1
2
3
4
5
6
7
8
9
10
11
12
13
14
15
16
17
18
19
20
21
22
23
24
25
26
27
28
29
30
31
32
33
34
35
36
37
38
39
40
41
42
43
44
45
46
47
48
49
50
51
52
53
54
55
56
57
58
59
60
61
62
63
64
65
66
67
68
69
70
71
72
73
74
75
76
77
78
79
80
81
82
83
84
85
86
87
88
89
90
91
92
93
94
95
96
97
98
99
100
101
102
103
104
105
106
107
108
109
110
111
112
113
114
115
116
117
118
119
120
121
122
123
124
125
126
127
128
129
130
131
132
133
134
135
136
137
138
139
140
141
142
143
144
145
146
147
148
149
150
151
152
153
154
155
156
157
158
159
160
161
162
163
164
165
166
167
168
169
170
171
172
173
174
175
176
177
178
179
180
181
182
183
184
185
186
187
188
189
190
191
192
193
194
195
196
197
198
199
200
201
202
203
204
205
206
207
208
209
210
211
212
213
214
215
216
217
218
219
220
221
222
223
224
225
226
227
228
229
230
231
232
233
234
235
236
237
238
239
240
241
242
243
244
245
246
247
248
249
250
251
252
253
254
255
256
257
258
259
260
261
262
263
264
265
266
267
268
269
270
271
272
273
274
275
276
277
278
279
280
281
282
283
284
285
286
287
288
289
290
291
292
293
294
295
296
297
298
299
300
301
302
303
304
305
306
307
308
309
310
311
312
313
314
315
316
317
318
319
320
321
322
323
324
325
326
327
328
329
330
331
332
333
334
335
336
337
338
339
340
341
342
343
344
345
346
347
348
349
350
351
352
353
354
355
356
357
358
359
360
361
362
363
364
365
366
367
368
369
370
371
372
373
374
375
376
377
378
379
380
381
382
383
384
385
386
387
388
389
390
391
392
393
394
395
396
397
398
399
400
401
402
403
404
405
406
407
408
409
410
411
412
413
414
415
416
417
418
419
420
421
422
423
424
425
426
427
428
429
430
431
432
433
434
435
436
437
438
439
440
441
442
443
444
445
446
447
448
449
450
451
452
453
454
455
456
457
458
459
460
461
462
463
464
465
466
467
468
469
470
471
472
473
474
475
476
477
478
479
480
481
482
483
484
485
486
487
488
489
490
491
492
493
494
495
496
497
498
499
500
501
502
503
504
505
506
507
508
509
510
511
512
513
514
515
516
517
518
519
520
521
522
523
524
525
526
527
528
529
530
531
532
533
534
535
536
537
538
539
540
541
542
543
544
545
546
547
548
549
550
551
552
553
554
555
556
557
558
559
560
561
562
563
564
565
566
567
568
569
570
571
572
573
574
575
576
577
578
579
580
581
582
583
584
585
586
587
588
589
590
591
592
593
594
595
596
597
598
599
600
601
602
603
604
605
606
607
608
609
610
611
612
613
614
615
616
617
618
619
620
621
622
623
624
625
626
627
628
629
630
631
632
633
634
635
636
637
638
639
640
641
642
643
644
645
646
647
648
649
650
651
652
653
654
655
656
657
658
659
660
661
662
663
664
665
666
667
668
669
670
671
672
673
674
675
676
677
678
679
680
681
682
683
684
685
686
687
688
689
690
691
692
693
694
695
696
697
698
699
700
701
702
703
704
705
706
707
708
709
710
711
712
713
714
715
716
717
718
719
720
721
722
723
724
725
726
727
728
729
730
731
732
733
734
735
736
737
738
739
740
741
742
743
744
745
746
747
748
749
750
751
752
753
754
755
756
757
758
759
760
761
762
763
764
765
766
767
768
769
770
771
772
773
774
775
776
777
778
779
780
781
782
783
784
785
786
787
788
789
790
791
792
793
794
795
796
797
798
799
800
801
802
803
804
805
806
807
808
809
810
811
812
813
814
815
816
817
818
819
820
821
822
823
824
825
826
827
828
829
830
831
832
833
834
835
836
837
838
839
840
841
842
843
844
845
846
847
848
849
850
851
852
853
854
855
856
857
858
859
860
861
862
863
864
865
866
867
868
869
870
871
872
873
874
875
876
877
878
879
880
881
882
883
884
885
886
887
888
889
890
891
892
893
894
895
896
897
898
899
900
901
902
903
904
905
906
907
908
909
910
911
912
913
914
915
916
917
918
919
920
921
922
923
924
925
926
927
928
929
930
931
932
933
934
935
936
937
938
939
940
941
942
943
944
945
946
947
948
949
950
951
952
953
954
955
956
957
958
959
960
961
962
963
964
965
966
967
968
969
970
971
972
973
974
975
976
977
978
979
980
981
982
983
984
985
986
987
988
989
990
991
992
993
994
995
996
997
998
999
1000
1001
1002
1003
1004
1005
1006
1007
1008
1009
1010
1011
1012
1013
1014
1015
1016
1017
1018
1019
1020
1021
1022
1023
1024
1025
1026
1027
1028
1029
1030
1031
1032
1033
1034
1035
1036
1037
1038
1039
1040
1041
1042
1043
1044
1045
1046
1047
1048
1049
1050
1051
1052
1053
1054
1055
1056
1057
1058
1059
1060
1061
1062
1063
1064
1065
1066
1067
1068
1069
1070
1071
1072
1073
1074
1075
1076
1077
1078
1079
1080
1081
1082
1083
1084
1085
1086
1087
1088
1089
1090
1091
1092
1093
1094
1095
1096
1097
1098
1099
1100
1101
1102
1103
1104
1105
1106
1107
1108
1109
1110
1111
1112
1113
1114
1115
1116
1117
1118
1119
1120
1121
1122
1123
1124
1125
1126
1127
1128
1129
1130
1131
1132
1133
1134
1135
1136
1137
1138
1139
1140
1141
1142
1143
1144
1145
1146
1147
1148
1149
1150
1151
1152
1153
1154
1155
1156
1157
1158
1159
1160
1161
1162
1163
1164
1165
1166
1167
1168
1169
1170
1171
1172
1173
1174
1175
1176
1177
1178
1179
1180
1181
1182
1183
1184
1185
1186
1187
1188
1189
1190
1191
1192
1193
1194
1195
1196
1197
1198
1199
1200
1201
1202
1203
1204
1205
1206
1207
1208
1209
1210
1211
1212
1213
1214
1215
1216
1217
1218
1219
1220
1221
1222
1223
1224
1225
1226
1227
1228
1229
1230
1231
1232
1233
1234
1235
1236
1237
1238
1239
1240
1241
1242
1243
1244
1245
1246
1247
1248
1249
1250
1251
1252
1253
1254
1255
1256
1257
1258
1259
1260
1261
1262
1263
1264
1265
1266
1267
1268
1269
1270
1271
1272
1273
1274
1275
1276
1277
1278
1279
1280
1281
1282
1283
1284
1285
1286
1287
1288
1289
1290
1291
1292
1293
1294
1295
1296
1297
1298
1299
1300
1301
1302
1303
1304
1305
1306
1307
1308
1309
1310
1311
1312
1313
1314
1315
1316
1317
1318
1319
1320
1321
1322
1323
1324
1325
1326
1327
1328
1329
1330
1331
1332
1333
1334
1335
1336
1337
1338
1339
1340
1341
1342
1343
1344
1345
1346
1347
1348
1349
1350
1351
1352
1353
1354
1355
1356
1357
1358
1359
1360
1361
1362
1363
1364
1365
1366
1367
1368
1369
1370
1371
1372
1373
1374
1375
1376
1377
1378
1379
1380
1381
1382
1383
1384
1385
1386
1387
1388
1389
1390
1391
1392
1393
1394
1395
1396
1397
1398
1399
1400
1401
1402
1403
1404
1405
1406
1407
1408
1409
1410
1411
1412
1413
1414
1415
1416
1417
1418
1419
1420
1421
1422
1423
1424
1425
1426
1427
1428
1429
1430
1431
1432
1433
1434
1435
1436
1437
1438
1439
1440
1441
1442
1443
1444
1445
1446
1447
1448
1449
1450
1451
1452
1453
1454
1455
1456
1457
1458
1459
1460
1461
1462
1463
1464
1465
1466
1467
1468
1469
1470
1471
1472
1473
1474
1475
1476
1477
1478
1479
1480
1481
1482
1483
1484
1485
1486
1487
1488
1489
1490
1491
1492
1493
1494
1495
1496
1497
1498
1499
1500
1501
1502
1503
1504
1505
1506
1507
1508
1509
1510
1511
1512
1513
1514
1515
1516
1517
1518
1519
1520
1521
1522
1523
1524
1525
1526
1527
1528
1529
1530
1531
1532
1533
1534
1535
1536
1537
1538
1539
1540
1541
1542
1543
1544
1545
1546
1547
1548
1549
1550
1551
1552
1553
1554
1555
1556
1557
1558
1559
1560
1561
1562
1563
1564
1565
1566
1567
1568
1569
1570
1571
1572
1573
1574
1575
1576
1577
1578
1579
1580
1581
1582
1583
1584
1585
1586
1587
1588
1589
1590
1591
1592
1593
1594
1595
1596
1597
1598
1599
1600
1601
1602
1603
1604
1605
1606
1607
1608
1609
1610
1611
1612
1613
1614
1615
1616
1617
1618
1619
1620
1621
1622
1623
1624
1625
1626
1627
1628
1629
1630
1631
1632
1633
1634
1635
1636
1637
1638
1639
1640
1641
1642
1643
1644
1645
1646
1647
1648
1649
1650
1651
1652
1653
1654
1655
1656
1657
1658
1659
1660
1661
1662
1663
1664
1665
1666
1667
1668
1669
1670
1671
1672
1673
1674
1675
1676
1677
1678
1679
1680
1681
1682
1683
1684
1685
1686
1687
1688
1689
1690
1691
1692
1693
1694
1695
1696
1697
1698
1699
1700
1701
1702
1703
1704
1705
1706
1707
1708
1709
1710
1711
1712
1713
1714
1715
1716
1717
1718
1719
1720
1721
1722
1723
1724
1725
1726
1727
1728
1729
1730
1731
1732
1733
1734
1735
1736
1737
1738
1739
1740
1741
1742
1743
1744
1745
1746
1747
1748
1749
1750
1751
1752
1753
1754
1755
1756
1757
1758
1759
1760
1761
1762
1763
1764
1765
1766
1767
1768
1769
1770
1771
1772
1773
1774
1775
1776
1777
1778
1779
1780
1781
1782
1783
1784
1785
1786
1787
1788
1789
1790
1791
1792
1793
1794
1795
1796
1797
1798
1799
1800
1801
1802
1803
1804
1805
1806
1807
1808
1809
1810
1811
1812
1813
1814
1815
1816
1817
1818
1819
1820
1821
1822
1823
1824
1825
1826
1827
1828
1829
1830
1831
1832
1833
1834
1835
1836
1837
1838
1839
1840
1841
1842
1843
1844
1845
1846
1847
1848
1849
1850
1851
1852
1853
1854
1855
1856
1857
1858
1859
1860
1861
1862
1863
1864
1865
1866
1867
1868
1869
1870
1871
1872
1873
1874
1875
1876
1877
1878
1879
1880
1881
1882
1883
1884
1885
1886
1887
1888
1889
1890
1891
1892
1893
1894
1895
1896
1897
1898
1899
1900
1901
1902
1903
1904
1905
1906
1907
1908
1909
1910
1911
1912
1913
1914
1915
1916
1917
1918
1919
1920
1921
1922
1923
1924
1925
1926
1927
1928
1929
1930
1931
1932
1933
1934
1935
1936
1937
1938
1939
1940
1941
1942
1943
1944
1945
1946
1947
1948
1949
1950
1951
1952
1953
1954
1955
1956
1957
1958
1959
1960
1961
1962
1963
1964
1965
1966
1967
1968
1969
1970
1971
1972
1973
1974
1975
1976
1977
1978
1979
1980
1981
1982
1983
1984
1985
1986
1987
1988
1989
1990
1991
1992
1993
1994
1995
1996
1997
1998
1999
2000
2001
2002
2003
2004
2005
2006
2007
2008
2009
2010
2011
2012
2013
2014
2015
2016
2017
2018
2019
2020
2021
2022
2023
2024
2025
2026
2027
2028
2029
2030
2031
2032
2033
2034
2035
2036
2037
2038
2039
2040
2041
2042
2043
2044
2045
2046
2047
2048
2049
2050
2051
2052
2053
2054
2055
2056
2057
2058
2059
2060
2061
2062
2063
2064
2065
2066
2067
2068
2069
2070
2071
2072
2073
2074
2075
2076
2077
2078
2079
2080
2081
2082
2083
2084
2085
2086
2087
2088
2089
2090
2091
2092
2093
2094
2095
2096
2097
2098
2099
2100
2101
2102
2103
2104
2105
2106
2107
2108
2109
2110
2111
2112
2113
2114
2115
2116
2117
2118
2119
2120
2121
2122
2123
2124
2125
2126
2127
2128
2129
2130
2131
2132
2133
2134
2135
2136
2137
2138
2139
2140
2141
2142
2143
2144
2145
2146
2147
2148
2149
2150
2151
2152
2153
2154
2155
2156
2157
2158
2159
2160
2161
2162
2163
2164
2165
2166
2167
2168
2169
2170
2171
2172
2173
2174
2175
2176
2177
2178
2179
2180
2181
2182
2183
2184
2185
2186
2187
2188
2189
2190
2191
2192
2193
2194
2195
2196
2197
2198
2199
2200
2201
2202
2203
2204
2205
2206
2207
2208
2209
2210
2211
2212
2213
2214
2215
2216
2217
2218
2219
2220
2221
2222
2223
2224
2225
2226
2227
2228
2229
2230
2231
2232
2233
2234
2235
2236
2237
2238
2239
2240
2241
2242
2243
2244
2245
2246
2247
2248
2249
2250
2251
2252
2253
2254
2255
2256
2257
2258
2259
2260
2261
2262
2263
2264
2265
2266
2267
2268
2269
2270
2271
2272
2273
2274
2275
2276
2277
2278
2279
2280
2281
2282
2283
2284
2285
2286
2287
2288
2289
2290
2291
2292
2293
2294
2295
2296
2297
2298
2299
2300
2301
2302
2303
2304
2305
2306
2307
2308
2309
2310
2311
2312
2313
2314
2315
2316
2317
2318
2319
2320
2321
2322
2323
2324
2325
2326
2327
2328
2329
2330
2331
2332
2333
2334
2335
2336
2337
2338
2339
2340
2341
2342
2343
2344
2345
2346
2347
2348
2349
2350
2351
2352
2353
2354
2355
2356
2357
2358
2359
2360
2361
2362
2363
2364
2365
2366
2367
2368
2369
2370
2371
2372
2373
2374
2375
2376
2377
2378
2379
2380
2381
2382
2383
2384
2385
2386
2387
2388
2389
2390
2391
2392
2393
2394
2395
2396
2397
2398
2399
2400
2401
2402
2403
2404
2405
2406
2407
2408
2409
2410
2411
2412
2413
2414
2415
2416
2417
2418
2419
2420
2421
2422
2423
2424
2425
2426
2427
2428
2429
2430
2431
2432
2433
2434
2435
2436
2437
2438
2439
2440
2441
2442
2443
2444
2445
2446
2447
2448
2449
2450
2451
2452
2453
2454
2455
2456
2457
2458
2459
2460
2461
2462
2463
2464
2465
2466
2467
2468
2469
2470
2471
2472
2473
2474
2475
2476
2477
2478
2479
2480
2481
2482
2483
2484
2485
2486
2487
2488
2489
2490
2491
2492
2493
2494
2495
2496
2497
2498
2499
2500
2501
2502
2503
2504
2505
2506
2507
2508
2509
2510
2511
2512
2513
2514
2515
2516
2517
