Fedora desktop Planet

New LVFS redirect behavior

Posted by Richard Hughes on November 25, 2021 04:20 PM

tl;dr: if you’re using libfwupd to download firmware, nothing changes and everything continues as before. If you’re using something like wget that doesn’t follow redirects by default you might need to add a command line argument to download firmware from the LVFS.

Just a quick note to explain something that some people might have noticed; if you’re using fwupd >= 1.6.1 or >= 1.5.10 when you connect to the LVFS to download a firmware file you actually get redirected to the same file on the CDN. e.g. downloading https://fwupd/download/foo.cab gets a redirect to https://cdn.fwupd/download/foo.cab which is then streamed to the user. Why this insanity?

As some of you know, egress charges from AWS are insanely high. The Linux Foundation are the kind people that kindly pay the LVFS bill every month, and a 4 years ago that was just a few hundred dollars and that was a rounding error to them. Last year we again grew at more than 100% and the projection for next year is going to surpass even that; the average size of firmware files has gone from ~30MB to ~50MB with much, much, larger server firmware in the pipeline. We certainly can’t watch the egress bill scale linearly with the LVFS popularity, else some accountant at the Linux Foundation is going to start asking questions – especially when Fastly provides the LF a geo-replicated CDN – which we’re not using.

So why don’t we put the CDN URL in the XML metadata directly, and then avoid all this redirect complexity altogether? This time the lawyers get us, as we’re required by US law to restrict distribution of some firmware to some countries on an embargo list. It’s very complicated, and it varies by vendor, but it’s not something we can avoid. So for this reason, the LVFS does a GeoIP lookup on the client IP, and if it’s all okay we then redirect the client to the CDN-cached version. It also lets us tell the vendor how many times the firmware has been downloaded without importing the CDN logs every 24 hours – which would be even harder as we only keep them for a short time for privacy reasons.

Introducing GNOME Crosswords

Posted by Jonathan Blandford on November 18, 2021 08:00 AM

GNOME Crosswords

Howdy folks! I want to announce a game for GNOME that I’ve been working on for a few months.

I’ve always enjoyed solving Crossword puzzles. It’s something I grew up doing as a kid, and we continue to do them as a family at the dinner table at night. I’ve wanted to try my hand at writing crosswords for a while, but there isn’t really a good tool available for doing so, and certainly no free software ones that work well with a recent GNOME release. I recently bought myself a lovely new Fedora-loaded Lenovo, and after it arrived, I thought I’d take a shot at writing such a tool.

Over the past four months or so I managed to get something worth releasing. The code is available here. It should build on relatively recent Linux distributions, though it does need libadwaita from git (toasts!). I also put together a flatpak file for testing here (no repo yet, as getting that set up defeated me). Once I’m more confident that the puzzles are solvable and fun I plan to publish it to flathub.

<figure aria-describedby="caption-attachment-6652" class="wp-caption aligncenter" id="attachment_6652" style="width: 480px">A dog's day<figcaption class="wp-caption-text" id="caption-attachment-6652">Non-traditional grid</figcaption></figure> <figure aria-describedby="caption-attachment-6649" class="wp-caption aligncenter" id="attachment_6649" style="width: 480px">Guardian cryptic No 28,605<figcaption class="wp-caption-text" id="caption-attachment-6649">The Guardian Daily Cryptic with reveal answers enabled</figcaption></figure>

Features:

It’s still early, but it already has some fun features:

  • Puzzle Sets. The heart of the game is the Puzzle Set. It’s a collection of crossword puzzles that are tied together by a theme. Solving a puzzle unlocks more puzzles. I currently have one puzzle set (“Cats and Dogs”) with nine puzzles in it, but I have a few more puzzle sets planned. It contains mostly traditional puzzles, but I threw in a cryptic to keep people on their toes.
  • Nontraditional shapes and styles: I wanted to make something a bit little more whimsical and fun, as well as the more traditional puzzle grids. So I added support for colors and shapes as well. My son had fun doing pixel art to create some of the grids.
  • Reveal mistakes: For when you get stuck! It also supports checksums for puzzles that don’t include the solution.
  • Scalable grid: Currently the UI only exposes four sizes, but we have all the pieces to scale crosswords to different sizes.
  • Support for the .ipuz spec: This spec supports a ton of things, and I don’t support it fully yet, but most of the crossword part of the spec is included. There aren’t a ton of .ipuz files floating around, but you can use puzzlepull to download the Guardian Daily puzzle if you want to try some other examples.
<figure aria-describedby="caption-attachment-6691" class="wp-caption aligncenter" id="attachment_6691" style="width: 261px">Puzzle Set<figcaption class="wp-caption-text" id="caption-attachment-6691">The first Puzzle Set</figcaption></figure>

Crossword Editor

GNOME Crosswords Editor

As part of building this app, I realized that creating grids was as big a part of the app as writing the actual game itself. To facilitate that, I started writing a crossword editor as well. It’s in the early stages, but it already has one of the most important features: a tool to create the initial grid. Making puzzles that fit well together is surprisingly hard. To make it easier, I wrote a crossword solver that quickly suggests words to fill in the grid. I’m proud of the design – it’s able to efficiently suggest options out of a list of 500K words really quickly (<1 μs on my machine). I was able to use it to build an autofill dialog that can recursively fill in a section of the puzzle when making a grid.

I still have more work to do on the editor and it’s clear that the autofill dialog isn’t a panacea, but it helped me figure out out some tricky corners. Here’s a video of the autofill dialog in action:

<video class="wp-video-shortcode" controls="controls" height="473" id="video-6632-1" preload="metadata" width="840"><source src="https://blogs.gnome.org/jrb/files/2021/11/Screencast-from-11-16-2021-071758-PM.webm?_=1" type="video/webm">https://blogs.gnome.org/jrb/files/2021/11/Screencast-from-11-16-2021-071758-PM.webm</video>

Thanks

I especially want to thank Rosanna and my kids for play-testing this and suggesting clues, as well as their patience while I was writing it. Thanks also to Federico for giving great advice, great code, and for being a star. Matthias for helping me relearn GTK and explaining GtkIMContext. Also, the example code in GNOME Builder was immensely helpful for getting this started.

What’s next?

There are a ton of features I’d like to add to this game. It really needs printing support, which should be relatively easy. I’d also love to see it get internationalized (and not just translated) – are crosswords in non-Latin languages a thing? And I’ve seen enough of Benjamin‘s GUADEC presentations over the years to know GTK can do something cooler than popup a dialog when you finish a crossword.

But the most important thing is that the game needs to be fun! For that, we need more puzzles and the existing puzzles need to be better. If you’re interested in joining me in creating a good set of puzzles for Linux, try the game out and let me know.

Pango Updates

Posted by Matthias Clasen on November 10, 2021 04:07 AM

Here is another update on what will appear in Pango 1.50 soon. Quite a few things, as it turns out. (This is a continuation of my last Pango update)

Bidi Improvements

Pango has long claimed to have good support for bidirectional text rendering and editing. But until recently, there were some annoying bugs: you could end up in loops when moving the cursor through mixed-direction text.

The relevant code has been rewritten, closing a very old bug (#157).

Useful Information

Modern fonts can contain a lot of different information – there can be colormaps, and glyphs can be represented not just as splines, but also as pngs or svgs.

For GTK, an important piece of information for each glyph is whether it is colored or monochrome. With the new PangoGlyphVisAttr.is_color field,  GTK no longer needs to poke directly at the font data to find out.

Another case where GTK needed data that wasn’t available through Pango APIs has been closed with pango_font_get_languages(). This data is used in the font chooser filter popup:

Superscripts and subscripts

Superscripts and subscripts now use font metrics information for size changes and baseline shift.  They can also be properly nested, so markup like

2<sup>2<sup>2</sup></sup>

yields the expected rendering:

Customizing segmentation

Pango determines word and sentence boundaries according to the Unicode Text Segmentation Spec (TR29). The specification can’t deal with all the complexities of formatted text, so sometimes its results need tailoring, which can now be done with the new word and sentence attributes:

$ pango-segmentation --kind=word --text=\
    "<span segment='word'>1-based</span> index"

|1-based| |index|

Better Markup

Besides the new attributes that have been mentioned, Pango markup now lets you specify many attributes in more natural units.

For example,  you can now say

<span font_size='12.5pt'>

instead of

<span font_size='12800'>

Small Caps

Pango has had the PangoVariant enumeration with its PANGO_VARIANT_SMALL_CAPS value since forever, but it has never done anything. Since we’ve added support for OpenType features (in Pango 1.37), it has been possible to produce Small Caps by using the smcp=1 feature, if the font supports it. Sadly, most fonts don’t.

With the text transformation and font scaling infrastructure now in place, it was easy to emulate Small Caps for fonts that don’t  support this natively:

While at it, we’ve expanded the PangoVariant enumeration to cover all the CSS casing variants, and made the GTK CSS engine use them instead of OpenType features.

Better debugging

Pango ships with a versatile test utility called pango-view, which has options to test many of Pangos layout features. It has recently learned to show more auxiliary information in visual form, like glyph extents, caret positions and slopes:

$ pango-view --text Boxes \ --annotate=glyph,caret,slope

Enjoy!

Firefox 94 comes with EGL on X11

Posted by Martin Stransky on October 30, 2021 12:05 PM
<figure class="wp-block-image size-large"><figcaption>(In)Famous WebGL Aquarium demo. EGL brings you more fish 🙂</figcaption></figure>

Firefox 94 is coming out next week and brings awesome news. OpenGL EGL backend its enabled by default on Intel/AMD and recent Mesa for users on X11.

This project has been driven by Robert Mader (most of the EGL work), Andrew Osmond (glxtest fixes and config), Jamie Nicol (EGL/Android and partial damage), Greg V (partial damage support) and Jan Ikenmeyer (Darkspirit) (help with issues, testing).

Historically Linux comes with GLX (OpenGL X11 extension) but that era is finally ending and we’re moving forward to EGL which promises all the goodness you can expect from modern graphics subsystem…or you can create a texture over graphics memory at least 🙂

I’ll keep aside all EGL / GLX difference and focus to changes from user perspective. GLX is old, well debugged and tied closely to X11 which means seamless experience and wide support by gfx drivers (like proprietary NVIDIA ones). It’s used by most X11 applications and ‘just works’.

EGL is ‘new’ from Linux desktop perspective and used mainly by Wayland, Android and various small devices. It’s not fully supported by all desktop drivers and has glitches (broken rendering of transparent windows for instance). But as Wayland is gaining momentum also EGL is getting more attention and fixes on Linux desktop.

And why we actually want EGL? Because it gives us a cool toy – EGLImages (and EGLFence). EGLImage is an object which is created over a piece of GPU memory (which can be DMA-Buf), shared with different process, used as a frame buffer (target of GL rendering) or a texture (source of GL rendering).

EGLImage allows to use GPU memory in a very creative way. VA-API decoded video frames or WebGL scenes can be mapped as EGLImages, moved from decode process to rendering process and used as a texture. EGLFence allows to lock EGLImages across process so we don’t re-paint WebGL scene while it’s used in a different process or recycle VA-API video frames too early.

And what you can expect from EGL in Firefox? Faster WebGL rendering (used on Google Maps for instance), more effective rendering due to to partial damage support and potential VA-API video decoding (that’s blocked by Bug 1698778 on both Wayland and X11). It also unifies rendering path for Wayland and X11, which means X11 will gain features done for Wayland (suspended rendering for invisible windows, better VSync support and more).

NVIDIA is also working on EGL & DMA-Buf support in their proprietary drivers so there’s a hope for owners of such FOSS unfriendly hardware.

So give Firefox 94 a try. If anything goes wrong, please file a bug. You can also disable EGL and switch back to GLX. Go to about:config page and flip gfx.x11-egl.force-disabled preference and restart browser.

GTK4: Toolbars in Sidebar

Posted by Caolán McNamara on October 28, 2021 04:28 PM

 

GTK4 port of Libreoffice now supports the "widebutton" Toolbar MenuButtons that show a preview of the selected color.

PSA: gnome-settings-daemon's MediaKeys API is going away

Posted by Bastien Nocera on October 20, 2021 12:12 PM

 In 2007, Jan Arne Petersen added a D-Bus API to what was still pretty much an import into gnome-control-center of the "acme" utility I wrote to have all the keys on my iBook working.

It switched the code away from remapping keyboard keys to "XF86Audio*", to expecting players to contact the D-Bus daemon and ask to be forwarded key events.

 

Multimedia keys circa 2003

In 2013, we added support for controlling media players using MPRIS, as another interface. Fast-forward to 2021, and MPRIS support is ubiquitous, whether in free software, proprietary applications or even browsers. So we'll be parting with the "org.gnome.SettingsDaemon.MediaKeys" D-Bus API. If your application still wants to work with older versions of GNOME, it is recommended to at least quiet the MediaKeys API's unavailability.

 

Multimedia keys in 2021
 

TL;DR: Remove code that relies on gnome-settings-daemon's MediaKeys API, make sure to add MPRIS support to your app.

Firefox Wayland development in 2021

Posted by Martin Stransky on October 01, 2021 08:45 AM
<figure class="aligncenter size-large"><figcaption>I swear, no more crashes on Wayland!</figcaption></figure>

It’s been long time from my last update about Firefox news on Linux and I’ve finally got some time to sum up what we’ve been working on for last year and what’s coming. There haven’t been introduced any new exciting features (from Linux perspective) for the last year but rather a hidden but important changes.

From Linux desktop developers perspective 2021 is a year of Wayland. KDE has been shipping decent Wayland compositor which becomes default for Fedora 34. It’s actually pretty fast and gives you smooth feeling of “good old times” with X11/Gtk2/name-your-favorite environment where any graphics change was just instant without lags or slow transitions. I must mention Robert Mader who created a new Firefox Wayland SW backend for the KDE.

As major desktop distro (Ubuntu) is slowly moving to Wayland we’re getting more and more Wayland/Firefox users. Even notorious troublemaker (NVIDIA) decided to step in and support it.

What’s done for next releases?

It’s good that Wayland market share is rising but we also need to make sure that Firefox is ready to run there without any major issues and matches its X11 variant. There are two major areas where Firefox is behind its X11 counterpart – clipboard and popup handling. It’s given by some Wayland protocol features where we can’t simply duplicate the X11 code here.

Clipboard on Wayland is similar to X11 one but we need to translate Wayland asynchronous clipboard to Firefox/Web synced one. I tried various approaches but the best one seems to be just use the asynchronous Wayland clipboard as is and implement some kind of abstraction over it. That was implemented in Firefox 93 an it’s going to be shipped by default in Firefox 94.

On the other hand popups are the most annoying thing we have to implement on Wayland. Firefox just expect any popup can be created any time without its parent (or use main window as a parent) but Wayland requests strict popup hierarchy. It means every window can have only one child popup. When more than one popup is opened it has to be attached to the previously opened popup which becomes a parent for it. And when any popup in the chain is closed the popups must be rearranged to keep the chain connected. This involves all kind of popups like content menus, tooltips, extensions popups, permissions popups and so on. Plus there are some interesting bugs in Wayland protocol or Gtk so excitement/frustration is guaranteed and basic popup implementation becomes extraordinary challenge where small changes can introduce various regressions. Despite the ‘fun with popups’ the popup tracker is almost clear and we’ll ship it in Firefox 94.

One of main Wayland feature is support of monitors with various DPI/scale factor together. Fedora default compositor Mutter shows here a creative approach and reports screen sizes differently than other compositors. As we really want to know screen sizes Firefox tracks monitor changes from Wayland directly and find correct screen by matching screen left top corner point – which fortunately stays stays same for all compositor. We also stop painting Firefox window when screen scale changes so you should enjoy seamless experience on systems with various screen sizes with Firefox 93.

Future plans for Firefox 95

Firefox 95 development cycle begins next ween and I’m going to look to drag and drop which has been partially broken for long long time. Some Wayland specific fixes are already in Firefox 94 but we need to rework some parts to correctly copy files from remote destinations (like inbox) to local filesystems, fix names of dropped files or do tab preview of moved tabs. There are also new interesting compositors bugs as usually 🙂

Future plans for Firefox 96

Firefox Wayland port is generally done and there isn’t any big difference between X11 and Wayland variant at least on GNOME which Fedora uses as default environment. We’re fixing minor bugs and keep eye on user reports.

For next quarter I’d like to look at GPU process for Wayland. GPU process is running tasks related to graphics hardware and shields browser from HW drivers crashes. It’s also place where VAAPI video decoding should be run and will be properly sandboxed there (right now VAAPI is run in content process along general Firefox code, it’s restricted by content sanbox which leads to various VAAPI code crashes and failures).

Reducing the effectiveness of a safety feature

Posted by Richard Hughes on September 23, 2021 03:37 PM

We just purchased a 2021 KIA eNIRO to use as our family car. As typical with EVs, this has to produce a fake engine noise to avoid squashing the hard of hearing (or animals) not expecting two tons of metal to be silently moving. When we test drove a 2020 car last year, there was a physical button on the dash to turn this off, on the premise that the noise sometimes isn’t required or appropriate, but it always defaulted to “on” for every start. As the car gets faster the noise also increases in volume, until after about 30km/h fading to nothing. In reverse the same thing happens with some additional beeps. Getting our 2021 car this year, the button was no longer present and the less-than-affectionately known VESS module cannot be muted. I can guess why, someone probably turned it off and squashed something or someone and someone in the UK/US/EU government understandably freaked out. KIA also removed the wire from the wiring loom, and won’t sell the 2020 button cluster, so it’s not even like you can retrofit a new car to act like the old one.

To be super clear: I don’t have a problem with the VESS noise, but because the “speaker” is in the front bumper the solution for going backwards is “turn up the volume”. Living in London means that houses are pretty close together and me reversing into the drive at 2mph shouldn’t submit the house opposite with a noise several times louder than a huge garbage truck. The solution in the various KIA owner forums seems to be “just unplug the VESS module” but this seems at best unethical and probably borderline illegal given it’s a device with the express purpose of trying to avoid hurting someone with your 2 ton lump of metal.

Given VESS is, as you might expect, just another device on the CAN bus and people have reverse engineered the command stream so you can actually just plug the VESS module into a USB device (with a CAN converter) and play with it yourself. My idea would be to make a modchip-like device that plugs into the VESS module using the existing plug, and basically MITM the CAN messages. All messages going from VESS back to the ECU get allow-listed (even though the ECU doesn’t seem to care if the VESS goes AWOL…) and any speed measurements going forward also get passed straight through. The clever part would be to MITM the speed when the “reverse gear” command has been issued, so that the car thinks it’s going about 20km/h backwards. This makes the VESS still make the engine and beeping noise but it’s only about as loud as the VESS module when going forwards when outside the car.

Technically this is quite easy, VESS->txrxCAN->MCU->txrxCAN->ECU and you can probably use an inexpensive Microchip reference board for a prototype. My question is more if this would:

  1. Be ethical
  2. Be legal
  3. Invalidate my insurance
  4. Invalidate the warranty of my nice new shiny car

Feedback very welcome!

What's new in XI 2.4 - touchpad gestures

Posted by Peter Hutterer on September 23, 2021 05:26 AM

After a nine year hiatus, a new version of the X Input Protocol is out. Credit for the work goes to Povilas Kanapickas, who also implemented support for XI 2.4 in the various pieces of the stack [0]. So let's have a look.

X has had touch events since XI 2.2 (2012) but those were only really useful for direct touch devices (read: touchscreens). There were accommodations for indirect touch devices like touchpads but they were never used. The synaptics driver set the required bits for a while but it was dropped in 2015 because ... it was complicated to make use of and no-one seemed to actually use it anyway. Meanwhile, the rest of the world moved on and touchpad gestures are now prevalent. They've been standard in MacOS for ages, in Windows for almost ages and - with recent GNOME releases - now feature prominently on the Linux desktop as well. They have been part of libinput and the Wayland protocol for years (and even recently gained a new set of "hold" gestures). Meanwhile, X was left behind in the dust or mud, depending on your local climate.

XI 2.4 fixes this, it adds pinch and swipe gestures to the XI2 protocol and makes those available to supporting clients [2]. Notably here is that the interpretation of gestures is left to the driver [1]. The server takes the gestures and does the required state handling but otherwise has no decision into what constitutes a gesture. This is of course no different to e.g. 2-finger scrolling on a touchpad where the server just receives scroll events and passes them on accordingly.

XI 2.4 gesture events are quite similar to touch events in that they are processed as a sequence of begin/update/end with both types having their own event types. So the events you will receive are e.g. XIGesturePinchBegin or XIGestureSwipeUpdate. As with touch events, a client must select for all three (begin/update/end) on a window. Only one gesture can exist at any time, so if you are a multi-tasking octopus prepare to be disappointed.

Because gestures are tied to an indirect-touch device, the location they apply at is wherever the cursor is currently positioned. In that, they work similar to button presses, and passive grabs apply as expected too. So long-term the window manager will likely want a passive grab on the root window for swipe gestures while applications will implement pinch-to-zoom as you'd expect.

In terms of API there are no suprises. libXi 1.8 is the version to implement the new features and there we have a new XIGestureClassInfo returned by XIQueryDevice and of course the two events: XIGesturePinchEvent and XIGestureSwipeEvent. Grabbing is done via e.g. XIGrabSwipeGestureBegin, so for those of you with XI2 experience this will all look familiar. For those of you without - it's probably no longer worth investing time into becoming an XI2 expert.

Overall, it's a nice addition to the protocol and it will help getting the X server slightly closer to Wayland for a widely-used feature. Once GTK, mutter and all the other pieces in the stack are in place, it will just work for any (GTK) application that supports gestures under Wayland already. The same will be true for Qt I expect.

X server 21.1 will be out in a few weeks, xf86-input-libinput 1.2.0 is already out and so are xorgproto 2021.5 and libXi 1.8.

[0] In addition to taking on the Xorg release, so clearly there are no limits here
[1] More specifically: it's done by libinput since neither xf86-input-evdev nor xf86-input-synaptics will ever see gestures being implemented
[2] Hold gestures missed out on the various deadlines

An Xorg release without Xwayland

Posted by Peter Hutterer on September 22, 2021 11:00 PM

Xorg is about to released.

And it's a release without Xwayland.

And... wait, what?

Let's unwind this a bit, and ideally you should come away with a better understanding of Xorg vs Xwayland, and possibly even Wayland itself.

Heads up: if you are familiar with X, the below is simplified to the point it hurts. Sorry about that, but as an X developer you're probably good at coping with pain.

Let's go back to the 1980s, when fashion was weird and there were still reasons to be optimistic about the future. Because this is a thought exercise, we go back with full hindsight 20/20 vision and, ideally, the winning Lotto numbers in case we have some time for some self-indulgence.

If we were to implement an X server from scratch, we'd come away with a set of components. libxprotocol that handles the actual protocol wire format parsing and provides a C api to access that (quite like libxcb, actually). That one will just be the protocol-to-code conversion layer.

We'd have a libxserver component which handles all the state management required for an X server to actually behave like an X server (nothing in the X protocol require an X server to display anything). That library has a few entry points for abstract input events (pointer and keyboard, because this is the 80s after all) and a few exit points for rendered output.

libxserver uses libxprotocol but that's an implementation detail, we can ignore the protocol for the rest of the post.

Let's create a github organisation and host those two libraries. We now have: http://github.com/x/libxserver and http://github.com/x/libxprotocol [1].

Now, to actually implement a working functional X server, our new project would link against libxserver hook into this library's API points. For input, you'd use libinput and pass those events through, for output you'd use the modesetting driver that knows how to scream at the hardware until something finally shows up. This is somewhere between outrageously simplified and unacceptably wrong but it'll do for this post.

Your X server has to handle a lot of the hardware-specifics but other than that it's a wrapper around libxserver which does the work of ... well, being an X server.

Our stack looks like this:


+------------------------+
| xserver [libxserver]|--------[ X client ]
| |
|[libinput] [modesetting]|
+------------------------+
| kernel |
+------------------------+
Hooray, we have re-implemented Xorg. Or rather, XFree86 because we're 20 years from all the pent-up frustratrion that caused the Xorg fork. Let's host this project on http://github.com/x/xorg

Now, let's say instead of physical display devices, we want to render into an framebuffer, and we have no input devices.


+------------------------+
| xserver [libxserver]|--------[ X client ]
| |
| [write()] |
+------------------------+
| some buffer |
+------------------------+
This is basically Xvfb or, if you are writing out PostScript, Xprint. Let's host those on github too, we're accumulating quite a set of projects here.

Now, let's say those buffers are allocated elsewhere and we're just rendering to them. And those buffer are passed to us via an IPC protocol, like... Wayland!


+------------------------+
| xserver [libxserver]|--------[ X client ]
| |
|input events [render]|
+------------------------+
| |
+------------------------+
| Wayland compositor |
+------------------------+
And voila, we have Xwayland. If you swap out the protocol you can have Xquartz (X on Macos) or Xwin (X on Windows) or Xnext/Xephyr (X on X) or Xvnc (X over VNC). The principle is always the same.

Fun fact: the Wayland compositor doesn't need to run on the hardware, you can play display server matryoshka until you run out of turtles.

In our glorious revisioned past all these are distinct projects, re-using libxserver and some external libraries where needed. Depending on the projects things may be very simple or get very complex, it depends on how we render things.

But in the end, we have several independent projects all providing us with an X server process - the specific X bits are done in libxserver though. We can release Xwayland without having to release Xorg or Xvfb.

libxserver won't need a lot of releases, the behaviour is largely specified by the protocol requirements and once you're done implementing it, it'll be quite a slow-moving project.

Ok, now, fast forward to 2021, lose some hindsight, hope, and attitude and - oh, we have exactly the above structure. Except that it's not spread across multiple independent repos on github, it's all sitting in the same git directory: our Xorg, Xwayland, Xvfb, etc. are all sitting in hw/$name, and libxserver is basically the rest of the repo.

A traditional X server release was a tag in that git directory. An XWayland-only release is basically an rm -rf hw/*-but-not-xwayland followed by a tag, an Xorg-only release is basically an rm -rf hw/*-but-not-xfree86 [2].

In theory, we could've moved all these out into separate projects a while ago but the benefits are small and no-one has the time for that anyway.

So there you have it - you can have Xorg-only or XWayland-only releases without the world coming to an end.

Now, for the "Xorg is dead" claims - it's very likely that the current release will be the last Xorg release. [3] There is little interest in an X server that runs on hardware, or rather: there's little interest in the effort required to push out releases. Povilas did a great job in getting this one out but again, it's likely this is the last release. [4]

Xwayland - very different, it'll hang around for a long time because it's "just" a protocol translation layer. And of course the interest is there, so we have volunteers to do the releases.

So basically: expecting Xwayland releases, be surprised (but not confused) by Xorg releases.

[1] Github of course doesn't exist yet because we're in the 80s. Time-travelling is complicated.
[2] Historical directory name, just accept it.
[3] Just like the previous release...
[4] At least until the next volunteer steps ups. Turns out the problem "no-one wants to work on this" is easily fixed by "me! me! I want to work on this". A concept that is apparently quite hard to understand in the peanut gallery.

