The ability to port drivers from Linux is very important, because there are very many drivers. The ability to port them with minimal changes is important because they are said to have a high rate of change.
Device drivers, in most cases, have the same security properties as the kernel - they must be trusted by everyone, because they direct hardware that on most architectures has write access to all of physical memory. In KeyKOS and early versions of EROS, device drivers executed as part of the kernel.
However, there are several advantages to having drivers execute in user processes:
There are also some disadvantages:
The decision in CapROS is to run virtually all drivers in user processes.
Drivers will of course be granted whatever privileges are needed to do their job (this is architecture-dependent). The DevicePrivs key allows drivers to field interrupts and (together with a Range key) access memory containing device registers. The DevicePrivs key will also make it possible for a driver to obtain contiguous physical memory for use with DMA.
The Page Fault Handler (PFH) is defined as all the user objects which could be needed to service requests (from the kernel) to fetch and write pages and nodes from or to backing store. This includes the process that fields those requests, any other processes that service those requests, the process(es) that service backing store device interrupts, all the code and data used by those processes, and any other pages and nodes used by those processes.
Backing store is usually disk, but the PFH is not exactly the same as the disk driver. The PFH includes any objects used to handle the buses by which the backing store is accessed, such as PCI, USB, Firewire, or even Ethernet. The PFH excludes any part of the disk driver that is used to service user requests (for example, querying statistics) but not used to service kernel requests.
The CapROS model is that user data and capabilities are persistent, meaning their state is preserved even when the system is rebooted. Persistence is accomplished by checkpointing the state to nonvolatile backing storage. By checkpointing all user state, we assure that the security state (the state of the capabilities) is valid and consistent.
The problem is that the PFH, though part of the user state, cannot be persistent, for several reasons:
Therefore, we must declare that some user state is non-persistent. Non-persistent objects will live in OID ranges that are marked as such. These ranges will be loaded by the boot loader (in GRUB, they are considered to be "modules"). Non-persistent objects must never be cleaned, checkpointed, or written back to disk. They must never be removed from memory. (It may prove useful to allow non-persistent objects to be explicitly released from memory, after a configuration procedure has determined they are not needed.)
We can now identify four different classes of objects:
Note that objects move between classes 3 and 4 as they are pinned and unpinned. And an object in class 2 effectively moves to class 1 while it holds a lock that is also used by the PFH.
Also note that if the system is configured as a diskless or "embedded" system, the entire system is non-persistent. Such systems need not be considered further here.
We said above that the system takes checkpoints to ensure the consistency (including security) of the user data. Now the question is, what is the consistency contract between the persistent and non-persistent worlds? For example, we want to avoid scenarios such as the following :
Here are the design constraints for the interaction of the four classes of objects.
The PFH must never reference unpinned objects, because that would lead to a deadlock. (It is permissible to have a capability to such an object, as long as it is never used.)
The PFH may read persistent pinned objects, but must never write them. The reason is that at the instant a checkpoint is taken, persistent objects are made temporarily read-only until their state can be saved. Saving the state may require using the PFH.
It follows that the PFH cannot use a gate key to any persistent process. A process should not CALL a start key to the PFH, because the PFH must not receive and use the resume key.
More generally, the PFH could be designed to handle its own page faults recursively. There needs to be some limit on the depth of recursion. For simplicity, we will design with no recursion.
Another constraint applies to locks (such as semaphores, mutexes, or spinlocks) that are used by the PFH. While holding such a lock, a process must obey the same rules as the PFH.
When the kernel allocates a page frame, if space is not available, it begins cleaning pages. Meanwhile the process requesting the space waits for the cleaning to finish. Similarly, when a node frame is needed, it may require cleaning nodes into a node pot, which in turn may require a page frame, which may require page cleaning.
If the PFH needs to allocate a page frame, this could lead to deadlock. The PFH may allocate pages or nodes to expand heap space, or to build a mapping table (mapping tables are allocated from page frame space).