2518
2519
2520
2521
2522
2523
2524
2525
2526
2527
2528
2529
2530
2531
2532
2533
2534
2535
2536
2537
2538
2539
2540
2541
2542
2543
2544
2545
2546
2547
2548
2549
2550
2551
2552
2553
2554
2555
2556
2557
2558
2559
2560
2561
2562
2563
2564
2565
2566
2567
2568
2569
2570
2571
2572
2573
2574
2575
2576
2577
2578
2579
2580
2581
2582
2583
2584
2585
2586
2587
2588
2589
2590
2591
2592
2593
2594
2595
2596
2597
2598
2599
2600
2601
2602
2603
2604
2605
2606
2607
2608
2609
2610
2611
2612
2613
2614
2615
2616
2617
2618
2619
2620
2621
2622
2623
2624
2625
2626
2627
2628
2629
2630
2631
2632
2633
2634
2635
2636
2637
2638
2639
2640
2641
2642
2643
2644
2645
2646
2647
2648
2649
2650
2651
2652
2653
2654
2655
2656
2657
2658
2659
2660
2661
2662
2663
2664
2665
2666
2667
2668
2669
2670
2671
2672
2673
2674
2675
2676
2677
2678
2679
2680
2681
2682
2683
2684
2685
2686
2687
2688
2689
2690
2691
2692
2693
2694
2695
2696
2697
2698
2699
2700
2701
2702
2703
2704
2705
2706
2707
2708
2709
2710
2711
2712
2713
2714
2715
2716
2717
2718
2719
2720
2721
2722
2723
2724
2725
2726
2727
2728
2729
2730
2731
2732
2733
2734
2735
2736
2737
2738
2739
2740
2741
2742
2743
2744
2745
2746
2747
2748
2749
2750
2751
2752
2753
2754
2755
2756
2757
2758
2759
2760
2761
2762
2763
2764
2765
2766
2767
2768
2769
2770
2771
2772
2773
2774
2775
2776
2777
2778
2779
2780
2781
2782
2783
2784
2785
2786
2787
2788
2789
2790
2791
2792
2793
2794
2795
2796
2797
2798
2799
2800
2801
2802
2803
2804
2805
2806
2807
2808
2809
2810
2811
2812
2813
2814
2815
2816
2817
2818
2819
2820
2821
2822
2823
2824
2825
2826
2827
2828
2829
2830
2831
2832
2833
2834
2835
2836
2837
2838
2839
2840
2841
2842
2843
2844
2845
2846
2847
2848
2849
2850
2851
2852
2853
2854
2855
2856
2857
2858
2859
2860
2861
2862
2863
2864
2865
2866
2867
2868
2869
2870
2871
2872
2873
2874
2875
2876
2877
2878
2879
2880
2881
2882
2883
2884
2885
2886
2887
2888
2889
2890
2891
2892
2893
2894
2895
2896
2897
2898
2899
2900
2901
2902
2903
2904
2905
2906
2907
2908
2909
2910
2911
2912
2913
2914
2915
2916
2917
2918
2919
2920
2921
2922
2923
2924
2925
2926
2927
2928
2929
2930
2931
2932
2933
2934
2935
2936
2937
2938
2939
2940
2941
2942
2943
2944
2945
2946
2947
2948
2949
2950
2951
2952
2953
2954
2955
2956
2957
2958
2959
2960
2961
2962
2963
2964
2965
2966
2967
2968
2969
2970
2971
2972
2973
2974
2975
2976
2977
2978
2979
2980
2981
2982
2983
2984
2985
2986
2987
2988
2989
2990
2991
2992
2993
2994
2995
2996
2997
2998
2999
3000
3001
3002
3003
3004
3005
3006
3007
3008
3009
3010
3011
3012
3013
3014
3015
3016
3017
3018
3019
3020
3021
3022
3023
3024
3025
3026
3027
3028
3029
3030
3031
3032
3033
3034
3035
3036
3037
3038
3039
3040
3041
3042
3043
3044
3045
3046
3047
3048
3049
3050
3051
3052
3053
3054
3055
3056
3057
3058
3059
3060
3061
3062
3063
3064
3065
3066
3067
3068
3069
3070
3071
3072
3073
3074
3075
3076
3077
3078
3079
3080
3081
3082
3083
3084
3085
3086
3087
3088
3089
3090
3091
3092
3093
3094
3095
3096
3097
3098
3099
3100
3101
3102
3103
3104
|
<!DOCTYPE HTML PUBLIC "-//W3C//DTD HTML 4.01 Transitional//EN"
"http://www.w3.org/TR/html4/loose.dtd">
<html>
<head><title>A Tour Through RCU's Requirements [LWN.net]</title>
<meta HTTP-EQUIV="Content-Type" CONTENT="text/html; charset=utf-8">
<h1>A Tour Through RCU's Requirements</h1>
<p>Copyright IBM Corporation, 2015</p>
<p>Author: Paul E. McKenney</p>
<p><i>The initial version of this document appeared in the
<a href="https://lwn.net/">LWN</a> articles
<a href="https://lwn.net/Articles/652156/">here</a>,
<a href="https://lwn.net/Articles/652677/">here</a>, and
<a href="https://lwn.net/Articles/653326/">here</a>.</i></p>
<h2>Introduction</h2>
<p>
Read-copy update (RCU) is a synchronization mechanism that is often
used as a replacement for reader-writer locking.
RCU is unusual in that updaters do not block readers,
which means that RCU's read-side primitives can be exceedingly fast
and scalable.
In addition, updaters can make useful forward progress concurrently
with readers.
However, all this concurrency between RCU readers and updaters does raise
the question of exactly what RCU readers are doing, which in turn
raises the question of exactly what RCU's requirements are.
<p>
This document therefore summarizes RCU's requirements, and can be thought
of as an informal, high-level specification for RCU.
It is important to understand that RCU's specification is primarily
empirical in nature;
in fact, I learned about many of these requirements the hard way.
This situation might cause some consternation, however, not only
has this learning process been a lot of fun, but it has also been
a great privilege to work with so many people willing to apply
technologies in interesting new ways.
<p>
All that aside, here are the categories of currently known RCU requirements:
</p>
<ol>
<li> <a href="#Fundamental Requirements">
Fundamental Requirements</a>
<li> <a href="#Fundamental Non-Requirements">Fundamental Non-Requirements</a>
<li> <a href="#Parallelism Facts of Life">
Parallelism Facts of Life</a>
<li> <a href="#Quality-of-Implementation Requirements">
Quality-of-Implementation Requirements</a>
<li> <a href="#Linux Kernel Complications">
Linux Kernel Complications</a>
<li> <a href="#Software-Engineering Requirements">
Software-Engineering Requirements</a>
<li> <a href="#Other RCU Flavors">
Other RCU Flavors</a>
<li> <a href="#Possible Future Changes">
Possible Future Changes</a>
</ol>
<p>
This is followed by a <a href="#Summary">summary</a>,
however, the answers to each quick quiz immediately follows the quiz.
Select the big white space with your mouse to see the answer.
<h2><a name="Fundamental Requirements">Fundamental Requirements</a></h2>
<p>
RCU's fundamental requirements are the closest thing RCU has to hard
mathematical requirements.
These are:
<ol>
<li> <a href="#Grace-Period Guarantee">
Grace-Period Guarantee</a>
<li> <a href="#Publish-Subscribe Guarantee">
Publish-Subscribe Guarantee</a>
<li> <a href="#Memory-Barrier Guarantees">
Memory-Barrier Guarantees</a>
<li> <a href="#RCU Primitives Guaranteed to Execute Unconditionally">
RCU Primitives Guaranteed to Execute Unconditionally</a>
<li> <a href="#Guaranteed Read-to-Write Upgrade">
Guaranteed Read-to-Write Upgrade</a>
</ol>
<h3><a name="Grace-Period Guarantee">Grace-Period Guarantee</a></h3>
<p>
RCU's grace-period guarantee is unusual in being premeditated:
Jack Slingwine and I had this guarantee firmly in mind when we started
work on RCU (then called “rclock”) in the early 1990s.
That said, the past two decades of experience with RCU have produced
a much more detailed understanding of this guarantee.
<p>
RCU's grace-period guarantee allows updaters to wait for the completion
of all pre-existing RCU read-side critical sections.
An RCU read-side critical section
begins with the marker <tt>rcu_read_lock()</tt> and ends with
the marker <tt>rcu_read_unlock()</tt>.
These markers may be nested, and RCU treats a nested set as one
big RCU read-side critical section.
Production-quality implementations of <tt>rcu_read_lock()</tt> and
<tt>rcu_read_unlock()</tt> are extremely lightweight, and in
fact have exactly zero overhead in Linux kernels built for production
use with <tt>CONFIG_PREEMPT=n</tt>.
<p>
This guarantee allows ordering to be enforced with extremely low
overhead to readers, for example:
<blockquote>
<pre>
1 int x, y;
2
3 void thread0(void)
4 {
5 rcu_read_lock();
6 r1 = READ_ONCE(x);
7 r2 = READ_ONCE(y);
8 rcu_read_unlock();
9 }
10
11 void thread1(void)
12 {
13 WRITE_ONCE(x, 1);
14 synchronize_rcu();
15 WRITE_ONCE(y, 1);
16 }
</pre>
</blockquote>
<p>
Because the <tt>synchronize_rcu()</tt> on line 14 waits for
all pre-existing readers, any instance of <tt>thread0()</tt> that
loads a value of zero from <tt>x</tt> must complete before
<tt>thread1()</tt> stores to <tt>y</tt>, so that instance must
also load a value of zero from <tt>y</tt>.
Similarly, any instance of <tt>thread0()</tt> that loads a value of
one from <tt>y</tt> must have started after the
<tt>synchronize_rcu()</tt> started, and must therefore also load
a value of one from <tt>x</tt>.
Therefore, the outcome:
<blockquote>
<pre>
(r1 == 0 && r2 == 1)
</pre>
</blockquote>
cannot happen.
<table>
<tr><th> </th></tr>
<tr><th align="left">Quick Quiz:</th></tr>
<tr><td>
Wait a minute!
You said that updaters can make useful forward progress concurrently
with readers, but pre-existing readers will block
<tt>synchronize_rcu()</tt>!!!
Just who are you trying to fool???
</td></tr>
<tr><th align="left">Answer:</th></tr>
<tr><td bgcolor="#ffffff"><font color="ffffff">
First, if updaters do not wish to be blocked by readers, they can use
<tt>call_rcu()</tt> or <tt>kfree_rcu()</tt>, which will
be discussed later.
Second, even when using <tt>synchronize_rcu()</tt>, the other
update-side code does run concurrently with readers, whether
pre-existing or not.
</font></td></tr>
<tr><td> </td></tr>
</table>
<p>
This scenario resembles one of the first uses of RCU in
<a href="https://en.wikipedia.org/wiki/DYNIX">DYNIX/ptx</a>,
which managed a distributed lock manager's transition into
a state suitable for handling recovery from node failure,
more or less as follows:
<blockquote>
<pre>
1 #define STATE_NORMAL 0
2 #define STATE_WANT_RECOVERY 1
3 #define STATE_RECOVERING 2
4 #define STATE_WANT_NORMAL 3
5
6 int state = STATE_NORMAL;
7
8 void do_something_dlm(void)
9 {
10 int state_snap;
11
12 rcu_read_lock();
13 state_snap = READ_ONCE(state);
14 if (state_snap == STATE_NORMAL)
15 do_something();
16 else
17 do_something_carefully();
18 rcu_read_unlock();
19 }
20
21 void start_recovery(void)
22 {
23 WRITE_ONCE(state, STATE_WANT_RECOVERY);
24 synchronize_rcu();
25 WRITE_ONCE(state, STATE_RECOVERING);
26 recovery();
27 WRITE_ONCE(state, STATE_WANT_NORMAL);
28 synchronize_rcu();
29 WRITE_ONCE(state, STATE_NORMAL);
30 }
</pre>
</blockquote>
<p>
The RCU read-side critical section in <tt>do_something_dlm()</tt>
works with the <tt>synchronize_rcu()</tt> in <tt>start_recovery()</tt>
to guarantee that <tt>do_something()</tt> never runs concurrently
with <tt>recovery()</tt>, but with little or no synchronization
overhead in <tt>do_something_dlm()</tt>.
<table>
<tr><th> </th></tr>
<tr><th align="left">Quick Quiz:</th></tr>
<tr><td>
Why is the <tt>synchronize_rcu()</tt> on line 28 needed?
</td></tr>
<tr><th align="left">Answer:</th></tr>
<tr><td bgcolor="#ffffff"><font color="ffffff">
Without that extra grace period, memory reordering could result in
<tt>do_something_dlm()</tt> executing <tt>do_something()</tt>
concurrently with the last bits of <tt>recovery()</tt>.
</font></td></tr>
<tr><td> </td></tr>
</table>
<p>
In order to avoid fatal problems such as deadlocks,
an RCU read-side critical section must not contain calls to
<tt>synchronize_rcu()</tt>.
Similarly, an RCU read-side critical section must not
contain anything that waits, directly or indirectly, on completion of
an invocation of <tt>synchronize_rcu()</tt>.
<p>
Although RCU's grace-period guarantee is useful in and of itself, with
<a href="https://lwn.net/Articles/573497/">quite a few use cases</a>,
it would be good to be able to use RCU to coordinate read-side
access to linked data structures.
For this, the grace-period guarantee is not sufficient, as can
be seen in function <tt>add_gp_buggy()</tt> below.
We will look at the reader's code later, but in the meantime, just think of
the reader as locklessly picking up the <tt>gp</tt> pointer,
and, if the value loaded is non-<tt>NULL</tt>, locklessly accessing the
<tt>->a</tt> and <tt>->b</tt> fields.