Authenticated Boot and Disk Encryption on Linux

Posted by Lennart Poettering on September 22, 2021 10:00 PM

The Strange State of Authenticated Boot and Disk Encryption on Generic Linux Distributions

TL;DR: Linux has been supporting Full Disk Encryption (FDE) and technologies such as UEFI SecureBoot and TPMs for a long time. However, the way they are set up by most distributions is not as secure as they should be, and in some ways quite frankly weird. In fact, right now, your data is probably more secure if stored on current ChromeOS, Android, Windows or MacOS devices, than it is on typical Linux distributions.

Generic Linux distributions (i.e. Debian, Fedora, Ubuntu, …) adopted Full Disk Encryption (FDE) more than 15 years ago, with the LUKS/cryptsetup infrastructure. It was a big step forward to a more secure environment. Almost ten years ago the big distributions started adding UEFI SecureBoot to their boot process. Support for Trusted Platform Modules (TPMs) has been added to the distributions a long time ago as well — but even though many PCs/laptops these days have TPM chips on-board it's generally not used in the default setup of generic Linux distributions.

How these technologies currently fit together on generic Linux distributions doesn't really make too much sense to me — and falls short of what they could actually deliver. In this story I'd like to have a closer look at why I think that, and what I propose to do about it.

The Basic Technologies

Let's have a closer look what these technologies actually deliver:

  1. LUKS/dm-crypt/cryptsetup provide disk encryption, and optionally data authentication. Disk encryption means that reading the data in clear-text form is only possible if you possess a secret of some form, usually a password/passphrase. Data authentication means that no one can make changes to the data on disk unless they possess a secret of some form. Most distributions only enable the former though — the latter is a more recent addition to LUKS/cryptsetup, and is not used by default on most distributions (though it probably should be). Closely related to LUKS/dm-crypt is dm-verity (which can authenticate immutable volumes) and dm-integrity (which can authenticate writable volumes, among other things).

  2. UEFI SecureBoot provides mechanisms for authenticating boot loaders and other pre-OS binaries before they are invoked. If those boot loaders then authenticate the next step of booting in a similar fashion there's a chain of trust which can ensure that only code that has some level of trust associated with it will run on the system. Authentication of boot loaders is done via cryptographic signatures: the OS/boot loader vendors cryptographically sign their boot loader binaries. The cryptographic certificates that may be used to validate these signatures are then signed by Microsoft, and since Microsoft's certificates are basically built into all of today's PCs and laptops this will provide some basic trust chain: if you want to modify the boot loader of a system you must have access to the private key used to sign the code (or to the private keys further up the certificate chain).

  3. TPMs do many things. For this text we'll focus one facet: they can be used to protect secrets (for example for use in disk encryption, see above), that are released only if the code that booted the host can be authenticated in some form. This works roughly like this: every component that is used during the boot process (i.e. code, certificates, configuration, …) is hashed with a cryptographic hash function before it is used. The resulting hash is written to some small volatile memory the TPM maintains that is write-only (the so called Platform Configuration Registers, "PCRs"): each step of the boot process will write hashes of the resources needed by the next part of the boot process into these PCRs. The PCRs cannot be written freely: the hashes written are combined with what is already stored in the PCRs — also through hashing and the result of that then replaces the previous value. Effectively this means: only if every component involved in the boot matches expectations the hash values exposed in the TPM PCRs match the expected values too. And if you then use those values to unlock the secrets you want to protect you can guarantee that the key is only released to the OS if the expected OS and configuration is booted. The process of hashing the components of the boot process and writing that to the TPM PCRs is called "measuring". What's also important to mention is that the secrets are not only protected by these PCR values but encrypted with a "seed key" that is generated on the TPM chip itself, and cannot leave the TPM (at least so goes the theory). The idea is that you cannot read out a TPM's seed key, and thus you cannot duplicate the chip: unless you possess the original, physical chip you cannot retrieve the secret it might be able to unlock for you. Finally, TPMs can enforce a limit on unlock attempts per time ("anti-hammering"): this makes it hard to brute force things: if you can only execute a certain number of unlock attempts within some specific time then brute forcing will be prohibitively slow.

How Linux Distributions use these Technologies

As mentioned already, Linux distributions adopted the first two of these technologies widely, the third one not so much.

So typically, here's how the boot process of Linux distributions works these days:

  1. The UEFI firmware invokes a piece of code called "shim" (which is stored in the EFI System Partition — the "ESP" — of your system), that more or less is just a list of certificates compiled into code form. The shim is signed with the aforementioned Microsoft key, that is built into all PCs/laptops. This list of certificates then can be used to validate the next step of the boot process. The shim is measured by the firmware into the TPM. (Well, the shim can do a bit more than what I describe here, but this is outside of the focus of this article.)

  2. The shim then invokes a boot loader (often Grub) that is signed by a private key owned by the distribution vendor. The boot loader is stored in the ESP as well, plus some other places (i.e. possibly a separate boot partition). The corresponding certificate is included in the list of certificates built into the shim. The boot loader components are also measured into the TPM.

  3. The boot loader then invokes the kernel and passes it an initial RAM disk image (initrd), which contains initial userspace code. The kernel itself is signed by the distribution vendor too. It's also validated via the shim. The initrd is not validated, though (!). The kernel is measured into the TPM, the initrd sometimes too.

  4. The kernel unpacks the initrd image, and invokes what is contained in it. Typically, the initrd then asks the user for a password for the encrypted root file system. The initrd then uses that to set up the encrypted volume. No code authentication or TPM measurements take place.

  5. The initrd then transitions into the root file system. No code authentication or TPM measurements take place.

  6. When the OS itself is up the user is prompted for their user name, and their password. If correct, this will unlock the user account: the system is now ready to use. At this point no code authentication, no TPM measurements take place. Moreover, the user's password is not used to unlock any data, it's used only to allow or deny the login attempt — the user's data has already been decrypted a long time ago, by the initrd, as mentioned above.

What you'll notice here of course is that code validation happens for the shim, the boot loader and the kernel, but not for the initrd or the main OS code anymore. TPM measurements might go one step further: the initrd is measured sometimes too, if you are lucky. Moreover, you might notice that the disk encryption password and the user password are inquired by code that is not validated, and is thus not safe from external manipulation. You might also notice that even though TPM measurements of boot loader/OS components are done nothing actually ever makes use of the resulting PCRs in the typical setup.

Attack Scenarios

Of course, before determining whether the setup described above makes sense or not, one should have an idea what one actually intends to protect against.

The most basic attack scenario to focus on is probably that you want to be reasonably sure that if someone steals your laptop that contains all your data then this data remains confidential. The model described above probably delivers that to some degree: the full disk encryption when used with a reasonably strong password should make it hard for the laptop thief to access the data. The data is as secure as the password used is strong. The attacker might attempt to brute force the password, thus if the password is not chosen carefully the attacker might be successful.

Two more interesting attack scenarios go something like this:

  1. Instead of stealing your laptop the attacker takes the harddisk from your laptop while you aren't watching (e.g. while you went for a walk and left it at home or in your hotel room), makes a copy of it, and then puts it back. You'll never notice they did that. The attacker then analyzes the data in their lab, maybe trying to brute force the password. In this scenario you won't even know that your data is at risk, because for you nothing changed — unlike in the basic scenario above. If the attacker manages to break your password they have full access to the data included on it, i.e. everything you so far stored on it, but not necessarily on what you are going to store on it later. This scenario is worse than the basic one mentioned above, for the simple fact that you won't know that you might be attacked. (This scenario could be extended further: maybe the attacker has a chance to watch you type in your password or so, effectively lowering the password strength.)

  2. Instead of stealing your laptop the attacker takes the harddisk from your laptop while you aren't watching, inserts backdoor code on it, and puts it back. In this scenario you won't know your data is at risk, because physically everything is as before. What's really bad though is that the attacker gets access to anything you do on your laptop, both the data already on it, and whatever you will do in the future.

I think in particular this backdoor attack scenario is something we should be concerned about. We know for a fact that attacks like that happen all the time (Pegasus, industry espionage, …), hence we should make them hard.

Are we Safe?

So, does the scheme so far implemented by generic Linux distributions protect us against the latter two scenarios? Unfortunately not at all. Because distributions set up disk encryption the way they do, and only bind it to a user password, an attacker can easily duplicate the disk, and then attempt to brute force your password. What's worse: since code authentication ends at the kernel — and the initrd is not authenticated anymore —, backdooring is trivially easy: an attacker can change the initrd any way they want, without having to fight any kind of protections. And given that FDE unlocking is implemented in the initrd, and it's the initrd that asks for the encryption password things are just too easy: an attacker could trivially easily insert some code that picks up the FDE password as you type it in and send it wherever they want. And not just that: since once they are in they are in, they can do anything they like for the rest of the system's lifecycle, with full privileges — including installing backdoors for versions of the OS or kernel that are installed on the device in the future, so that their backdoor remains open for as long as they like.

That is sad of course. It's particular sad given that the other popular OSes all address this much better. ChromeOS, Android, Windows and MacOS all have way better built-in protections against attacks like this. And it's why one can certainly claim that your data is probably better protected right now if you store it on those OSes then it is on generic Linux distributions.

(Yeah, I know that there are some niche distros which do this better, and some hackers hack their own. But I care about general purpose distros here, i.e. the big ones, that most people base their work on.)

Note that there are more problems with the current setup. For example, it's really weird that during boot the user is queried for an FDE password which actually protects their data, and then once the system is up they are queried again – now asking for a username, and another password. And the weird thing is that this second authentication that appears to be user-focused doesn't really protect the user's data anymore — at that moment the data is already unlocked and accessible. The username/password query is supposed to be useful in multi-user scenarios of course, but how does that make any sense, given that these multiple users would all have to know a disk encryption password that unlocks the whole thing during the FDE step, and thus they have access to every user's data anyway if they make an offline copy of the harddisk?

Can we do better?

Of course we can, and that is what this story is actually supposed to be about.

Let's first figure out what the minimal issues we should fix are (at least in my humble opinion):

  1. The initrd must be authenticated before being booted into. (And measured unconditionally.)

  2. The OS binary resources (i.e. /usr/) must be authenticated before being booted into. (But don't need to be encrypted, since everyone has the same anyway, there's nothing to hide here.)

  3. The OS configuration and state (i.e. /etc/ and /var/) must be encrypted, and authenticated before they are used. The encryption key should be bound to the TPM device; i.e system data should be locked to a security concept belonging to the system, not the user.

  4. The user's home directory (i.e. /home/lennart/ and similar) must be encrypted and authenticated. The unlocking key should be bound to a user password or user security token (FIDO2 or PKCS#11 token); i.e. user data should be locked to a security concept belonging to the user, not the system.

Or to summarize this differently:

  1. Every single component of the boot process and OS needs to be authenticated, i.e. all of shim (done), boot loader (done), kernel (done), initrd (missing so far), OS binary resources (missing so far), OS configuration and state (missing so far), the user's home (missing so far).

  2. Encryption is necessary for the OS configuration and state (bound to TPM), and for the user's home directory (bound to a user password or user security token).

In Detail

Let's see how we can achieve the above in more detail.

How to Authenticate the initrd

At the moment initrds are generated on the installed host via scripts (dracut and similar) that try to figure out a minimal set of binaries and configuration data to build an initrd that contains just enough to be able to find and set up the root file system. What is included in the initrd hence depends highly on the individual installation and its configuration. Pretty likely no two initrds generated that way will be fully identical due to this. This model clearly has benefits: the initrds generated this way are very small and minimal, and support exactly what is necessary for the system to boot, and not less or more. It comes with serious drawbacks too though: the generation process is fragile and sometimes more akin to black magic than following clear rules: the generator script natively has to understand a myriad of storage stacks to determine what needs to be included and what not. It also means that authenticating the image is hard: given that each individual host gets a different specialized initrd, it means we cannot just sign the initrd with the vendor key like we sign the kernel. If we want to keep this design we'd have to figure out some other mechanism (e.g. a per-host signature key – that is generated locally; or by authenticating it with a message authentication code bound to the TPM). While these approaches are certainly thinkable, I am not convinced they actually are a good idea though: locally and dynamically generated per-host initrds is something we probably should move away from.

If we move away from locally generated initrds, things become a lot simpler. If the distribution vendor generates the initrds on their build systems then it can be attached to the kernel image itself, and thus be signed and measured along with the kernel image, without any further work. This simplicity is simply lovely. Besides robustness and reproducibility this gives us an easy route to authenticated initrds.

But of course, nothing is really that simple: working with vendor-generated initrds means that we can't adjust them anymore to the specifics of the individual host: if we pre-build the initrds and include them in the kernel image in immutable fashion then it becomes harder to support complex, more exotic storage or to parameterize it with local network server information, credentials, passwords, and so on. Now, for my simple laptop use-case these things don't matter, there's no need to extend/parameterize things, laptops and their setups are not that wildly different. But what to do about the cases where we want both: extensibility to cover for less common storage subsystems (iscsi, LVM, multipath, drivers for exotic hardware…) and parameterization?

Here's a proposal how to achieve that: let's build a basic initrd into the kernel as suggested, but then do two things to make this scheme both extensible and parameterizable, without compromising security.

  1. Let's define a way how the basic initrd can be extended with additional files, which are stored in separate "extension images". The basic initrd should be able to discover these extension images, authenticate them and then activate them, thus extending the initrd with additional resources on-the-fly.

  2. Let's define a way how we can safely pass additional parameters to the kernel/initrd (and actually the rest of the OS, too) in an authenticated (and possibly encrypted) fashion. Parameters in this context can be anything specific to the local installation, i.e. server information, security credentials, certificates, SSH server keys, or even just the root password that shall be able to unlock the root account in the initrd …

In such a scheme we should be able to deliver everything we are looking for:

  1. We'll have a full trust chain for the code: the boot loader will authenticate and measure the kernel and basic initrd. The initrd extension images will then be authenticated by the basic initrd image.

  2. We'll have authentication for all the parameters passed to the initrd.

This so far sounds very unspecific? Let's make it more specific by looking closer at the components I'd suggest to be used for this logic:

  1. The systemd suite since a few months contains a subsystem implementing system extensions (v248). System extensions are ultimately just disk images (for example a squashfs file system in a GPT envelope) that can extend an underlying OS tree. Extending in this regard means they simply add additional files and directories into the OS tree, i.e. below /usr/. For a longer explanation see systemd-sysext(8). When a system extension is activated it is simply mounted and then merged into the main /usr/ tree via a read-only overlayfs mount. Now what's particularly nice about them in this context we are talking about here is that the extension images may carry dm-verity authentication data, and PKCS#7 signatures (once this is merged, that is, i.e. v250).

  2. The systemd suite also contains a concept called service "credentials". These are small pieces of information passed to services in a secure way. One key feature of these credentials is that they can be encrypted and authenticated in a very simple way with a key bound to the TPM (v250). See LoadCredentialEncrypted= and systemd-creds(1) for details. They are great for safely storing SSL private keys and similar on your system, but they also come handy for parameterizing initrds: an encrypted credential is just a file that can only be decoded if the right TPM is around with the right PCR values set.

  3. The systemd suite contains a component called systemd-stub(7). It's an EFI stub, i.e. a small piece of code that is attached to a kernel image, and turns the kernel image into a regular EFI binary that can be directly executed by the firmware (or a boot loader). This stub has a number of nice features (for example, it can show a boot splash before invoking the Linux kernel itself and such). Once this work is merged (v250) the stub will support one more feature: it will automatically search for system extension image files and credential files next to the kernel image file, measure them and pass them on to the main initrd of the host.

Putting this together we have nice way to provide fully authenticated kernel images, initrd images and initrd extension images; as well as encrypted and authenticated parameters via the credentials logic.

How would a distribution actually make us of this? A distribution vendor would pre-build the basic initrd, and glue it into the kernel image, and sign that as a whole. Then, for each supposed extension of the basic initrd (e.g. one for iscsi support, one for LVM, one for multipath, …), the vendor would use a tool such as mkosi to build an extension image, i.e. a GPT disk image containing the files in squashfs format, a Verity partition that authenticates it, plus a PKCS#7 signature partition that validates the root hash for the dm-verity partition, and that can be checked against a key provided by the boot loader or main initrd. Then, any parameters for the initrd will be encrypted using systemd-creds encrypt -T. The resulting encrypted credentials and the initrd extension images are then simply placed next to the kernel image in the ESP (or boot partition). Done.

This checks all boxes: everything is authenticated and measured, the credentials also encrypted. Things remain extensible and modular, can be pre-built by the vendor, and installation is as simple as dropping in one file for each extension and/or credential.

How to Authenticate the Binary OS Resources

Let's now have a look how to authenticate the Binary OS resources, i.e. the stuff you find in /usr/, i.e. the stuff traditionally shipped to the user's system via RPMs or DEBs.

I think there are three relevant ways how to authenticate this:

  1. Make /usr/ a dm-verity volume. dm-verity is a concept implemented in the Linux kernel that provides authenticity to read-only block devices: every read access is cryptographically verified against a top-level hash value. This top-level hash is typically a 256bit value that you can either encode in the kernel image you are using, or cryptographically sign (which is particularly nice once this is merged). I think this is actually the best approach since it makes the /usr/ tree entirely immutable in a very simple way. However, this also means that the whole of /usr/ needs to be updated as once, i.e. the traditional rpm/apt based update logic cannot work in this mode.

  2. Make /usr/ a dm-integrity volume. dm-integrity is a concept provided by the Linux kernel that offers integrity guarantees to writable block devices, i.e. in some ways it can be considered to be a bit like dm-verity while permitting write access. It can be used in three ways, one of which I think is particularly relevant here. The first way is with a simple hash function in "stand-alone" mode: this is not too interesting here, it just provides greater data safety for file systems that don't hash check their files' data on their own. The second way is in combination with dm-crypt, i.e. with disk encryption. In this case it adds authenticity to confidentiality: only if you know the right secret you can read and make changes to the data, and any attempt to make changes without knowing this secret key will be detected as IO error on next read by those in possession of the secret (more about this below). The third way is the one I think is most interesting here: in "stand-alone" mode, but with a keyed hash function (e.g. HMAC). What's this good for? This provides authenticity without encryption: if you make changes to the disk without knowing the secret this will be noticed on the next read attempt of the data and result in IO errors. This mode provides what we want (authenticity) and doesn't do what we don't need (encryption). Of course, the secret key for the HMAC must be provided somehow, I think ideally by the TPM.

  3. Make /usr/ a dm-crypt (LUKS) + dm-integrity volume. This provides both authenticity and encryption. The latter isn't typically needed for /usr/ given that it generally contains no secret data: anyone can download the binaries off the Internet anyway, and the sources too. By encrypting this you'll waste CPU cycles, but beyond that it doesn't hurt much. (Admittedly, some people might want to hide the precise set of packages they have installed, since it of course does reveal a bit of information about you: i.e. what you are working on, maybe what your job is – think: if you are a hacker you have hacking tools installed – and similar). Going this way might simplify things in some cases, as it means you don't have to distinguish "OS binary resources" (i.e /usr/) and "OS configuration and state" (i.e. /etc/ + /var/, see below), and just make it the same volume. Here too, the secret key must be provided somehow, I think ideally by the TPM.

All three approach are valid. The first approach has my primary sympathies, but for distributions not willing to abandon client-side updates via RPM/dpkg this is not an option, in which case I would propose the other two approaches for these cases.

The LUKS encryption key (and in case of dm-integrity standalone mode the key for the keyed hash function) should be bound to the TPM. Why the TPM for this? You could also use a user password, a FIDO2 or PKCS#11 security token — but I think TPM is the right choice: why that? To reduce the requirement for repeated authentication, i.e. that you first have to provide the disk encryption password, and then you have to login, providing another password. It should be possible that the system boots up unattended and then only one authentication prompt is needed to unlock the user's data properly. The TPM provides a way to do this in a reasonably safe and fully unattended way. Also, when we stop considering just the laptop use-case for a moment: on servers interactive disk encryption prompts don't make much sense — the fact that TPMs can provide secrets without this requiring user interaction and thus the ability to work in entirely unattended environments is quite desirable. Note that crypttab(5) as implemented by systemd (v248) provides native support for authentication via password, via TPM2, via PKCS#11 or via FIDO2, so the choice is ultimately all yours.

How to Encrypt/Authenticate OS Configuration and State

Let's now look at the OS configuration and state, i.e. the stuff in /etc/ and /var/. It probably makes sense to not consider these two hierarchies independently but instead just consider this to be the root file system. If the OS binary resources are in a separate file system it is then mounted onto the /usr/ sub-directory of the root file system.

The OS configuration and state (or: root file system) should be both encrypted and authenticated: it might contain secret keys, user passwords, privileged logs and similar. This data matters and contains plenty data that should remain confidential.

The encryption of choice here is dm-crypt (LUKS) + dm-integrity similar as discussed above, again with the key bound to the TPM.

If the OS binary resources are protected the same way it is safe to merge these two volumes and have a single partition for both (see above)

How to Encrypt/Authenticate the User's Home Directory

The data in the user's home directory should be encrypted, and bound to the user's preferred token of authentication (i.e. a password or FIDO2/PKCS#11 security token). As mentioned, in the traditional mode of operation the user's home directory is not individually encrypted, but only encrypted because FDE is in use. The encryption key for that is a system wide key though, not a per-user key. And I think that's problem, as mentioned (and probably not even generally understood by our users). We should correct that and ensure that the user's password is what unlocks the user's data.

In the systemd suite we provide a service systemd-homed(8) (v245) that implements this in a safe way: each user gets its own LUKS volume stored in a loopback file in /home/, and this is enough to synthesize a user account. The encryption password for this volume is the user's account password, thus it's really the password provided at login time that unlocks the user's data. systemd-homed also supports other mechanisms of authentication, in particular PKCS#11/FIDO2 security tokens. It also provides support for other storage back-ends (such as fscrypt), but I'd always suggest to use the LUKS back-end since it's the only one providing the comprehensive confidentiality guarantees one wants for a UNIX-style home directory.

Note that there's one special caveat here: if the user's home directory (e.g. /home/lennart/) is encrypted and authenticated, what about the file system this data is stored on, i.e. /home/ itself? If that dir is part of the the root file system this would result in double encryption: first the data is encrypted with the TPM root file system key, and then again with the per-user key. Such double encryption is a waste of resources, and unnecessary. I'd thus suggest to make /home/ its own dm-integrity volume with a HMAC, keyed by the TPM. This means the data stored directly in /home/ will be authenticated but not encrypted. That's good not only for performance, but also has practical benefits: it allows extracting the encrypted volume of the various users in case the TPM key is lost, as a way to recover from dead laptops or similar.

Why authenticate /home/, if it only contains per-user home directories that are authenticated on their own anyway? That's a valid question: it's because the kernel file system maintainers made clear that Linux file system code is not considered safe against rogue disk images, and is not tested for that; this means before you mount anything you need to establish trust in some way because otherwise there's a risk that the act of mounting might exploit your kernel.

Summary of Resources and their Protections

So, let's now put this all together. Here's a table showing the various resources we deal with, and how I think they should be protected (in my idealized world).

Resource Needs Authentication Needs Encryption Suggested Technology Validation/Encryption Keys/Certificates acquired via Stored where
Shim yes no SecureBoot signature verification firmware certificate database ESP
Boot loader yes no ditto firmware certificate database/shim ESP/boot partition
Kernel yes no ditto ditto ditto
initrd yes no ditto ditto ditto
initrd parameters yes yes systemd TPM encrypted credentials TPM ditto
initrd extensions yes no systemd-sysext with Verity+PKCS#7 signatures firmware/initrd certificate database ditto
OS binary resources yes no dm-verity root hash linked into kernel image, or firmware/initrd certificate database top-level partition
OS configuration and state yes yes dm-crypt (LUKS) + dm-integrity TPM top-level partition
/home/ itself yes no dm-integrity with HMAC TPM top-level partition
User home directories yes yes dm-crypt (LUKS) + dm-integrity in loopback files User password/FIDO2/PKCS#11 security token loopback file inside /home partition

This should provide all the desired guarantees: everything is authenticated, and the individualized per-host or per-user data is also encrypted. No double encryption takes place. The encryption keys/verification certificates are stored/bound to the most appropriate infrastructure.

Does this address the three attack scenarios mentioned earlier? I think so, yes. The basic attack scenario I described is addressed by the fact that /var/, /etc/ and /home/*/ are encrypted. Brute forcing the former two is harder than in the status quo ante model, since a high entropy key is used instead of one derived from a user provided password. Moreover, the "anti-hammering" logic of the TPM will make brute forcing prohibitively slow. The home directories are protected by the user's password or ideally a personal FIDO2/PKCS#11 security token in this model. Of course, a password isn't better security-wise then the status quo ante. But given the FIDO2/PKCS#11 support built into systemd-homed it should be easier to lock down the home directories securely.

Binding encryption of /var/ and /etc/ to the TPM also addresses the first of the two more advanced attack scenarios: a copy of the harddisk is useless without the physical TPM chip, since the seed key is sealed into that. (And even if the attacker had the chance to watch you type in your password, it won't help unless they possess access to to the TPM chip.) For the home directory this attack is not addressed as long as a plain password is used. However, since binding home directories to FIDO2/PKCS#11 tokens is built into systemd-homed things should be safe here too — provided the user actually possesses and uses such a device.

The backdoor attack scenario is addressed by the fact that every resource in play now is authenticated: it's hard to backdoor the OS if there's no component that isn't verified by signature keys or TPM secrets the attacker hopefully doesn't know.

For general purpose distributions that focus on updating the OS per RPM/dpkg the idealized model above won't work out, since (as mentioned) this implies an immutable /usr/, and thus requires updating /usr/ via an atomic update operation. For such distros a setup like the following is probably more realistic, but see above.

Resource Needs Authentication Needs Encryption Suggested Technology Validation/Encryption Keys/Certificates acquired via Stored where
Shim yes no SecureBoot signature verification firmware certificate database ESP
Boot loader yes no ditto firmware certificate database/shim ESP/boot partition
Kernel yes no ditto ditto ditto
initrd yes no ditto ditto ditto
initrd parameters yes yes systemd TPM encrypted credentials TPM ditto
initrd extensions yes no systemd-sysext with Verity+PKCS#7 signatures firmware/initrd certificate database ditto
OS binary resources, configuration and state yes yes dm-crypt (LUKS) + dm-integrity TPM top-level partition
/home/ itself yes no dm-integrity with HMAC TPM top-level partition
User home directories yes yes dm-crypt (LUKS) + dm-integrity in loopback files User password/FIDO2/PKCS#11 security token loopback file inside /home partition

This means there's only one root file system that contains all of /etc/, /var/ and /usr/.

Recovery Keys

When binding encryption to TPMs one problem that arises is what strategy to adopt if the TPM is lost, due to hardware failure: if I need the TPM to unlock my encrypted volume, what do I do if I need the data but lost the TPM?

The answer here is supporting recovery keys (this is similar to how other OSes approach this). Recovery keys are pretty much the same concept as passwords. The main difference being that they are computer generated rather than user-chosen. Because of that they typically have much higher entropy (which makes them more annoying to type in, i.e you want to use them only when you must, not day-to-day). By having higher entropy they are useful in combination with TPM, FIDO2 or PKCS#11 based unlocking: unlike a combination with passwords they do not compromise the higher strength of protection that TPM/FIDO2/PKCS#11 based unlocking is supposed to provide.