To address this problem, the kernel reserves a number of page frames (KTUNE_MapTabReserve) for mapping tables and for allocations by the PFH. If this number is large enough to guarantee that the PFH can finish cleaning, then the deadlock is eliminated.
On a restart, the state of the non-persistent world is reset to its initial state. This means that any capabilities acquired during execution will be lost. This should not be a problem for the non-persistent world, because it will be initialized to a consistent state. However, a persistent object that expected to be invoked by a non-persistent object may wait forever.
We will solve this problem for resume keys. When the system takes a checkpoint, it will examine all non-persistent nodes. For each resume key it finds to a persistent process, it will save the state of the waiting process such that if restarted from that state, the process will find that the CALL has returned with a distinctive reply code (RC_capros_key_Restart). Alternatively, it could save the state such that if restarted, the process will run and re-execute its CALL invocation. Assuming that the CALL was directly to a non-persistent process, that cap will have been rescinded (see below), and the response from the now-void cap will signal the restart.
In that case, the operation that was requested by the CALL invocation may have partially completed before the last checkpoint, and partially or fully completed after the last checkpoint and before the restart. This behavior is unavoidable, because the system cannot in general checkpoint and roll back the state of I/O devices. Programs that invoke I/O devices must be prepared to cope with this possibility. In some cases, the driver can respond to the re-executed CALL by saying "the device is not now known to be in a state where you can perform that request (for example, not "open").
Any keys other than resume keys will be dropped on restart. If a non-persistent object held a start capability to a persistent process, the persistent process will not be notified that the start capability has been dropped. As for security, it should be clear that dropping capabilities does not open any vulnerabilities.
Non-persistent objects in class 2 could be initialized with capabilities to persistent objects, but they must recognize that the persistent object's state may get out of sync.
Persistent objects may hold capabilities to non-persistent objects. A persistent process may hold and call a start key to an object in class 2. When the system is restarted, from the point of view of the persistent object, the non-persistent objects appear to travel backwards in time to their initial state. To alert the persistent object that this has happened, on restart we will rescind all capabilites in persistent objects to non-persistent objects.
(The effect of this is similar, from the point of view of the persistent world, to what would happen if the non-persistent world were created out of a single space bank, and on restart that bank were asked to reclaim all the space created from it. Note that on restart, unlike the SpaceBank.destroySpace operation, resume caps held in objects created under the bank will be effectively invoked with an error code. The equivalent KeyKOS space bank operation did invoke such resume caps, and we believe the CapROS bank should also.)
This can be implemented as follows. During execution, non-persistent objects may be allocated and rescinded as normal, from a space bank that allocates from a non-persistent range. This may result in incrementing the object's allocation count. We will keep track of the maximum allocation count or call count of any non-persistent object, and save that on a checkpoint. On restart, all non-persistent objects are given an allocation count and call count equal to that maximum plus one. All capabilities in persistent objects to non-persistent objects will thus be rescinded. Capabilities to non-persistent objects in non-persistent objects will be initialized to contain the new allocation count, so those capabilities will be valid.
There are two ways to calculate the maximum allocation/call count.
There is an implementation wrinkle. On restart, we load and start the non-persistent processes, which then read the checkpoint headers. We will have to initially load the non-persistent objects with zero allocation and call counts. Once the needed counts are read from the checkpoint headers, and before any persistent objects are loaded, we must adjust the counts of the already-loaded non-persistent objects.
The design for communication between the persistent and non-persistent worlds respects the severe limitations described above.
There is a persistent node called (for historical reasons) Volsize with a well-known OID (PVOLSIZE_OID which is 0x0100000000000000). A slot (volsize_nplinkCap) of this node contains a start capability to a persistent bridge process called NPLink (Non-Persistent Link). Non-persistent drivers access the persistent Volsize node to obtain the capability to NPLink.
When a device, to which access is required from persistent objects, is recognized, either when the system is booted, or when the device is later attached to the system, the driver sends to the start cap to the NPLINK process a message containing a cap (the "device cap") (usually a start cap) to the non-persistent object (which must be in class 2) that provides access to the device. The NPLink process is responsible for granting this cap to the appropriate client of the device.