<blockquote>
<pre>
1 bool add_gp_buggy(int a, int b)
2 {
3 p = kmalloc(sizeof(*p), GFP_KERNEL);
4 if (!p)
5 return -ENOMEM;
6 spin_lock(&gp_lock);
7 if (rcu_access_pointer(gp)) {
8 spin_unlock(&gp_lock);
9 return false;
10 }
11 p->a = a;
12 p->b = a;
13 gp = p; /* ORDERING BUG */
14 spin_unlock(&gp_lock);
15 return true;
16 }
</pre>
</blockquote>
<p>
The problem is that both the compiler and weakly ordered CPUs are within
their rights to reorder this code as follows:
<blockquote>
<pre>
1 bool add_gp_buggy_optimized(int a, int b)
2 {
3 p = kmalloc(sizeof(*p), GFP_KERNEL);
4 if (!p)
5 return -ENOMEM;
6 spin_lock(&gp_lock);
7 if (rcu_access_pointer(gp)) {
8 spin_unlock(&gp_lock);
9 return false;
10 }
<b>11 gp = p; /* ORDERING BUG */
12 p->a = a;
13 p->b = a;</b>
14 spin_unlock(&gp_lock);
15 return true;
16 }
</pre>
</blockquote>
<p>
If an RCU reader fetches <tt>gp</tt> just after
<tt>add_gp_buggy_optimized</tt> executes line 11,
it will see garbage in the <tt>->a</tt> and <tt>->b</tt>
fields.
And this is but one of many ways in which compiler and hardware optimizations
could cause trouble.
Therefore, we clearly need some way to prevent the compiler and the CPU from
reordering in this manner, which brings us to the publish-subscribe
guarantee discussed in the next section.
<h3><a name="Publish-Subscribe Guarantee">Publish/Subscribe Guarantee</a></h3>
<p>
RCU's publish-subscribe guarantee allows data to be inserted
into a linked data structure without disrupting RCU readers.
The updater uses <tt>rcu_assign_pointer()</tt> to insert the
new data, and readers use <tt>rcu_dereference()</tt> to
access data, whether new or old.
The following shows an example of insertion:
<blockquote>
<pre>
1 bool add_gp(int a, int b)
2 {
3 p = kmalloc(sizeof(*p), GFP_KERNEL);
4 if (!p)
5 return -ENOMEM;
6 spin_lock(&gp_lock);
7 if (rcu_access_pointer(gp)) {
8 spin_unlock(&gp_lock);
9 return false;
10 }
11 p->a = a;
12 p->b = a;
13 rcu_assign_pointer(gp, p);
14 spin_unlock(&gp_lock);
15 return true;
16 }
</pre>
</blockquote>
<p>
The <tt>rcu_assign_pointer()</tt> on line 13 is conceptually
equivalent to a simple assignment statement, but also guarantees
that its assignment will
happen after the two assignments in lines 11 and 12,
similar to the C11 <tt>memory_order_release</tt> store operation.
It also prevents any number of “interesting” compiler
optimizations, for example, the use of <tt>gp</tt> as a scratch
location immediately preceding the assignment.
<table>
<tr><th> </th></tr>
<tr><th align="left">Quick Quiz:</th></tr>
<tr><td>
But <tt>rcu_assign_pointer()</tt> does nothing to prevent the
two assignments to <tt>p->a</tt> and <tt>p->b</tt>
from being reordered.
Can't that also cause problems?
</td></tr>
<tr><th align="left">Answer:</th></tr>
<tr><td bgcolor="#ffffff"><font color="ffffff">
No, it cannot.
The readers cannot see either of these two fields until
the assignment to <tt>gp</tt>, by which time both fields are
fully initialized.
So reordering the assignments
to <tt>p->a</tt> and <tt>p->b</tt> cannot possibly
cause any problems.
</font></td></tr>
<tr><td> </td></tr>
</table>
<p>
It is tempting to assume that the reader need not do anything special
to control its accesses to the RCU-protected data,
as shown in <tt>do_something_gp_buggy()</tt> below:
<blockquote>
<pre>
1 bool do_something_gp_buggy(void)
2 {
3 rcu_read_lock();
4 p = gp; /* OPTIMIZATIONS GALORE!!! */
5 if (p) {
6 do_something(p->a, p->b);
7 rcu_read_unlock();
8 return true;
9 }
10 rcu_read_unlock();
11 return false;
12 }
</pre>
</blockquote>
<p>
However, this temptation must be resisted because there are a
surprisingly large number of ways that the compiler
(to say nothing of
<a href="https://h71000.www7.hp.com/wizard/wiz_2637.html">DEC Alpha CPUs</a>)
can trip this code up.
For but one example, if the compiler were short of registers, it
might choose to refetch from <tt>gp</tt> rather than keeping
a separate copy in <tt>p</tt> as follows:
<blockquote>
<pre>
1 bool do_something_gp_buggy_optimized(void)
2 {
3 rcu_read_lock();
4 if (gp) { /* OPTIMIZATIONS GALORE!!! */
<b> 5 do_something(gp->a, gp->b);</b>
6 rcu_read_unlock();
7 return true;
8 }
9 rcu_read_unlock();
10 return false;
11 }
</pre>
</blockquote>
<p>
If this function ran concurrently with a series of updates that
replaced the current structure with a new one,
the fetches of <tt>gp->a</tt>
and <tt>gp->b</tt> might well come from two different structures,
which could cause serious confusion.
To prevent this (and much else besides), <tt>do_something_gp()</tt> uses
<tt>rcu_dereference()</tt> to fetch from <tt>gp</tt>:
<blockquote>
<pre>
1 bool do_something_gp(void)
2 {
3 rcu_read_lock();
4 p = rcu_dereference(gp);
5 if (p) {
6 do_something(p->a, p->b);
7 rcu_read_unlock();
8 return true;
9 }
10 rcu_read_unlock();
11 return false;
12 }
</pre>
</blockquote>
<p>
The <tt>rcu_dereference()</tt> uses volatile casts and (for DEC Alpha)
memory barriers in the Linux kernel.
Should a
<a href="http://www.rdrop.com/users/paulmck/RCU/consume.2015.07.13a.pdf">high-quality implementation of C11 <tt>memory_order_consume</tt> [PDF]</a>
ever appear, then <tt>rcu_dereference()</tt> could be implemented
as a <tt>memory_order_consume</tt> load.
Regardless of the exact implementation, a pointer fetched by
<tt>rcu_dereference()</tt> may not be used outside of the
outermost RCU read-side critical section containing that
<tt>rcu_dereference()</tt>, unless protection of
the corresponding data element has been passed from RCU to some
other synchronization mechanism, most commonly locking or
<a href="https://www.kernel.org/doc/Documentation/RCU/rcuref.txt">reference counting</a>.
<p>
In short, updaters use <tt>rcu_assign_pointer()</tt> and readers
use <tt>rcu_dereference()</tt>, and these two RCU API elements
work together to ensure that readers have a consistent view of
newly added data elements.
<p>
Of course, it is also necessary to remove elements from RCU-protected
data structures, for example, using the following process:
<ol>
<li> Remove the data element from the enclosing structure.
<li> Wait for all pre-existing RCU read-side critical sections
to complete (because only pre-existing readers can possibly have
a reference to the newly removed data element).
<li> At this point, only the updater has a reference to the
newly removed data element, so it can safely reclaim
the data element, for example, by passing it to <tt>kfree()</tt>.
</ol>
This process is implemented by <tt>remove_gp_synchronous()</tt>:
<blockquote>
<pre>
1 bool remove_gp_synchronous(void)
2 {
3 struct foo *p;
4
5 spin_lock(&gp_lock);
6 p = rcu_access_pointer(gp);
7 if (!p) {
8 spin_unlock(&gp_lock);
9 return false;
10 }
11 rcu_assign_pointer(gp, NULL);
12 spin_unlock(&gp_lock);
13 synchronize_rcu();
14 kfree(p);
15 return true;
16 }
</pre>
</blockquote>
<p>
This function is straightforward, with line 13 waiting for a grace
period before line 14 frees the old data element.
This waiting ensures that readers will reach line 7 of
<tt>do_something_gp()</tt> before the data element referenced by
<tt>p</tt> is freed.
The <tt>rcu_access_pointer()</tt> on line 6 is similar to
<tt>rcu_dereference()</tt>, except that:
<ol>
<li> The value returned by <tt>rcu_access_pointer()</tt>
cannot be dereferenced.
If you want to access the value pointed to as well as
the pointer itself, use <tt>rcu_dereference()</tt>
instead of <tt>rcu_access_pointer()</tt>.
<li> The call to <tt>rcu_access_pointer()</tt> need not be
protected.
In contrast, <tt>rcu_dereference()</tt> must either be
within an RCU read-side critical section or in a code
segment where the pointer cannot change, for example, in
code protected by the corresponding update-side lock.
</ol>
<table>
<tr><th> </th></tr>
<tr><th align="left">Quick Quiz:</th></tr>
<tr><td>
Without the <tt>rcu_dereference()</tt> or the
<tt>rcu_access_pointer()</tt>, what destructive optimizations
might the compiler make use of?
</td></tr>
<tr><th align="left">Answer:</th></tr>
<tr><td bgcolor="#ffffff"><font color="ffffff">
Let's start with what happens to <tt>do_something_gp()</tt>
if it fails to use <tt>rcu_dereference()</tt>.
It could reuse a value formerly fetched from this same pointer.
It could also fetch the pointer from <tt>gp</tt> in a byte-at-a-time
manner, resulting in <i>load tearing</i>, in turn resulting a bytewise
mash-up of two distinct pointer values.
It might even use value-speculation optimizations, where it makes
a wrong guess, but by the time it gets around to checking the
value, an update has changed the pointer to match the wrong guess.
Too bad about any dereferences that returned pre-initialization garbage
in the meantime!
</font>
<p><font color="ffffff">
For <tt>remove_gp_synchronous()</tt>, as long as all modifications
to <tt>gp</tt> are carried out while holding <tt>gp_lock</tt>,
the above optimizations are harmless.
However,
with <tt>CONFIG_SPARSE_RCU_POINTER=y</tt>,
<tt>sparse</tt> will complain if you
define <tt>gp</tt> with <tt>__rcu</tt> and then
access it without using
either <tt>rcu_access_pointer()</tt> or <tt>rcu_dereference()</tt>.
</font></td></tr>
<tr><td> </td></tr>
</table>
<p>
In short, RCU's publish-subscribe guarantee is provided by the combination
of <tt>rcu_assign_pointer()</tt> and <tt>rcu_dereference()</tt>.
This guarantee allows data elements to be safely added to RCU-protected
linked data structures without disrupting RCU readers.
This guarantee can be used in combination with the grace-period
guarantee to also allow data elements to be removed from RCU-protected
linked data structures, again without disrupting RCU readers.
<p>
This guarantee was only partially premeditated.
DYNIX/ptx used an explicit memory barrier for publication, but had nothing
resembling <tt>rcu_dereference()</tt> for subscription, nor did it
have anything resembling the <tt>smp_read_barrier_depends()</tt>
that was later subsumed into <tt>rcu_dereference()</tt>.
The need for these operations made itself known quite suddenly at a
late-1990s meeting with the DEC Alpha architects, back in the days when
DEC was still a free-standing company.
It took the Alpha architects a good hour to convince me that any sort
of barrier would ever be needed, and it then took me a good <i>two</i> hours
to convince them that their documentation did not make this point clear.
More recent work with the C and C++ standards committees have provided
much education on tricks and traps from the compiler.
In short, compilers were much less tricky in the early 1990s, but in
2015, don't even think about omitting <tt>rcu_dereference()</tt>!
<h3><a name="Memory-Barrier Guarantees">Memory-Barrier Guarantees</a></h3>
<p>
The previous section's simple linked-data-structure scenario clearly
demonstrates the need for RCU's stringent memory-ordering guarantees on
systems with more than one CPU:
<ol>
<li> Each CPU that has an RCU read-side critical section that
begins before <tt>synchronize_rcu()</tt> starts is
guaranteed to execute a full memory barrier between the time
that the RCU read-side critical section ends and the time that
<tt>synchronize_rcu()</tt> returns.
Without this guarantee, a pre-existing RCU read-side critical section
might hold a reference to the newly removed <tt>struct foo</tt>
after the <tt>kfree()</tt> on line 14 of
<tt>remove_gp_synchronous()</tt>.
<li> Each CPU that has an RCU read-side critical section that ends
after <tt>synchronize_rcu()</tt> returns is guaranteed
to execute a full memory barrier between the time that
<tt>synchronize_rcu()</tt> begins and the time that the RCU
read-side critical section begins.
Without this guarantee, a later RCU read-side critical section
running after the <tt>kfree()</tt> on line 14 of
<tt>remove_gp_synchronous()</tt> might
later run <tt>do_something_gp()</tt> and find the
newly deleted <tt>struct foo</tt>.
<li> If the task invoking <tt>synchronize_rcu()</tt> remains
on a given CPU, then that CPU is guaranteed to execute a full
memory barrier sometime during the execution of
<tt>synchronize_rcu()</tt>.
This guarantee ensures that the <tt>kfree()</tt> on
line 14 of <tt>remove_gp_synchronous()</tt> really does
execute after the removal on line 11.
<li> If the task invoking <tt>synchronize_rcu()</tt> migrates
among a group of CPUs during that invocation, then each of the
CPUs in that group is guaranteed to execute a full memory barrier
sometime during the execution of <tt>synchronize_rcu()</tt>.
This guarantee also ensures that the <tt>kfree()</tt> on
line 14 of <tt>remove_gp_synchronous()</tt> really does
execute after the removal on
line 11, but also in the case where the thread executing the
<tt>synchronize_rcu()</tt> migrates in the meantime.
</ol>
<table>
<tr><th> </th></tr>
<tr><th align="left">Quick Quiz:</th></tr>
<tr><td>
Given that multiple CPUs can start RCU read-side critical sections
at any time without any ordering whatsoever, how can RCU possibly
tell whether or not a given RCU read-side critical section starts
before a given instance of <tt>synchronize_rcu()</tt>?
</td></tr>
<tr><th align="left">Answer:</th></tr>
<tr><td bgcolor="#ffffff"><font color="ffffff">
If RCU cannot tell whether or not a given
RCU read-side critical section starts before a
given instance of <tt>synchronize_rcu()</tt>,
then it must assume that the RCU read-side critical section
started first.
In other words, a given instance of <tt>synchronize_rcu()</tt>
can avoid waiting on a given RCU read-side critical section only
if it can prove that <tt>synchronize_rcu()</tt> started first.
<p>
A related question is “When <tt>rcu_read_lock()</tt>
doesn't generate any code, why does it matter how it relates
to a grace period?”