Current versions of systemd-cryptenroll(1) implement a recovery key concept in an attempt to address this problem. You may enroll any combination of TPM chips, PKCS#11 tokens, FIDO2 tokens, recovery keys and passwords on the same LUKS volume. When enrolling a recovery key it is generated and shown on screen both in text form and as QR code you can scan off screen if you like. The idea is write down/store this recovery key at a safe place so that you can use it when you need it. Note that such recovery keys can be entered wherever a LUKS password is requested, i.e. after generation they behave pretty much the same as a regular password.

TPM PCR Brittleness

Locking devices to TPMs and enforcing a PCR policy with this (i.e. configuring the TPM key to be unlockable only if certain PCRs match certain values, and thus requiring the OS to be in a certain state) brings a problem with it: TPM PCR brittleness. If the key you want to unlock with the TPM requires the OS to be in a specific state (i.e. that all OS components' hashes match certain expectations or similar) then doing OS updates might have the affect of making your key inaccessible: the OS updates will cause the code to change, and thus the hashes of the code, and thus certain PCRs. (Thankfully, you unrolled a recovery key, as described above, so this doesn't mean you lost your data, right?).

To address this I'd suggest three strategies:

  1. Most importantly: don't actually use the TPM PCRs that contain code hashes. There are actually multiple PCRs defined, each containing measurements of different aspects of the boot process. My recommendation is to bind keys to PCR 7 only, a PCR that contains measurements of the UEFI SecureBoot certificate databases. Thus, the keys will remain accessible as long as these databases remain the same, and updates to code will not affect it (updates to the certificate databases will, and they do happen too, though hopefully much less frequent then code updates). Does this reduce security? Not much, no, because the code that's run is after all not just measured but also validated via code signatures, and those signatures are validated with the aforementioned certificate databases. Thus binding an encrypted TPM key to PCR 7 should enforce a similar level of trust in the boot/OS code as binding it to a PCR with hashes of specific versions of that code. i.e. using PCR 7 means you say "every code signed by these vendors is allowed to unlock my key" while using a PCR that contains code hashes means "only this exact version of my code may access my key".

  2. Use LUKS key management to enroll multiple versions of the TPM keys in relevant volumes, to support multiple versions of the OS code (or multiple versions of the certificate database, as discussed above). Specifically: whenever an update is done that might result changing the relevant PCRs, pre-calculate the new PCRs, and enroll them in an additional LUKS slot on the relevant volumes. This means that the unlocking keys tied to the TPM remain accessible in both states of the system. Eventually, once rebooted after the update, remove the old slots.

  3. If these two strategies didn't work out (maybe because the OS/firmware was updated outside of OS control, or the update mechanism was aborted at the wrong time) and the TPM PCRs changed unexpectedly, and the user now needs to use their recovery key to get access to the OS back, let's handle this gracefully and automatically reenroll the current TPM PCRs at boot, after the recovery key checked out, so that for future boots everything is in order again.

Other approaches can work too: for example, some OSes simply remove TPM PCR policy protection of disk encryption keys altogether immediately before OS or firmware updates, and then reenable it right after. Of course, this opens a time window where the key bound to the TPM is much less protected than people might assume. I'd try to avoid such a scheme if possible.

Anything Else?

So, given that we are talking about idealized systems: I personally actually think the ideal OS would be much simpler, and thus more secure than this:

I'd try to ditch the Shim, and instead focus on enrolling the distribution vendor keys directly in the UEFI firmware certificate list. This is actually supported by all firmwares too. This has various benefits: it's no longer necessary to bind everything to Microsoft's root key, you can just enroll your own stuff and thus make sure only what you want to trust is trusted and nothing else. To make an approach like this easier, we have been working on doing automatic enrollment of these keys from the systemd-boot boot loader, see this work in progress for details. This way the Firmware will authenticate the boot loader/kernel/initrd without any further component for this in place.

I'd also not bother with a separate boot partition, and just use the ESP for everything. The ESP is required anyway by the firmware, and is good enough for storing the few files we need.

FAQ

Can I implement all of this in my distribution today?

Probably not. While the big issues have mostly been addressed there's a lot of integration work still missing. As you might have seen I linked some PRs that haven't even been merged into our tree yet, and definitely not been released yet or even entered the distributions.

Will this show up in Fedora/Debian/Ubuntu soon?

I don't know. I am making a proposal how these things might work, and am working on getting various building blocks for this into shape. What the distributions do is up to them. But even if they don't follow the recommendations I make 100%, or don't want to use the building blocks I propose I think it's important they start thinking about this, and yes, I think they should be thinking about defaulting to setups like this.

Work for measuring/signing initrds on Fedora has been started, here's a slide deck with some information about it.

But isn't a TPM evil?

Some corners of the community tried (unfortunately successfully to some degree) to paint TPMs/Trusted Computing/SecureBoot as generally evil technologies that stop us from using our systems the way we want. That idea is rubbish though, I think. We should focus on what it can deliver for us (and that's a lot I think, see above), and appreciate the fact we can actually use it to kick out perceived evil empires from our devices instead of being subjected to them. Yes, the way SecureBoot/TPMs are defined puts you in the driver seat if you want — and you may enroll your own certificates to keep out everything you don't like.

What if my system doesn't have a TPM?

TPMs are becoming quite ubiquitous, in particular as the upcoming Windows versions will require them. In general I think we should focus on modern, fully equipped systems when designing all this, and then find fall-backs for more limited systems. Frankly it feels as if so far the design approach for all this was the other way round: try to make the new stuff work like the old rather than the old like the new (I mean, to me it appears this thinking is the main raison d'être for the Grub boot loader).

More specifically, on the systems where we have no TPM we ultimately cannot provide the same security guarantees as for those which have. So depending on the resource to protect we should fall back to different TPM-less mechanisms. For example, if we have no TPM then the root file system should probably be encrypted with a user provided password, typed in at boot as before. And for the encrypted boot credentials we probably should simply not encrypt them, and place them in the ESP unencrypted.

Effectively this means: without TPM you'll still get protection regarding the basic attack scenario, as before, but not the other two.

What if my system doesn't have UEFI?

Many of the mechanisms explained above taken individually do not require UEFI. But of course the chain of trust suggested above requires something like UEFI SecureBoot. If your system lacks UEFI it's probably best to find work-alikes to the technologies suggested above, but I doubt I'll be able to help you there.

rpm/dpkg already cryptographically validates all packages at installation time (gpg), why would I need more than that?

This type of package validation happens once: at the moment of installation (or update) of the package, but not anymore when the data installed is actually used. Thus when an attacker manages to modify the package data after installation and before use they can make any change they like without this ever being noticed. Such package download validation does address certain attack scenarios (i.e. man-in-the-middle attacks on network downloads), but it doesn't protect you from attackers with physical access, as described in the attack scenarios above.

Systems such as ostree aren't better than rpm/dpkg regarding this BTW, their data is not validated on use either, but only during download or when processing tree checkouts.

Key really here is that the scheme explained here provides offline protection for the data "at rest" — even someone with physical access to your device cannot easily make changes that aren't noticed on next use. rpm/dpkg/ostree provide online protection only: as long as the system remains up, and all OS changes are done through the intended program code-paths, and no one has physical access everything should be good. In today's world I am sure this is not good enough though. As mentioned most modern OSes provide offline protection for the data at rest in one way or another. Generic Linux distributions are terribly behind on this.

This is all so desktop/laptop focused, what about servers?

I am pretty sure servers should provide similar security guarantees as outlined above. In a way servers are a much simpler case: there are no users and no interactivity. Thus the discussion of /home/ and what it contains and of user passwords doesn't matter. However, the authenticated initrd and the unattended TPM-based encryption I think are very important for servers too, in a trusted data center environment. It provides security guarantees so far not given by Linux server OSes.

I'd like to help with this, or discuss/comment on this

Submit patches or reviews through GitHub. General discussion about this is best done on the systemd mailing list.

Creating Quality Backtraces for Crash Reports

Posted by Michael Catanzaro on September 19, 2021 02:41 AM

Hello Linux users! Help developers help you: include a quality backtraces taken with gdb each and every time you create an issue report for a crash. If you don’t, most developers will request that you provide a backtrace, then ignore your issue until you manage to figure out how to do so. Save us the trouble and just provide the backtrace with your initial report, so everything goes smoother. (Backtraces are often called “stack traces.” They are the same thing.)

Don’t just copy the lower-quality backtrace you see in your system journal into your issue report. That’s a lot better than nothing, but if you really want the crash to be fixed, you should provide the developers with a higher-quality backtrace from gdb. Don’t know how to get a quality backtrace with gdb? Read on.

Future of Crash Reporting

Here are instructions for getting a quality backtrace for a crashing process on Fedora 35, which is scheduled to be released in October:

$ coredumpctl gdb
(gdb) bt full

Press ‘c’ (continue) when required. When it’s done printing, press ‘q’ to quit. That’s it! That’s all you need to know. You’re done. Two points of note:

  • When a process crashes, a core dump is caught by systemd-coredump and stored for future use. The coredumpctl gdb command opens the most recent core dump in gdb. systemd-coredump has been enabled by default in Fedora since Fedora 26, which was released four years ago. (It’s also enabled by default in RHEL 8.)
  • After opening the core dump, gdb uses debuginfod to automatically download all required debuginfo packages, ensuring the generated backtrace is useful. debuginfod is a new feature in Fedora 35.

If you’re not an inhabitant of the future, you are probably missing at least debuginfod today, and possibly also systemd-coredump depending on which operating system you are using, so we will also learn how to take a backtrace without these tools. It will be more complicated, of course.

systemd-coredump

If your operating system enables systemd-coredump by default, then congratulations! This makes reporting crashes much easier because you can easily retrieve a core dump for any recent crash using the coredumpctl command. For example, coredumpctl alone will list all available core dumps. coredumpctl gdb will open the core dump of the most recent crash in gdb. coredumpctl gdb 1234 will open the core dump corresponding to the most recent crash of a process with pid 1234. It doesn’t get easier than this.

Core dumps are stored under /var/lib/systemd/coredump. systemd-coredump will automatically delete core dumps that exceed configurable size limits (2 GB by default). It also deletes core dumps if your free disk space falls below a configurable threshold (15% free by default). Additionally, systemd-tmpfiles will delete core dumps automatically after some time has passed (three days by default). This ensures your disk doesn’t fill up with old core dumps. Although most of these settings seem good to me, the default 2 GB size limit is way too low in my opinion, as it causes systemd to immediately discard crashes of any application that uses WebKit. I recommend raising this limit to 20 GB by creating an /etc/systemd/systemd.conf.d/50-coredump.conf drop-in containing the following:

[Coredump]
ProcessSizeMax=20G
ExternalSizeMax=20G

The other settings are likely sufficient to prevent your disk from filling up with core dumps.

Sadly, although systemd-coredump has been around for a good while now and many Linux operating systems have it enabled by default, many still do not. Most notably, the Debian ecosystem is still not yet on board. To check if systemd-coredump is enabled on your system:

$ cat /proc/sys/kernel/core_pattern

If you see systemd-coredump, then you’re good.

To enable it in Debian or Ubuntu, just install it:

# apt install systemd-coredump

Ubuntu users, note this will cause apport to be uninstalled, since it is currently incompatible. Also note that I switched from $ (which indicates a normal prompt) to # (which indicates a root prompt).

In other operating systems, you may have to manually enable it:

# echo "kernel.core_pattern=|/usr/lib/systemd/systemd-coredump %P %u %g %s %t %c %h" > /etc/sysctl.d/50-coredump.conf
# /usr/lib/systemd/systemd-sysctl --prefix kernel.core_pattern

Note the exact core pattern to use changes occasionally in newer versions of systemd, so these instructions may not work everywhere.

Detour: Manual Core Dump Handling

If you don’t want to enable systemd-coredump, life is harder and you should probably reconsider, but it’s still possible to debug most crashes. First, enable core dump creation by removing the default 0-byte size limit on core files:

$ ulimit -c unlimited

This change is temporary and only affects the current instance of your shell. For example, if you open a new tab in your terminal, you will need to set the ulimit again in the new tab.

Next, run your program in the terminal and try to make it crash. A core file will be generated in the current directory. Open it by starting the program that crashed in gdb and passing the filename of the core file that was created. For example:

$ gdb gnome-chess ./core

This is downright primitive, though:

  • You’re going to have a hard time getting backtraces for services that are crashing, for starters. If starting the service normally, how do you set the ulimit? I’m sure there’s a way to do it, but I don’t know how! It’s probably easier to start the service manually, but then what command line flags are needed to properly do so? It will be different for each service, and you have to figure this all out for yourself.
  • Special situations become very difficult. For example, if a service is crashing only when run early during boot, or only during an initial setup session, you are going to have an especially hard time.
  • If you don’t know how to reproduce a crash that occurs only rarely, it’s inevitably going to crash when you’re not prepared to manually catch the core dump. Sadly, not all crashes will occur on demand when you happen to be running the software from a terminal with the right ulimit configured.
  • Lastly, you have to remember to delete that core file when you’re done, because otherwise it will take up space on your disk space until you do. You’ll probably notice if you leave core files scattered in ~/home, but you might not notice if you’re working someplace else.

Seriously, just enable systemd-coredump. It solves all of these problems and guarantees you will always have easy access to a core dump when something crashes, even for crashes that occur only rarely.

Debuginfo Installation

Now that we know how to open a core dump in gdb, let’s talk about debuginfo. When you don’t have the right debuginfo packages installed, the backtrace generated by gdb will be low-quality. Almost all Linux software developers deal with low-quality backtraces on a regular basis, because most users are not very good at installing debuginfo. Again, if you’re reading this in the future using Fedora 35 or newer, you don’t have to worry about this anymore because debuginfod will take care of everything for you. I would be thrilled if other Linux operating systems would quickly adopt debuginfod so we can put the era of low-quality crash reports behind us. But since most readers of this blog today will not yet have debuginfod enabled, let’s learn how to install debuginfo manually.

As an example, I decided to force gnome-chess to crash using the command killall -SEGV gnome-chess, then I ran coredumpctl gdb to open the resulting core dump in gdb. After a bunch of spam, I saw this:

Missing separate debuginfos, use: dnf debuginfo-install gnome-chess-40.1-1.fc34.x86_64
--Type <RET> for more, q to quit, c to continue without paging--
Core was generated by `/usr/bin/gnome-chess --gapplication-service'.
Program terminated with signal SIGSEGV, Segmentation fault.
#0  0x00007fa23d8b55bf in __GI___poll (fds=0x5636deb06930, nfds=2, timeout=2830)
    at ../sysdeps/unix/sysv/linux/poll.c:29
29  return SYSCALL_CANCEL (poll, fds, nfds, timeout);
[Current thread is 1 (Thread 0x7fa23ca0cd00 (LWP 140177))]
(gdb)

If you are using Fedora, RHEL, or related operating systems, the line “missing separate debuginfos” is a good hint that debuginfo is missing. It even tells you exactly which dnf debuginfo-install command to run to remedy this problem! But this is a Fedora ecosystem feature, and you won’t see this on most other operating systems. Usually, you’ll need to manually locate the right debuginfo packages to install. Debian and Ubuntu users can do this by searching for and installing -dbg or -dbgsym packages until each frame in the backtrace looks good. You’ll just have to manually guess the names of which debuginfo packages you need to install based on the names of the libraries in the backtrace. Look here for instructions for popular operating systems.

How do you know when the backtrace looks good? When each frame has file names, line numbers, function parameters, and local variables! Here is an example of a bad backtrace, if I continue the gnome-chess example above without properly installing the required debuginfo:

(gdb) bt full
#0 0x00007fa23d8b55bf in __GI___poll (fds=0x5636deb06930, nfds=2, timeout=2830)
    at ../sysdeps/unix/sysv/linux/poll.c:29
        sc_ret = -516
        sc_cancel_oldtype = 0
#1 0x00007fa23eee648c in g_main_context_iterate.constprop () at /lib64/libglib-2.0.so.0
#2 0x00007fa23ee8fc03 in g_main_context_iteration () at /lib64/libglib-2.0.so.0
#3 0x00007fa23e4b599d in g_application_run () at /lib64/libgio-2.0.so.0
#4 0x00005636dd7b79a2 in chess_application_main ()
#5 0x00007fa23d7e7b75 in __libc_start_main (main=0x5636dd7aaa50 <main>, argc=2, argv=0x7fff827b6438, init=<optimized out>, fini=<optimized out>, rtld_fini=<optimized out>, stack_end=0x7fff827b6428)
    at ../csu/libc-start.c:332
        self = <optimized out>
        result = <optimized out>
        unwind_buf = 
              {cancel_jmp_buf = {{jmp_buf = {94793644186304, 829313697107602221, 94793644026480, 0, 0, 0, -829413713854928083, -808912263273321683}, mask_was_saved = 0}}, priv = {pad = {0x0, 0x0, 0x2, 0x7fff827b6438}, data = {prev = 0x0, cleanup = 0x0, canceltype = 2}}}
        not_first_call = <optimized out>
#6 0x00005636dd7aaa9e in _start ()

This backtrace has six frames, which shows where the code was during program execution when the crash occurred. You can see line numbers for frame #0 (poll.c:29) and #5 (libc-start.c:332), and these frames also show the values of function parameters and variables on the stack, which are often useful for figuring out what went wrong. These frames have good debuginfo because I already had debuginfo installed for glibc. But frames #1 through #4 do not look so useful, showing only function names and the library and nothing else. This is because I’m using Fedora 34 rather than Fedora 35, so I don’t have debuginfod yet, and I did not install proper debuginfo for libgio, libglib, and gnome-chess. (The function names are actually only there because programs in Fedora include some limited debuginfo by default. In many operating systems, you will see ??? instead of function names.) A developer looking at this backtrace is not going to know what went wrong.

Now, let’s run the recommended debuginfo-install command:

# dnf debuginfo-install gnome-chess-40.1-1.fc34.x86_64

When the command finishes, we’ll start gdb again, using coredumpctl gdb just like before. This time, we see this:

Missing separate debuginfos, use: dnf debuginfo-install avahi-libs-0.8-14.fc34.x86_64 colord-libs-1.4.5-2.fc34.x86_64 cups-libs-2.3.3op2-7.fc34.x86_64 fontconfig-2.13.94-2.fc34.x86_64 glib2-2.68.4-1.fc34.x86_64 graphene-1.10.6-2.fc34.x86_64 gstreamer1-1.19.1-2.1.18.4.fc34.x86_64 gstreamer1-plugins-bad-free-1.19.1-3.1.18.4.fc34.x86_64 gstreamer1-plugins-base-1.19.1-2.1.18.4.fc34.x86_64 gtk4-4.2.1-1.fc34.x86_64 json-glib-1.6.6-1.fc34.x86_64 krb5-libs-1.19.2-2.fc34.x86_64 libX11-1.7.2-3.fc34.x86_64 libX11-xcb-1.7.2-3.fc34.x86_64 libXfixes-6.0.0-1.fc34.x86_64 libdrm-2.4.107-1.fc34.x86_64 libedit-3.1-38.20210714cvs.fc34.x86_64 libepoxy-1.5.9-1.fc34.x86_64 libgcc-11.2.1-1.fc34.x86_64 libidn2-2.3.2-1.fc34.x86_64 librsvg2-2.50.7-1.fc34.x86_64 libstdc++-11.2.1-1.fc34.x86_64 libxcrypt-4.4.25-1.fc34.x86_64 llvm-libs-12.0.1-1.fc34.x86_64 mesa-dri-drivers-21.1.8-1.fc34.x86_64 mesa-libEGL-21.1.8-1.fc34.x86_64 mesa-libgbm-21.1.8-1.fc34.x86_64 mesa-libglapi-21.1.8-1.fc34.x86_64 nettle-3.7.3-1.fc34.x86_64 openldap-2.4.57-5.fc34.x86_64 openssl-libs-1.1.1l-1.fc34.x86_64 pango-1.48.9-2.fc34.x86_64

Yup, Fedora ecosystem users will need to run dnf debuginfo-install twice to install everything required, because gdb doesn’t list all required packages until the second time. Next, we’ll run coredumpctl gdb one last time. There will usually be a few more debuginfo packages that are still missing because they’re not available in the Fedora repositories, but now you’ll probably have enough to get a quality backtrace:

(gdb) bt full
#0  0x00007fa23d8b55bf in __GI___poll (fds=0x5636deb06930, nfds=2, timeout=2830)
    at ../sysdeps/unix/sysv/linux/poll.c:29
        sc_ret = -516
        sc_cancel_oldtype = 0
#1  0x00007fa23eee648c in g_main_context_poll
    (priority=, n_fds=2, fds=0x5636deb06930, timeout=, context=0x5636de7b24a0)
    at ../glib/gmain.c:4434
        ret = 
        errsv = 
        poll_func = 0x7fa23ee97c90 
        max_priority = 2147483647
        timeout = 2830
        some_ready = 
        nfds = 2
        allocated_nfds = 2
        fds = 0x5636deb06930
        begin_time_nsec = 30619110638882
#2  g_main_context_iterate.constprop.0
    (context=context@entry=0x5636de7b24a0, block=block@entry=1, dispatch=dispatch@entry=1, self=)
    at ../glib/gmain.c:4126
        max_priority = 2147483647
        timeout = 2830
        some_ready = 
        nfds = 2
        allocated_nfds = 2
        fds = 0x5636deb06930
        begin_time_nsec = 30619110638882
#3  0x00007fa23ee8fc03 in g_main_context_iteration
    (context=context@entry=0x5636de7b24a0, may_block=may_block@entry=1) at ../glib/gmain.c:4196
        retval = 
#4  0x00007fa23e4b599d in g_application_run
    (application=0x5636de7ae260 [ChessApplication], argc=-2105843004, argv=)
    at ../gio/gapplication.c:2560
        arguments = 0x5636de7b2400
        status = 0
        context = 0x5636de7b24a0
        acquired_context = 
        __func__ = "g_application_run"
#5  0x00005636dd7b79a2 in chess_application_main (args=0x7fff827b6438, args_length1=2)
    at src/gnome-chess.p/gnome-chess.c:5623
        _tmp0_ = 0x5636de7ae260 [ChessApplication]
        _tmp1_ = 0x5636de7ae260 [ChessApplication]
        _tmp2_ = 
        result = 0
...

I removed the last two frames because they are triggering a strange WordPress bug, but that’s enough to get the point. It looks much better! Now the developer can see exactly where the program crashed, including filenames, line numbers, and the values of function parameters and variables on the stack. This is as good as a crash report is normally going to get. In this case, it crashed when running poll() because gnome-chess was not actually doing anything at the time of the crash, since we crashed it by manually sending a SIGSEGV signal. Normally the backtrace will look more interesting.

Special Note for Arch Linux Users

Arch does not ship debuginfo packages, btw. To prepare a quality backtrace, you need to rebuild each package yourself with debuginfo enabled. This is a chore, to say the least. You will probably need to rebuild several system libraries in addition to the application itself if you want to get a quality backtrace. This is hardly impossible, but it’s a lot of work, too much to expect users to do for a typical bug report. You might want to consider attempting to reproduce your crash on another operating system in order to make it easier to get the backtrace. What a shame!

debuginfod for Fedora 34

Again, all of that manual debuginfo installation is no longer required as of Fedora 35, where debuginfod is enabled by default. It’s actually all ready to go in Fedora 34, just not enabled by default yet. You can try it early using the DEBUGINFOD_URLS environment variable:

$ DEBUGINFOD_URLS=https://debuginfod.fedoraproject.org/ coredumpctl gdb

Then you can watch gdb download the required debuginfo for you! Again, this environment variable will no longer be necessary in Fedora 35. (Technically, it will still be needed, but it will be configured by default.)

debuginfod for Debian Users

Debian users can use debuginfod, but it has to be enabled manually:

$ DEBUGINFOD_URLS=https://debuginfod.debian.net/ gdb

See here for more information. This requires Debian 11 “bullseye” or newer. If you’re using Ubuntu or other operating systems derived from Debian, you’ll need to wait until a debuginfod server for your operating system is available.

Flatpak

If your application uses Flatpak, you can use the flatpak-coredumpctl script to open core dumps in gdb. For most runtimes, including those distributed by GNOME or Flathub, you will need to manually install (a) the debug extension for your app, (b) the SDK runtime corresponding to the platform runtime that you are using, and (c) the debug extension for the SDK runtime. For example, to install everything required to debug Epiphany 40 from Flathub, you would run:

$ flatpak install org.gnome.Epiphany.Debug//stable
$ flatpak install org.gnome.Sdk//40
$ flatpak install org.gnome.Sdk.Debug//40

(flatpak-coredumpctl will fail to start if you don’t have the correct SDK runtime installed, but it will not fail if you’re missing the debug extensions. You’ll just wind up with a bad backtrace.)

The debug extensions need to exactly match the versions of the app and runtime that crashed, so backtrace generation may be unreliable after you install them for the very first time, because you would have installed the latest versions of the extensions, but your core dump might correspond to an older app or runtime version. If the crash is reproducible, it’s a good idea to run flatpak update after installing to ensure you have the latest version of everything, then reproduce the crash again.

Once your debuginfo is installed, you can open the backtrace in gdb using flatpak-coredumpctl. You just have to tell flatpak-coredumpctl the app ID to use:

$ flatpak-coredumpctl org.gnome.Epiphany

You can pass matches to coredumpctl using -m. For example, to open the core dump corresponding to a crashed process with pid 1234:

$ flatpak-coredumpctl -m 1234 org.gnome.Epiphany

Thibault Saunier wrote flatpak-coredumpctl because I complained about how hard it used to be to debug crashed Flatpak applications. Clearly it is no longer hard. Thanks Thibault!

Update: On newer versions of Debian and Ubuntu, flatpak-coredumpctl is included in the libflatpak-dev subpackage rather than the base flatpak package, so you’ll have to install libflatpak-dev first. But on older OS versions, including Debian 10 “buster” and Ubuntu 20.04, it is unfortunately installed as /usr/share/doc/flatpak/examples/flatpak-coredumpctl rather than /usr/bin/flatpak-coredumpctl due to a regrettable packaging choice that has been corrected in newer package versions. As a workaround, you can simply copy it to /usr/local/bin. Don’t forget to delete your copy after upgrading to a newer OS version, or it will shadow the packaged version.

Fedora Flatpaks

Flatpaks distributed by Fedora are different than those distributed by GNOME or by Flathub because they do not have debug extensions. Historically, this has meant that debugging crashes was impractical. The best solution was to give up.

Good news! Fedora’s Flatpaks are compatible with debuginfod, which means debug extensions will no longer be missed. You do still need to manually install the org.fedoraproject.Sdk runtime corresponding to the version of the org.fedoraproject.Platform runtime that the application uses, because this is required for flatpak-coredumpctl to work, but nothing else is required. For example, to get a backtrace for Fedora’s Epiphany Flatpak using a Fedora 35 host system, I ran:

$ flatpak install org.fedoraproject.Sdk//f34
$ flatpak-coredumpctl org.gnome.Epiphany
(gdb) bt full

(The f34 is not a typo. Epiphany currently uses the Fedora 34 runtime regardless of what host system you are using.)

