- staffKeymasterSeptember 22, 2017 at 3:48 pmPost count: 47
Talk by Ivan Godard – October 4, 2017, at the
Silicon Valley Linux Users Group
Slides: 2017-10-04-IPC.4 (.pptx)
This was the twelfth topic publicly presented related to the Mill general-purpose CPU architecture. It covers Inter-Process Communication for the Mill CPU architecture family. The talk assumes a familiarity with aspects of CPU architecture in general and C++ programming in particular.
The Mill is a new general-purpose architectural family, with an emphasis on secure and inexpensive communication across protection boundaries. The large (page) granularity of protection on conventional architectures makes such communication difficult compared to communication within a protection boundary, such as a function call. As a result, the large granularity has forced communication protocols on conventional architectures into two models: pass-by-sharing (using shared pages), and pass-by-copy (using the OS kernel for files/message passing). Both have drawbacks: sharing requires difficult-to-get-right synchronization, while copy involves kernel transitions as well as the costs of the copy itself.
The Mill supports both these protocols, for use by legacy code. However, the Mill hardware also supports inter-process communication using the same program protocols as for intra-process communication and function call: pass-by-value, pass-by-copy, pass-by-reference, and pass-by-name, but all without kernel involvement or overhead. The protocols are secure: neither party can see anything of the other except the explicit arguments to the communication. Neither caller nor callee codes need source changes to replace intra-process communication with Mill inter-process argument passing. However, the pass-by-reference protocol may require use of shims to delimit the extent of sharing in some languages. And granularity is no longer an issue: arguments can be of any size down to the byte.
The talk describes the machinery behind the Mill IPC protocols, together with suggestions as to how the hardware facilities may be integrated with representative language runtime systems such as those found in Linux.
Ivan Godard is CTO and a founder of Mill Computing, Inc., developer of the Mill family of general-purpose CPUs. He has written or led the development team for a dozen compilers, an OS, an OODBMS, and much other software. Ivan has been active in the field of computers since the 1960s.
- Christophe BioccaParticipantJanuary 3, 2018 at 1:29 pmPost count: 3
As far as the wiki is concerned, whether a call is a portal call or a regular call is not determined by the caller (both use the same
callopcode). Portal-call vs. in-turf execution vs. fault is determined by the
permissions of the calling turf on the entry point of the called code.
In addition to the above, a turf can grant/revoke portal/execute permissions on their code to other turfs, as a way to achieve microkernel-style privilege separation.
This protects the caller, but the callee can’t know for certain that the foreign code they’re calling into will actually run in a different turf. I could write an evil portal service that grants plain execute to would be callers, and they wouldn’t be able to tell.
Am I missing something? Is this just a matter of policy (always have a separate, trusted, turf control the portal permissions so that you can trust that they’re set properly and never changed out from under you), or is there a way for a caller to say “this must be a portal call, fault otherwise”?
- This reply was modified 2 years ago by Christophe Biocca.
- Ivan GodardKeymasterJanuary 3, 2018 at 2:33 pmPost count: 558
A portal causes turf switch to a turf id contained in the portal structure. There are barriers to the vulnerability you suggest.
If the attacker gave the victim a code pointer that falsely purports to be a portal and the victim called it then the victim would still be in his original turf, executing the code referenced by the passed pointer. However, the victim must have execute rights for any code, so the substitute code must be executable by the victim’s turf; it can’t be attacker code because the victim does not have execute rights to attacker code. And the attacker cannot blindly give such rights to the victim; there is a check so that a suspicious victim must accept a proposed grant before it takes effect.
Thus the target address must thus be a valid entry point in the victims own code. Of course, getting the victim to call one of his own functions when he didn’t intend to is problematic too. There is a check, a bit more general than you suggest, that an untrusting program can use for this. It returns, for a given address, what permissions the caller has at that address. That check is necessary in a number of ways, but seems inelegant and we have been exploring alternatives, but with nothing entirely satisfactory yet.
Second, the portal structure itself is set up by trusted code, which always sets the associated turf to that of the thread creating the portal. That is, you can create portals into yourself, but not into anyone else.
- Christophe BioccaParticipantJanuary 4, 2018 at 5:38 amPost count: 3
And the attacker cannot blindly give such rights to the victim; there is a check so that a suspicious victim must accept a proposed grant before it takes effect.
That does solve the issue I had in mind: the attacker granting execute instead of portal permissions to its own code, the victim, unaware of this, calling into it and giving attacker-controlled code in-turf access.
Making each grant subject to approval does solve this issue, and means the program can simply check a purported portal call address once (to not have execute permissions at all) and use it safely afterwards. Then even revocation of the portal permission wouldn’t later cause problems for the program (beyond faulting).
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