The answer is that it is not the relationship of
<tt>rcu_read_lock()</tt> itself that is important, but rather
the relationship of the code within the enclosed RCU read-side
critical section to the code preceding and following the
grace period.
If we take this viewpoint, then a given RCU read-side critical
section begins before a given grace period when some access
preceding the grace period observes the effect of some access
within the critical section, in which case none of the accesses
within the critical section may observe the effects of any
access following the grace period.
<p>
As of late 2016, mathematical models of RCU take this
viewpoint, for example, see slides 62 and 63
of the
<a href="http://www2.rdrop.com/users/paulmck/scalability/paper/LinuxMM.2016.10.04c.LCE.pdf">2016 LinuxCon EU</a>
presentation.
</font></td></tr>
<tr><td> </td></tr>
</table>
<table>
<tr><th> </th></tr>
<tr><th align="left">Quick Quiz:</th></tr>
<tr><td>
The first and second guarantees require unbelievably strict ordering!
Are all these memory barriers <i> really</i> required?
</td></tr>
<tr><th align="left">Answer:</th></tr>
<tr><td bgcolor="#ffffff"><font color="ffffff">
Yes, they really are required.
To see why the first guarantee is required, consider the following
sequence of events:
</font>
<ol>
<li> <font color="ffffff">
CPU 1: <tt>rcu_read_lock()</tt>
</font>
<li> <font color="ffffff">
CPU 1: <tt>q = rcu_dereference(gp);
/* Very likely to return p. */</tt>
</font>
<li> <font color="ffffff">
CPU 0: <tt>list_del_rcu(p);</tt>
</font>
<li> <font color="ffffff">
CPU 0: <tt>synchronize_rcu()</tt> starts.
</font>
<li> <font color="ffffff">
CPU 1: <tt>do_something_with(q->a);
/* No smp_mb(), so might happen after kfree(). */</tt>
</font>
<li> <font color="ffffff">
CPU 1: <tt>rcu_read_unlock()</tt>
</font>
<li> <font color="ffffff">
CPU 0: <tt>synchronize_rcu()</tt> returns.
</font>
<li> <font color="ffffff">
CPU 0: <tt>kfree(p);</tt>
</font>
</ol>
<p><font color="ffffff">
Therefore, there absolutely must be a full memory barrier between the
end of the RCU read-side critical section and the end of the
grace period.
</font>
<p><font color="ffffff">
The sequence of events demonstrating the necessity of the second rule
is roughly similar:
</font>
<ol>
<li> <font color="ffffff">CPU 0: <tt>list_del_rcu(p);</tt>
</font>
<li> <font color="ffffff">CPU 0: <tt>synchronize_rcu()</tt> starts.
</font>
<li> <font color="ffffff">CPU 1: <tt>rcu_read_lock()</tt>
</font>
<li> <font color="ffffff">CPU 1: <tt>q = rcu_dereference(gp);
/* Might return p if no memory barrier. */</tt>
</font>
<li> <font color="ffffff">CPU 0: <tt>synchronize_rcu()</tt> returns.
</font>
<li> <font color="ffffff">CPU 0: <tt>kfree(p);</tt>
</font>
<li> <font color="ffffff">
CPU 1: <tt>do_something_with(q->a); /* Boom!!! */</tt>
</font>
<li> <font color="ffffff">CPU 1: <tt>rcu_read_unlock()</tt>
</font>
</ol>
<p><font color="ffffff">
And similarly, without a memory barrier between the beginning of the
grace period and the beginning of the RCU read-side critical section,
CPU 1 might end up accessing the freelist.
</font>
<p><font color="ffffff">
The “as if” rule of course applies, so that any
implementation that acts as if the appropriate memory barriers
were in place is a correct implementation.
That said, it is much easier to fool yourself into believing
that you have adhered to the as-if rule than it is to actually
adhere to it!
</font></td></tr>
<tr><td> </td></tr>
</table>
<table>
<tr><th> </th></tr>
<tr><th align="left">Quick Quiz:</th></tr>
<tr><td>
You claim that <tt>rcu_read_lock()</tt> and <tt>rcu_read_unlock()</tt>
generate absolutely no code in some kernel builds.
This means that the compiler might arbitrarily rearrange consecutive
RCU read-side critical sections.
Given such rearrangement, if a given RCU read-side critical section
is done, how can you be sure that all prior RCU read-side critical
sections are done?
Won't the compiler rearrangements make that impossible to determine?
</td></tr>
<tr><th align="left">Answer:</th></tr>
<tr><td bgcolor="#ffffff"><font color="ffffff">
In cases where <tt>rcu_read_lock()</tt> and <tt>rcu_read_unlock()</tt>
generate absolutely no code, RCU infers quiescent states only at
special locations, for example, within the scheduler.
Because calls to <tt>schedule()</tt> had better prevent calling-code
accesses to shared variables from being rearranged across the call to
<tt>schedule()</tt>, if RCU detects the end of a given RCU read-side
critical section, it will necessarily detect the end of all prior
RCU read-side critical sections, no matter how aggressively the
compiler scrambles the code.
</font>
<p><font color="ffffff">
Again, this all assumes that the compiler cannot scramble code across
calls to the scheduler, out of interrupt handlers, into the idle loop,
into user-mode code, and so on.
But if your kernel build allows that sort of scrambling, you have broken
far more than just RCU!
</font></td></tr>
<tr><td> </td></tr>
</table>
<p>
Note that these memory-barrier requirements do not replace the fundamental
RCU requirement that a grace period wait for all pre-existing readers.
On the contrary, the memory barriers called out in this section must operate in
such a way as to <i>enforce</i> this fundamental requirement.
Of course, different implementations enforce this requirement in different
ways, but enforce it they must.
<h3><a name="RCU Primitives Guaranteed to Execute Unconditionally">RCU Primitives Guaranteed to Execute Unconditionally</a></h3>
<p>
The common-case RCU primitives are unconditional.
They are invoked, they do their job, and they return, with no possibility
of error, and no need to retry.
This is a key RCU design philosophy.
<p>
However, this philosophy is pragmatic rather than pigheaded.
If someone comes up with a good justification for a particular conditional
RCU primitive, it might well be implemented and added.
After all, this guarantee was reverse-engineered, not premeditated.
The unconditional nature of the RCU primitives was initially an
accident of implementation, and later experience with synchronization
primitives with conditional primitives caused me to elevate this
accident to a guarantee.
Therefore, the justification for adding a conditional primitive to
RCU would need to be based on detailed and compelling use cases.
<h3><a name="Guaranteed Read-to-Write Upgrade">Guaranteed Read-to-Write Upgrade</a></h3>
<p>
As far as RCU is concerned, it is always possible to carry out an
update within an RCU read-side critical section.
For example, that RCU read-side critical section might search for
a given data element, and then might acquire the update-side
spinlock in order to update that element, all while remaining
in that RCU read-side critical section.
Of course, it is necessary to exit the RCU read-side critical section
before invoking <tt>synchronize_rcu()</tt>, however, this
inconvenience can be avoided through use of the
<tt>call_rcu()</tt> and <tt>kfree_rcu()</tt> API members
described later in this document.
<table>
<tr><th> </th></tr>
<tr><th align="left">Quick Quiz:</th></tr>
<tr><td>
But how does the upgrade-to-write operation exclude other readers?
</td></tr>
<tr><th align="left">Answer:</th></tr>
<tr><td bgcolor="#ffffff"><font color="ffffff">
It doesn't, just like normal RCU updates, which also do not exclude
RCU readers.
</font></td></tr>
<tr><td> </td></tr>
</table>
<p>
This guarantee allows lookup code to be shared between read-side
and update-side code, and was premeditated, appearing in the earliest
DYNIX/ptx RCU documentation.
<h2><a name="Fundamental Non-Requirements">Fundamental Non-Requirements</a></h2>
<p>
RCU provides extremely lightweight readers, and its read-side guarantees,
though quite useful, are correspondingly lightweight.
It is therefore all too easy to assume that RCU is guaranteeing more
than it really is.
Of course, the list of things that RCU does not guarantee is infinitely
long, however, the following sections list a few non-guarantees that
have caused confusion.
Except where otherwise noted, these non-guarantees were premeditated.
<ol>
<li> <a href="#Readers Impose Minimal Ordering">
Readers Impose Minimal Ordering</a>
<li> <a href="#Readers Do Not Exclude Updaters">
Readers Do Not Exclude Updaters</a>
<li> <a href="#Updaters Only Wait For Old Readers">
Updaters Only Wait For Old Readers</a>
<li> <a href="#Grace Periods Don't Partition Read-Side Critical Sections">
Grace Periods Don't Partition Read-Side Critical Sections</a>
<li> <a href="#Read-Side Critical Sections Don't Partition Grace Periods">
Read-Side Critical Sections Don't Partition Grace Periods</a>
<li> <a href="#Disabling Preemption Does Not Block Grace Periods">
Disabling Preemption Does Not Block Grace Periods</a>
</ol>
<h3><a name="Readers Impose Minimal Ordering">Readers Impose Minimal Ordering</a></h3>
<p>
Reader-side markers such as <tt>rcu_read_lock()</tt> and
<tt>rcu_read_unlock()</tt> provide absolutely no ordering guarantees
except through their interaction with the grace-period APIs such as
<tt>synchronize_rcu()</tt>.
To see this, consider the following pair of threads:
<blockquote>
<pre>
1 void thread0(void)
2 {
3 rcu_read_lock();
4 WRITE_ONCE(x, 1);
5 rcu_read_unlock();
6 rcu_read_lock();
7 WRITE_ONCE(y, 1);
8 rcu_read_unlock();
9 }
10
11 void thread1(void)
12 {
13 rcu_read_lock();
14 r1 = READ_ONCE(y);
15 rcu_read_unlock();
16 rcu_read_lock();
17 r2 = READ_ONCE(x);
18 rcu_read_unlock();
19 }
</pre>
</blockquote>
<p>
After <tt>thread0()</tt> and <tt>thread1()</tt> execute
concurrently, it is quite possible to have
<blockquote>
<pre>
(r1 == 1 && r2 == 0)
</pre>
</blockquote>
(that is, <tt>y</tt> appears to have been assigned before <tt>x</tt>),
which would not be possible if <tt>rcu_read_lock()</tt> and
<tt>rcu_read_unlock()</tt> had much in the way of ordering
properties.
But they do not, so the CPU is within its rights
to do significant reordering.
This is by design: Any significant ordering constraints would slow down
these fast-path APIs.
<table>
<tr><th> </th></tr>
<tr><th align="left">Quick Quiz:</th></tr>
<tr><td>
Can't the compiler also reorder this code?
</td></tr>
<tr><th align="left">Answer:</th></tr>
<tr><td bgcolor="#ffffff"><font color="ffffff">
No, the volatile casts in <tt>READ_ONCE()</tt> and
<tt>WRITE_ONCE()</tt> prevent the compiler from reordering in
this particular case.
</font></td></tr>
<tr><td> </td></tr>
</table>
<h3><a name="Readers Do Not Exclude Updaters">Readers Do Not Exclude Updaters</a></h3>
<p>
Neither <tt>rcu_read_lock()</tt> nor <tt>rcu_read_unlock()</tt>
exclude updates.
All they do is to prevent grace periods from ending.
The following example illustrates this:
<blockquote>
<pre>
1 void thread0(void)
2 {
3 rcu_read_lock();
4 r1 = READ_ONCE(y);
5 if (r1) {
6 do_something_with_nonzero_x();
7 r2 = READ_ONCE(x);
8 WARN_ON(!r2); /* BUG!!! */
9 }
10 rcu_read_unlock();
11 }
12
13 void thread1(void)
14 {
15 spin_lock(&my_lock);
16 WRITE_ONCE(x, 1);
17 WRITE_ONCE(y, 1);
18 spin_unlock(&my_lock);
19 }
</pre>
</blockquote>
<p>
If the <tt>thread0()</tt> function's <tt>rcu_read_lock()</tt>
excluded the <tt>thread1()</tt> function's update,
the <tt>WARN_ON()</tt> could never fire.
But the fact is that <tt>rcu_read_lock()</tt> does not exclude
much of anything aside from subsequent grace periods, of which
<tt>thread1()</tt> has none, so the
<tt>WARN_ON()</tt> can and does fire.
<h3><a name="Updaters Only Wait For Old Readers">Updaters Only Wait For Old Readers</a></h3>
<p>
It might be tempting to assume that after <tt>synchronize_rcu()</tt>
completes, there are no readers executing.
This temptation must be avoided because
new readers can start immediately after <tt>synchronize_rcu()</tt>
starts, and <tt>synchronize_rcu()</tt> is under no
obligation to wait for these new readers.
<table>
<tr><th> </th></tr>
<tr><th align="left">Quick Quiz:</th></tr>
<tr><td>
Suppose that synchronize_rcu() did wait until <i>all</i>
readers had completed instead of waiting only on
pre-existing readers.
For how long would the updater be able to rely on there
being no readers?
</td></tr>
<tr><th align="left">Answer:</th></tr>
<tr><td bgcolor="#ffffff"><font color="ffffff">
For no time at all.
Even if <tt>synchronize_rcu()</tt> were to wait until
all readers had completed, a new reader might start immediately after
<tt>synchronize_rcu()</tt> completed.
Therefore, the code following
<tt>synchronize_rcu()</tt> can <i>never</i> rely on there being
no readers.
</font></td></tr>
<tr><td> </td></tr>
</table>
<h3><a name="Grace Periods Don't Partition Read-Side Critical Sections">
Grace Periods Don't Partition Read-Side Critical Sections</a></h3>
<p>
It is tempting to assume that if any part of one RCU read-side critical
section precedes a given grace period, and if any part of another RCU
read-side critical section follows that same grace period, then all of
the first RCU read-side critical section must precede all of the second.
However, this just isn't the case: A single grace period does not
partition the set of RCU read-side critical sections.
An example of this situation can be illustrated as follows, where
<tt>x</tt>, <tt>y</tt>, and <tt>z</tt> are initially all zero:
<blockquote>
<pre>
1 void thread0(void)
2 {
3 rcu_read_lock();
4 WRITE_ONCE(a, 1);
5 WRITE_ONCE(b, 1);
6 rcu_read_unlock();
7 }
8
9 void thread1(void)
10 {
11 r1 = READ_ONCE(a);
12 synchronize_rcu();
13 WRITE_ONCE(c, 1);
14 }
15
16 void thread2(void)
17 {
18 rcu_read_lock();
19 r2 = READ_ONCE(b);
20 r3 = READ_ONCE(c);
21 rcu_read_unlock();
22 }
</pre>
</blockquote>
<p>
It turns out that the outcome:
<blockquote>
<pre>
(r1 == 1 && r2 == 0 && r3 == 1)
</pre>
</blockquote>
is entirely possible.