That’s it!

Miscellany

At this point, you should know enough to obtain a high-quality backtrace on most Linux systems. That will usually be all you really need, but it never hurts to know a little more, right?

Alternative Types of Backtraces

At the top of this blog post, I suggested using bt full to take the backtrace because this type of backtrace is the most useful to most developers. But there are other types of backtraces you might occasionally want to collect:

  • bt on its own without full prints a much shorter backtrace without stack variables or function parameters. This form of the backtrace is more useful for getting a quick feel for where the bug is occurring, because it is much shorter and easier to read than a full backtrace. But because there are no stack variables or function parameters, it might not contain enough information to solve the crash. I sometimes like to paste the first few lines of a bt backtrace directly into an issue report, then submit the bt full version of the backtrace as an attachment, since an entire bt full backtrace can be long and inconvenient if pasted directly into an issue report.
  • thread apply all bt prints a backtrace for every thread. Normally these backtraces are very long and noisy, so I don’t collect them very often, but when a threadsafety issue is suspected, this form of backtrace will sometimes be required.
  • thread apply all bt full prints a full backtrace for every thread. This is what automated bug report tools generally collect, because it provides the most information. But these backtraces are usually huge, and this level of detail is rarely needed, so I normally recommend starting with a normal bt full.

If in doubt, just use bt full like I showed at the top of this blog post. Developers will let you know if they want you to provide the backtrace in a different form.

gdb Logging

You can make gdb print your session to a file. For longer backtraces, this may be easier than copying the backtrace from a terminal:

(gdb) set logging on

Memory Corruption

While a backtrace taken with gdb is usually enough information for developers to debug crashes, memory corruption is an exception. Memory corruption is the absolute worst. When memory corruption occurs, the code will crash in a location that may be far removed from where the actual bug occurred, rendering gdb backtraces useless for tracking down the bug. As a general rule, if you see a crash inside a memory allocation routine like malloc() or g_slice_alloc(), you probably have memory corruption. If you see magazine_chain_pop_head(), that’s called by g_slice_alloc() and is a sure sign of memory corruption. Similarly, crashes in GTK’s CSS machinery are almost always caused by memory corruption somewhere else.

Memory corruption is generally impossible to debug unless you are able to reproduce the issue under valgrind. valgrind is extremely slow, so it’s impractical to use it on a regular basis, but it will get to the root of the problem where gdb cannot. As a general rule, you want to run valgrind with --track-origins=yes so that it shows you exactly what went wrong:

$ valgrind --track-origins=yes my_app

When valgrinding a GLib application, including all GNOME applications, always use the G_SLICE=always-malloc environment variable to disable GLib’s slice allocator, to ensure the highest-quality diagnostics. Correction: the slice allocator can now detect valgrind and disable itself automatically.

When valgrinding a WebKit application, there are some WebKit-specific environment variables to use. Malloc=1 will disable the bmalloc allocator, GIGACAGE_ENABLED=0 will disable JavaScriptCore’s Gigacage feature (Update: turns out this is actually implied by Malloc=1), and WEBKIT_FORCE_SANDBOX=0 will disable the web process sandbox used by WebKitGTK or WPE WebKit:

$ Malloc=1 WEBKIT_FORCE_SANDBOX=0 valgrind --track-origins=yes epiphany

If you cannot reproduce the issue under valgrind, you’re usually totally out of luck. Memory corruption that only occurs rarely or under unknown conditions will lurk in your code indefinitely and cause occasional crashes that are effectively impossible to fix.

Another good tool for debugging memory corruption is address sanitizer (asan), but this is more complicated to use. Experienced users who are comfortable with rebuilding applications using special compiler flags may find asan very useful. However, because it can be very difficult to use,  I recommend sticking with valgrind if you’re just trying to report a bug.

Apport and ABRT

There are two popular downstream bug reporting tools: Ubuntu has Apport, and Fedora has ABRT. These tools are relatively easy to use — no command line knowledge required — and produce quality crash reports. Unfortunately, while the tools are promising, the crash reports go to downstream packagers who are generally either not watching bug reports, or else not interested or capable of fixing upstream software problems. Since downstream reports are very often ignored, it’s better to report crashes directly to upstream if you want your issue to be seen by the right developers and actually fixed. Of course, only report issues upstream if you’re using a recent software version. Fedora and Arch users can pretty much always safely report directly to upstream, as can Ubuntu users who are using the very latest version of Ubuntu. If you are an Ubuntu LTS user, you should stick with reporting issues to downstream only, or at least take the time to verify that the issue still occurs with a more recent software version.

There are a couple more problems with these tools. As previously mentioned, Ubuntu’s apport is incompatible with systemd-coredump. If you’ve read this far, you know you really want systemd-coredump enabled, so I recommend disabling apport until it learns to play ball with systemd-coredump.

The technical design of Fedora’s ABRT is currently better because it actually retrieves your core dumps from systemd-coredump, so you don’t have to choose between one or the other. Unfortunately, ABRT has many serious user experience bugs and warts. I can’t recommend it for this reason, but it if it works well enough for you, it does create some great downstream crash reports. Whether a downstream package maintainer will look at those reports is hit or miss, though.

What is a crash, really?

Most developers consider crashes on Unix systems to be program termination via a Unix signal that triggers creation of a core dump. The most common of these are SIGSEGV (segmentation fault, “invalid memory reference”) or SIBABRT (usually an intentional crash due to an assertion failure). Less-common signals are SIGBUS (“bad memory access”) or SIGILL (“illegal instruction”). Sandboxed applications might occasionally see SIGSYS (“bad system call”). See the manpage signal(7) for a full list. These are cases where you can get a backtrace to help with tracking down the issues.

What is not a crash? If your application is hanging or just not behaving properly, that is not a crash. If your application is killed using SIGTERM or SIGKILL — this can happen when systemd-oomd determines you are low on memory,  or when a service is taking too long to stop — this is also not a crash in the usual sense of the word, because you’re not going to be able to get a backtrace for it. If a website is slow or unavailable, the news might say that it “crashed,” but it’s obviously not the same thing as what we’re talking about here. The techniques in this blog post are no use for these sorts of “crashes.”

Conclusion

If you have systemd-coredump enabled and debuginfod installed and working, most crash reports will be simple.  Memory corruption is a frustrating exception. Encourage your operating system to enable systemd-coredump and debuginfod if it doesn’t already.  Happy crash reporting!

Flatpak portals - how do they work?

Posted by Peter Hutterer on September 01, 2021 03:23 AM

I've been working on portals recently and one of the issues for me was that the documentation just didn't quite hit the sweet spot. At least the bits I found were either too high-level or too implementation-specific. So here's a set of notes on how a portal works, in the hope that this is actually correct.

First, Portals are supposed to be a way for sandboxed applications (flatpaks) to trigger functionality they don't have direct access too. The prime example: opening a file without the application having access to $HOME. This is done by the applications talking to portals instead of doing the functionality themselves.

There is really only one portal process: /usr/libexec/xdg-desktop-portal, started as a systemd user service. That process owns a DBus bus name (org.freedesktop.portal.Desktop) and an object on that name (/org/freedesktop/portal/desktop). You can see that bus name and object with D-Feet, from DBus' POV there's nothing special about it. What makes it the portal is simply that the application running inside the sandbox can talk to that DBus name and thus call the various methods. Obviously the xdg-desktop-portal needs to run outside the sandbox to do its things.

There are multiple portal interfaces, all available on that one object. Those interfaces have names like org.freedesktop.portal.FileChooser (to open/save files). The xdg-desktop-portal implements those interfaces and thus handles any method calls on those interfaces. So where an application is sandboxed, it doesn't implement the functionality itself, it instead calls e.g. the OpenFile() method on the org.freedesktop.portal.FileChooser interface. Then it gets an fd back and can read the content of that file without needing full access to the file system.

Some interfaces are fully handled within xdg-desktop-portal. For example, the Camera portal checks a few things internally, pops up a dialog for the user to confirm access to if needed [1] but otherwise there's nothing else involved with this specific method call.

Other interfaces have a backend "implementation" DBus interface. For example, the org.freedesktop.portal.FileChooser interface has a org.freedesktop.impl.portal.FileChooser (notice the "impl") counterpart. xdg-desktop-portal does not implement those impl.portals. xdg-desktop-portal instead routes the DBus calls to the respective "impl.portal". Your sandboxed application calls OpenFile(), xdg-desktop-portal now calls OpenFile() on org.freedesktop.impl.portal.FileChooser. That interface returns a value, xdg-desktop-portal extracts it and returns it back to the application in respones to the original OpenFile() call.

What provides those impl.portals doesn't matter to xdg-desktop-portal, and this is where things are hot-swappable. GTK and Qt both provide (some of) those impl portals, There are GTK and Qt-specific portals with xdg-desktop-portal-gtk and xdg-desktop-portal-kde but another one is provided by GNOME Shell directly. You can check the files in /usr/share/xdg-desktop-portal/portals/ and see which impl portal is provided on which bus name. The reason those impl.portals exist is so they can be native to the desktop environment - regardless what application you're running and with a generic xdg-desktop-portal, you see the native file chooser dialog for your desktop environment.

So the full call sequence is:

  • At startup, xdg-desktop-portal parses the /usr/libexec/xdg-desktop-portal/*.portal files to know which impl.portal interface is provided on which bus name
  • The application calls OpenFile() on the org.freedesktop.portal.FileChooser interface on the object path /org/freedesktop/portal/desktop. It can do so because the bus name this object sits on is not restricted by the sandbox
  • xdg-desktop-portal receives that call. This is portal with an impl.portal so xdg-desktop-portal calls OpenFile() on the bus name that provides the org.freedesktop.impl.portal.FileChooser interface (as previously established by reading the *.portal files)
  • Assuming xdg-desktop-portal-gtk provides that portal at the moment, that process now pops up a GTK FileChooser dialog that runs outside the sandbox. User selects a file
  • xdg-desktop-portal-gtk sends back the fd for the file to the xdg-desktop-portal, and the impl.portal parts are done
  • xdg-desktop-portal receives that fd and sends it back as reply to the OpenFile() method in the normal portal
  • The application receives the fd and can read the file now
A few details here aren't fully correct, but it's correct enough to understand the sequence - the exact details depend on the method call anyway.

Finally: because of DBus restrictions, the various methods in the portal interfaces don't just reply with values. Instead, the xdg-desktop-portal creates a new org.freedesktop.portal.Request object and returns the object path for that. Once that's done the method is complete from DBus' POV. When the actual return value arrives (e.g. the fd), that value is passed via a signal on that Request object, which is then destroyed. This roundabout way is done for purely technical reasons, regular DBus methods would time out while the user picks a file path.

Anyway. Maybe this helps someone understanding how the portal bits fit together.

[1] it does so using another portal but let's ignore that
[2] not really hot-swappable though. You need to restart xdg-desktop-portal but not your host. So luke-warm-swappable only

Edit Sep 01: clarify that it's not GTK/Qt providing the portals, but xdg-desktop-portal-gtk and -kde

libinput and high-resolution wheel scrolling

Posted by Peter Hutterer on August 31, 2021 07:50 AM

Gut Ding braucht Weile. Almost three years ago, we added high-resolution wheel scrolling to the kernel (v5.0). The desktop stack however was first lagging and eventually left behind (except for an update a year ago or so, see here). However, I'm happy to announce that thanks to José Expósito's efforts, we now pushed it across the line. So - in a socially distanced manner and masked up to your eyebrows - gather round children, for it is storytime.

Historical History

In the beginning, there was the wheel detent. Or rather there were 24 of them, dividing a 360 degree [1] movement of a wheel into a neat set of 15 clicks. libinput exposed those wheel clicks as part of the "pointer axis" namespace and you could get the click count with libinput_event_pointer_get_axis_discrete() (announced here). The degree value is exposed as libinput_event_pointer_get_axis_value(). Other scroll backends (finger-scrolling or button-based scrolling) expose the pixel-precise value via that same function.

In a "recent" Microsoft Windows version (Vista!), MS added the ability for wheels to trigger more than 24 clicks per rotation. The MS Windows API now treats one "traditional" wheel click as a value of 120, anything finer-grained will be a fraction thereof. You may have a mouse that triggers quarter-wheel clicks, each sending a value of 30. This makes for smoother scrolling and is supported(-ish) by a lot of mice introduced in the last 10 years [2]. Obviously, three small scrolls are nicer than one large scroll, so the UX is less bad than before.

Now it's time for libinput to catch up with Windows Vista! For $reasons, the existing pointer axis API could get changed to accommodate for the high-res values, so a new API was added for scroll events. Read on for the details, you will believe what happens next.

Out with the old, in with the new

As of libinput 1.19, libinput has three new events: LIBINPUT_EVENT_POINTER_SCROLL_WHEEL, LIBINPUT_EVENT_POINTER_SCROLL_FINGER, and LIBINPUT_EVENT_POINTER_SCROLL_CONTINUOUS. These events reflect, perhaps unsuprisingly, scroll movements of a wheel, a finger or along a continuous axis (e.g. button scrolling). And they replace the old event LIBINPUT_EVENT_POINTER_AXIS. Those familiar with libinput will notice that the new event names encode the scroll source in the event name now. This makes them slightly more flexible and saves callers an extra call.

In terms of actual API, the new events have two new functions: libinput_event_pointer_get_scroll_value(). For the FINGER and CONTINUOUS events, the value returned is in "pixels" [3]. For the new WHEEL events, the value is in degrees. IOW this is a drop-in replacement for the old libinput_event_pointer_get_axis_value() function. The second call is libinput_event_pointer_get_scroll_value_v120() which, for WHEEL events, returns the 120-based logical units the kernel uses as well. libinput_event_pointer_has_axis() returns true if the given axis has a value, just as before. With those three calls you now get the data for the new events.

Backwards compatibility

To ensure backwards compatibility, libinput generates both old and new events so the rule for callers is: if you want to support the new events, just ignore the old ones completely. libinput also guarantees new events even on pre-5.0 kernels. This makes the old and new code easy to ifdef out, and once you get past the immediate event handling the code paths are virtually identical.

When, oh when?

These changes have been merged into the libinput main branch and will be part of libinput 1.19. Which is due to be released over the next month or so, so feel free to work backwards from that for your favourite distribution.

Having said that, libinput is merely the lowest block in the Jenga tower that is the desktop stack. José linked to the various MRs in the upstream libinput MR, so if you're on your seat's edge waiting for e.g. GTK to get this, well, there's an MR for that.

[1] That's degrees of an angle, not Fahrenheit
[2] As usual, on a significant number of those you'll need to know whatever proprietary protocol the vendor deemed to be important IP. Older MS mice stand out here because they use straight HID.
[3] libinput doesn't really have a concept of pixels, but it has a normalized pixel that movements are defined as. Most callers take that as real pixels except for the high-resolution displays where it's appropriately scaled.

Pango updates

Posted by Matthias Clasen on August 26, 2021 05:15 PM

I’ve spent some time on Pango, recently. Here is a little update on the feature work that I’ve done there. All of these changes will appear in Pango 1.50 and GTK 4.6.

The general directions of this work are:

  • Take advantage of the fact that we are now using harfbuzz on all platforms. Among other things, this gives us much easier access to font information.
  • Match CSS where it makes sense. If nothing else, this makes it much easier to connect new Pango features to the CSS machinery in GTK.

CSS features

Lets start with the second point: matching CSS.

Line spacing has historically been a bit painful in GtkTextView. You can set distances before and after paragraphs, and between wrapped lines inside a paragraph. But this does not take font sizes into account—it is a fixed number of pixels.

A while ago, I added a line-spacing factor to Pango, which was meant to help with the font size dependency. You basically tell Pango: I want the baselines of this paragraph spaced apart 1.33 times as wide as they would normally be. The remaining problem is that Pango handles text one paragraph at a time. So as far as it is concerned, there is not previous baseline above the first line in a paragraph, and it does not increase the spacing between paragraphs.

The CSS solution to this problem is to just make the lines themselves taller, and place them flush next to each other.  With this approach, you need to be a little careful to make sure that you still get consistent baseline-to-baseline distances. But at least it solves the paragraph spacing issue.

Pango recently gained a line-height attribute that does just that, and GTK now supports the corresponding CSS property.

Another feature that has come to Pango from the CSS side is support for text transformation (also in the screenshot). This lets you change the capitalization of text. Note that this is just for presentation purposes—if you select STRASSE in the example and copy it to the clipboard,  you get the original straße.

And again, GTK supports the corresponding CSS property.

Font features

As I said, harfbuzz makes it much easier for us to access features of the fonts we use. One of these that I have recently looked into is caret metrics. High-quality italic fonts can contain information about the slope at which the text caret is best drawn to match the text:

I’ve added a new api to get this information, and made GTK use it, with this result:

<video class="wp-video-shortcode" controls="controls" height="267" id="video-9287-1" preload="metadata" width="474"><source src="https://blogs.gnome.org/mclasen/files/2021/08/Screencast-from-08-24-2021-035115-PM.webm?_=1" type="video/webm">https://blogs.gnome.org/mclasen/files/2021/08/Screencast-from-08-24-2021-035115-PM.webm</video>

Another useful bit of font information concerns placement or carets inside ligatures.

 Historically, Pango has just divided the width of the glyph evenly among the characters that are forming the ligature (w and i, in this example), but high-quality fonts can provide this information. This is most relevant for scripts using many ligatures, such as Arabic.I’ve made Pango use the ligature caret data if it is available from the font, and got this result:

<video class="wp-video-shortcode" controls="controls" height="267" id="video-9287-2" preload="metadata" width="474"><source src="https://blogs.gnome.org/mclasen/files/2021/08/Screencast-from-08-25-2021-105318-PM.webm?_=2" type="video/webm">https://blogs.gnome.org/mclasen/files/2021/08/Screencast-from-08-25-2021-105318-PM.webm</video>

The wi ligature in this test is what I could come up with after a struggling for a few hours with fontforge (clearly, font design is not in my future). The only other fonts I’ve found with ligature caret information are Arabic, and I sadly can’t read or write that script.

The last feature closed a 15 year old bug – not something you get to do every day!

libei - a status update

Posted by Peter Hutterer on August 25, 2021 05:29 AM

A year ago, I first announced libei - a library to support emulated input. After an initial spurt of development, it was left mostly untouched until a few weeks ago. Since then, another flurry of changes have been added, including some initial integration into GNOME's mutter. So, let's see what has changed.

A Recap

First, a short recap of what libei is: it's a transport layer for emulated input events to allow for any application to control the pointer, type, etc. But, unlike the XTEST extension in X, libei allows the compositor to be in control over clients, the devices they can emulate and the input events as well. So it's safer than XTEST but also a lot more flexible. libei already supports touch and smooth scrolling events, something XTest doesn't have or is struggling with.

Terminology refresher: libei is the client library (used by an application wanting to emulate input), EIS is the Emulated Input Server, i.e. the part that typically runs in the compositor.

Server-side Devices

So what has changed recently: first, the whole approach has flipped on its head - now a libei client connects to the EIS implementation and "binds" to the seats the EIS implementation provides. The EIS implementation then provides input devices to the client. In the simplest case, that's just a relative pointer but we have capabilities for absolute pointers, keyboards and touch as well. Plans for the future is to add gestures and tablet support too. Possibly joysticks, but I haven't really thought about that in detail yet.

So basically, the initial conversation with an EIS implementation goes like this:

  • Client: Hello, I am $NAME
  • Server: Hello, I have "seat0" and "seat1"
  • Client: Bind to "seat0" for pointer, keyboard and touch
  • Server: Here is a pointer device
  • Server: Here is a keyboard device
  • Client: Send relative motion event 10/2 through the pointer device
Notice how the touch device is missing? The capabilities the client binds to are just what the client wants, the server doesn't need to actually give the client a device for that capability.

One of the design choices for libei is that devices are effectively static. If something changes on the EIS side, the device is removed and a new device is created with the new data. This applies for example to regions and keymaps (see below), so libei clients need to be able to re-create their internal states whenever the screen or the keymap changes.

Device Regions

Devices can now have regions attached to them, also provided by the EIS implementation. These regions define areas reachable by the device and are required for clients such as Barrier. On a dual-monitor setup you may have one device with two regions or two devices with one region (representing one monitor), it depends on the EIS implementation. But either way, as libei client you will know that there is an area and you will know how to reach any given pixel on that area. Since the EIS implementation decides the regions, it's possible to have areas that are unreachable by emulated input (though I'm struggling a bit for a real-world use-case).

So basically, the conversation with an EIS implementation goes like this:

  • Client: Hello, I am $NAME
  • Server: Hello, I have "seat0" and "seat1"
  • Client: Bind to "seat0" for absolute pointer
  • Server: Here is an abs pointer device with regions 1920x1080@0,0, 1080x1920@1920,0
  • Server: Here is an abs pointer device with regions 1920x1080@0,0
  • Server: Here is an abs pointer device with regions 1080x1920@1920,0
  • Client: Send abs position 100/100 through the second device
Notice how we have three absolute devices? A client emulating a tablet that is mapped to a screen could just use the third device. As with everything, the server decides what devices are created and the clients have to figure out what they want to do and how to do it.

Perhaps unsurprisingly, the use of regions make libei clients windowing-system independent. The Barrier EI support WIP no longer has any Wayland-specific code in it. In theory, we could implement EIS in the X server and libei clients would work against that unmodified.

Keymap handling

The keymap handling has been changed so the keymap too is provided by the EIS implementation now, effectively in the same way as the Wayland compositor provides the keymap to Wayland clients. This means a client knows what keycodes to send, it can handle the state to keep track of things, etc. Using Barrier as an example again - if you want to generate an "a", you need to look up the keymap to figure out which keycode generates an A, then you can send that through libei to actually press the key.

Admittedly, this is quite messy. XKB (and specifically libxkbcommon) does not make it easy to go from a keysym to a key code. The existing Barrier X code is full of corner-cases with XKB already, I espect those to be necessary for the EI support as well.

Scrolling

Scroll events have four types: pixel-based scrolling, discrete scrolling, and scroll stop/cancel events. The first should be obvious, discrete scrolling is for mouse wheels. It uses the same 120-based API that Windows (and the kernel) use, so it's compatible with high-resolution wheel mice. The scroll stop event notifies an EIS implementation that the scroll interaction has stopped (e.g. lifting fingers off) which in turn may start kinetic scrolling - just like the libinput/Wayland scroll stop events. The scroll cancel event notifies the EIS implementation that scrolling really has stopped and no kinetic scrolling should be triggered. There's no equivalent in libinput/Wayland for this yet but it helps to get the hook in place.

Emulation "Transactions"

This has fairly little functional effect, but interactions with an EIS server are now sandwiched in a start/stop emulating pair. While this doesn't matter for one-shot tools like xdotool, it does matter for things like Barrier which can send the start emulating event when the pointer enters the local window. This again allows the EIS implementation to provide some visual feedback to the user. To correct the example from above, the sequence is actually:

  • ...
  • Server: Here is a pointer device
  • Client: Start emulating
  • Client: Send relative motion event 10/2 through the pointer device
  • Client: Send relative motion event 1/4 through the pointer device
  • Client: Stop emulating

Properties

Finally, there is now a generic property API, something copied from PipeWire. Properties are simple key/value string pairs and cover those things that aren't in the immediate API. One example here: the portal can set things like "ei.application.appid" to the Flatpak's appid. Properties can be locked down and only libei itself can set properties before the initial connection. This makes them reliable enough for the EIS implementation to make decisions based on their values. Just like with PipeWire, the list of useful properties will grow over time. it's too early to tell what is really needed.

Repositories

Now, for the actual demo bits: I've added enough support to Barrier, XWayland, Mutter and GNOME Shell that I can control a GNOME on Wayland session through Barrier (note: the controlling host still needs to run X since we don't have the ability to capture input events under Wayland yet). The keymap handling in Barrier is nasty but it's enough to show that it can work.

GNOME Shell has a rudimentary UI, again just to show what works:

The status icon shows ... if libei clients are connected, it changes to !!! while the clients are emulating events. Clients are listed by name and can be disconnected at will. I am not a designer, this is just a PoC to test the hooks.

Note how xdotool is listed in this screenshot: that tool is unmodified, it's the XWayland libei implementation that allows it to work and show up correctly

The various repositories are in the "wip/ei" branch of:

And of course libei itself.

Where to go from here? The last weeks were driven by rapid development, so there's plenty of test cases to be written to make sure the new code actually works as intended. That's easy enough. Looking at the Flatpak integration is another big ticket item, once the portal details are sorted all the pieces are (at least theoretically) in place. That aside, improving the integrations into the various systems above is obviously what's needed to get this working OOTB on the various distributions. Right now it's all very much in alpha stage and I could use help with all of those (unless you're happy to wait another year or so...). Do ping me if you're interested to work on any of this.

HighContrast variants for Adwaita-qt

Posted by Jan Grulich on August 24, 2021 10:23 AM

In the past we used to have a completely different project to cover HighContrast variants of GTK Adwaita theme. This was all implemented as Highcontrast-qt, a project nobody has touched for 6 years. You can imagine how it looks like these days when you compare it to what we have now. I think even GTK variant of HighContrast was a completely separate theme back then, while today days it’s just Adwaita with a different set of colors.

Since GTK made the new HighContrast theme with just few modifications to the original Adwaita theme, I decided to use same approach and have Adwaita-qt to provide all four variants as well (Adwaita, Adwaita-dark, HighContrast and HighContrastInverse). While this looks like a simple thing to do, as you just need to add additional color palette, it was a pain to do it in Adwaita-qt. The reason is that Adwaita-qt is full of hardcoded color definitions, where all of them were randomly taken from GTK Adwaita stylesheets. Everytime something changed in GTK Adwaita, we would have to manually pick the change and replace the changed color value on our side. This was not really sustainable, especially when I wanted to have four different variants.

To improve this situation and make my life easier, I decided to bundle GTK stylesheets for Adwaita theme, have it processed and write a simple parser to make everything automated. And I did exactly that. I included the stylesheets and wrote some definitions myself so I can let them processed with sassc (GTK uses SASS for theming) and have all my definitions substituted for simple parsing. I no longer have to pick all base colors manually, all of them are being parsed for all four variants and same goes for basic styling of Buttons, CheckBoxes and Radio Buttons, where each of them have all kind of possible states (active, hovered, checked etc.) and use not only a simple color, but also gradients. You can imagine how hard it was to hardcode all values for each state. The parser I wrote is really basic and simple with use of regular expressions as the code I’m trying to process is not that complex.

