the page at srcva in its address space,
but the receiver also obtains a mapping for this same physical page
at the dstva originally specified by the receiver,
in the receiver's address space.
As a result this page becomes shared between the sender and receiver.
If either the sender or the receiver does not indicate
that a page should be transferred,
then no page is transferred.
After any IPC
the kernel sets the new field env_ipc_perm
in the receiver's Env structure
to the permissions of the page received,
or zero if no page was received.
Exercise 10.
Implement sys_ipc_recv and
sys_ipc_try_send in kern/syscall.c.
Read the comments on both before implementing them, since they
have to work together.
When you call envid2env in these routines, you should
set the checkperm flag to 0,
meaning that any environment is allowed to send
IPC messages to any other environment,
and the kernel does no special permission checking
other than verifying that the target envid is valid.
Then implement
the ipc_recv and ipc_send functions
in lib/ipc.c.
Use the user/pingpong and user/primes functions to test your IPC mechanism. You might find it interesting to read user/primes.c to see all the forking and IPC going on behind the scenes.
Challenge!
The ipc_send function is not very fair.
Run three copies of user/fairness and you will
see this problem. The first two copies are both trying to send to
the third copy, but only one of them will ever succeed.
Make the IPC fair, so that each copy has approximately
equal chance of succeeding.
Challenge!
Why does ipc_send
have to loop? Change the system call interface so it
doesn't have to. Make sure you can handle multiple
environments trying to send to one environment at the
same time.
Challenge! The prime sieve is only one neat use of message passing between a large number of concurrent programs. Read C. A. R. Hoare, ``Communicating Sequential Processes,'' Communications of the ACM 21(8) (August 1978), 666-667, and implement the matrix multiplication example.
Challenge! Probably the most impressive example of the power of message passing is Doug McIlroy's power series calculator, described in M. Douglas McIlroy, ``Squinting at Power Series,'' Software--Practice and Experience, 20(7) (July 1990), 661-683. Implement his power series calculator and compute the power series for sin(x+x^3).
Challenge! Make JOS's IPC mechanism more efficient by applying some of the techniques from Liedtke's paper, "Improving IPC by Kernel Design", or any other tricks you may think of. Feel free to modify the kernel's system call API for this purpose, as long as your code is backwards compatible with what our grading scripts expect.
This ends part C. As usual, you can grade your submission with make grade and hand it in with make turnin-partC. If you are trying to figure out why a particular test case is failing, run sh grade-lab4.sh -v, which will show you the output of the kernel builds and QEMU runs for each test, until a test fails. When a test fails, the script will stop, and then you can inspect jos.out to see what the kernel actually printed.
Last updated: Tue Apr 05 19:50:09 -0500 2011 [validate xhtml]