The following figure show how this can happen, with each circled
<tt>QS</tt> indicating the point at which RCU recorded a
<i>quiescent state</i> for each thread, that is, a state in which
RCU knows that the thread cannot be in the midst of an RCU read-side
critical section that started before the current grace period:
<p><img src="GPpartitionReaders1.svg" alt="GPpartitionReaders1.svg" width="60%"></p>
<p>
If it is necessary to partition RCU read-side critical sections in this
manner, it is necessary to use two grace periods, where the first
grace period is known to end before the second grace period starts:
<blockquote>
<pre>
1 void thread0(void)
2 {
3 rcu_read_lock();
4 WRITE_ONCE(a, 1);
5 WRITE_ONCE(b, 1);
6 rcu_read_unlock();
7 }
8
9 void thread1(void)
10 {
11 r1 = READ_ONCE(a);
12 synchronize_rcu();
13 WRITE_ONCE(c, 1);
14 }
15
16 void thread2(void)
17 {
18 r2 = READ_ONCE(c);
19 synchronize_rcu();
20 WRITE_ONCE(d, 1);
21 }
22
23 void thread3(void)
24 {
25 rcu_read_lock();
26 r3 = READ_ONCE(b);
27 r4 = READ_ONCE(d);
28 rcu_read_unlock();
29 }
</pre>
</blockquote>
<p>
Here, if <tt>(r1 == 1)</tt>, then
<tt>thread0()</tt>'s write to <tt>b</tt> must happen
before the end of <tt>thread1()</tt>'s grace period.
If in addition <tt>(r4 == 1)</tt>, then
<tt>thread3()</tt>'s read from <tt>b</tt> must happen
after the beginning of <tt>thread2()</tt>'s grace period.
If it is also the case that <tt>(r2 == 1)</tt>, then the
end of <tt>thread1()</tt>'s grace period must precede the
beginning of <tt>thread2()</tt>'s grace period.
This mean that the two RCU read-side critical sections cannot overlap,
guaranteeing that <tt>(r3 == 1)</tt>.
As a result, the outcome:
<blockquote>
<pre>
(r1 == 1 && r2 == 1 && r3 == 0 && r4 == 1)
</pre>
</blockquote>
cannot happen.
<p>
This non-requirement was also non-premeditated, but became apparent
when studying RCU's interaction with memory ordering.
<h3><a name="Read-Side Critical Sections Don't Partition Grace Periods">
Read-Side Critical Sections Don't Partition Grace Periods</a></h3>
<p>
It is also tempting to assume that if an RCU read-side critical section
happens between a pair of grace periods, then those grace periods cannot
overlap.
However, this temptation leads nowhere good, as can be illustrated by
the following, with all variables initially zero:
<blockquote>
<pre>
1 void thread0(void)
2 {
3 rcu_read_lock();
4 WRITE_ONCE(a, 1);
5 WRITE_ONCE(b, 1);
6 rcu_read_unlock();
7 }
8
9 void thread1(void)
10 {
11 r1 = READ_ONCE(a);
12 synchronize_rcu();
13 WRITE_ONCE(c, 1);
14 }
15
16 void thread2(void)
17 {
18 rcu_read_lock();
19 WRITE_ONCE(d, 1);
20 r2 = READ_ONCE(c);
21 rcu_read_unlock();
22 }
23
24 void thread3(void)
25 {
26 r3 = READ_ONCE(d);
27 synchronize_rcu();
28 WRITE_ONCE(e, 1);
29 }
30
31 void thread4(void)
32 {
33 rcu_read_lock();
34 r4 = READ_ONCE(b);
35 r5 = READ_ONCE(e);
36 rcu_read_unlock();
37 }
</pre>
</blockquote>
<p>
In this case, the outcome:
<blockquote>
<pre>
(r1 == 1 && r2 == 1 && r3 == 1 && r4 == 0 && r5 == 1)
</pre>
</blockquote>
is entirely possible, as illustrated below:
<p><img src="ReadersPartitionGP1.svg" alt="ReadersPartitionGP1.svg" width="100%"></p>
<p>
Again, an RCU read-side critical section can overlap almost all of a
given grace period, just so long as it does not overlap the entire
grace period.
As a result, an RCU read-side critical section cannot partition a pair
of RCU grace periods.
<table>
<tr><th> </th></tr>
<tr><th align="left">Quick Quiz:</th></tr>
<tr><td>
How long a sequence of grace periods, each separated by an RCU
read-side critical section, would be required to partition the RCU
read-side critical sections at the beginning and end of the chain?
</td></tr>
<tr><th align="left">Answer:</th></tr>
<tr><td bgcolor="#ffffff"><font color="ffffff">
In theory, an infinite number.
In practice, an unknown number that is sensitive to both implementation
details and timing considerations.
Therefore, even in practice, RCU users must abide by the
theoretical rather than the practical answer.
</font></td></tr>
<tr><td> </td></tr>
</table>
<h3><a name="Disabling Preemption Does Not Block Grace Periods">
Disabling Preemption Does Not Block Grace Periods</a></h3>
<p>
There was a time when disabling preemption on any given CPU would block
subsequent grace periods.
However, this was an accident of implementation and is not a requirement.
And in the current Linux-kernel implementation, disabling preemption
on a given CPU in fact does not block grace periods, as Oleg Nesterov
<a href="https://lkml.kernel.org/g/20150614193825.GA19582@redhat.com">demonstrated</a>.
<p>
If you need a preempt-disable region to block grace periods, you need to add
<tt>rcu_read_lock()</tt> and <tt>rcu_read_unlock()</tt>, for example
as follows:
<blockquote>
<pre>
1 preempt_disable();
2 rcu_read_lock();
3 do_something();
4 rcu_read_unlock();
5 preempt_enable();
6
7 /* Spinlocks implicitly disable preemption. */
8 spin_lock(&mylock);
9 rcu_read_lock();
10 do_something();
11 rcu_read_unlock();
12 spin_unlock(&mylock);
</pre>
</blockquote>
<p>
In theory, you could enter the RCU read-side critical section first,
but it is more efficient to keep the entire RCU read-side critical
section contained in the preempt-disable region as shown above.
Of course, RCU read-side critical sections that extend outside of
preempt-disable regions will work correctly, but such critical sections
can be preempted, which forces <tt>rcu_read_unlock()</tt> to do
more work.
And no, this is <i>not</i> an invitation to enclose all of your RCU
read-side critical sections within preempt-disable regions, because
doing so would degrade real-time response.
<p>
This non-requirement appeared with preemptible RCU.
If you need a grace period that waits on non-preemptible code regions, use
<a href="#Sched Flavor">RCU-sched</a>.
<h2><a name="Parallelism Facts of Life">Parallelism Facts of Life</a></h2>
<p>
These parallelism facts of life are by no means specific to RCU, but
the RCU implementation must abide by them.
They therefore bear repeating:
<ol>
<li> Any CPU or task may be delayed at any time,
and any attempts to avoid these delays by disabling
preemption, interrupts, or whatever are completely futile.
This is most obvious in preemptible user-level
environments and in virtualized environments (where
a given guest OS's VCPUs can be preempted at any time by
the underlying hypervisor), but can also happen in bare-metal
environments due to ECC errors, NMIs, and other hardware
events.
Although a delay of more than about 20 seconds can result
in splats, the RCU implementation is obligated to use
algorithms that can tolerate extremely long delays, but where
“extremely long” is not long enough to allow
wrap-around when incrementing a 64-bit counter.
<li> Both the compiler and the CPU can reorder memory accesses.
Where it matters, RCU must use compiler directives and
memory-barrier instructions to preserve ordering.
<li> Conflicting writes to memory locations in any given cache line
will result in expensive cache misses.
Greater numbers of concurrent writes and more-frequent
concurrent writes will result in more dramatic slowdowns.
RCU is therefore obligated to use algorithms that have
sufficient locality to avoid significant performance and
scalability problems.
<li> As a rough rule of thumb, only one CPU's worth of processing
may be carried out under the protection of any given exclusive
lock.
RCU must therefore use scalable locking designs.
<li> Counters are finite, especially on 32-bit systems.
RCU's use of counters must therefore tolerate counter wrap,
or be designed such that counter wrap would take way more
time than a single system is likely to run.
An uptime of ten years is quite possible, a runtime
of a century much less so.
As an example of the latter, RCU's dyntick-idle nesting counter
allows 54 bits for interrupt nesting level (this counter
is 64 bits even on a 32-bit system).
Overflowing this counter requires 2<sup>54</sup>
half-interrupts on a given CPU without that CPU ever going idle.
If a half-interrupt happened every microsecond, it would take
570 years of runtime to overflow this counter, which is currently
believed to be an acceptably long time.
<li> Linux systems can have thousands of CPUs running a single
Linux kernel in a single shared-memory environment.
RCU must therefore pay close attention to high-end scalability.
</ol>
<p>
This last parallelism fact of life means that RCU must pay special
attention to the preceding facts of life.
The idea that Linux might scale to systems with thousands of CPUs would
have been met with some skepticism in the 1990s, but these requirements
would have otherwise have been unsurprising, even in the early 1990s.
<h2><a name="Quality-of-Implementation Requirements">Quality-of-Implementation Requirements</a></h2>
<p>
These sections list quality-of-implementation requirements.
Although an RCU implementation that ignores these requirements could
still be used, it would likely be subject to limitations that would
make it inappropriate for industrial-strength production use.
Classes of quality-of-implementation requirements are as follows:
<ol>
<li> <a href="#Specialization">Specialization</a>
<li> <a href="#Performance and Scalability">Performance and Scalability</a>
<li> <a href="#Composability">Composability</a>
<li> <a href="#Corner Cases">Corner Cases</a>
</ol>
<p>
These classes is covered in the following sections.
<h3><a name="Specialization">Specialization</a></h3>
<p>
RCU is and always has been intended primarily for read-mostly situations,
which means that RCU's read-side primitives are optimized, often at the
expense of its update-side primitives.
Experience thus far is captured by the following list of situations:
<ol>
<li> Read-mostly data, where stale and inconsistent data is not
a problem: RCU works great!
<li> Read-mostly data, where data must be consistent:
RCU works well.
<li> Read-write data, where data must be consistent:
RCU <i>might</i> work OK.
Or not.
<li> Write-mostly data, where data must be consistent:
RCU is very unlikely to be the right tool for the job,
with the following exceptions, where RCU can provide:
<ol type=a>
<li> Existence guarantees for update-friendly mechanisms.
<li> Wait-free read-side primitives for real-time use.
</ol>
</ol>
<p>
This focus on read-mostly situations means that RCU must interoperate
with other synchronization primitives.
For example, the <tt>add_gp()</tt> and <tt>remove_gp_synchronous()</tt>
examples discussed earlier use RCU to protect readers and locking to
coordinate updaters.
However, the need extends much farther, requiring that a variety of
synchronization primitives be legal within RCU read-side critical sections,
including spinlocks, sequence locks, atomic operations, reference
counters, and memory barriers.
<table>
<tr><th> </th></tr>
<tr><th align="left">Quick Quiz:</th></tr>
<tr><td>
What about sleeping locks?
</td></tr>
<tr><th align="left">Answer:</th></tr>
<tr><td bgcolor="#ffffff"><font color="ffffff">
These are forbidden within Linux-kernel RCU read-side critical
sections because it is not legal to place a quiescent state
(in this case, voluntary context switch) within an RCU read-side
critical section.
However, sleeping locks may be used within userspace RCU read-side
critical sections, and also within Linux-kernel sleepable RCU
<a href="#Sleepable RCU"><font color="ffffff">(SRCU)</font></a>
read-side critical sections.
In addition, the -rt patchset turns spinlocks into a
sleeping locks so that the corresponding critical sections
can be preempted, which also means that these sleeplockified
spinlocks (but not other sleeping locks!) may be acquire within
-rt-Linux-kernel RCU read-side critical sections.
</font>
<p><font color="ffffff">
Note that it <i>is</i> legal for a normal RCU read-side
critical section to conditionally acquire a sleeping locks
(as in <tt>mutex_trylock()</tt>), but only as long as it does
not loop indefinitely attempting to conditionally acquire that
sleeping locks.
The key point is that things like <tt>mutex_trylock()</tt>
either return with the mutex held, or return an error indication if
the mutex was not immediately available.
Either way, <tt>mutex_trylock()</tt> returns immediately without
sleeping.
</font></td></tr>
<tr><td> </td></tr>
</table>
<p>
It often comes as a surprise that many algorithms do not require a
consistent view of data, but many can function in that mode,
with network routing being the poster child.
Internet routing algorithms take significant time to propagate
updates, so that by the time an update arrives at a given system,
that system has been sending network traffic the wrong way for
a considerable length of time.
Having a few threads continue to send traffic the wrong way for a
few more milliseconds is clearly not a problem: In the worst case,
TCP retransmissions will eventually get the data where it needs to go.
In general, when tracking the state of the universe outside of the
computer, some level of inconsistency must be tolerated due to
speed-of-light delays if nothing else.
<p>
Furthermore, uncertainty about external state is inherent in many cases.
For example, a pair of veterinarians might use heartbeat to determine
whether or not a given cat was alive.
But how long should they wait after the last heartbeat to decide that
the cat is in fact dead?
Waiting less than 400 milliseconds makes no sense because this would
mean that a relaxed cat would be considered to cycle between death
and life more than 100 times per minute.
Moreover, just as with human beings, a cat's heart might stop for
some period of time, so the exact wait period is a judgment call.
One of our pair of veterinarians might wait 30 seconds before pronouncing
the cat dead, while the other might insist on waiting a full minute.
The two veterinarians would then disagree on the state of the cat during
the final 30 seconds of the minute following the last heartbeat.
<p>
Interestingly enough, this same situation applies to hardware.
When push comes to shove, how do we tell whether or not some
external server has failed?
We send messages to it periodically, and declare it failed if we
don't receive a response within a given period of time.
Policy decisions can usually tolerate short
periods of inconsistency.
The policy was decided some time ago, and is only now being put into
effect, so a few milliseconds of delay is normally inconsequential.
<p>
However, there are algorithms that absolutely must see consistent data.
For example, the translation between a user-level SystemV semaphore
ID to the corresponding in-kernel data structure is protected by RCU,
but it is absolutely forbidden to update a semaphore that has just been
removed.
In the Linux kernel, this need for consistency is accommodated by acquiring
spinlocks located in the in-kernel data structure from within
the RCU read-side critical section, and this is indicated by the
green box in the figure above.
Many other techniques may be used, and are in fact used within the
Linux kernel.
<p>
In short, RCU is not required to maintain consistency, and other
mechanisms may be used in concert with RCU when consistency is required.
RCU's specialization allows it to do its job extremely well, and its
ability to interoperate with other synchronization mechanisms allows
the right mix of synchronization tools to be used for a given job.
<h3><a name="Performance and Scalability">Performance and Scalability</a></h3>
<p>
Energy efficiency is a critical component of performance today,
and Linux-kernel RCU implementations must therefore avoid unnecessarily
awakening idle CPUs.