The code I try to process is either a simple definition:

@define-color base_color #ffffff;

Or widget definition in this form:

button:checked { color: #2e3436; border-color: #cdc7c2; background-image: image(#dad6d2); box-shadow: none; }

The result is:

<figure class="wp-block-image size-large"><figcaption>Adwaita-qt: Adwaita variant</figcaption></figure> <figure class="wp-block-image size-large"><figcaption>Adwaita-qt: HighContrast variant</figcaption></figure>

I think this is a big step forward for Adwaita-qt and it will allow me to more quicky respond on changes happening in GTK Adwaita/HighContrast theme. I can also imagine this being extended in the future to support some additional variants, like modified Adwaita theme you can find in Ubuntu (at least if the stylesheet is similar enough). As mentioned, Adwaita-qt now supports four variants and they should be on par with GTK 4.4, at least when it comes to colors and style for most used widgets, because there are still places in Adwaita-qt which need some extra work. Anyway, this all is now released as Adwaita-qt 1.4.0 and I will be updating Flathub and Fedora builds to it soon.

GTK4: Start Center

Posted by Caolán McNamara on August 11, 2021 02:54 PM

 

LibreOffice Start Center working with GTK4 widgets

power-profiles-daemon: Follow-up

Posted by Bastien Nocera on August 05, 2021 02:50 PM

Just about a year after the original announcement, I think it's time to see the progress on power-profiles-daemon.

Note that I would still recommend you read the up-to-date project README if you have questions about why this project was necessary, and why a new project was started rather than building on an existing one.

 The project was born out of the need to make a firmware feature available to end-users for a number of lines of Lenovo laptops for them to be fully usable on Fedora. For that, I worked with Mark Pearson from Lenovo, who wrote the initial kernel support for the feature and served as our link to the Lenovo firmware team, and Hans de Goede, who worked on making the kernel interfaces more generic.

More generic, but in a good way

 With the initial kernel support written for (select) Lenovo laptops, Hans implemented a more generic interface called platform_profile. This interface is now the one that power-profiles-daemon will integrate with, and means that it also supports a number of Microsoft Surface, HP, Lenovo's own Ideapad laptops, and maybe Razer laptops soon.

 The next item to make more generic is Lenovo's "lap detection" which still relies on a custom driver interface. This should be soon transformed into a generic proximity sensor, which will mean I get to work some more on iio-sensor-proxy.

Working those interactions

 power-profiles-dameon landed in a number of distributions, sometimes enabled by default, sometimes not enabled by default (sigh, the less said about that the better), which fortunately meant that we had some early feedback available.

 The goal was always to have the user in control, but we still needed to think carefully about how the UI would look and how users would interact with it when a profile was temporarily unavailable, or the system started a "power saver" mode because battery was running out.

 The latter is something that David Redondo's work on the "HoldProfile" API made possible. Software can programmatically switch to the power-saver or performance profile for the duration of a command. This is useful to switch to the Performance profile when running a compilation (eg. powerprofilesctl jhbuild --no-interact build gnome-shell), or for gnome-settings-daemon to set the power-saver profile when low on battery.

 The aforementioned David Redondo and Kai Uwe Broulik also worked on the KDE interface to power-profiles-daemon, as Florian Müllner implemented the gnome-shell equivalent.

Promised by me, delivered by somebody else :)

 I took this opportunity to update the Power panel in Settings, which shows off the temporary switch to the performance mode, and the setting to automatically switch to power-saver when low on battery.

Low-Power, everywhere

 Talking of which, while it's important for the system to know that they're targetting a power saving behaviour, it's also pretty useful for applications to try and behave better.
 
 Maybe you've already integrated with "low memory" events using GLib, but thanks to Patrick Griffis you can be an event better ecosystem citizen and monitor whether the system is in "Power Saver" mode and adjust your application's behaviour.
 
 This feature will be available in GLib 2.70 along with documentation of useful steps to take. GNOME Software will already be using this functionality to avoid large automated downloads when energy saving is needed.

Availability

 The majority of the above features are available in the GNOME 41 development branches and should get to your favourite GNOME-friendly distribution for their next release, such as Fedora 35.

It's templates all the way down - part 4

Posted by Peter Hutterer on July 28, 2021 03:22 AM

Part 1, Part 2, Part 3

After getting thouroughly nerd-sniped a few weeks back, we now have FreeBSD support through qemu in the freedesktop.org ci-templates. This is possible through the qemu image generation we have had for quite a while now. So let's see how we can easily add a FreeBSD VM (or other distributions) to our gitlab CI pipeline:


.freebsd:
variables:
FDO_DISTRIBUTION_VERSION: '13.0'
FDO_DISTRIBUTION_TAG: 'freebsd.0' # some value for humans to read

build-image:
extends:
- .freebsd
- .fdo.qemu-build@freebsd
variables:
FDO_DISTRIBUTION_PACKAGES: "curl wget"
Now, so far this may all seem quite familiar. And indeed, this is almost exactly the same process as for normal containers (see Part 1), the only difference is the .fdo.qemu-build base template. Using this template means we build an image babushka: our desired BSD image is actual a QEMU RAW image sitting inside another generic container image. That latter image only exists to start the QEMU image and set up the environment if need be, you don't need to care what distribution it runs out (Fedora for now).

Because of the nesting, we need to handle this accordingly in our script: tag for the actual test job - we need to start the image and make sure our jobs are actually built within. The templates set up an ssh alias "vm" for this and the vmctl script helps to do things on the vm:


test-build:
extends:
- .freebsd
- .fdo.distribution-image@freebsd
script:
# start our QEMU image
- /app/vmctl start

# copy our current working directory to the VM
# (this is a yaml multiline command to work around the colon)
- |
scp -r $PWD vm:

# Run the build commands on the VM and if they succeed, create a .success file
- /app/vmctl exec "cd $CI_PROJECT_NAME; meson builddir; ninja -C builddir" && touch .success || true

# Copy results back to our run container so we can include them in artifacts:
- |
scp -r vm:$CI_PROJECT_NAME/builddir .

# kill the VM
- /app/vmctl stop

# Now that we have cleaned up: if our build job before
# failed, exit with an error
- [[ -e .success ]] || exit 1
Now, there's a bit to unpack but with the comments above it should be fairly obvious what is happening. We start the VM, copy our working directory over and then run a command on the VM before cleaning up. The reason we use touch .success is simple: it allows us to copy things out and clean up before actually failing the job.

Obviously, if you want to build any other distribution you just swap the freebsd out for fedora or whatever - the process is the same. libinput has been using fedora qemu images for ages now.

libinput and hold gestures

Posted by Peter Hutterer on July 27, 2021 05:58 AM

Thanks to the work done by <https:>Josè Expòsito, libinput 1.19 will ship with a new type of gesture: Hold Gestures. So far libinput supported swipe (moving multiple fingers in the same direction) and pinch (moving fingers towards each other or away from each other). These gestures are well-known, commonly used, and familiar to most users. For example, GNOME 40 recently has increased its use of touchpad gestures to switch between workspaces, etc. Swipe and pinch gestures require movement, it was not possible (for callers) to detect fingers on the touchpad that don't move.

This gap is now filled by Hold gestures. These are triggered when a user puts fingers down on the touchpad, without moving the fingers. This allows for some new interactions and we had two specific ones in mind: hold-to-click, a common interaction on older touchscreen interfaces where holding a finger in place eventually triggers the context menu. On a touchpad, a three-finger hold could zoom in, or do dictionary lookups, or kill a kitten. Whatever matches your user interface most, I guess.

The second interaction was the ability to stop kinetic scrolling. libinput does not actually provide kinetic scrolling, it merely provides the information needed in the client to do it there: specifically, it tells the caller when a finger was lifted off a touchpad at the end of a scroll movement. It's up to the caller (usually: the toolkit) to implement the kinetic scrolling effects. One missing piece was that while libinput provided information about lifting the fingers, it didn't provide information about putting fingers down again later - a common way to stop scrolling on other systems.

Hold gestures are intended to address this: a hold gesture triggered after a flick with two fingers can now be used by callers (read: toolkits) to stop scrolling.

Now, one important thing about hold gestures is that they will generate a lot of false positives, so be careful how you implement them. The vast majority of interactions with the touchpad will trigger some movement - once that movement hits a certain threshold the hold gesture will be cancelled and libinput sends out the movement events. Those events may be tiny (depending on touchpad sensitivity) so getting the balance right for the aforementioned hold-to-click gesture is up to the caller.

As usual, the required bits to get hold gestures into the wayland protocol are either in the works, mid-flight or merge-ready so expect this to hit the various repositories over the medium-term future.

threads and libxcb: problems now we have two

Posted by Adam Jackson on July 22, 2021 08:59 PM

If you want to write an X application, you need to use some library that speaks the X11 protocol. For a long time this meant libX11, often called xlib, which - like most things about X - is a fantastic bit of engineering that is very much a product of its time with some confusing baroque bits. Overall it does a very nice job of hiding the icky details of the protocol from the application developer.

One of the details it hides has to do with how resource IDs are allocated in X. A resource ID (an XID, in the jargon) is a 32 29-bit integer that names a resource - window, colormap, what have you. Those 29 bits are split up netmask/hostmask style, where the top 8 or so uniquely identify the client, and the rest identify the resource belonging to that client. When you create a window in X, what you really tell the server is "I want a window that's initially this size, this background color (etc.) and from now on when I say (my client id + 17) I mean that window." This is great for performance because it means resource allocation is assumed to succeed and you don't have to wait for a reply from the server.

Key to all this is that in xlib the XID is the return value from the call that issues the resource creation request. Internally the request gets queued into the protocol's write buffer, but the client can march ahead and issue the next few commands as if creation had succeeded - because it probably did, and if it didn't you're probably going to crash anyway.

So to allocate XIDs the client just marches forward through its XID range. What happens when you hit the end of the range? Before X11R4, you'd crash, because xlib doesn't keep track of which XIDs it's allocated, just the lowest one it hasn't allocated yet. Starting in R4 the server added an extension called XC-MISC that lets the client ask the server for a list of unused XIDs, so when xlib hits the end of the range it can request a new range from the server.

But. UI programming tends to want threads, and xlib is perhaps not the most thread-friendly. So XCB was invented, which sacrifices some of xlib's ease of use for a more direct binding to the protocol and (in theory) an explicitly thread-safe design. We then modified xlib and XCB to coexist in the same process, using the same I/O buffers, reply and event management, etc.

This literal reflection of the protocol into the API has consequences. In XCB, unlike xlib, XID generation is an explicit step. The client first calls into XCB to allocate the XID, and then passes that XID to the creation request in order to give the resource a name.

Which... sorta ruins that whole thread-safety thing.

Let's say you call xcb_generate_id in thread A and the XID it returns is the last one in your range. Then thread B schedules in and tries to allocate another XID. You'll ask the server for a new range, but since thread A hasn't called its resource creation request yet, from the server's perspective that "allocated" XID looks like it's still free! So now, whichever thread issues their resource creation request second will get BadIDChoice thrown at them if the other thread's resource hasn't been destroyed in the interim.

A library that was supposed to be about thread safety baked a thread safety hazard into the API. Good work, team.

How do you fix this without changing the API? Maybe you could keep a bitmap on the client side that tracks XID allocation, that's only like 256KB worst case, you can grow it dynamically and most clients don't create more than a few dozen resources anyway. Make xcb_generate_id consult that bitmap for the first unallocated ID, and mark it used when it returns. Then track every resource destruction request and zero it back out of the bitmap. You'd only need XC-MISC if some other client destroyed one of your resources and you were completely out of XIDs otherwise.

And you can implement this, except. One, XCB has zero idea what a resource destruction request is, that's simply not in the protocol description. Not a big deal, you can fix that, there's only like forty destructors you'd need to annotate. But then two, that would only catch resource destruction calls that flow through XCB's protocol binding API, which xlib does not, xlib instead pushes raw data through xcb_writev. So now you need to modify every client library (libXext, libGL, ...) to inform XCB about resource destruction.

Which is doable. Tedious. But doable.

I think.

I feel a little weird writing about this because: surely I can't be the first person to notice this.

Does free software benefit from ML models being derived works of training data?

Posted by Matthew Garrett on July 13, 2021 01:57 AM
Github recently announced Copilot, a machine learning system that makes suggestions for you when you're writing code. It's apparently trained on all public code hosted on Github, which means there's a lot of free software in its training set. Github assert that the output of Copilot belongs to the user, although they admit that it may occasionally produce output that is identical to content from the training set.

Unsurprisingly, this has led to a number of questions along the lines of "If Copilot embeds code that is identical to GPLed training data, is my code now GPLed?". This is extremely understandable, but the underlying issue is actually more general than that. Even code under permissive licenses like BSD requires retention of copyright notices and disclaimers, and failing to include them is just as much a copyright violation as incorporating GPLed code into a work and not abiding by the terms of the GPL is.

But free software licenses only have power to the extent that copyright permits them to. If your code isn't a derived work of GPLed material, you have no obligation to follow the terms of the GPL. Github clearly believe that Copilot's output doesn't count as a derived work as far as US copyright law goes, and as a result the licenses on the training data don't apply to the output. Some people have interpreted this as an attack on free software - Copilot may insert code that's either identical or extremely similar to GPLed code, and claim that there are no license obligations created as a result, effectively allowing the laundering of GPLed code into proprietary software.

I'm completely unqualified to hold a strong opinion on whether Github's legal position is justifiable or not, and right now I'm also not interested in thinking about it too much. What I think is more interesting is what the impact of either position has on free software. Do we benefit more from a future where the output of Copilot (or similar projects) is considered a derived work of the training data, or one where it isn't? Having been involved in a bunch of GPL enforcement activities, it's very easy to think of this as something that weakens the GPL and, as a result, weakens free software. That was my initial reaction, but that's shifted over the past few days.

Let's look at the GNU manifesto, specifically this section:

The fact that the easiest way to copy a program is from one neighbor to another, the fact that a program has both source code and object code which are distinct, and the fact that a program is used rather than read and enjoyed, combine to create a situation in which a person who enforces a copyright is harming society as a whole both materially and spiritually; in which a person should not do so regardless of whether the law enables him to.

The GPL makes use of copyright law to ensure that GPLed work can't be taken from the commons. Anyone who produces a derived work of GPLed code is obliged to provide that work under the same terms. If software weren't copyrightable, the GPL would have no power. But this is the outcome Stallman wanted! The GPL doesn't exist because copyright is good, it exists because software being copyrightable is what enables the concept of proprietary software in the first place.

The powers that the GPL uses to enforce sharing of code are used by the authors of proprietary software to reduce that sharing. They attempt to forbid us from examining their code to determine how it works - they argue that anyone who does so is tainted, unable to contribute similar code to free software projects in case they produce a derived work of the original. Broadly speaking, the further the definition of a derived work reaches, the greater the power of proprietary software authors. If Oracle's argument that APIs are copyrightable had prevailed, it would have been disastrous for free software. If the Apple look and feel suit had established that Microsoft infringed Apple's copyright, we might be living in a future where we had no free software desktop environments.

When we argue for an interpretation of copyright law that enhances the power of the GPL, we're also enhancing the power of giant corporations with a lot of lawyers on hand. So let's look at this another way. If Github's interpretation of copyright law holds, we can train a model on proprietary code and extract concepts without having to worry about being tainted. The proprietary code itself won't enter the commons, but the ideas it embodies will. No more worries about whether you're literally copying the code that implements an algorithm you want to duplicate - simply start typing and let the model remove the risk for you.

There's a reasonable counter argument about equality here. How much GPL-influenced code is going to end up in proprietary projects when compared to the reverse? It's not an easy question to answer, but we should bear in mind that the majority of public repositories on Github aren't under an open source license. Copilot is already claiming to give us access to the concepts embodied in those repositories. Do these provide more value than is given up? I honestly don't know how to measure that. But what I do know is that free software was founded in a belief that software shouldn't be constrained by copyright, and our default stance shouldn't be to argue against the idea that copyright is weaker than we imagined.

(Edit: this post by Julia Reda makes some of the same arguments, but spends some more time focusing on a legal analysis of why having copyright cover the output of Copilot would be a problem)

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GTK4: Adapting to Toolbar changes

Posted by Caolán McNamara on July 06, 2021 10:53 AM

 

In GTK4 the dedicated GtkToolbar is gone, replaced by GtkBox with a style of "toolbar". The spelling dialog now supports this conversion.

GTK4 Port: DoubleDecker Notebooks

Posted by Caolán McNamara on June 17, 2021 04:15 PM

 

Double-decker mode, for notebooks with excessive numbers of tabs, works again in the GTK4 Port.

The Wondrous World of Discoverable GPT Disk Images

Posted by Lennart Poettering on June 10, 2021 10:00 PM

TL;DR: Tag your GPT partitions with the right, descriptive partition types, and the world will become a better place.

A number of years ago we started the Discoverable Partitions Specification which defines GPT partition type UUIDs and partition flags for the various partitions Linux systems typically deal with. Before the specification all Linux partitions usually just used the same type, basically saying "Hey, I am a Linux partition" and not much else. With this specification the GPT partition type, flags and label system becomes a lot more expressive, as it can tell you:

  1. What kind of data a partition contains (i.e. is this swap data, a file system or Verity data?)
  2. What the purpose/mount point of a partition is (i.e. is this a /home/ partition or a root file system?)
  3. What CPU architecture a partition is intended for (i.e. is this a root partition for x86-64 or for aarch64?)
  4. Shall this partition be mounted automatically? (i.e. without specifically be configured via /etc/fstab)
  5. And if so, shall it be mounted read-only?
  6. And if so, shall the file system be grown to its enclosing partition size, if smaller?
  7. Which partition contains the newer version of the same data (i.e. multiple root file systems, with different versions)

By embedding all of this information inside the GPT partition table disk images become self-descriptive: without requiring any other source of information (such as /etc/fstab) if you look at a compliant GPT disk image it is clear how an image is put together and how it should be used and mounted. This self-descriptiveness in particular breaks one philosophical weirdness of traditional Linux installations: the original source of information which file system the root file system is, typically is embedded in the root file system itself, in /etc/fstab. Thus, in a way, in order to know what the root file system is you need to know what the root file system is. 🤯 🤯 🤯

(Of course, the way this recursion is traditionally broken up is by then copying the root file system information from /etc/fstab into the boot loader configuration, resulting in a situation where the primary source of information for this — i.e. /etc/fstab — is actually mostly irrelevant, and the secondary source — i.e. the copy in the boot loader — becomes the configuration that actually matters.)

Today, the GPT partition type UUIDs defined by the specification have been adopted quite widely, by distributions and their installers, as well as a variety of partitioning tools and other tools.

In this article I want to highlight how the various tools the systemd project provides make use of the concepts the specification introduces.

But before we start with that, let's underline why tagging partitions with these descriptive partition type UUIDs (and the associated partition flags) is a good thing, besides the philosophical points made above.

  1. Simplicity: in particular OS installers become simpler — adjusting /etc/fstab as part of the installation is not necessary anymore, as the partitioning step already put all information into place for assembling the system properly at boot. i.e. installing doesn't mean that you always have to get fdisk and /etc/fstab into place, the former suffices entirely.

  2. Robustness: since partition tables mostly remain static after installation the chance of corruption is much lower than if the data is stored in file systems (e.g. in /etc/fstab). Moreover by associating the metadata directly with the objects it describes the chance of things getting out of sync is reduced. (i.e. if you lose /etc/fstab, or forget to rerun your initrd builder you still know what a partition is supposed to be just by looking at it.)

  3. Programmability: if partitions are self-descriptive it's much easier to automatically process them with various tools. In fact, this blog story is mostly about that: various systemd tools can naturally process disk images prepared like this.

  4. Alternative entry points: on traditional disk images, the boot loader needs to be told which kernel command line option root= to use, which then provides access to the root file system, where /etc/fstab is then found which describes the rest of the file systems. Where precisely root= is configured for the boot loader highly depends on the boot loader and distribution used, and is typically encoded in a Turing complete programming language (Grub…). This makes it very hard to automatically determine the right root file system to use, to implement alternative entry points to the system. By alternative entry points I mean other ways to boot the disk image, specifically for running it as a systemd-nspawn container — but this extends to other mechanisms where the boot loader may be bypassed to boot up the system, for example qemu when configured without a boot loader.

  5. User friendliness: it's simply a lot nicer for the user looking at a partition table if the partition table explains what is what, instead of just saying "Hey, this is a Linux partition!" and nothing else.

Uses for the concept

Now that we cleared up the Why?, lets have a closer look how this is currently used and exposed in systemd's various components.

Use #1: Running a disk image in a container

If a disk image follows the Discoverable Partition Specification then systemd-nspawn has all it needs to just boot it up. Specifically, if you have a GPT disk image in a file foobar.raw and you want to boot it up in a container, just run systemd-nspawn -i foobar.raw -b, and that's it (you can specify a block device like /dev/sdb too if you like). It becomes easy and natural to prepare disk images that can be booted either on a physical machine, inside a virtual machine manager or inside such a container manager: the necessary meta-information is included in the image, easily accessible before actually looking into its file systems.

Use #2: Booting an OS image on bare-metal without /etc/fstab or kernel command line root=

If a disk image follows the specification in many cases you can remove /etc/fstab (or never even install it) — as the basic information needed is already included in the partition table. The systemd-gpt-auto-generator logic implements automatic discovery of the root file system as well as all auxiliary file systems. (Note that the former requires an initrd that uses systemd, some more conservative distributions do not support that yet, unfortunately). Effectively this means you can boot up a kernel/initrd with an entirely empty kernel command line, and the initrd will automatically find the root file system (by looking for a suitably marked partition on the same drive the EFI System Partition was found on).

(Note, if /etc/fstab or root= exist and contain relevant information they always takes precedence over the automatic logic. This is in particular useful to tweaks thing by specifying additional mount options and such.)

Use #3: Mounting a complex disk image for introspection or manipulation

The systemd-dissect tool may be used to introspect and manipulate OS disk images that implement the specification. If you pass the path to a disk image (or block device) it will extract various bits of useful information from the image (e.g. what OS is this? what partitions to mount?) and display it.

With the --mount switch a disk image (or block device) can be mounted to some location. This is useful for looking what is inside it, or changing its contents. This will dissect the image and then automatically mount all contained file systems matching their GPT partition description to the right places, so that you subsequently could chroot into it. (But why chroot if you can just use systemd-nspawn? 😎)

Use #4: Copying files in and out of a disk image

The systemd-dissect tool also has two switches --copy-from and --copy-to which allow copying files out of or into a compliant disk image, taking all included file systems and the resulting mount hierarchy into account.

Use #5: Running services directly off a disk image

The RootImage= setting in service unit files accepts paths to compliant disk images (or block device nodes), and can mount them automatically, running service binaries directly off them (in chroot() style). In fact, this is the base for the Portable Service concept of systemd.

Use #6: Provisioning disk images

systemd provides various tools that can run operations provisioning disk images in an "offline" mode. Specifically:

systemd-tmpfiles

With the --image= switch systemd-tmpfiles can directly operate on a disk image, and for example create all directories and other inodes defined in its declarative configuration files included in the image. This can be useful for example to set up the /var/ or /etc/ tree according to such configuration before first boot.

systemd-sysusers

Similar, the --image= switch of systemd-sysusers tells the tool to read the declarative system user specifications included in the image and synthesizes system users from it, writing them to the /etc/passwd (and related) files in the image. This is useful for provisioning these users before the first boot, for example to ensure UID/GID numbers are pre-allocated, and such allocations not delayed until first boot.

systemd-machine-id-setup

The --image= switch of systemd-machine-id-setup may be used to provision a fresh machine ID into /etc/machine-id of a disk image, before first boot.

systemd-firstboot

The --image= switch of systemd-firstboot may be used to set various basic system setting (such as root password, locale information, hostname, …) on the specified disk image, before booting it up.

Use #7: Extracting log information

The journalctl switch --image= may be used to show the journal log data included in a disk image (or, as usual, the specified block device). This is very useful for analyzing failed systems offline, as it gives direct access to the logs without any further, manual analysis.

Use #8: Automatic repartitioning/growing of file systems

The systemd-repart tool may be used to repartition a disk or image in an declarative and additive way. One primary use-case for it is to run during boot on physical or VM systems to grow the root file system to the disk size, or to add in, format, encrypt, populate additional partitions at boot.

With its --image= switch it the tool may operate on compliant disk images in offline mode of operation: it will then read the partition definitions that shall be grown or created off the image itself, and then apply them to the image. This is particularly useful in combination with the --size= which allows growing disk images to the specified size.

Specifically, consider the following work-flow: you download a minimized disk image foobar.raw that contains only the minimized root file system (and maybe an ESP, if you want to boot it on bare-metal, too). You then run systemd-repart --image=foo.raw --size=15G to enlarge the image to the 15G, based on the declarative rules defined in the repart.d/ drop-in files included in the image (this means this can grow the root partition, and/or add in more partitions, for example for /srv or so, maybe encrypted with a locally generated key or so). Then, you proceed to boot it up with systemd-nspawn --image=foo.raw -b, making use of the full 15G.

Versioning + Multi-Arch

Disk images implementing this specifications can carry OS executables in one of three ways:

  1. Only a root file system

  2. Only a /usr/ file system (in which case the root file system is automatically picked as tmpfs).

  3. Both a root and a /usr/file system (in which case the two are combined, the /usr/ file system mounted into the root file system, and the former possibly in read-only fashion`)

They may also contain OS executables for different architectures, permitting "multi-arch" disk images that can safely boot up on multiple CPU architectures. As the root and /usr/ partition type UUIDs are specific to architectures this is easily done by including one such partition for x86-64, and another for aarch64. If the image is now used on an x86-64 system automatically the former partition is used, on aarch64 the latter.

Moreover, these OS executables may be contained in different versions, to implement a simple versioning scheme: when tools such as systemd-nspawn or systemd-gpt-auto-generator dissect a disk image, and they find two or more root or /usr/ partitions of the same type UUID, they will automatically pick the one whose GPT partition label (a 36 character free-form string every GPT partition may have) is the newest according to strverscmp() (OK, truth be told, we don't use strverscmp() as-is, but a modified version with some more modern syntax and semantics, but conceptually identical).

This logic allows to implement a very simple and natural A/B update scheme: an updater can drop multiple versions of the OS into separate root or /usr/ partitions, always updating the partition label to the version included there-in once the download is complete. All of the tools described here will then honour this, and always automatically pick the newest version of the OS.

Verity

When building modern OS appliances, security is highly relevant. Specifically, offline security matters: an attacker with physical access should have a difficult time modifying the OS in a way that isn't noticed. i.e. think of a car or a cell network base station: these appliances are usually parked/deployed in environments attackers can get physical access to: it's essential that in this case the OS itself sufficiently protected, so that the attacker cannot just mount the OS file system image, make modifications (inserting a backdoor, spying software or similar) and the system otherwise continues to run without this being immediately detected.

A great way to implement offline security is via Linux' dm-verity subsystem: it allows to securely bind immutable disk IO to a single, short trusted hash value: if an attacker manages to offline modify the disk image the modified disk image won't match the trusted hash anymore, and will not be trusted anymore (depending on policy this then just result in IO errors being generated, or automatic reboot/power-off).

The Discoverable Partitions Specification declares how to include Verity validation data in disk images, and how to relate them to the file systems they protect, thus making if very easy to deploy and work with such protected images. For example systemd-nspawn supports a --root-hash= switch, which accepts the Verity root hash and then will automatically assemble dm-verity with this, automatically matching up the payload and verity partitions. (Alternatively, just place a .roothash file next to the image file).