Communication between the device and the client proceeds normally until the system restarts. Then, as described above, the device cap, and any other caps to non-persistent objects, which may have been derived from the device cap, are rescinded. The client will discover the rescission the next time it tries to use one of these caps, and will use that information to initiate whatever restart recovery is appropriate. If and when the device is present after the restart, a new device cap to it will be created and distributed once again as described above.
This design allows either side of the interface to be upgraded, while protecting against the possibility of an unscheduled restart at any time. To upgrade the non-persistent side:
Upgrading the persistent side can be done as follows:
Consider the driver for a USB mouse.
The mouse is obviously not a paging device, but it interfaces with the USB bus, which may also have paging devices. In the current implementation, the USB host controller driver (which manages the bus) has dependencies on the USB device drivers. For example, if the mouse is unplugged, the mouse driver receives a notification, to which it must respond promptly. While processing such a notification, the mouse driver is in class 1, because a page fault may depend on the USB host controller driver which depends on the mouse driver.
Because USB devices are hot-plugged, the driver is constructed as needed. The constructor has as a constituent a capability to the non-persistent Volsize node, which has a capability to the persistent Volsize node, which has a start capability to the NPLink process. When a mouse is recognized (either when the system is booted, or when a device is plugged in), a driver is constructed, consisting of two processes (threads). One thread responds to requests from the USB bus handler, which it must do as class 1. The second thread will respond to requests from persistent processes for mouse I/O. The second thread will be in class 2 while it is receiving or replying to requests, and must not hold any locks needed by the first thread at that time.
The second thread passes to the NPLink process a start capability to itself, which we will call a mouse capability. If the mouse is unplugged, the driver rescinds the mouse cap. If the system restarts, the system rescinds the mouse cap.
Some drivers, such as the serial port driver, are totally separate from the PFH. They can be entirely in class 2.
A device driver could be divided into different parts, each in a different one of the four classes of objects. In many cases a driver will need a persistent front end that presents a more usable interface to clients, and can handle multiple independent clients. The only requirement is that any part used by the kernel for paging is part of the PFH and therefore must be non-persistent.
Perhaps the simplest design is to have a part of the driver in class 2 that is the user interface, and the rest in class 1 acting as a server for both the class 2 part and the kernel. The user interface can pin (and unpin) user objects if it is necessary for the class 1 part to read them.
Another design issue concerns protection boundaries. Class 3 objects have more authority than class 4 (because they can pin), and class 2 have more than class 3 (because they can cause time travel). We want to keep different authorities in different processes, and use gate keys to transfer control between them. For efficiency, we want to avoid the need for one gate jump from 4 to 3, and a second jump from 3 to 2. (And possibly more from 2 to 1.) In other words, we want to traverse the most authority change with the fewest gate jumps. From class 4, we cannot call class 1, so calling class 2 makes sense.
Support for device drivers in user processes was added to the Linux kernel in version 2.6.23-rc4 (drivers/uio). We took a quick look at this feature and discovered the following:
Apparently kernel code is also needed to support open, release, and mmap.
The interface between the kernel and the disk driver(s) has yet to be fully designed, but should be straightforward. For paging in, the driver periodically calls a special key to learn of any OIDs (object identifiers) that need to be fetched. The driver also receives a handle for the Activity that needs to be awakened when the OID is available. This handle could be a special key that the driver invokes when it is done. Alternatively, the kernel could enqueue the Activity on a wait queue based on a hash of the OID, and when done the driver could instruct the kernel to wake up that queue.
|Copyright 2006, 2007, 2008, 2009 by Strawberry Development Group. All rights reserved. For terms of redistribution, see the GNU General Public License This material is based upon work supported by the US Defense Advanced Research Projects Agency under Contract No. W31P4Q-07-C-0070. Approved for public release, distribution unlimited.|