I cannot claim that this requirement was premeditated.
In fact, I learned of it during a telephone conversation in which I
was given “frank and open” feedback on the importance
of energy efficiency in battery-powered systems and on specific
energy-efficiency shortcomings of the Linux-kernel RCU implementation.
In my experience, the battery-powered embedded community will consider
any unnecessary wakeups to be extremely unfriendly acts.
So much so that mere Linux-kernel-mailing-list posts are
insufficient to vent their ire.
<p>
Memory consumption is not particularly important for in most
situations, and has become decreasingly
so as memory sizes have expanded and memory
costs have plummeted.
However, as I learned from Matt Mackall's
<a href="http://elinux.org/Linux_Tiny-FAQ">bloatwatch</a>
efforts, memory footprint is critically important on single-CPU systems with
non-preemptible (<tt>CONFIG_PREEMPT=n</tt>) kernels, and thus
<a href="https://lkml.kernel.org/g/20090113221724.GA15307@linux.vnet.ibm.com">tiny RCU</a>
was born.
Josh Triplett has since taken over the small-memory banner with his
<a href="https://tiny.wiki.kernel.org/">Linux kernel tinification</a>
project, which resulted in
<a href="#Sleepable RCU">SRCU</a>
becoming optional for those kernels not needing it.
<p>
The remaining performance requirements are, for the most part,
unsurprising.
For example, in keeping with RCU's read-side specialization,
<tt>rcu_dereference()</tt> should have negligible overhead (for
example, suppression of a few minor compiler optimizations).
Similarly, in non-preemptible environments, <tt>rcu_read_lock()</tt> and
<tt>rcu_read_unlock()</tt> should have exactly zero overhead.
<p>
In preemptible environments, in the case where the RCU read-side
critical section was not preempted (as will be the case for the
highest-priority real-time process), <tt>rcu_read_lock()</tt> and
<tt>rcu_read_unlock()</tt> should have minimal overhead.
In particular, they should not contain atomic read-modify-write
operations, memory-barrier instructions, preemption disabling,
interrupt disabling, or backwards branches.
However, in the case where the RCU read-side critical section was preempted,
<tt>rcu_read_unlock()</tt> may acquire spinlocks and disable interrupts.
This is why it is better to nest an RCU read-side critical section
within a preempt-disable region than vice versa, at least in cases
where that critical section is short enough to avoid unduly degrading
real-time latencies.
<p>
The <tt>synchronize_rcu()</tt> grace-period-wait primitive is
optimized for throughput.
It may therefore incur several milliseconds of latency in addition to
the duration of the longest RCU read-side critical section.
On the other hand, multiple concurrent invocations of
<tt>synchronize_rcu()</tt> are required to use batching optimizations
so that they can be satisfied by a single underlying grace-period-wait
operation.
For example, in the Linux kernel, it is not unusual for a single
grace-period-wait operation to serve more than
<a href="https://www.usenix.org/conference/2004-usenix-annual-technical-conference/making-rcu-safe-deep-sub-millisecond-response">1,000 separate invocations</a>
of <tt>synchronize_rcu()</tt>, thus amortizing the per-invocation
overhead down to nearly zero.
However, the grace-period optimization is also required to avoid
measurable degradation of real-time scheduling and interrupt latencies.
<p>
In some cases, the multi-millisecond <tt>synchronize_rcu()</tt>
latencies are unacceptable.
In these cases, <tt>synchronize_rcu_expedited()</tt> may be used
instead, reducing the grace-period latency down to a few tens of
microseconds on small systems, at least in cases where the RCU read-side
critical sections are short.
There are currently no special latency requirements for
<tt>synchronize_rcu_expedited()</tt> on large systems, but,
consistent with the empirical nature of the RCU specification,
that is subject to change.
However, there most definitely are scalability requirements:
A storm of <tt>synchronize_rcu_expedited()</tt> invocations on 4096
CPUs should at least make reasonable forward progress.
In return for its shorter latencies, <tt>synchronize_rcu_expedited()</tt>
is permitted to impose modest degradation of real-time latency
on non-idle online CPUs.
That said, it will likely be necessary to take further steps to reduce this
degradation, hopefully to roughly that of a scheduling-clock interrupt.
<p>
There are a number of situations where even
<tt>synchronize_rcu_expedited()</tt>'s reduced grace-period
latency is unacceptable.
In these situations, the asynchronous <tt>call_rcu()</tt> can be
used in place of <tt>synchronize_rcu()</tt> as follows:
<blockquote>
<pre>
1 struct foo {
2 int a;
3 int b;
4 struct rcu_head rh;
5 };
6
7 static void remove_gp_cb(struct rcu_head *rhp)
8 {
9 struct foo *p = container_of(rhp, struct foo, rh);
10
11 kfree(p);
12 }
13
14 bool remove_gp_asynchronous(void)
15 {
16 struct foo *p;
17
18 spin_lock(&gp_lock);
19 p = rcu_dereference(gp);
20 if (!p) {
21 spin_unlock(&gp_lock);
22 return false;
23 }
24 rcu_assign_pointer(gp, NULL);
25 call_rcu(&p->rh, remove_gp_cb);
26 spin_unlock(&gp_lock);
27 return true;
28 }
</pre>
</blockquote>
<p>
A definition of <tt>struct foo</tt> is finally needed, and appears
on lines 1-5.
The function <tt>remove_gp_cb()</tt> is passed to <tt>call_rcu()</tt>
on line 25, and will be invoked after the end of a subsequent
grace period.
This gets the same effect as <tt>remove_gp_synchronous()</tt>,
but without forcing the updater to wait for a grace period to elapse.
The <tt>call_rcu()</tt> function may be used in a number of
situations where neither <tt>synchronize_rcu()</tt> nor
<tt>synchronize_rcu_expedited()</tt> would be legal,
including within preempt-disable code, <tt>local_bh_disable()</tt> code,
interrupt-disable code, and interrupt handlers.
However, even <tt>call_rcu()</tt> is illegal within NMI handlers
and from idle and offline CPUs.
The callback function (<tt>remove_gp_cb()</tt> in this case) will be
executed within softirq (software interrupt) environment within the
Linux kernel,
either within a real softirq handler or under the protection
of <tt>local_bh_disable()</tt>.
In both the Linux kernel and in userspace, it is bad practice to
write an RCU callback function that takes too long.
Long-running operations should be relegated to separate threads or
(in the Linux kernel) workqueues.
<table>
<tr><th> </th></tr>
<tr><th align="left">Quick Quiz:</th></tr>
<tr><td>
Why does line 19 use <tt>rcu_access_pointer()</tt>?
After all, <tt>call_rcu()</tt> on line 25 stores into the
structure, which would interact badly with concurrent insertions.
Doesn't this mean that <tt>rcu_dereference()</tt> is required?
</td></tr>
<tr><th align="left">Answer:</th></tr>
<tr><td bgcolor="#ffffff"><font color="ffffff">
Presumably the <tt>->gp_lock</tt> acquired on line 18 excludes
any changes, including any insertions that <tt>rcu_dereference()</tt>
would protect against.
Therefore, any insertions will be delayed until after
<tt>->gp_lock</tt>
is released on line 25, which in turn means that
<tt>rcu_access_pointer()</tt> suffices.
</font></td></tr>
<tr><td> </td></tr>
</table>
<p>
However, all that <tt>remove_gp_cb()</tt> is doing is
invoking <tt>kfree()</tt> on the data element.
This is a common idiom, and is supported by <tt>kfree_rcu()</tt>,
which allows “fire and forget” operation as shown below:
<blockquote>
<pre>
1 struct foo {
2 int a;
3 int b;
4 struct rcu_head rh;
5 };
6
7 bool remove_gp_faf(void)
8 {
9 struct foo *p;
10
11 spin_lock(&gp_lock);
12 p = rcu_dereference(gp);
13 if (!p) {
14 spin_unlock(&gp_lock);
15 return false;
16 }
17 rcu_assign_pointer(gp, NULL);
18 kfree_rcu(p, rh);
19 spin_unlock(&gp_lock);
20 return true;
21 }
</pre>
</blockquote>
<p>
Note that <tt>remove_gp_faf()</tt> simply invokes
<tt>kfree_rcu()</tt> and proceeds, without any need to pay any
further attention to the subsequent grace period and <tt>kfree()</tt>.
It is permissible to invoke <tt>kfree_rcu()</tt> from the same
environments as for <tt>call_rcu()</tt>.
Interestingly enough, DYNIX/ptx had the equivalents of
<tt>call_rcu()</tt> and <tt>kfree_rcu()</tt>, but not
<tt>synchronize_rcu()</tt>.
This was due to the fact that RCU was not heavily used within DYNIX/ptx,
so the very few places that needed something like
<tt>synchronize_rcu()</tt> simply open-coded it.
<table>
<tr><th> </th></tr>
<tr><th align="left">Quick Quiz:</th></tr>
<tr><td>
Earlier it was claimed that <tt>call_rcu()</tt> and
<tt>kfree_rcu()</tt> allowed updaters to avoid being blocked
by readers.
But how can that be correct, given that the invocation of the callback
and the freeing of the memory (respectively) must still wait for
a grace period to elapse?
</td></tr>
<tr><th align="left">Answer:</th></tr>
<tr><td bgcolor="#ffffff"><font color="ffffff">
We could define things this way, but keep in mind that this sort of
definition would say that updates in garbage-collected languages
cannot complete until the next time the garbage collector runs,
which does not seem at all reasonable.
The key point is that in most cases, an updater using either
<tt>call_rcu()</tt> or <tt>kfree_rcu()</tt> can proceed to the
next update as soon as it has invoked <tt>call_rcu()</tt> or
<tt>kfree_rcu()</tt>, without having to wait for a subsequent
grace period.
</font></td></tr>
<tr><td> </td></tr>
</table>
<p>
But what if the updater must wait for the completion of code to be
executed after the end of the grace period, but has other tasks
that can be carried out in the meantime?
The polling-style <tt>get_state_synchronize_rcu()</tt> and
<tt>cond_synchronize_rcu()</tt> functions may be used for this
purpose, as shown below:
<blockquote>
<pre>
1 bool remove_gp_poll(void)
2 {
3 struct foo *p;
4 unsigned long s;
5
6 spin_lock(&gp_lock);
7 p = rcu_access_pointer(gp);
8 if (!p) {
9 spin_unlock(&gp_lock);
10 return false;
11 }
12 rcu_assign_pointer(gp, NULL);
13 spin_unlock(&gp_lock);
14 s = get_state_synchronize_rcu();
15 do_something_while_waiting();
16 cond_synchronize_rcu(s);
17 kfree(p);
18 return true;
19 }
</pre>
</blockquote>
<p>
On line 14, <tt>get_state_synchronize_rcu()</tt> obtains a
“cookie” from RCU,
then line 15 carries out other tasks,
and finally, line 16 returns immediately if a grace period has
elapsed in the meantime, but otherwise waits as required.
The need for <tt>get_state_synchronize_rcu</tt> and
<tt>cond_synchronize_rcu()</tt> has appeared quite recently,
so it is too early to tell whether they will stand the test of time.
<p>
RCU thus provides a range of tools to allow updaters to strike the
required tradeoff between latency, flexibility and CPU overhead.
<h3><a name="Composability">Composability</a></h3>
<p>
Composability has received much attention in recent years, perhaps in part
due to the collision of multicore hardware with object-oriented techniques
designed in single-threaded environments for single-threaded use.
And in theory, RCU read-side critical sections may be composed, and in
fact may be nested arbitrarily deeply.
In practice, as with all real-world implementations of composable
constructs, there are limitations.
<p>
Implementations of RCU for which <tt>rcu_read_lock()</tt>
and <tt>rcu_read_unlock()</tt> generate no code, such as
Linux-kernel RCU when <tt>CONFIG_PREEMPT=n</tt>, can be
nested arbitrarily deeply.
After all, there is no overhead.
Except that if all these instances of <tt>rcu_read_lock()</tt>
and <tt>rcu_read_unlock()</tt> are visible to the compiler,
compilation will eventually fail due to exhausting memory,
mass storage, or user patience, whichever comes first.
If the nesting is not visible to the compiler, as is the case with
mutually recursive functions each in its own translation unit,
stack overflow will result.
If the nesting takes the form of loops, either the control variable
will overflow or (in the Linux kernel) you will get an RCU CPU stall warning.
Nevertheless, this class of RCU implementations is one
of the most composable constructs in existence.
<p>
RCU implementations that explicitly track nesting depth
are limited by the nesting-depth counter.
For example, the Linux kernel's preemptible RCU limits nesting to
<tt>INT_MAX</tt>.
This should suffice for almost all practical purposes.
That said, a consecutive pair of RCU read-side critical sections
between which there is an operation that waits for a grace period
cannot be enclosed in another RCU read-side critical section.
This is because it is not legal to wait for a grace period within
an RCU read-side critical section: To do so would result either
in deadlock or
in RCU implicitly splitting the enclosing RCU read-side critical
section, neither of which is conducive to a long-lived and prosperous
kernel.
<p>
It is worth noting that RCU is not alone in limiting composability.
For example, many transactional-memory implementations prohibit
composing a pair of transactions separated by an irrevocable
operation (for example, a network receive operation).
For another example, lock-based critical sections can be composed
surprisingly freely, but only if deadlock is avoided.
<p>
In short, although RCU read-side critical sections are highly composable,
care is required in some situations, just as is the case for any other
composable synchronization mechanism.
<h3><a name="Corner Cases">Corner Cases</a></h3>
<p>
A given RCU workload might have an endless and intense stream of
RCU read-side critical sections, perhaps even so intense that there
was never a point in time during which there was not at least one
RCU read-side critical section in flight.
RCU cannot allow this situation to block grace periods: As long as
all the RCU read-side critical sections are finite, grace periods
must also be finite.
<p>
That said, preemptible RCU implementations could potentially result
in RCU read-side critical sections being preempted for long durations,
which has the effect of creating a long-duration RCU read-side
critical section.
This situation can arise only in heavily loaded systems, but systems using
real-time priorities are of course more vulnerable.
Therefore, RCU priority boosting is provided to help deal with this
case.
That said, the exact requirements on RCU priority boosting will likely
evolve as more experience accumulates.
<p>
Other workloads might have very high update rates.
Although one can argue that such workloads should instead use
something other than RCU, the fact remains that RCU must
handle such workloads gracefully.
This requirement is another factor driving batching of grace periods,
but it is also the driving force behind the checks for large numbers
of queued RCU callbacks in the <tt>call_rcu()</tt> code path.
Finally, high update rates should not delay RCU read-side critical
sections, although some read-side delays can occur when using
<tt>synchronize_rcu_expedited()</tt>, courtesy of this function's use
of <tt>try_stop_cpus()</tt>.
(In the future, <tt>synchronize_rcu_expedited()</tt> will be
converted to use lighter-weight inter-processor interrupts (IPIs),
but this will still disturb readers, though to a much smaller degree.)