Future

The above already is a powerful tool set for working with disk images. However, there are some more areas I'd like to extend this logic to:

bootctl

Similar to the other tools mentioned above, bootctl (which is a tool to interface with the boot loader, and install/update systemd's own EFI boot loader sd-boot) should learn a --image= switch, to make installation of the boot loader on disk images easy and natural. It would automatically find the ESP and other relevant partitions in the image, and copy the boot loader binaries into them (or update them).

coredumpctl

Similar to the existing journalctl --image= logic the coredumpctl tool should also gain an --image= switch for extracting coredumps from compliant disk images. The combination of journalctl --image= and coredumpctl --image= would make it exceptionally easy to work with OS disk images of appliances and extracting logging and debugging information from them after failures.

And that's all for now. Please refer to the specification and the man pages for further details. If your distribution's installer does not yet tag the GPT partition it creates with the right GPT type UUIDs, consider asking them to do so.

Thank you for your time.

LibreOffice GTK4 Port: More MenuButtons

Posted by Caolán McNamara on June 04, 2021 04:02 PM

 

More MenuButton features working under GTK4. Now with working radio entries in GtkPopoverMenu dropdowns.

Mike Lindell's Cyber "Evidence"

Posted by Matthew Garrett on June 04, 2021 05:14 AM
Mike Lindell, notable for absolutely nothing relevant in this field, today filed a lawsuit against a couple of voting machine manufacturers in response to them suing him for defamation after he claimed that they were covering up hacks that had altered the course of the US election. Paragraph 104 of his suit asserts that he has evidence of at least 20 documented hacks, including the number of votes that were changed. The citation is just a link to a video called Absolute 9-0, which claims to present sufficient evidence that the US supreme court will come to a 9-0 decision that the election was tampered with.

The claim is that Lindell was provided with a set of files on the 9th of January, and gave these to some cyber experts to verify. These experts identified them as packet captures. The video contains scrolling hex, and we are told that this is the raw encrypted data from the files. In reality, the hex values correspond very clearly to printable ASCII, and appear to just be the Pennsylvania voter roll. They're not encrypted, and they're not packet captures (they contain no packet headers).

20 of these packet captures were then selected and analysed, giving us the tables contained within Exhibit 12. The alleged source IPs appear to correspond to the networks the tables claim, and the latitude and longitude presumably just come from a geoip lookup of some sort (although clearly those values are far too precise to be accurate). But if we look at the target IPs, we find something interesting. Most of them resolve to the website for the county that was the nominal target (eg, 198.108.253.104 is www.deltacountymi.org). So, we're supposed to believe that in many cases, the county voting infrastructure was hosted on the county website.

Unfortunately we're not given the destination port, but 198.108.253.104 isn't listening on anything other than 80 and 443. We're told that the packet data is encrypted, so presumably it's over HTTPS. So, uh, how did they decrypt this to figure out how many votes were switched? If Mike's hackers have broken TLS, they really don't need to be dealing with this.

We're also given some background information on how it's impossible to reconstruct packet captures after the fact (untrue), or that modifying them would change their hashes (true, but in the absence of known good hash values that tells us nothing), but it's pretty clear that nothing we're shown actually demonstrates what we're told it does.

In summary: yes, any supreme court decision on this would be 9-0, just not the way he's hoping for.

Update: It was pointed out that this data appears to be part of a larger dataset. This one is even more dubious - it somehow has MAC addresses for both the source and destination (which is impossible), and almost none of these addresses are in actual issued ranges.

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LibreOffice GTK4 Port: MenuButtons

Posted by Caolán McNamara on June 03, 2021 04:18 PM

 

Initial MenuButton and Popover support under GTK4 now working

GTK4 LibreOffice Port: Print Dialog

Posted by Caolán McNamara on June 02, 2021 07:04 PM

 

LibreOffice's Print Dialog in GTK4 Port with fancy "suggested-action" blue "Print" button.

Producing a trustworthy x86-based Linux appliance

Posted by Matthew Garrett on June 02, 2021 04:21 AM
Let's say you're building some form of appliance on top of general purpose x86 hardware. You want to be able to verify the software it's running hasn't been tampered with. What's the best approach with existing technology?

Let's split this into two separate problems. The first is to do as much as we can to ensure that the software can't be modified without our consent[1]. This requires that each component in the boot chain verify that the next component is legitimate. We call the first component in this chain the root of trust, and in the x86 world this is the system firmware[2]. This firmware is responsible for verifying the bootloader, and the easiest way to do this on x86 is to use UEFI Secure Boot. In this setup the firmware contains a set of trusted signing certificates and will only boot executables with a chain of trust to one of these certificates. Switching the system into setup mode from the firmware menu will allow you to remove the existing keys and install new ones.

(Note: You shouldn't use the trusted certificate directly for signing bootloaders - instead, the trusted certificate should be used to sign another certificate and the key for that certificate used to sign your bootloader. This way, if you ever need to revoke the signing certificate, you can simply sign a new one with the trusted parent and push out a revocation update instead of having to provision new keys)

But what do you want to sign? In the general purpose Linux world, we use an intermediate bootloader called Shim to bridge from the Microsoft signing authority to a distribution one. Shim then verifies the signature on grub, and grub in turn verifies the signature on the kernel. This is a large body of code that exists because of the use cases that general purpose distributions need to support - primarily, booting on arbitrary off the shelf hardware, and allowing arbitrary and complicated boot setups. This is unnecessary in the appliance case, where the hardware target can be well defined, where there's no need for interoperability with the Microsoft signing authority, and where the boot configuration can be extremely static.

We can skip all of this complexity using systemd-boot's unified Linux image support. This has the format described here, but the short version is that it's simply a kernel and initramfs linked into a small EFI executable that will run them. Instructions for generating such an image are here, and if you follow them you'll end up with a single static image that can be directly executed by the firmware. Signing this avoids dealing with a whole host of problems associated with relying on shim and grub, but note that you'll be embedding the initramfs as well. Again, this should be fine for appliance use-cases, but you'll need your build system to support building the initramfs at image creation time rather than relying on it being generated on the host.

At this point we have a single image that can be verified by the firmware and will get us to the point of a running kernel and initramfs. Unless you've got enough RAM that you can put your entire workload in the initramfs, you're going to want a filesystem as well, and you're going to want to verify that that filesystem hasn't been tampered with. The easiest approach to this is to use dm-verity, a device-mapper layer that uses a hash tree to verify that the filesystem contents haven't been modified. The kernel needs to know what the root hash is, so this can either be embedded into your initramfs image or into the kernel command line. Either way, it'll end up in the signed boot image, so nobody will be able to tamper with it.

It's important to note that a dm-verity partition is read-only - the kernel doesn't have the cryptographic secret that would be required to generate a new hash tree if the partition is modified. So if you require the ability to write data or logs anywhere, you'll need to add a new partition for that. If this partition is unencrypted, an attacker with access to the device will be able to put whatever they want on there. You should treat any data you read from there as untrusted, and ensure that it's validated before use (ie, don't just feed it to a random parser written in C and expect that everything's going to be ok). On the other hand, if it's encrypted, remember that you can't just put the encryption key in the boot image - an attacker with access to the device is going to be able to dump that and extract it. You'll probably want to use a TPM-sealed encryption secret, which will be discussed later on.

At this point everything in the boot process is cryptographically verified, and so should be difficult to tamper with. Unfortunately this isn't really sufficient - on x86 systems there's typically no verification of the integrity of the secure boot database. An attacker with physical access to the system could attach a programmer directly to the firmware flash and rewrite the secure boot database to include keys they control. They could then replace the boot image with one that they've signed, and the machine would happily boot code that the attacker controlled. We need to be able to demonstrate that the system booted using the correct secure boot keys, and the only way we can do that is to use the TPM.

I wrote an introduction to TPMs a while back. The important thing to know here is that the TPM contains a set of Platform Configuration Registers that are large enough to contain a cryptographic hash. During boot, each component of the boot process will generate a "measurement" of other security critical components, including the next component to be booted. These measurements are a representation of the data in question - they may simply be a hash of the object being measured, or the hash of a structure containing various pieces of metadata. Each measurement is passed to the TPM, along with the PCR it should be measured into. The TPM takes the new measurement, appends it to the existing value, and then stores the hash of this concatenated data in the PCR. This means that the final PCR value depends not only on the measurement, but also on every previous measurement. Without breaking the hash algorithm, there's no way to set the PCR to an arbitrary value. The hash values and some associated data are stored in a log that's kept in system RAM, which we'll come back to later.

Different PCRs store different pieces of information, but the one that's most interesting to us is PCR 7. Its use is documented in the TCG PC Client Platform Firmware Profile (section 3.3.4.8), but the short version is that the firmware will measure the secure boot keys that are used to boot the system. If the secure boot keys are altered (such as by an attacker flashing new ones), the PCR 7 value will change.

What can we do with this? There's a couple of choices. For devices that are online, we can perform remote attestation, a process where the device can provide a signed copy of the PCR values to another system. If the system also provides a copy of the TPM event log, the individual events in the log can be replayed in the same way that the TPM would use to calculate the PCR values, and then compared to the actual PCR values. If they match, that implies that the log values are correct, and we can then analyse individual log entries to make assumptions about system state. If a device has been tampered with, the PCR 7 values and associated log entries won't match the expected values, and we can detect the tampering.

If a device is offline, or if there's a need to permit local verification of the device state, we still have options. First, we can perform remote attestation to a local device. I demonstrated doing this over Bluetooth at LCA back in 2020. Alternatively, we can take advantage of other TPM features. TPMs can be configured to store secrets or keys in a way that renders them inaccessible unless a chosen set of PCRs have specific values. This is used in tpm2-totp, which uses a secret stored in the TPM to generate a TOTP value. If the same secret is enrolled in any standard TOTP app, the value generated by the machine can be compared to the value in the app. If they match, the PCR values the secret was sealed to are unmodified. If they don't, or if no numbers are generated at all, that demonstrates that PCR 7 is no longer the same value, and that the system has been tampered with.

Unfortunately, TOTP requires that both sides have possession of the same secret. This is fine when a user is making that association themselves, but works less well if you need some way to ship the secret on a machine and then separately ship the secret to a user. If the user can simply download the secret via some API, so can an attacker. If an attacker has the secret, they can modify the secure boot database and re-seal the secret to the new PCR 7 value. That means having to add some form of authentication, along with a strong binding of machine serial number to a user (in order to avoid someone with valid credentials simply downloading all the secrets).

Instead, we probably want some mechanism that uses asymmetric cryptography. A keypair can be generated on the TPM, which will refuse to release an unencrypted copy of the private key. The public key, however, can be exported and stored. If it's acceptable for a verification app to connect to the internet then the public key can simply be obtained that way - if not, a certificate can be issued to the key, and this exposed to the verifier via a QR code. The app then verifies that the certificate is signed by the vendor, and if so extracts the public key from that. The private key can have an associated policy that only permits its use when PCR 7 has an appropriate value, so the app then generates a nonce and asks the user to type that into the device. The device generates a signature over that nonce and displays that as a QR code. The app verifies the signature matches, and can then assert that PCR 7 has the expected value.

Once we can assert that PCR 7 has the expected value, we can assert that the system booted something signed by us and thus infer that the rest of the boot chain is also secure. But this is still dependent on the TPM obtaining trustworthy information, and unfortunately the bus that the TPM sits on isn't really terribly secure (TPM Genie is an example of an interposer for i2c-connected TPMs, but there's no reason an LPC one can't be constructed to attack the sort usually used on PCs). TPMs do support encrypted communication channels, but bootstrapping those isn't straightforward without firmware support. The easiest way around this is to make use of a firmware-based TPM, where the TPM is implemented in software running on an ancillary controller. Intel's solution is part of their Platform Trust Technology and runs on the Management Engine, AMD run it on the Platform Security Processor. In both cases it's not terribly feasible to intercept the communications, so we avoid this attack. The downside is that we're then placing more trust in components that are running much more code than a TPM would and which have a correspondingly larger attack surface. Which is preferable is going to depend on your threat model.

Most of this should be achievable using Yocto, which now has support for dm-verity built in. It's almost certainly going to be easier using this than trying to base on top of a general purpose distribution. I'd love to see this become a largely push button receive secure image process, so might take a go at that if I have some free time in the near future.

[1] Obviously technologies that can be used to ensure nobody other than me is able to modify the software on devices I own can also be used to ensure that nobody other than the manufacturer is able to modify the software on devices that they sell to third parties. There's no real technological solution to this problem, but we shouldn't allow the fact that a technology can be used in ways that are hostile to user freedom to cause us to reject that technology outright.
[2] This is slightly complicated due to the interactions with the Management Engine (on Intel) or the Platform Security Processor (on AMD). Here's a good writeup on the Intel side of things.

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LibreOffice GTK4 Port: Some more progress

Posted by Caolán McNamara on May 26, 2021 08:01 PM

More progress, Entries, SpinButtons, PushButtons, CheckBoxes/RadioButtons, TreeViews and basic Notebooks and DrawingAreas working in the GTK4 port


Reminder: SoupSessionSync and SoupSessionAsync default to no TLS certificate verification

Posted by Michael Catanzaro on May 25, 2021 02:12 PM

This is a public service announcement! The modern SoupSession class is secure by default, but the older, deprecated SoupSessionSync and SoupSessionAsync subclasses of SoupSession are not. If your code uses SoupSessionSync or SoupSessionAsync and does not set SoupSession:tls-database, SoupSession:ssl-use-system-ca-file, or SoupSession:ssl-ca-file, then you get no TLS certificate verification. This is almost always worth requesting a CVE.

SoupSessionSync and SoupSessionAsync have both been deprecated since libsoup 2.42 was released back in 2013, so surely they are no longer commonly used, right? Right? Of course not, so please check your applications and make sure they’re actually doing TLS certificate verification. Thanks!

Update: we decided to change this behavior.

File Descriptor Limits

Posted by Lennart Poettering on May 18, 2021 10:00 PM

TL;DR: don't use select() + bump the RLIMIT_NOFILE soft limit to the hard limit in your modern programs.

The primary way to reference, allocate and pin runtime OS resources on Linux today are file descriptors ("fds"). Originally they were used to reference open files and directories and maybe a bit more, but today they may be used to reference almost any kind of runtime resource in Linux userspace, including open devices, memory (memfd_create(2)), timers (timefd_create(2)) and even processes (with the new pidfd_open(2) system call). In a way, the philosophically skewed UNIX concept of "everything is a file" through the proliferation of fds actually acquires a bit of sensible meaning: "everything has a file descriptor" is certainly a much better motto to adopt.

Because of this proliferation of fds, non-trivial modern programs tend to have to deal with substantially more fds at the same time than they traditionally did. Today, you'll often encounter real-life programs that have a few thousand fds open at the same time.

Like on most runtime resources on Linux limits are enforced on file descriptors: once you hit the resource limit configured via RLIMIT_NOFILE any attempt to allocate more is refused with the EMFILE error — until you close a couple of those you already have open.

Because fds weren't such a universal concept traditionally, the limit of RLIMIT_NOFILE used to be quite low. Specifically, when the Linux kernel first invokes userspace it still sets RLIMIT_NOFILE to a low value of 1024 (soft) and 4096 (hard). (Quick explanation: the soft limit is what matters and causes the EMFILE issues, the hard limit is a secondary limit that processes may bump their soft limit to — if they like — without requiring further privileges to do so. Bumping the limit further would require privileges however.). A limit of 1024 fds made fds a scarce resource: APIs tried to be careful with using fds, since you simply couldn't have that many of them at the same time. This resulted in some questionable coding decisions and concepts at various places: often secondary descriptors that are very similar to fds — but were not actually fds — were introduced (e.g. inotify watch descriptors), simply to avoid for them the low limits enforced on true fds. Or code tried to aggressively close fds when not absolutely needing them (e.g. ftw()/nftw()), losing the nice + stable "pinning" effect of open fds.

Worse though is that certain OS level APIs were designed having only the low limits in mind. The worst offender being the BSD/POSIX select(2) system call: it only works with fds in the numeric range of 0…1023 (aka FD_SETSIZE-1). If you have an fd outside of this range, tough luck: select() won't work, and only if you are lucky you'll detect that and can handle it somehow.

Linux fds are exposed as simple integers, and for most calls it is guaranteed that the lowest unused integer is allocated for new fds. Thus, as long as the RLIMIT_NOFILE soft limit is set to 1024 everything remains compatible with select(): the resulting fds will also be below 1024. Yay. If we'd bump the soft limit above this threshold though and at some point in time an fd higher than the threshold is allocated, this fd would not be compatible with select() anymore.

Because of that, indiscriminately increasing the soft RLIMIT_NOFILE resource limit today for every userspace process is problematic: as long as there's userspace code still using select() doing so will risk triggering hard-to-handle, hard-to-debug errors all over the place.

However, given the nowadays ubiquitous use of fds for all kinds of resources (did you know, an eBPF program is an fd? and a cgroup too? and attaching an eBPF program to cgroup is another fd? …), we'd really like to raise the limit anyway. 🤔

So before we continue thinking about this problem, let's make the problem more complex (…uh, I mean… "more exciting") first. Having just one hard and one soft per-process limit on fds is boring. Let's add more limits on fds to the mix. Specifically on Linux there are two system-wide sysctls: fs.nr_open and fs.file-max. (Don't ask me why one uses a dash and the other an underscore, or why there are two of them...) On today's kernels they kinda lost their relevance. They had some originally, because fds weren't accounted by any other counter. But today, the kernel tracks fds mostly as small pieces of memory allocated on userspace requests — because that's ultimately what they are —, and thus charges them to the memory accounting done anyway.

So now, we have four limits (actually: five if you count the memory accounting) on the same kind of resource, and all of them make a resource artificially scarce that we don't want to be scarce. So what to do?

Back in systemd v240 already (i.e. 2019) we decided to do something about it. Specifically:

  • Automatically at boot we'll now bump the two sysctls to their maximum, making them effectively ineffective. This one was easy. We got rid of two pretty much redundant knobs. Nice!

  • The RLIMIT_NOFILE hard limit is bumped substantially to 512K. Yay, cheap fds! You may have an fd, and you, and you as well, everyone may have an fd!

  • But … we left the soft RLIMIT_NOFILE limit at 1024. We weren't quite ready to break all programs still using select() in 2019 yet. But it's not as bad as it might sound I think: given the hard limit is bumped every program can easily opt-in to a larger number of fds, by setting the soft limit to the hard limit early on — without requiring privileges.

So effectively, with this approach fds should be much less scarce (at least for programs that opt into that), and the limits should be much easier to configure, since there are only two knobs now one really needs to care about:

  • Configure the RLIMIT_NOFILE hard limit to the maximum number of fds you actually want to allow a process.

  • In the program code then either bump the soft to the hard limit, or not. If you do, you basically declare "I understood the problem, I promise to not use select(), drown me fds please!". If you don't then effectively everything remains as it always was.

Apparently this approach worked, since the negative feedback on change was even scarcer than fds traditionally were (ha, fun!). We got reports from pretty much only two projects that were bitten by the change (one being a JVM implementation): they already bumped their soft limit automatically to their hard limit during program initialization, and then allocated an array with one entry per possible fd. With the new high limit this resulted in one massive allocation that traditionally was just a few K, and this caused memory checks to be hit.

Anyway, here's the take away of this blog story:

  • Don't use select() anymore in 2021. Use poll(), epoll, iouring, …, but for heaven's sake don't use select(). It might have been all the rage in the 1990s but it doesn't scale and is simply not designed for today's programs. I wished the man page of select() would make clearer how icky it is and that there are plenty of more preferably APIs.

  • If you hack on a program that potentially uses a lot of fds, add some simple code somewhere to its start-up that bumps the RLIMIT_NOFILE soft limit to the hard limit. But if you do this, you have to make sure your code (and any code that you link to from it) refrains from using select(). (Note: there's at least one glibc NSS plugin using select() internally. Given that NSS modules can end up being loaded into pretty much any process such modules should probably be considered just buggy.)

  • If said program you hack on forks off foreign programs, make sure to reset the RLIMIT_NOFILE soft limit back to 1024 for them. Just because your program might be fine with fds >= 1024 it doesn't mean that those foreign programs might. And unfortunately RLIMIT_NOFILE is inherited down the process tree unless explicitly set.

And that's all I have for today. I hope this was enlightening.

First Cut at a GTK4 Port of LibreOffice

Posted by Caolán McNamara on May 14, 2021 03:57 PM

Early days yet, but here's some screenshots of an experimental port of LibreOffice to GTK4 I've commited upstream during the week. It starts without crashing, renders, accepts user keystrokes and mouse events, basic Input Method support also works.

Standard GTK file dialog functionality works.

Standard GTK message dialogs work.

Context menus, now using Popover Menus, work.

Toplevel Menubar and menus, via gtk_popover_menu_bar_new_from_model, are also functional.

Working sample native GTK dialog. A work in progress, my current thinking is with GtkButtonBox gone that moving the old action area contents to the dialogs GtkHeaderBar seems the easiest route.

More doorbell adventures

Posted by Matthew Garrett on May 06, 2021 06:26 AM
Back in my last post on this topic, I'd got shell on my doorbell but hadn't figured out why the HTTP callbacks weren't always firing. I still haven't, but I have learned some more things.

Doorbird sell a chime, a network connected device that is signalled by the doorbell when someone pushes a button. It costs about $150, which seems excessive, but would solve my problem (ie, that if someone pushes the doorbell and I'm not paying attention to my phone, I miss it entirely). But given a shell on the doorbell, how hard could it be to figure out how to mimic the behaviour of one?

Configuration for the doorbell is all stored under /mnt/flash, and there's a bunch of files prefixed 1000eyes that contain config (1000eyes is the German company that seems to be behind Doorbird). One of these was called 1000eyes.peripherals, which seemed like a good starting point. The initial contents were {"Peripherals":[]}, so it seemed likely that it was intended to be JSON. Unfortunately, since I had no access to any of the peripherals, I had no idea what the format was. I threw the main application into Ghidra and found a function that had debug statements referencing "initPeripherals and read a bunch of JSON keys out of the file, so I could simply look at the keys it referenced and write out a file based on that. I did so, and it didn't work - the app stubbornly refused to believe that there were any defined peripherals. The check that was failing was pcVar4 = strstr(local_50[0],PTR_s_"type":"_0007c980);, which made no sense, since I very definitely had a type key in there. And then I read it more closely. strstr() wasn't being asked to look for "type":, it was being asked to look for "type":". I'd left a space between the : and the opening " in the value, which meant it wasn't matching. The rest of the function seems to call an actual JSON parser, so I have no idea why it doesn't just use that for this part as well, but deleting the space and restarting the service meant it now believed I had a peripheral attached.

The mobile app that's used for configuring the doorbell now showed a device in the peripherals tab, but it had a weird corrupted name. Tapping it resulted in an error telling me that the device was unavailable, and on the doorbell itself generated a log message showing it was trying to reach a device with the hostname bha-04f0212c5cca and (unsurprisingly) failing. The hostname was being generated from the MAC address field in the peripherals file and was presumably supposed to be resolved using mDNS, but for now I just threw a static entry in /etc/hosts pointing at my Home Assistant device. That was enough to show that when I opened the app the doorbell was trying to call a CGI script called peripherals.cgi on my fake chime. When that failed, it called out to the cloud API to ask it to ask the chime[1] instead. Since the cloud was completely unaware of my fake device, this didn't work either. I hacked together a simple server using Python's HTTPServer and was able to return data (another block of JSON). This got me to the point where the app would now let me get to the chime config, but would then immediately exit. adb logcat showed a traceback in the app caused by a failed assertion due to a missing key in the JSON, so I ran the app through jadx, found the assertion and from there figured out what keys I needed. Once that was done, the app opened the config page just fine.

Unfortunately, though, I couldn't edit the config. Whenever I hit "save" the app would tell me that the peripheral wasn't responding. This was strange, since the doorbell wasn't even trying to hit my fake chime. It turned out that the app was making a CGI call to the doorbell, and the thread handling that call was segfaulting just after reading the peripheral config file. This suggested that the format of my JSON was probably wrong and that the doorbell was not handling that gracefully, but trying to figure out what the format should actually be didn't seem easy and none of my attempts improved things.

So, new approach. Rather than writing the config myself, why not let the doorbell do it? I should be able to use the genuine pairing process if I could mimic the chime sufficiently well. Hitting the "add" button in the app asked me for the username and password for the chime, so I typed in something random in the expected format (six characters followed by four zeroes) and a sufficiently long password and hit ok. A few seconds later it told me it couldn't find the device, which wasn't unexpected. What was a little more unexpected was that the log on the doorbell was showing it trying to hit another bha-prefixed hostname (and, obviously, failing). The hostname contains the MAC address, but I hadn't told the doorbell the MAC address of the chime, just its username. Some more digging showed that the doorbell was calling out to the cloud API, giving it the 6 character prefix from the username and getting a MAC address back. Doing the same myself revealed that there was a straightforward mapping from the prefix to the mac address - changing the final character from "a" to "b" incremented the MAC by one. It's actually just a base 26 encoding of the MAC, with aaaaaa translating to 00408C000000.

That explained how the hostname was being generated, and in return I was able to work backwards to figure out which username I should use to generate the hostname I was already using. Attempting to add it now resulted in the doorbell making another CGI call to my fake chime in order to query its feature set, and by mocking that up as well I was able to send back a file containing X-Intercom-Type, X-Intercom-TypeId and X-Intercom-Class fields that made the doorbell happy. I now had a valid JSON file, which cleared up a couple of mysteries. The corrupt name was because the name field isn't supposed to be ASCII - it's base64 encoded UTF16-BE. And the reason I hadn't been able to figure out the JSON format correctly was because it looked something like this:

{"Peripherals":[]{"prefix":{"type":"DoorChime","name":"AEQAbwBvAHIAYwBoAGkAbQBlACAAVABlAHMAdA==","mac":"04f0212c5cca","user":"username","password":"password"}}]}


Note that there's a total of one [ in this file, but two ]s? Awesome. Anyway, I could now modify the config in the app and hit save, and the doorbell would then call out to my fake chime to push config to it. Weirdly, the association between the chime and a specific button on the doorbell is only stored on the chime, not on the doorbell. Further, hitting the doorbell didn't result in any more HTTP traffic to my fake chime. However, it did result in some broadcast UDP traffic being generated. Searching for the port number led me to the Doorbird LAN API and a complete description of the format and encryption mechanism in use. Argon2I is used to turn the first five characters of the chime's password (which is also stored on the doorbell itself) into a 256-bit key, and this is used with ChaCha20 to decrypt the payload. The payload then contains a six character field describing the device sending the event, and then another field describing the event itself. Some more scrappy Python and I could pick up these packets and decrypt them, which showed that they were being sent whenever any event occurred on the doorbell. This explained why there was no storage of the button/chime association on the doorbell itself - the doorbell sends packets for all events, and the chime is responsible for deciding whether to act on them or not.

On closer examination, it turns out that these packets aren't just sent if there's a configured chime. One is sent for each configured user, avoiding the need for a cloud round trip if your phone is on the same network as the doorbell at the time. There was literally no need for me to mimic the chime at all, suitable events were already being sent.