<p>
Although all three of these corner cases were understood in the early
1990s, a simple user-level test consisting of <tt>close(open(path))</tt>
in a tight loop
in the early 2000s suddenly provided a much deeper appreciation of the
high-update-rate corner case.
This test also motivated addition of some RCU code to react to high update
rates, for example, if a given CPU finds itself with more than 10,000
RCU callbacks queued, it will cause RCU to take evasive action by
more aggressively starting grace periods and more aggressively forcing
completion of grace-period processing.
This evasive action causes the grace period to complete more quickly,
but at the cost of restricting RCU's batching optimizations, thus
increasing the CPU overhead incurred by that grace period.
<h2><a name="Software-Engineering Requirements">
Software-Engineering Requirements</a></h2>
<p>
Between Murphy's Law and “To err is human”, it is necessary to
guard against mishaps and misuse:
<ol>
<li> It is all too easy to forget to use <tt>rcu_read_lock()</tt>
everywhere that it is needed, so kernels built with
<tt>CONFIG_PROVE_RCU=y</tt> will splat if
<tt>rcu_dereference()</tt> is used outside of an
RCU read-side critical section.
Update-side code can use <tt>rcu_dereference_protected()</tt>,
which takes a
<a href="https://lwn.net/Articles/371986/">lockdep expression</a>
to indicate what is providing the protection.
If the indicated protection is not provided, a lockdep splat
is emitted.
<p>
Code shared between readers and updaters can use
<tt>rcu_dereference_check()</tt>, which also takes a
lockdep expression, and emits a lockdep splat if neither
<tt>rcu_read_lock()</tt> nor the indicated protection
is in place.
In addition, <tt>rcu_dereference_raw()</tt> is used in those
(hopefully rare) cases where the required protection cannot
be easily described.
Finally, <tt>rcu_read_lock_held()</tt> is provided to
allow a function to verify that it has been invoked within
an RCU read-side critical section.
I was made aware of this set of requirements shortly after Thomas
Gleixner audited a number of RCU uses.
<li> A given function might wish to check for RCU-related preconditions
upon entry, before using any other RCU API.
The <tt>rcu_lockdep_assert()</tt> does this job,
asserting the expression in kernels having lockdep enabled
and doing nothing otherwise.
<li> It is also easy to forget to use <tt>rcu_assign_pointer()</tt>
and <tt>rcu_dereference()</tt>, perhaps (incorrectly)
substituting a simple assignment.
To catch this sort of error, a given RCU-protected pointer may be
tagged with <tt>__rcu</tt>, after which running sparse
with <tt>CONFIG_SPARSE_RCU_POINTER=y</tt> will complain
about simple-assignment accesses to that pointer.
Arnd Bergmann made me aware of this requirement, and also
supplied the needed
<a href="https://lwn.net/Articles/376011/">patch series</a>.
<li> Kernels built with <tt>CONFIG_DEBUG_OBJECTS_RCU_HEAD=y</tt>
will splat if a data element is passed to <tt>call_rcu()</tt>
twice in a row, without a grace period in between.
(This error is similar to a double free.)
The corresponding <tt>rcu_head</tt> structures that are
dynamically allocated are automatically tracked, but
<tt>rcu_head</tt> structures allocated on the stack
must be initialized with <tt>init_rcu_head_on_stack()</tt>
and cleaned up with <tt>destroy_rcu_head_on_stack()</tt>.
Similarly, statically allocated non-stack <tt>rcu_head</tt>
structures must be initialized with <tt>init_rcu_head()</tt>
and cleaned up with <tt>destroy_rcu_head()</tt>.
Mathieu Desnoyers made me aware of this requirement, and also
supplied the needed
<a href="https://lkml.kernel.org/g/20100319013024.GA28456@Krystal">patch</a>.
<li> An infinite loop in an RCU read-side critical section will
eventually trigger an RCU CPU stall warning splat, with
the duration of “eventually” being controlled by the
<tt>RCU_CPU_STALL_TIMEOUT</tt> <tt>Kconfig</tt> option, or,
alternatively, by the
<tt>rcupdate.rcu_cpu_stall_timeout</tt> boot/sysfs
parameter.
However, RCU is not obligated to produce this splat
unless there is a grace period waiting on that particular
RCU read-side critical section.
<p>
Some extreme workloads might intentionally delay
RCU grace periods, and systems running those workloads can
be booted with <tt>rcupdate.rcu_cpu_stall_suppress</tt>
to suppress the splats.
This kernel parameter may also be set via <tt>sysfs</tt>.
Furthermore, RCU CPU stall warnings are counter-productive
during sysrq dumps and during panics.
RCU therefore supplies the <tt>rcu_sysrq_start()</tt> and
<tt>rcu_sysrq_end()</tt> API members to be called before
and after long sysrq dumps.
RCU also supplies the <tt>rcu_panic()</tt> notifier that is
automatically invoked at the beginning of a panic to suppress
further RCU CPU stall warnings.
<p>
This requirement made itself known in the early 1990s, pretty
much the first time that it was necessary to debug a CPU stall.
That said, the initial implementation in DYNIX/ptx was quite
generic in comparison with that of Linux.
<li> Although it would be very good to detect pointers leaking out
of RCU read-side critical sections, there is currently no
good way of doing this.
One complication is the need to distinguish between pointers
leaking and pointers that have been handed off from RCU to
some other synchronization mechanism, for example, reference
counting.
<li> In kernels built with <tt>CONFIG_RCU_TRACE=y</tt>, RCU-related
information is provided via both debugfs and event tracing.
<li> Open-coded use of <tt>rcu_assign_pointer()</tt> and
<tt>rcu_dereference()</tt> to create typical linked
data structures can be surprisingly error-prone.
Therefore, RCU-protected
<a href="https://lwn.net/Articles/609973/#RCU List APIs">linked lists</a>
and, more recently, RCU-protected
<a href="https://lwn.net/Articles/612100/">hash tables</a>
are available.
Many other special-purpose RCU-protected data structures are
available in the Linux kernel and the userspace RCU library.
<li> Some linked structures are created at compile time, but still
require <tt>__rcu</tt> checking.
The <tt>RCU_POINTER_INITIALIZER()</tt> macro serves this
purpose.
<li> It is not necessary to use <tt>rcu_assign_pointer()</tt>
when creating linked structures that are to be published via
a single external pointer.
The <tt>RCU_INIT_POINTER()</tt> macro is provided for
this task and also for assigning <tt>NULL</tt> pointers
at runtime.
</ol>
<p>
This not a hard-and-fast list: RCU's diagnostic capabilities will
continue to be guided by the number and type of usage bugs found
in real-world RCU usage.
<h2><a name="Linux Kernel Complications">Linux Kernel Complications</a></h2>
<p>
The Linux kernel provides an interesting environment for all kinds of
software, including RCU.
Some of the relevant points of interest are as follows:
<ol>
<li> <a href="#Configuration">Configuration</a>.
<li> <a href="#Firmware Interface">Firmware Interface</a>.
<li> <a href="#Early Boot">Early Boot</a>.
<li> <a href="#Interrupts and NMIs">
Interrupts and non-maskable interrupts (NMIs)</a>.
<li> <a href="#Loadable Modules">Loadable Modules</a>.
<li> <a href="#Hotplug CPU">Hotplug CPU</a>.
<li> <a href="#Scheduler and RCU">Scheduler and RCU</a>.
<li> <a href="#Tracing and RCU">Tracing and RCU</a>.
<li> <a href="#Energy Efficiency">Energy Efficiency</a>.
<li> <a href="#Memory Efficiency">Memory Efficiency</a>.
<li> <a href="#Performance, Scalability, Response Time, and Reliability">
Performance, Scalability, Response Time, and Reliability</a>.
</ol>
<p>
This list is probably incomplete, but it does give a feel for the
most notable Linux-kernel complications.
Each of the following sections covers one of the above topics.
<h3><a name="Configuration">Configuration</a></h3>
<p>
RCU's goal is automatic configuration, so that almost nobody
needs to worry about RCU's <tt>Kconfig</tt> options.
And for almost all users, RCU does in fact work well
“out of the box.”
<p>
However, there are specialized use cases that are handled by
kernel boot parameters and <tt>Kconfig</tt> options.
Unfortunately, the <tt>Kconfig</tt> system will explicitly ask users
about new <tt>Kconfig</tt> options, which requires almost all of them
be hidden behind a <tt>CONFIG_RCU_EXPERT</tt> <tt>Kconfig</tt> option.
<p>
This all should be quite obvious, but the fact remains that
Linus Torvalds recently had to
<a href="https://lkml.kernel.org/g/CA+55aFy4wcCwaL4okTs8wXhGZ5h-ibecy_Meg9C4MNQrUnwMcg@mail.gmail.com">remind</a>
me of this requirement.
<h3><a name="Firmware Interface">Firmware Interface</a></h3>
<p>
In many cases, kernel obtains information about the system from the
firmware, and sometimes things are lost in translation.
Or the translation is accurate, but the original message is bogus.
<p>
For example, some systems' firmware overreports the number of CPUs,
sometimes by a large factor.
If RCU naively believed the firmware, as it used to do,
it would create too many per-CPU kthreads.
Although the resulting system will still run correctly, the extra
kthreads needlessly consume memory and can cause confusion
when they show up in <tt>ps</tt> listings.
<p>
RCU must therefore wait for a given CPU to actually come online before
it can allow itself to believe that the CPU actually exists.
The resulting “ghost CPUs” (which are never going to
come online) cause a number of
<a href="https://paulmck.livejournal.com/37494.html">interesting complications</a>.
<h3><a name="Early Boot">Early Boot</a></h3>
<p>
The Linux kernel's boot sequence is an interesting process,
and RCU is used early, even before <tt>rcu_init()</tt>
is invoked.
In fact, a number of RCU's primitives can be used as soon as the
initial task's <tt>task_struct</tt> is available and the
boot CPU's per-CPU variables are set up.
The read-side primitives (<tt>rcu_read_lock()</tt>,
<tt>rcu_read_unlock()</tt>, <tt>rcu_dereference()</tt>,
and <tt>rcu_access_pointer()</tt>) will operate normally very early on,
as will <tt>rcu_assign_pointer()</tt>.
<p>
Although <tt>call_rcu()</tt> may be invoked at any
time during boot, callbacks are not guaranteed to be invoked until after
the scheduler is fully up and running.
This delay in callback invocation is due to the fact that RCU does not
invoke callbacks until it is fully initialized, and this full initialization
cannot occur until after the scheduler has initialized itself to the
point where RCU can spawn and run its kthreads.
In theory, it would be possible to invoke callbacks earlier,
however, this is not a panacea because there would be severe restrictions
on what operations those callbacks could invoke.
<p>
Perhaps surprisingly, <tt>synchronize_rcu()</tt>,
<a href="#Bottom-Half Flavor"><tt>synchronize_rcu_bh()</tt></a>
(<a href="#Bottom-Half Flavor">discussed below</a>),
and
<a href="#Sched Flavor"><tt>synchronize_sched()</tt></a>
will all operate normally
during very early boot, the reason being that there is only one CPU
and preemption is disabled.
This means that the call <tt>synchronize_rcu()</tt> (or friends)
itself is a quiescent
state and thus a grace period, so the early-boot implementation can
be a no-op.
<p>
Both <tt>synchronize_rcu_bh()</tt> and <tt>synchronize_sched()</tt>
continue to operate normally through the remainder of boot, courtesy
of the fact that preemption is disabled across their RCU read-side
critical sections and also courtesy of the fact that there is still
only one CPU.
However, once the scheduler starts initializing, preemption is enabled.
There is still only a single CPU, but the fact that preemption is enabled
means that the no-op implementation of <tt>synchronize_rcu()</tt> no
longer works in <tt>CONFIG_PREEMPT=y</tt> kernels.
Therefore, as soon as the scheduler starts initializing, the early-boot
fastpath is disabled.
This means that <tt>synchronize_rcu()</tt> switches to its runtime
mode of operation where it posts callbacks, which in turn means that
any call to <tt>synchronize_rcu()</tt> will block until the corresponding
callback is invoked.
Unfortunately, the callback cannot be invoked until RCU's runtime
grace-period machinery is up and running, which cannot happen until
the scheduler has initialized itself sufficiently to allow RCU's
kthreads to be spawned.
Therefore, invoking <tt>synchronize_rcu()</tt> during scheduler
initialization can result in deadlock.
<table>
<tr><th> </th></tr>
<tr><th align="left">Quick Quiz:</th></tr>
<tr><td>
So what happens with <tt>synchronize_rcu()</tt> during
scheduler initialization for <tt>CONFIG_PREEMPT=n</tt>
kernels?
</td></tr>
<tr><th align="left">Answer:</th></tr>
<tr><td bgcolor="#ffffff"><font color="ffffff">
In <tt>CONFIG_PREEMPT=n</tt> kernel, <tt>synchronize_rcu()</tt>
maps directly to <tt>synchronize_sched()</tt>.
Therefore, <tt>synchronize_rcu()</tt> works normally throughout
boot in <tt>CONFIG_PREEMPT=n</tt> kernels.
However, your code must also work in <tt>CONFIG_PREEMPT=y</tt> kernels,
so it is still necessary to avoid invoking <tt>synchronize_rcu()</tt>
during scheduler initialization.
</font></td></tr>
<tr><td> </td></tr>
</table>
<p>
I learned of these boot-time requirements as a result of a series of
system hangs.
<h3><a name="Interrupts and NMIs">Interrupts and NMIs</a></h3>
<p>
The Linux kernel has interrupts, and RCU read-side critical sections are
legal within interrupt handlers and within interrupt-disabled regions
of code, as are invocations of <tt>call_rcu()</tt>.
<p>
Some Linux-kernel architectures can enter an interrupt handler from
non-idle process context, and then just never leave it, instead stealthily
transitioning back to process context.
This trick is sometimes used to invoke system calls from inside the kernel.
These “half-interrupts” mean that RCU has to be very careful
about how it counts interrupt nesting levels.
I learned of this requirement the hard way during a rewrite
of RCU's dyntick-idle code.
<p>
The Linux kernel has non-maskable interrupts (NMIs), and
RCU read-side critical sections are legal within NMI handlers.
Thankfully, RCU update-side primitives, including
<tt>call_rcu()</tt>, are prohibited within NMI handlers.
<p>
The name notwithstanding, some Linux-kernel architectures
can have nested NMIs, which RCU must handle correctly.
Andy Lutomirski
<a href="https://lkml.kernel.org/g/CALCETrXLq1y7e_dKFPgou-FKHB6Pu-r8+t-6Ds+8=va7anBWDA@mail.gmail.com">surprised me</a>
with this requirement;
he also kindly surprised me with
<a href="https://lkml.kernel.org/g/CALCETrXSY9JpW3uE6H8WYk81sg56qasA2aqmjMPsq5dOtzso=g@mail.gmail.com">an algorithm</a>
that meets this requirement.