Still. There's a fair amount of WTFery here, ranging from the strstr() based JSON parsing, the invalid JSON, the symmetric encryption that uses device passwords as the key (requiring the doorbell to be aware of the chime's password) and the use of only the first five characters of the password as input to the KDF. It doesn't give me a great deal of confidence in the rest of the device's security, so I'm going to keep playing.

[1] This seems to be to handle the case where the chime isn't on the same network as the doorbell

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Untitled Post

Posted by Caolán McNamara on April 30, 2021 04:07 PM

Math Selection Rendering

Towards 7.2 the Math edit window text selection is now drawn the same as the selection in the main applications. This affects the selection of similar uses of this EditView in LibreOffice such as the writer comments in sidebar.

fwupd 1.6.0

Posted by Richard Hughes on April 26, 2021 10:12 AM

I’ve just released the first release of the 1.6.x series, and since 1.5.x some internal plugin API has been changed and removed. Although we’ve tested this release on all the hardware we have regression tests for, bugs may have crept in; please report failures to the issue tracker as required.

There are several new plugins adding support for new hardware and a lot of code has been migrated to the new plugin API. The public libfwupd API also has some trivial additions, although no action is required.

This release adds the following notable features:

  • Add a composite ID that is used to identify dock device components
  • Add an Intel Flash Descriptor parser
  • Add API to allow the device to report its own battery level
  • Add API to refount why the the device is non-updatable
  • Add lspcon-i2c-spi programmer support
  • Add more hardware support to the pixart-rf plugin
  • Add some more new category types for firmware to use
  • Add support for downloading the SPI image from the Intel eSPI device
  • Add support for some Analogix hardware
  • Add support for writing SREC firmware
  • Add the firmware-sign command to fwupdtool to allow resigning archives
  • Split UEFI EFI binary into a subproject
  • Use an OFD or Unix lock to prevent more than one fwupdtool process

This release fixes the following bugs:

  • Actually write the bcm57xx stage1 version into the file
  • Add option to disable the UEFI capsule splash screen generation
  • Avoid use-after-free when specifying the VID/PID in dfu-tool
  • Cancel the GDBusObjectManager operation to fix a potential crash
  • Check PixArt firmware compatibility with hardware before flashing
  • Do not check for native dependencies as target dependencies
  • Do not use help2man to build manual pages
  • Fix a crash when shutting down the daemon
  • Fix build on musl
  • Fix build when using BSD
  • Fix /etc/os-release ID_LIKE field parsing
  • Force the synaptics-rmi hardware into IEP mode as required
  • Never allow D-Bus replacement when a firmware update is in operation
  • Offer the user to refresh the remote after enabling
  • Remove unused, unsafe and deprecated functions from libfwupdplugin
  • Simplify asking the user about reviews
  • Write BMP data directly without using PIL
  • Write synaptics-rmi files with valid checksum data

From a Fedora point of view, I’m waiting for the fwupd-efi new package to be reviewed and approved, and then I’ll upload the new version for Rawhide only. I might put 1.6.x into Fedora 34 after a couple more minor releases, but at the moment I’m keeping it with the 1.5.x stable branch which has all the important fixes backported. There’s a lot of new code in 1.6.x which needs to settle.

An accidental bootsplash

Posted by Matthew Garrett on April 23, 2021 11:21 AM
Back in 2005 we had Debconf in Helsinki. Earlier in the year I'd ended up invited to Canonical's Ubuntu Down Under event in Sydney, and one of the things we'd tried to design was a reasonable graphical boot environment that could also display status messages. The design constraints were awkward - we wanted it to be entirely in userland (so we didn't need to carry kernel patches), and we didn't want to rely on vesafb[1] (because at the time we needed to reinitialise graphics hardware from userland on suspend/resume[2], and vesa was not super compatible with that). Nothing currently met our requirements, but by the time we'd got to Helsinki there was a general understanding that Paul Sladen was going to implement this.

The Helsinki Debconf ended being an extremely strange event, involving me having to explain to Mark Shuttleworth what the physics of a bomb exploding on a bus were, many people being traumatised by the whole sauna situation, and the whole unfortunate water balloon incident, but it also involved Sladen spending a bunch of time trying to produce an SVG of a London bus as a D-Bus logo and not really writing our hypothetical userland bootsplash program, so on the last night, fueled by Koff that we'd bought by just collecting all the discarded empty bottles and returning them for the deposits, I started writing one.

I knew that Debian was already using graphics mode for installation despite having a textual installer, because they needed to deal with more complex fonts than VGA could manage. Digging into the code, I found that it used BOGL - a graphics library that made use of the VGA framebuffer to draw things. VGA had a pre-allocated memory range for the framebuffer[3], which meant the firmware probably wouldn't map anything else there any hitting those addresses probably wouldn't break anything. This seemed safe.

A few hours later, I had some code that could use BOGL to print status messages to the screen of a machine booted with vga16fb. I woke up some time later, somehow found myself in an airport, and while sitting at the departure gate[4] I spent a while staring at VGA documentation and worked out which magical calls I needed to make to have it behave roughly like a linear framebuffer. Shortly before I got on my flight back to the UK, I had something that could also draw a graphical picture.

Usplash shipped shortly afterwards. We hit various issues - vga16fb produced a 640x480 mode, and some laptops were not inclined to do that without a BIOS call first. 640x400 worked basically everywhere, but meant we had to redraw the art because circles don't work the same way if you change the resolution. My brief "UBUNTU BETA" artwork that was me literally writing "UBUNTU BETA" on an HP TC1100 shortly after I'd got the Wacom screen working did not go down well, and thankfully we had better artwork before release.

But 16 colours is somewhat limiting. SVGALib offered a way to get more colours and better resolution in userland, retaining our prerequisites. Unfortunately it relied on VM86, which doesn't exist in 64-bit mode on Intel systems. I ended up hacking the X.org x86emu into a thunk library that exposed the same API as LRMI, so we could run it without needing VM86. Shockingly, it worked - we had support for 256 colour bootsplashes in any supported resolution on 64 bit systems as well as 32 bit ones.

But by now it was obvious that the future was having the kernel manage graphics support, both in terms of native programming and in supporting suspend/resume. Plymouth is much more fully featured than Usplash ever was, but relies on functionality that simply didn't exist when we started this adventure. There's certainly an argument that we'd have been better off making reasonable kernel modesetting support happen faster, but at this point I had literally no idea how to write decent kernel code and everyone should be happy I kept this to userland.

Anyway. The moral of all of this is that sometimes history works out such that you write some software that a huge number of people run without any idea of who you are, and also that this can happen without you having any fucking idea what you're doing.

Write code. Do crimes.

[1] vesafb relied on either the bootloader or the early stage kernel performing a VBE call to set a mode, and then just drawing directly into that framebuffer. When we were doing GPU reinitialisation in userland we couldn't guarantee that we'd run before the kernel tried to draw stuff into that framebuffer, and there was a risk that that was mapped to something dangerous if the GPU hadn't been reprogrammed into the same state. It turns out that having GPU modesetting in the kernel is a Good Thing.

[2] ACPI didn't guarantee that the firmware would reinitialise the graphics hardware, and as a result most machines didn't. At this point Linux didn't have native support for initialising most graphics hardware, so we fell back to doing it from userland. VBEtool was a terrible hack I wrote to try to re-execute the system's graphics hardware through a range of mechanisms, and it worked in a surprising number of cases.

[3] As long as you were willing to deal with 640x480 in 16 colours

[4] Helsinki-Vantaan had astonishingly comfortable seating for time

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Free Software Charities

Posted by Michael Catanzaro on March 25, 2021 01:32 PM

I believe we have reached a point where it is time to discontinue donations to the Free Software Foundation, in light of the outrageously poor judgment shown by its board of directors in reinstating Richard Stallman to the board. I haven’t seen other free software community members calling for cutting off donations yet. Even the open letter doesn’t call for this. Edit: I’ve been pointed to the line “We urge those in a position to do so to stop supporting the Free Software Foundation,” which surely implies a call to stop donating, so I was wrong about that.

I have no doubt there will be follow-up blog posts explaining why cutting off donations is harmful to the community and the FSF’s staff, and will hurt the FSF in the long-term… but seriously, enough is enough. If we don’t draw the line here, there will never be any line anywhere. Continued support for the FSF is continued complicity, and is harming rather than helping advance the ideals of the free software community. So, where should you donate instead?

The Software Freedom Conservancy is a great choice. It hosts many member projects you’re probably familiar with, including Outreachy and git, among many others. It does some GPL compliance work and takes a strict view on software freedom. For US donors, it is a 501(c)(3), so your donations may be tax-deductible. This is where I send my free software donations not directed to GNOME. Read its statement on recent events at the FSF.

Another good option, especially for people in the EU, is the Free Software Foundation Europe, an independently-run sister organization of the FSF. If you’re not familiar with FSFE, think of it as a more moderate version of the FSF, with a special focus on Europe. It shares the same goals as the FSF, but with more reasonable leadership and much less popcorn. Read its statement on recent events at the FSF.

Most people reading this blog are likely GNOME users. Contributing to GNOME is a great way to support the desktop you use to get things done. Monthly sustaining donations are especially appreciated. Despite the historical association between GNOME and GNU, GNOME has had little to do with GNU for a long time now. The GNOME Foundation is run independently and has formally signed the open letter calling for the resignation of the FSF’s board of directors. For US donors, it is a 501(c)(3). (If you use KDE, donate to KDE here.)

It should be obvious, but this is a personal blog. None of the above organizations have endorsed cutting off donations to the FSF, to my knowledge. They would probably find it to be in poor taste to abuse a crisis to solicit funds. But I doubt they’d object if you send some money their way.

Exploring my doorbell

Posted by Matthew Garrett on March 15, 2021 06:58 PM
I've talked about my doorbell before, but started looking at it again this week because sometimes it simply doesn't send notifications to my Home Assistant setup - the push notifications appear on my phone, but the doorbell simply doesn't trigger the HTTP callback it's meant to[1]. This is obviously suboptimal, but it's also tricky to debug a device when you have no access to it.

Normally I'd just head straight in with a screwdriver, but the doorbell is shared with the other units in this building and it seemed a little anti-social to interfere with a shared resource. So I bought some broken units from ebay and pulled one of them apart. There's several boards inside, but one of them had a conveniently empty connector at the top with "TX", "RX" and "GND" labelled. Sticking a USB-serial converter on this gave me output from U-Boot, and then kernel output. Confirmation that my doorbell runs Linux, but unfortunately it didn't give me a shell prompt. My next approach would often me to just dump the flash and look for vulnerabilities that way, but this device uses TSOP-48 packaged NAND flash rather than the more convenient SPI NOR flash that I already have adapters to access. Dumping this sort of NAND isn't terribly hard, but the easiest way to do it involves desoldering it from the board and plugging it into something like a Flashcat USB adapter, and my soldering's not good enough to put it back on the board afterwards. So I wanted another approach.

U-Boot gave a short countdown to hit a key before continuing with boot, and for once hitting a key actually did something. Unfortunately it then prompted for a password, and giving the wrong one resulted in boot continuing[2]. In the past I've had good luck forcing U-Boot to drop to a prompt by simply connecting one of the data lines on SPI flash to ground while it's trying to read the kernel - the failed read causes U-Boot to error out. It turns out the same works fine on raw NAND, so I just edited the kernel boot arguments to append "init=/bin/sh" and soon I had a shell.

From here on, things were made easier by virtue of the device using the YAFFS filesystem. Unlike many flash filesystems, it's read/write, so I could make changes that would persist through to the running system. There was a convenient copy of telnetd included, but it segfaulted on startup, which reduced its usefulness. Fortunately there was also a copy of Netcat[3]. If you make a fifo somewhere on the filesystem, you can cat the fifo to a shell, pipe the shell to a netcat listener, and then pipe netcat's output back to the fifo. The shell's output all gets passed to whatever connects to netcat, and whatever's sent to netcat gets passed through the fifo back to the shell. This is, obviously, horribly insecure, but it was enough to get a root shell over the network on the running device.

The doorbell runs various bits of software, one of which is Lighttpd to provide a local API and access to the device. Another component ("nxp-client") connects to the vendor's cloud infrastructure and passes cloud commands back to the local webserver. This is where I found something strange. Lighttpd was refusing to start because its modules wanted library symbols that simply weren't present on the device. My best guess is that a firmware update went wrong and left the device in a partially upgraded state - and without a working local webserver, there was no way to perform any further updates. This may explain why this doorbell was sitting on ebay.

Anyway. Now that I had shell, I could simply dump the flash by copying it directly off the /dev/mtdblock devices - since I had netcat, I could just pipe stuff through that back to my actual computer. Now I had access to the filesystem I could extract that locally and start digging into it more deeply. One incredibly useful tool for this is qemu-user. qemu is a general purpose hardware emulation platform, usually used to emulate entire systems. But in qemu-user mode, it instead only emulates the CPU. When a piece of code tries to make a system call to access the kernel, qemu-user translates that to the appropriate calling convention for the host kernel and makes that call instead. Combined with binfmt_misc, you can configure a Linux system to be able to run Linux binaries from other architectures. One of the best things about this is that, because they're still using the host convention for making syscalls, you can run the host strace on them and see what they're doing.

What I found was that nxp-client was calling back to the cloud platform, setting up an encrypted communication channel (using ChaCha20 and a bunch of key setup stuff I couldn't be bothered picking apart) and then waiting for commands from the cloud. It would then proxy those through to the local webserver. Since I couldn't run the local lighttpd, I just wrote a trivial Python app using http.server and waited to see what requests I got. The first was a GET to a CGI script called editcgi.cgi, along with a path name. I mocked up the GET request to respond with what was on the actual filesystem. The cloud then proceeded to POST to editcgi.cgi, with the same pathname and with new file contents. editcgi.cgi is apparently able to read and write to files on the filesystem.

But this is on the interface that's exposed to the cloud client, so this didn't appear immediately useful - and, indeed, trying to hit the same CGI binary over the local network gave me a 401 unauthorized error. There's a local API spec for these doorbells, but they all refer to scripts in the bha-api namespace, and this script was in the plain cgi-bin namespace. But then I noticed that the bha-api namespace didn't actually exist in the filesystem - instead, lighttpd's mod_alias was configured to rewrite requests to bha-api through to files in cgi-bin. And by using the documented API to get a session token, I could call editcgi.cgi to read and write arbitrary files on the doorbell. Which means I can drop an extra script in /etc/rc.d/rc3.d and get a shell on my doorbell.

This all requires the ability to have local authentication credentials, so it's not a big security deal other than it allowing you to retain access to a monitoring device even after you've moved out and had your credentials revoked. I'm sure it's all fine.

[1] I can ping the doorbell from the Home Assistant machine, so it's not that the network is flaky
[2] The password appears to be hy9$gnhw0z6@ if anyone else ends up in this situation
[3] https://twitter.com/mjg59/status/654578208545751040

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Unauthenticated MQTT endpoints on Linksys Velop routers enable local DoS

Posted by Matthew Garrett on March 09, 2021 07:55 PM
(Edit: this is CVE-2021-1000002)

Linksys produces a series of wifi mesh routers under the Velop line. These routers use MQTT to send messages to each other for coordination purposes. In the version I tested against, there was zero authentication on this - anyone on the local network is able to connect to the MQTT interface on a router and send commands. As an example:
mosquitto_pub -h 192.168.1.1 -t "network/master/cmd/nodes_temporary_blacklist" -m '{"data": {"client": "f8:16:54:43:e2:0c", "duration": "3600", "action": "start"}}'
will ask the router to block the client with MAC address f8:16:54:43:e2:0c from the network for an hour. Various other MQTT topics pass parameters to shell scripts without quoting them or escaping metacharacters, so more serious outcomes may be possible.

The vendor has released two firmware updates since report - I have not verified whether either fixes this, but the changelog does not indicate any security issues were addressed.

Timeline:

2020-07-30: Submitted through the vendor's security vulnerability report form, indicating that I plan to disclose in either 90 days or after a fix is released. The form turns out to file a Bugcrowd submission.
2020-07-30: I claim the Bugcrowd submission.
2020-08-19: Vendor acknowledges the issue, is able to reproduce and assigns it a P3 priority.
2020-12-15: I ask if there's an update.
2021-02-02: I ask if there's an update.
2021-02-03: Bugcrowd raise a blocker on the issue, asking the vendor to respond.
2021-02-17: I ask for permission to disclose.
2021-03-09: In the absence of any response from the vendor since 2020-08-19, I violate Bugcrowd disclosure policies and unilaterally disclose.

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Making hibernation work under Linux Lockdown

Posted by Matthew Garrett on February 21, 2021 08:37 AM
Linux draws a distinction between code running in kernel (kernel space) and applications running in userland (user space). This is enforced at the hardware level - in x86-speak[1], kernel space code runs in ring 0 and user space code runs in ring 3[2]. If you're running in ring 3 and you attempt to touch memory that's only accessible in ring 0, the hardware will raise a fault. No matter how privileged your ring 3 code, you don't get to touch ring 0.

Kind of. In theory. Traditionally this wasn't well enforced. At the most basic level, since root can load kernel modules, you could just build a kernel module that performed any kernel modifications you wanted and then have root load it. Technically user space code wasn't modifying kernel space code, but the difference was pretty semantic rather than useful. But it got worse - root could also map memory ranges belonging to PCI devices[3], and if the device could perform DMA you could just ask the device to overwrite bits of the kernel[4]. Or root could modify special CPU registers ("Model Specific Registers", or MSRs) that alter CPU behaviour via the /dev/msr interface, and compromise the kernel boundary that way.

It turns out that there were a number of ways root was effectively equivalent to ring 0, and the boundary was more about reliability (ie, a process running as root that ends up misbehaving should still only be able to crash itself rather than taking down the kernel with it) than security. After all, if you were root you could just replace the on-disk kernel with a backdoored one and reboot. Going deeper, you could replace the bootloader with one that automatically injected backdoors into a legitimate kernel image. We didn't have any way to prevent this sort of thing, so attempting to harden the root/kernel boundary wasn't especially interesting.

In 2012 Microsoft started requiring vendors ship systems with UEFI Secure Boot, a firmware feature that allowed[5] systems to refuse to boot anything without an appropriate signature. This not only enabled the creation of a system that drew a strong boundary between root and kernel, it arguably required one - what's the point of restricting what the firmware will stick in ring 0 if root can just throw more code in there afterwards? What ended up as the Lockdown Linux Security Module provides the tooling for this, blocking userspace interfaces that can be used to modify the kernel and enforcing that any modules have a trusted signature.

But that comes at something of a cost. Most of the features that Lockdown blocks are fairly niche, so the direct impact of having it enabled is small. Except that it also blocks hibernation[6], and it turns out some people were using that. The obvious question is "what does hibernation have to do with keeping root out of kernel space", and the answer is a little convoluted and is tied into how Linux implements hibernation. Basically, Linux saves system state into the swap partition and modifies the header to indicate that there's a hibernation image there instead of swap. On the next boot, the kernel sees the header indicating that it's a hibernation image, copies the contents of the swap partition back into RAM, and then jumps back into the old kernel code. What ensures that the hibernation image was actually written out by the kernel? Absolutely nothing, which means a motivated attacker with root access could turn off swap, write a hibernation image to the swap partition themselves, and then reboot. The kernel would happily resume into the attacker's image, giving the attacker control over what gets copied back into kernel space.

This is annoying, because normally when we think about attacks on swap we mitigate it by requiring an encrypted swap partition. But in this case, our attacker is root, and so already has access to the plaintext version of the swap partition. Disk encryption doesn't save us here. We need some way to verify that the hibernation image was written out by the kernel, not by root. And thankfully we have some tools for that.

Trusted Platform Modules (TPMs) are cryptographic coprocessors[7] capable of doing things like generating encryption keys and then encrypting things with them. You can ask a TPM to encrypt something with a key that's tied to that specific TPM - the OS has no access to the decryption key, and nor does any other TPM. So we can have the kernel generate an encryption key, encrypt part of the hibernation image with it, and then have the TPM encrypt it. We store the encrypted copy of the key in the hibernation image as well. On resume, the kernel reads the encrypted copy of the key, passes it to the TPM, gets the decrypted copy back and is able to verify the hibernation image.

That's great! Except root can do exactly the same thing. This tells us the hibernation image was generated on this machine, but doesn't tell us that it was done by the kernel. We need some way to be able to differentiate between keys that were generated in kernel and ones that were generated in userland. TPMs have the concept of "localities" (effectively privilege levels) that would be perfect for this. Userland is only able to access locality 0, so the kernel could simply use locality 1 to encrypt the key. Unfortunately, despite trying pretty hard, I've been unable to get localities to work. The motherboard chipset on my test machines simply doesn't forward any accesses to the TPM unless they're for locality 0. I needed another approach.

TPMs have a set of Platform Configuration Registers (PCRs), intended for keeping a record of system state. The OS isn't able to modify the PCRs directly. Instead, the OS provides a cryptographic hash of some material to the TPM. The TPM takes the existing PCR value, appends the new hash to that, and then stores the hash of the combination in the PCR - a process called "extension". This means that the new value of the TPM depends not only on the value of the new data, it depends on the previous value of the PCR - and, in turn, that previous value depended on its previous value, and so on. The only way to get to a specific PCR value is to either (a) break the hash algorithm, or (b) perform exactly the same sequence of writes. On system reset the PCRs go back to a known value, and the entire process starts again.

Some PCRs are different. PCR 23, for example, can be reset back to its original value without resetting the system. We can make use of that. The first thing we need to do is to prevent userland from being able to reset or extend PCR 23 itself. All TPM accesses go through the kernel, so this is a simple matter of parsing the write before it's sent to the TPM and returning an error if it's a sensitive command that would touch PCR 23. We now know that any change in PCR 23's state will be restricted to the kernel.

When we encrypt material with the TPM, we can ask it to record the PCR state. This is given back to us as metadata accompanying the encrypted secret. Along with the metadata is an additional signature created by the TPM, which can be used to prove that the metadata is both legitimate and associated with this specific encrypted data. In our case, that means we know what the value of PCR 23 was when we encrypted the key. That means that if we simply extend PCR 23 with a known value in-kernel before encrypting our key, we can look at the value of PCR 23 in the metadata. If it matches, the key was encrypted by the kernel - userland can create its own key, but it has no way to extend PCR 23 to the appropriate value first. We now know that the key was generated by the kernel.

But what if the attacker is able to gain access to the encrypted key? Let's say a kernel bug is hit that prevents hibernation from resuming, and you boot back up without wiping the hibernation image. Root can then read the key from the partition, ask the TPM to decrypt it, and then use that to create a new hibernation image. We probably want to prevent that as well. Fortunately, when you ask the TPM to encrypt something, you can ask that the TPM only decrypt it if the PCRs have specific values. "Sealing" material to the TPM in this way allows you to block decryption if the system isn't in the desired state. So, we define a policy that says that PCR 23 must have the same value at resume as it did on hibernation. On resume, the kernel resets PCR 23, extends it to the same value it did during hibernation, and then attempts to decrypt the key. Afterwards, it resets PCR 23 back to the initial value. Even if an attacker gains access to the encrypted copy of the key, the TPM will refuse to decrypt it.

And that's what this patchset implements. There's one fairly significant flaw at the moment, which is simply that an attacker can just reboot into an older kernel that doesn't implement the PCR 23 blocking and set up state by hand. Fortunately, this can be avoided using another aspect of the boot process. When you boot something via UEFI Secure Boot, the signing key used to verify the booted code is measured into PCR 7 by the system firmware. In the Linux world, the Shim bootloader then measures any additional keys that are used. By either using a new key to tag kernels that have support for the PCR 23 restrictions, or by embedding some additional metadata in the kernel that indicates the presence of this feature and measuring that, we can have a PCR 7 value that verifies that the PCR 23 restrictions are present. We then seal the key to PCR 7 as well as PCR 23, and if an attacker boots into a kernel that doesn't have this feature the PCR 7 value will be different and the TPM will refuse to decrypt the secret.

While there's a whole bunch of complexity here, the process should be entirely transparent to the user. The current implementation requires a TPM 2, and I'm not certain whether TPM 1.2 provides all the features necessary to do this properly - if so, extending it shouldn't be hard, but also all systems shipped in the past few years should have a TPM 2, so that's going to depend on whether there's sufficient interest to justify the work. And we're also at the early days of review, so there's always the risk that I've missed something obvious and there are terrible holes in this. And, well, given that it took almost 8 years to get the Lockdown patchset into mainline, let's not assume that I'm good at landing security code.

[1] Other architectures use different terminology here, such as "supervisor" and "user" mode, but it's broadly equivalent
[2] In theory rings 1 and 2 would allow you to run drivers with privileges somewhere between full kernel access and userland applications, but in reality we just don't talk about them in polite company
[3] This is how graphics worked in Linux before kernel modesetting turned up. XFree86 would just map your GPU's registers into userland and poke them directly. This was not a huge win for stability
[4] IOMMUs can help you here, by restricting the memory PCI devices can DMA to or from. The kernel then gets to allocate ranges for device buffers and configure the IOMMU such that the device can't DMA to anything else. Except that region of memory may still contain sensitive material such as function pointers, and attacks like this can still cause you problems as a result.
[5] This describes why I'm using "allowed" rather than "required" here
[6] Saving the system state to disk and powering down the platform entirely - significantly slower than suspending the system while keeping state in RAM, but also resilient against the system losing power.
[7] With some handwaving around "coprocessor". TPMs can't be part of the OS or the system firmware, but they don't technically need to be an independent component. Intel have a TPM implementation that runs on the Management Engine, a separate processor built into the motherboard chipset. AMD have one that runs on the Platform Security Processor, a small ARM core built into their CPU. Various ARM implementations run a TPM in Trustzone, a special CPU mode that (in theory) is able to access resources that are entirely blocked off from anything running in the OS, kernel or otherwise.

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A pre-supplied "custom" keyboard layout for X11

Posted by Peter Hutterer on February 18, 2021 01:57 AM

Last year I wrote about how to create a user-specific XKB layout, followed by a post explaining that this won't work in X. But there's a pandemic going on, which is presumably the only reason people haven't all switched to Wayland yet. So it was time to figure out a workaround for those still running X.

This Merge Request (scheduled for xkeyboard-config 2.33) adds a "custom" layout to the evdev.xml and base.xml files. These XML files are parsed by the various GUI tools to display the selection of available layouts. An entry in there will thus show up in the GUI tool.

Our rulesets, i.e. the files that convert a layout/variant configuration into the components to actually load already have wildcard matching [1]. So the custom layout will resolve to the symbols/custom file in your XKB data dir - usually /usr/share/X11/xkb/symbols/custom.

This file is not provided by xkeyboard-config. It can be created by the user though and whatever configuration is in there will be the "custom" keyboard layout. Because xkeyboard-config does not supply this file, it will not get overwritten on update.

From XKB's POV it is just another layout and it thus uses the same syntax. For example, to override the +/* key on the German keyboard layout with a key that produces a/b/c/d on the various Shift/Alt combinations, use this:


default
xkb_symbols "basic" {
include "de(basic)"
key <AD12> { [ a, b, c, d ] };
};
This example includes the "basic" section from the symbols/de file (i.e. the default German layout), then overrides the 12th alphanumeric key from left in the 4th row from bottom (D) with the given symbols. I'll leave it up to the reader to come up with a less useful example.

There are a few drawbacks:

  • If the file is missing and the user selects the custom layout, the results are... undefined. For run-time configuration like GNOME it doesn't really matter - the layout compilation fails and you end up with the one the device already had (i.e. the default one built into X, usually the US layout).
  • If the file is missing and the custom layout is selected in the xorg.conf, the results are... undefined. I tested it and ended up with the US layout but that seems more by accident than design. My recommendation is to not do that.
  • No variants are available in the XML files, so the only accessible section is the one marked default.
  • If a commandline tool uses a variant of custom, the GUI will not reflect this. If the GUI goes boom, that's a bug in the GUI.

So overall, it's a hack[2]. But it's a hack that fixes real user issues and given we're talking about X, I doubt anyone notices another hack anyway.

[1] If you don't care about GUIs, setxkbmap -layout custom -variant foobar has been possible for years.
[2] Sticking with the UNIX principle, it's a hack that fixes the issue at hand, is badly integrated, and weird to configure.

fwupd 1.5.6

Posted by Richard Hughes on February 16, 2021 12:16 PM

Today I released fwupd 1.5.6 which the usual smattering of new features and bugfixes. These are some of the more interesting ones:

With the help of a lot of people we added support for quite a bit of new hardware. The slightly odd GD32VF103 as found in the Longan Nano is now supported, and more of the DFU ST devices with huge amounts of flash. The former should enable us to support the Pinecil device soon and the latter will be a nice vendor announcement in the future. We’ve also added support for RMI PS2 devices as found in some newer Lenovo ThinkPads, the Starlabs LabTop L4 and the new System76 Keyboard. We’ve refactored the udev and usb backends into self contained modules, allowing someone else to contribute new bluetooth peripheral functionality in the future. There are more than a dozen teams of people all working on fwupd features at the moment. Exciting times!

One problem that has been reported was that downloads from the datacenter in the US were really slow from China, specifically because the firewall was deliberately dropping packets. I assume compressed firmware looks quite a lot like a large encrypted message from a firewalls’ point of view, and thus it was only letting through ~20% of the traffic. All non-export controlled public firmware is now also mirrored onto the IPFS, and we experimentally fall back to peer-to-peer downloads where the HTTP download failed. You can prefer IPFS downloads using fwupdmgr --ipfs update although you need to have a running ipfs daemon on your local computer. If this works well for you, let me know and we might add support for downloading metadata in the future too.

We’ve fully integrated the fwupd CI with oss-fuzz, a popular fuzzing service from Google. Generating horribly corrupt firmware files has found a few little memory leaks, files that cause fwupd to spin in a loop and even the odd crash. It was a lot of work to build each fuzzer into a small static binary using a 16.04-based container but it was well worth all the hard work. All new PRs will run the same fuzzers checking for regressions which also means new plugins now also have to implement building new firmware (so the test payload can be a few tens of bytes, not 32kB), rather than just parsing it.

On some Lenovo hardware there’s a “useful” feature called Boot Order Lock that means whatever the OS adds as a BootXXXX entry the old bootlist gets restored on next boot. This breaks firmware updates using fwupdx64.efi and until we can detect BOL from a kernel interface we also check if our EFI entry has been deleted by the firmware on next boot and give the user a more helpful message than just “it failed”. Also, on some Lenovo hardware we’re limiting the number of UEFI updates to be deployed on one reboot as they appear to have slightly quirky capsule coalesce behavior. In the same vein we’re also checking the system clock is set approximately correct (as in, not set to before 2020…) so we can tell the user to check the clock on the machine rather than just failing with a obscure certificate error.

Now there are systems that can be switched to coreboot (and back to EDK2 again) we’ve polished up the “switch-branch” feature. We’re also checking both BIOSWE and BLE before identifying systems that can be supported. We’re also including the lockdown status in uploaded UEFI reports and added SBAT metadata to the fwupd EFI binary, which will be required for future versions of shim and grub – so for distro fwupd binaries the packager will need to set meson build options like -Defi_sbat_distro_id=. There are examples in the fwupd source tree.

Auto-updating XKB for new kernel keycodes

Posted by Peter Hutterer on January 22, 2021 12:58 AM

Your XKB keymap contains two important parts. One is the mapping from the hardware scancode to some internal representation, for example:

  <AB10> = 61;  

Which basically means Alphanumeric key in row B (from bottom), 10th key from the left. In other words: the /? key on a US keyboard.

The second part is mapping that internal representation to a keysym, for example:

  key <AB10> {        [     slash,    question        ]       }; 

This is the actual layout mapping - once in place this key really produces a slash or question mark (on level2, i.e. when Shift is down).

This two-part approach exists so either part can be swapped without affecting the other. Swap the second part to an exclamation mark and paragraph symbol and you have the French version of this key, swap it to dash/underscore and you have the German version of the key - all without having to change the keycode.

Back in the golden days of everyone-does-what-they-feel-like, keyboard manufacturers (presumably happily so) changed the key codes and we needed model-specific keycodes in XKB. The XkbModel configuration is a leftover from these trying times.

The Linux kernel's evdev API has largely done away with this. It provides a standardised set of keycodes, defined in linux/input-event-codes.h, and ensures, with the help of udev [0], that all keyboards actually conform to that. An evdev XKB keycode is a simple "kernel keycode + 8" [1] and that applies to all keyboards. On top of that, the kernel uses semantic definitions for the keys as they'd be in the US layout. KEY_Q is the key that would, behold!, produce a Q. Or an A in the French layout because they just have to be different, don't they? Either way, with evdev the Xkb Model configuration largely points to nothing and only wastes a few cycles with string parsing.

The second part, the keysym mapping, uses two approaches. One is to use a named #define like the "slash", "question" outlined above (see X11/keysymdef.h for the defines). The other is to use unicode directly like this example from  the Devangari layout:

  key <AB10> { [ U092f, U095f, slash, question ] };

As you can see, mix and match is available too. Using Unicode code points of course makes the layouts less immediately readable but on the other hand we don't need to #define the whole of Unicode. So from a maintenance perspective it's a win.

However, there's a third type of key that we care about: functional keys. Those are the multimedia (historically: "internet") keys that most devices have these days. Volume up, touchpad on/off, cycle display connectors, etc. Those keys are special in that they don't have a Unicode representation and they are always mapped to the same fixed functionality. Even Dvorak users want their volume keys to do what it says on the key.

Because they have no Unicode code points, those keys are defined, historically, in XF86keysyms.h:

  #define XF86XK_MonBrightnessUp    0x1008FF02  /* Monitor/panel brightness */

And mapping a key like this looks like this [2]:

  key <I21>   {       [ XF86Calculator        ] };

The only drawback: every key needs to be added manually. This has been done for some, but not for others. And some keys were added with different names than what the kernel uses [3].

So we're in this weird situation where we have a flexible keymap system  but the kernel already tells us what a key does anyway and we don't want to change that. Virtually all keys added in the last decade or so falls into that group of keys, but to actually make use of them requires a #define in xorgproto and an update to the keycodes and symbols in xkeyboard-config. That again introduces discrepancies and we end up in the situation where we're at right now: some keys don't work until someone files a bug, and then the users still need to wait for several components to be released and those releases trickle into the distributions.

10 years ago would've been a good time to make this more efficient. The situation wasn't that urgent then, most of the kernel keycodes added are >255 which means they cannot be used in X anyway. [4] The second best time to do it is now. What we need is basically a pass-through from kernel code to symbol and that's currently sitting in various MRs:

- xkeyboard-config can generate the keycodes/evdev file based on the list of kernel keycodes, so all kernel keycodes are mapped to internal representations by default

- xorgproto has reserved a range within the XF86 keysym reserved range for pass-through mappings, i.e. any KEY_FOO define from the kernel is mapped to XF86XK_Foo with a specific value [5]. The #define format is fixed so it can be parsed.

- xkeyboard-config parses theses XF86 keysyms and sets up a keysym mapping in the default keymap.

This is semi-automatic, i.e. there are helper scripts that detect changes and notify us, hooked into the CI, but the actual work must be done manually. These keysyms immediately become set-in-stone API so we don't want some unsupervised script to go wild on them.

There's a huge backlog of keys to be added (dating to kernels pre-v3.18) and I'll go through them one-by-one over the next weeks to make sure they're correct. But eventually they'll be done and we have a full keymap for all kernel keys to be immediately available in the XKB layout.

The last part of all of this is a calendar reminder for me to do this after every new kernel release. Let's hope this crucial part isn't the first to fail.

[0] 60-keyboard.hwdb has a mere ~1800 lines!
[1] Historical reasons, you don't want to know. *jedi wave*
[2] the XK_ part of the key name is dropped, implementation detail.
[3] This can also happen when a kernel define is renamed/aliased but we cannot easily do so for this header.
[4] X has an 8 bit keycode limit and that won't change until someone develops XKB2 with support for 32-bit keycodes, i.e. never.

[5] The actual value is an implementation detail and no client must care


Parsing HID Unit Items

Posted by Peter Hutterer on January 13, 2021 11:28 AM

This post explains how to parse the HID Unit Global Item as explained by the HID Specification, page 37. The table there is quite confusing and it took me a while to fully understand it (Benjamin Tissoires was really the one who cracked it). I couldn't find any better explanation online which means either I'm incredibly dense and everyone's figured it out or no-one has posted a better explanation. On the off-chance it's the latter [1], here are the instructions on how to parse this item.

We know a HID Report Descriptor consists of a number of items that describe the content of each HID Report (read: an event from a device). These Items include things like Logical Minimum/Maximum for axis ranges, etc. A HID Unit item specifies the physical unit to apply. For example, a Report Descriptor may specify that X and Y axes are in mm which can be quite useful for all the obvious reasons.

Like most HID items, a HID Unit Item consists of a one-byte item tag and 1, 2 or 4 byte payload. The Unit item in the Report Descriptor itself has the binary value 0110 01nn where the nn is either 1, 2, or 3 indicating 1, 2 or 4 bytes of payload, respectively. That's standard HID.

The payload is divided into nibbles (4-bit units) and goes from LSB to MSB. The lowest-order 4 bits (first byte & 0xf) define the unit System to apply: one of SI Linear, SI Rotation, English Linear or English Rotation (well, or None/Reserved). The rest of the nibbles are in this order: "length", "mass", "time", "temperature", "current", "luminous intensity". In something resembling code this means:


system = value & 0xf
length_exponent = (value & 0xf0) >> 4
mass_exponent = (value & 0xf00) >> 8
time_exponent = (value & 0xf000) >> 12
...
The System defines which unit is used for length (e.g. SILinear means length is in cm). The actual value of each nibble is the exponent for the unit in use [2]. In something resembling code:

switch (system)
case SILinear:
print("length is in cm^{length_exponent}");
break;
case SIRotation:
print("length is in rad^{length_exponent}");
break;
case EnglishLinear:
print("length is in in^{length_exponent}");
break;
case EnglishRotation:
print("length is in deg^{length_exponent}");
break;
case None:
case Reserved"
print("boo!");
break;

For example, the value 0x321 means "SI Linear" (0x1) so the remaining nibbles represent, in ascending nibble order: Centimeters, Grams, Seconds, Kelvin, Ampere, Candela. The length nibble has a value of 0x2 so it's square cm, the mass nibble has a value of 0x3 so it is cubic grams (well, it's just an example, so...). This means that any report containing this item comes in cm²g³. As a more realistic example: 0xF011 would be cm/s.

If we changed the lowest nibble to English Rotation (0x4), i.e. our value is now 0x324, the units represent: Degrees, Slug, Seconds, F, Ampere, Candela [3]. The length nibble 0x2 means square degrees, the mass nibble is cubic slugs. As a more realistic example, 0xF014 would be degrees/s.

Any nibble with value 0 means the unit isn't in use, so the example from the spec with value 0x00F0D121 is SI linear, units cm² g s⁻³ A⁻¹, which is... Voltage! Of course you knew that and totally didn't have to double-check with wikipedia.

Because bits are expensive and the base units are of course either too big or too small or otherwise not quite right, HID also provides a Unit Exponent item. The Unit Exponent item (a separate item to Unit in the Report Descriptor) then describes the exponent to be applied to the actual value in the report. For example, a Unit Eponent of -3 means 10⁻³ to be applied to the value. If the report descriptor specifies an item of Unit 0x00F0D121 (i.e. V) and Unit Exponent -3, the value of this item is mV (milliVolt), Unit Exponent of 3 would be kV (kiloVolt).

Now, in hindsight all this is pretty obvious and maybe even sensible. It'd have been nice if the spec would've explained it a bit clearer but then I would have nothing to write about, so I guess overall I call it a draw.

[1] This whole adventure was started because there's a touchpad out there that measures touch pressure in radians, so at least one other person out there struggled with the docs...
[2] The nibble value is twos complement (i.e. it's a signed 4-bit integer). Values 0x1-0x7 are exponents 1 to 7, values 0x8-0xf are exponents -8 to -1.
[3] English Linear should've trolled everyone and use Centimetres instead of Centimeters in SI Linear.

Unlocking LUKS2 volumes with TPM2, FIDO2, PKCS#11 Security Hardware on systemd 248

Posted by Lennart Poettering on January 12, 2021 11:00 PM

TL;DR: It's now easy to unlock your LUKS2 volume with a FIDO2 security token (e.g. YubiKey, Nitrokey FIDO2, AuthenTrend ATKey.Pro). And TPM2 unlocking is easy now too.

Blogging is a lot of work, and a lot less fun than hacking. I mostly focus on the latter because of that, but from time to time I guess stuff is just too interesting to not be blogged about. Hence here, finally, another blog story about exciting new features in systemd.

With the upcoming systemd v248 the systemd-cryptsetup component of systemd (which is responsible for assembling encrypted volumes during boot) gained direct support for unlocking encrypted storage with three types of security hardware:

  1. Unlocking with FIDO2 security tokens (well, at least with those which implement the hmac-secret extension; most do). i.e. your YubiKeys (series 5 and above), Nitrokey FIDO2, AuthenTrend ATKey.Pro and such.

  2. Unlocking with TPM2 security chips (pretty ubiquitous on non-budget PCs/laptops/…)

  3. Unlocking with PKCS#11 security tokens, i.e. your smartcards and older YubiKeys (the ones that implement PIV). (Strictly speaking this was supported on older systemd already, but was a lot more "manual".)

For completeness' sake, let's keep in mind that the component also allows unlocking with these more traditional mechanisms:

  1. Unlocking interactively with a user-entered passphrase (i.e. the way most people probably already deploy it, supported since about forever)

  2. Unlocking via key file on disk (optionally on removable media plugged in at boot), supported since forever.

  3. Unlocking via a key acquired through trivial AF_UNIX/SOCK_STREAM socket IPC. (Also new in v248)

  4. Unlocking via recovery keys. These are pretty much the same thing as a regular passphrase (and in fact can be entered wherever a passphrase is requested) — the main difference being that they are always generated by the computer, and thus have guaranteed high entropy, typically higher than user-chosen passphrases. They are generated in a way they are easy to type, in many cases even if the local key map is misconfigured. (Also new in v248)

In this blog story, let's focus on the first three items, i.e. those that talk to specific types of hardware for implementing unlocking.

To make working with security tokens and TPM2 easy, a new, small tool was added to the systemd tool set: systemd-cryptenroll. It's only purpose is to make it easy to enroll your security token/chip of choice into an encrypted volume. It works with any LUKS2 volume, and embeds a tiny bit of meta-information into the LUKS2 header with parameters necessary for the unlock operation.

Unlocking with FIDO2

So, let's see how this fits together in the FIDO2 case. Most likely this is what you want to use if you have one of these fancy FIDO2 tokens (which need to implement the hmac-secret extension, as mentioned). Let's say you already have your LUKS2 volume set up, and previously unlocked it with a simple passphrase. Plug in your token, and run:

# systemd-cryptenroll --fido2-device=auto /dev/sda5

(Replace /dev/sda5 with the underlying block device of your volume).

This will enroll the key as an additional way to unlock the volume, and embeds all necessary information for it in the LUKS2 volume header. Before we can unlock the volume with this at boot, we need to allow FIDO2 unlocking via /etc/crypttab. For that, find the right entry for your volume in that file, and edit it like so:

myvolume /dev/sda5 - fido2-device=auto

Replace myvolume and /dev/sda5 with the right volume name, and underlying device of course. Key here is the fido2-device=auto option you need to add to the fourth column in the file. It tells systemd-cryptsetup to use the FIDO2 metadata now embedded in the LUKS2 header, wait for the FIDO2 token to be plugged in at boot (utilizing systemd-udevd, …) and unlock the volume with it.

And that's it already. Easy-peasy, no?

Note that all of this doesn't modify the FIDO2 token itself in any way. Moreover you can enroll the same token in as many volumes as you like. Since all enrollment information is stored in the LUKS2 header (and not on the token) there are no bounds on any of this. (OK, well, admittedly, there's a cap on LUKS2 key slots per volume, i.e. you can't enroll more than a bunch of keys per volume.)

Unlocking with PKCS#11

Let's now have a closer look how the same works with a PKCS#11 compatible security token or smartcard. For this to work, you need a device that can store an RSA key pair. I figure most security tokens/smartcards that implement PIV qualify. How you actually get the keys onto the device might differ though. Here's how you do this for any YubiKey that implements the PIV feature:

# ykman piv reset
# ykman piv generate-key -a RSA2048 9d pubkey.pem
# ykman piv generate-certificate --subject "Knobelei" 9d pubkey.pem
# rm pubkey.pem

(This chain of commands erases what was stored in PIV feature of your token before, be careful!)

For tokens/smartcards from other vendors a different series of commands might work. Once you have a key pair on it, you can enroll it with a LUKS2 volume like so:

# systemd-cryptenroll --pkcs11-token-uri=auto /dev/sda5

Just like the same command's invocation in the FIDO2 case this enrolls the security token as an additional way to unlock the volume, any passphrases you already have enrolled remain enrolled.

For the PKCS#11 case you need to edit your /etc/crypttab entry like this:

myvolume /dev/sda5 - pkcs11-uri=auto

If you have a security token that implements both PKCS#11 PIV and FIDO2 I'd probably enroll it as FIDO2 device, given it's the more contemporary, future-proof standard. Moreover, it requires no special preparation in order to get an RSA key onto the device: FIDO2 keys typically just work.

Unlocking with TPM2

Most modern (non-budget) PC hardware (and other kind of hardware too) nowadays comes with a TPM2 security chip. In many ways a TPM2 chip is a smartcard that is soldered onto the mainboard of your system. Unlike your usual USB-connected security tokens you thus cannot remove them from your PC, which means they address quite a different security scenario: they aren't immediately comparable to a physical key you can take with you that unlocks some door, but they are a key you leave at the door, but that refuses to be turned by anyone but you.

Even though this sounds a lot weaker than the FIDO2/PKCS#11 model TPM2 still bring benefits for securing your systems: because the cryptographic key material stored in TPM2 devices cannot be extracted (at least that's the theory), if you bind your hard disk encryption to it, it means attackers cannot just copy your disk and analyze it offline — they always need access to the TPM2 chip too to have a chance to acquire the necessary cryptographic keys. Thus, they can still steal your whole PC and analyze it, but they cannot just copy the disk without you noticing and analyze the copy.

Moreover, you can bind the ability to unlock the harddisk to specific software versions: for example you could say that only your trusted Fedora Linux can unlock the device, but not any arbitrary OS some hacker might boot from a USB stick they plugged in. Thus, if you trust your OS vendor, you can entrust storage unlocking to the vendor's OS together with your TPM2 device, and thus can be reasonably sure intruders cannot decrypt your data unless they both hack your OS vendor and steal/break your TPM2 chip.

Here's how you enroll your LUKS2 volume with your TPM2 chip:

# systemd-cryptenroll --tpm2-device=auto --tpm2-pcrs=7 /dev/sda5

This looks almost as straightforward as the two earlier sytemd-cryptenroll command lines — if it wasn't for the --tpm2-pcrs= part. With that option you can specify to which TPM2 PCRs you want to bind the enrollment. TPM2 PCRs are a set of (typically 24) hash values that every TPM2 equipped system at boot calculates from all the software that is invoked during the boot sequence, in a secure, unfakable way (this is called "measurement"). If you bind unlocking to a specific value of a specific PCR you thus require the system has to follow the same sequence of software at boot to re-acquire the disk encryption key. Sounds complex? Well, that's because it is.

For now, let's see how we have to modify your /etc/crypttab to unlock via TPM2:

myvolume /dev/sda5 - tpm2-device=auto

This part is easy again: the tpm2-device= option is what tells systemd-cryptsetup to use the TPM2 metadata from the LUKS2 header and to wait for the TPM2 device to show up.

Bonus: Recovery Key Enrollment

FIDO2, PKCS#11 and TPM2 security tokens and chips pair well with recovery keys: since you don't need to type in your password everyday anymore it makes sense to get rid of it, and instead enroll a high-entropy recovery key you then print out or scan off screen and store a safe, physical location. i.e. forget about good ol' passphrase-based unlocking, go for FIDO2 plus recovery key instead! Here's how you do it:

# systemd-cryptenroll --recovery-key /dev/sda5

This will generate a key, enroll it in the LUKS2 volume, show it to you on screen and generate a QR code you may scan off screen if you like. The key has highest entropy, and can be entered wherever you can enter a passphrase. Because of that you don't have to modify /etc/crypttab to make the recovery key work.

Future

There's still plenty room for further improvement in all of this. In particular for the TPM2 case: what the text above doesn't really mention is that binding your encrypted volume unlocking to specific software versions (i.e. kernel + initrd + OS versions) actually sucks hard: if you naively update your system to newer versions you might lose access to your TPM2 enrolled keys (which isn't terrible, after all you did enroll a recovery key — right? — which you then can use to regain access). To solve this some more integration with distributions would be necessary: whenever they upgrade the system they'd have to make sure to enroll the TPM2 again — with the PCR hashes matching the new version. And whenever they remove an old version of the system they need to remove the old TPM2 enrollment. Alternatively TPM2 also knows a concept of signed PCR hash values. In this mode the distro could just ship a set of PCR signatures which would unlock the TPM2 keys. (But quite frankly I don't really see the point: whether you drop in a signature file on each system update, or enroll a new set of PCR hashes in the LUKS2 header doesn't make much of a difference). Either way, to make TPM2 enrollment smooth some more integration work with your distribution's system update mechanisms need to happen. And yes, because of this OS updating complexity the example above — where I referenced your trusty Fedora Linux — doesn't actually work IRL (yet? hopefully…). Nothing updates the enrollment automatically after you initially enrolled it, hence after the first kernel/initrd update you have to manually re-enroll things again, and again, and again … after every update.

The TPM2 could also be used for other kinds of key policies, we might look into adding later too. For example, Windows uses TPM2 stuff to allow short (4 digits or so) "PINs" for unlocking the harddisk, i.e. kind of a low-entropy password you type in. The reason this is reasonably safe is that in this case the PIN is passed to the TPM2 which enforces that not more than some limited amount of unlock attempts may be made within some time frame, and that after too many attempts the PIN is invalidated altogether. Thus making dictionary attacks harder (which would normally be easier given the short length of the PINs).

Postscript

(BTW: Yubico sent me two YubiKeys for testing, Nitrokey a Nitrokey FIDO2, and AuthenTrend three ATKey.Pro tokens, thank you! — That's why you see all those references to YubiKey/Nitrokey/AuthenTrend devices in the text above: it's the hardware I had to test this with. That said, I also tested the FIDO2 stuff with a SoloKey I bought, where it also worked fine. And yes, you!, other vendors!, who might be reading this, please send me your security tokens for free, too, and I might test things with them as well. No promises though. And I am not going to give them back, if you do, sorry. ;-))

fwupd 1.5.5

Posted by Richard Hughes on January 11, 2021 10:39 AM

I’ve just released fwupd 1.5.5 with the following new features:

  • Add a plugin to update PixArt RF devices; the hardware this enables we’ll announce in a few weeks hopefully
  • Add new hardware to use the elantp (for TouchPads) and rts54hid (for USB Hubs) plugins
  • Allow specifying more than one VendorID for a device, which allows ATA devices to use the OUI-assigned vendor if set
  • Detect the AMD TSME encryption state for HSI-4 — use fwupdmgr security --force to help test
  • Detect the AMI PK test key is not installed for HSI-1 — a failure here is very serious
  • As usual, this release fixes quite a few bugs too:

  • Fix flashing a fingerprint reader that is in use; in theory the window to hit this is vanishingly small, but on some hardware we ask the user to authorise the request using the very device that we’re trying to update…
  • Fix several critical warnings when parsing invalid firmware, found using hongfuzz, warming my office on these cold winter days
  • Fix updating DFU devices that use DNLOAD_BUSY which fixes fwupd on some other future hardware support
  • Ignore the legacy UEFI OVMF dummy GUID so that we can test the dbx updates using qemu on older releases like RHEL
  • Make libfwupd more thread safe to fix a crash in gnome-software — many thanks to Philip Withnall for explaining a lot of the GMainContext threading complexities to me
  • We now never show unprintable chars from invalid firmware in the logs — as a result of fuzzing insane things the logs would often be full of gobbledygook, but no longer
  • I’m now building 1.5.5 into Fedora 33 and Fedora 32, packages should appear soon.

    Firefox – we’re finally getting HW acceleration on Linux

    Posted by Martin Stransky on January 10, 2021 08:40 PM
    <figure class="wp-block-image size-large"><figcaption>A first image from original WebRender article. Published three years ago.</figcaption></figure>

    Firefox 84.0 is a big milestone for Firefox Linux development as it comes with HW acceleration by default for some Linux users. Stock Mozilla Firefox 84.0 enables WebRender (HW accelerated backend) for Gnome/X.org and Gnome/Wayland will be supported in Firefox 85.0. Fedora is bit ahead and enables WebRender for Gnome/Wayland in Firefox 84.0 too.

    WebRender by default is restricted to AMD/Intel graphics cards as NVIDIA is known for various issues – both proprietary and Noveau drivers.

    And why it’s enabled in Gnome only for now? For instance KDE is also a popular desktop environment. I think it’s because Gnome utilizes HW acceleration so when Gnome works on your box there’s assumption that Firefox will work too. KDE provides choices how to disable/restrict HW acceleration setup (for instance it supports disabled screen compositing) and it’s more difficult to cover various scenarios.

    Another excluded group are XWayland users. It means you have Wayland as a desktop compositor but for some reasons you use X11 emulation layer and run Firefox as X11 application. It’s a valid scenario, Firefox with Wayland backend still suffers from some annoying bug, mostly related to popup windows.

    But don’t worry, Mozilla folks are going to bring WebRender to the most Linux users on various desktops and graphics. Jan created a brief Linux WebRender state overview. And you can help with it! Please check if you have WebRender enabled and eventually try to enable it. Test various web pages, video playback, WebGL and report your experience. You can use comments below or drop me a mail at stransky@redhat.com.

    <figure class="wp-block-image size-large"></figure>