<h3><a name="Loadable Modules">Loadable Modules</a></h3>
<p>
The Linux kernel has loadable modules, and these modules can
also be unloaded.
After a given module has been unloaded, any attempt to call
one of its functions results in a segmentation fault.
The module-unload functions must therefore cancel any
delayed calls to loadable-module functions, for example,
any outstanding <tt>mod_timer()</tt> must be dealt with
via <tt>del_timer_sync()</tt> or similar.
<p>
Unfortunately, there is no way to cancel an RCU callback;
once you invoke <tt>call_rcu()</tt>, the callback function is
going to eventually be invoked, unless the system goes down first.
Because it is normally considered socially irresponsible to crash the system
in response to a module unload request, we need some other way
to deal with in-flight RCU callbacks.
<p>
RCU therefore provides
<tt><a href="https://lwn.net/Articles/217484/">rcu_barrier()</a></tt>,
which waits until all in-flight RCU callbacks have been invoked.
If a module uses <tt>call_rcu()</tt>, its exit function should therefore
prevent any future invocation of <tt>call_rcu()</tt>, then invoke
<tt>rcu_barrier()</tt>.
In theory, the underlying module-unload code could invoke
<tt>rcu_barrier()</tt> unconditionally, but in practice this would
incur unacceptable latencies.
<p>
Nikita Danilov noted this requirement for an analogous filesystem-unmount
situation, and Dipankar Sarma incorporated <tt>rcu_barrier()</tt> into RCU.
The need for <tt>rcu_barrier()</tt> for module unloading became
apparent later.
<h3><a name="Hotplug CPU">Hotplug CPU</a></h3>
<p>
The Linux kernel supports CPU hotplug, which means that CPUs
can come and go.
It is of course illegal to use any RCU API member from an offline CPU.
This requirement was present from day one in DYNIX/ptx, but
on the other hand, the Linux kernel's CPU-hotplug implementation
is “interesting.”
<p>
The Linux-kernel CPU-hotplug implementation has notifiers that
are used to allow the various kernel subsystems (including RCU)
to respond appropriately to a given CPU-hotplug operation.
Most RCU operations may be invoked from CPU-hotplug notifiers,
including even normal synchronous grace-period operations
such as <tt>synchronize_rcu()</tt>.
However, expedited grace-period operations such as
<tt>synchronize_rcu_expedited()</tt> are not supported,
due to the fact that current implementations block CPU-hotplug
operations, which could result in deadlock.
<p>
In addition, all-callback-wait operations such as
<tt>rcu_barrier()</tt> are also not supported, due to the
fact that there are phases of CPU-hotplug operations where
the outgoing CPU's callbacks will not be invoked until after
the CPU-hotplug operation ends, which could also result in deadlock.
<h3><a name="Scheduler and RCU">Scheduler and RCU</a></h3>
<p>
RCU depends on the scheduler, and the scheduler uses RCU to
protect some of its data structures.
This means the scheduler is forbidden from acquiring
the runqueue locks and the priority-inheritance locks
in the middle of an outermost RCU read-side critical section unless either
(1) it releases them before exiting that same
RCU read-side critical section, or
(2) interrupts are disabled across
that entire RCU read-side critical section.
This same prohibition also applies (recursively!) to any lock that is acquired
while holding any lock to which this prohibition applies.
Adhering to this rule prevents preemptible RCU from invoking
<tt>rcu_read_unlock_special()</tt> while either runqueue or
priority-inheritance locks are held, thus avoiding deadlock.
<p>
Prior to v4.4, it was only necessary to disable preemption across
RCU read-side critical sections that acquired scheduler locks.
In v4.4, expedited grace periods started using IPIs, and these
IPIs could force a <tt>rcu_read_unlock()</tt> to take the slowpath.
Therefore, this expedited-grace-period change required disabling of
interrupts, not just preemption.
<p>
For RCU's part, the preemptible-RCU <tt>rcu_read_unlock()</tt>
implementation must be written carefully to avoid similar deadlocks.
In particular, <tt>rcu_read_unlock()</tt> must tolerate an
interrupt where the interrupt handler invokes both
<tt>rcu_read_lock()</tt> and <tt>rcu_read_unlock()</tt>.
This possibility requires <tt>rcu_read_unlock()</tt> to use
negative nesting levels to avoid destructive recursion via
interrupt handler's use of RCU.
<p>
This pair of mutual scheduler-RCU requirements came as a
<a href="https://lwn.net/Articles/453002/">complete surprise</a>.
<p>
As noted above, RCU makes use of kthreads, and it is necessary to
avoid excessive CPU-time accumulation by these kthreads.
This requirement was no surprise, but RCU's violation of it
when running context-switch-heavy workloads when built with
<tt>CONFIG_NO_HZ_FULL=y</tt>
<a href="http://www.rdrop.com/users/paulmck/scalability/paper/BareMetal.2015.01.15b.pdf">did come as a surprise [PDF]</a>.
RCU has made good progress towards meeting this requirement, even
for context-switch-have <tt>CONFIG_NO_HZ_FULL=y</tt> workloads,
but there is room for further improvement.
<h3><a name="Tracing and RCU">Tracing and RCU</a></h3>
<p>
It is possible to use tracing on RCU code, but tracing itself
uses RCU.
For this reason, <tt>rcu_dereference_raw_notrace()</tt>
is provided for use by tracing, which avoids the destructive
recursion that could otherwise ensue.
This API is also used by virtualization in some architectures,
where RCU readers execute in environments in which tracing
cannot be used.
The tracing folks both located the requirement and provided the
needed fix, so this surprise requirement was relatively painless.
<h3><a name="Energy Efficiency">Energy Efficiency</a></h3>
<p>
Interrupting idle CPUs is considered socially unacceptable,
especially by people with battery-powered embedded systems.
RCU therefore conserves energy by detecting which CPUs are
idle, including tracking CPUs that have been interrupted from idle.
This is a large part of the energy-efficiency requirement,
so I learned of this via an irate phone call.
<p>
Because RCU avoids interrupting idle CPUs, it is illegal to
execute an RCU read-side critical section on an idle CPU.
(Kernels built with <tt>CONFIG_PROVE_RCU=y</tt> will splat
if you try it.)
The <tt>RCU_NONIDLE()</tt> macro and <tt>_rcuidle</tt>
event tracing is provided to work around this restriction.
In addition, <tt>rcu_is_watching()</tt> may be used to
test whether or not it is currently legal to run RCU read-side
critical sections on this CPU.
I learned of the need for diagnostics on the one hand
and <tt>RCU_NONIDLE()</tt> on the other while inspecting
idle-loop code.
Steven Rostedt supplied <tt>_rcuidle</tt> event tracing,
which is used quite heavily in the idle loop.
However, there are some restrictions on the code placed within
<tt>RCU_NONIDLE()</tt>:
<ol>
<li> Blocking is prohibited.
In practice, this is not a serious restriction given that idle
tasks are prohibited from blocking to begin with.
<li> Although nesting <tt>RCU_NONIDLE()</tt> is permitted, they cannot
nest indefinitely deeply.
However, given that they can be nested on the order of a million
deep, even on 32-bit systems, this should not be a serious
restriction.
This nesting limit would probably be reached long after the
compiler OOMed or the stack overflowed.
<li> Any code path that enters <tt>RCU_NONIDLE()</tt> must sequence
out of that same <tt>RCU_NONIDLE()</tt>.
For example, the following is grossly illegal:
<blockquote>
<pre>
1 RCU_NONIDLE({
2 do_something();
3 goto bad_idea; /* BUG!!! */
4 do_something_else();});
5 bad_idea:
</pre>
</blockquote>
<p>
It is just as illegal to transfer control into the middle of
<tt>RCU_NONIDLE()</tt>'s argument.
Yes, in theory, you could transfer in as long as you also
transferred out, but in practice you could also expect to get sharply
worded review comments.
</ol>
<p>
It is similarly socially unacceptable to interrupt an
<tt>nohz_full</tt> CPU running in userspace.
RCU must therefore track <tt>nohz_full</tt> userspace
execution.
And in
<a href="https://lwn.net/Articles/558284/"><tt>CONFIG_NO_HZ_FULL_SYSIDLE=y</tt></a>
kernels, RCU must separately track idle CPUs on the one hand and
CPUs that are either idle or executing in userspace on the other.
In both cases, RCU must be able to sample state at two points in
time, and be able to determine whether or not some other CPU spent
any time idle and/or executing in userspace.
<p>
These energy-efficiency requirements have proven quite difficult to
understand and to meet, for example, there have been more than five
clean-sheet rewrites of RCU's energy-efficiency code, the last of
which was finally able to demonstrate
<a href="http://www.rdrop.com/users/paulmck/realtime/paper/AMPenergy.2013.04.19a.pdf">real energy savings running on real hardware [PDF]</a>.
As noted earlier,
I learned of many of these requirements via angry phone calls:
Flaming me on the Linux-kernel mailing list was apparently not
sufficient to fully vent their ire at RCU's energy-efficiency bugs!
<h3><a name="Memory Efficiency">Memory Efficiency</a></h3>
<p>
Although small-memory non-realtime systems can simply use Tiny RCU,
code size is only one aspect of memory efficiency.
Another aspect is the size of the <tt>rcu_head</tt> structure
used by <tt>call_rcu()</tt> and <tt>kfree_rcu()</tt>.
Although this structure contains nothing more than a pair of pointers,
it does appear in many RCU-protected data structures, including
some that are size critical.
The <tt>page</tt> structure is a case in point, as evidenced by
the many occurrences of the <tt>union</tt> keyword within that structure.
<p>
This need for memory efficiency is one reason that RCU uses hand-crafted
singly linked lists to track the <tt>rcu_head</tt> structures that
are waiting for a grace period to elapse.
It is also the reason why <tt>rcu_head</tt> structures do not contain
debug information, such as fields tracking the file and line of the
<tt>call_rcu()</tt> or <tt>kfree_rcu()</tt> that posted them.
Although this information might appear in debug-only kernel builds at some
point, in the meantime, the <tt>->func</tt> field will often provide
the needed debug information.
<p>
However, in some cases, the need for memory efficiency leads to even
more extreme measures.
Returning to the <tt>page</tt> structure, the <tt>rcu_head</tt> field
shares storage with a great many other structures that are used at
various points in the corresponding page's lifetime.
In order to correctly resolve certain
<a href="https://lkml.kernel.org/g/1439976106-137226-1-git-send-email-kirill.shutemov@linux.intel.com">race conditions</a>,
the Linux kernel's memory-management subsystem needs a particular bit
to remain zero during all phases of grace-period processing,
and that bit happens to map to the bottom bit of the
<tt>rcu_head</tt> structure's <tt>->next</tt> field.
RCU makes this guarantee as long as <tt>call_rcu()</tt>
is used to post the callback, as opposed to <tt>kfree_rcu()</tt>
or some future “lazy”
variant of <tt>call_rcu()</tt> that might one day be created for
energy-efficiency purposes.
<p>
That said, there are limits.
RCU requires that the <tt>rcu_head</tt> structure be aligned to a
two-byte boundary, and passing a misaligned <tt>rcu_head</tt>
structure to one of the <tt>call_rcu()</tt> family of functions
will result in a splat.
It is therefore necessary to exercise caution when packing
structures containing fields of type <tt>rcu_head</tt>.
Why not a four-byte or even eight-byte alignment requirement?
Because the m68k architecture provides only two-byte alignment,
and thus acts as alignment's least common denominator.
<p>
The reason for reserving the bottom bit of pointers to
<tt>rcu_head</tt> structures is to leave the door open to
“lazy” callbacks whose invocations can safely be deferred.
Deferring invocation could potentially have energy-efficiency
benefits, but only if the rate of non-lazy callbacks decreases
significantly for some important workload.
In the meantime, reserving the bottom bit keeps this option open
in case it one day becomes useful.
<h3><a name="Performance, Scalability, Response Time, and Reliability">
Performance, Scalability, Response Time, and Reliability</a></h3>
<p>
Expanding on the
<a href="#Performance and Scalability">earlier discussion</a>,
RCU is used heavily by hot code paths in performance-critical
portions of the Linux kernel's networking, security, virtualization,
and scheduling code paths.
RCU must therefore use efficient implementations, especially in its
read-side primitives.
To that end, it would be good if preemptible RCU's implementation
of <tt>rcu_read_lock()</tt> could be inlined, however, doing
this requires resolving <tt>#include</tt> issues with the
<tt>task_struct</tt> structure.
<p>
The Linux kernel supports hardware configurations with up to
4096 CPUs, which means that RCU must be extremely scalable.
Algorithms that involve frequent acquisitions of global locks or
frequent atomic operations on global variables simply cannot be
tolerated within the RCU implementation.
RCU therefore makes heavy use of a combining tree based on the
<tt>rcu_node</tt> structure.
RCU is required to tolerate all CPUs continuously invoking any
combination of RCU's runtime primitives with minimal per-operation
overhead.
In fact, in many cases, increasing load must <i>decrease</i> the
per-operation overhead, witness the batching optimizations for
<tt>synchronize_rcu()</tt>, <tt>call_rcu()</tt>,
<tt>synchronize_rcu_expedited()</tt>, and <tt>rcu_barrier()</tt>.
As a general rule, RCU must cheerfully accept whatever the
rest of the Linux kernel decides to throw at it.
<p>
The Linux kernel is used for real-time workloads, especially
in conjunction with the
<a href="https://rt.wiki.kernel.org/index.php/Main_Page">-rt patchset</a>.
The real-time-latency response requirements are such that the
traditional approach of disabling preemption across RCU
read-side critical sections is inappropriate.
Kernels built with <tt>CONFIG_PREEMPT=y</tt> therefore
use an RCU implementation that allows RCU read-side critical
sections to be preempted.
This requirement made its presence known after users made it
clear that an earlier
<a href="https://lwn.net/Articles/107930/">real-time patch</a>
did not meet their needs, in conjunction with some
<a href="https://lkml.kernel.org/g/20050318002026.GA2693@us.ibm.com">RCU issues</a>
encountered by a very early version of the -rt patchset.
<p>
In addition, RCU must make do with a sub-100-microsecond real-time latency
budget.
In fact, on smaller systems with the -rt patchset, the Linux kernel
provides sub-20-microsecond real-time latencies for the whole kernel,
including RCU.
RCU's scalability and latency must therefore be sufficient for
these sorts of configurations.
To my surprise, the sub-100-microsecond real-time latency budget
<a href="http://www.rdrop.com/users/paulmck/realtime/paper/bigrt.2013.01.31a.LCA.pdf">
applies to even the largest systems [PDF]</a>,
up to and including systems with 4096 CPUs.
This real-time requirement motivated the grace-period kthread, which
also simplified handling of a number of race conditions.
<p>
RCU must avoid degrading real-time response for CPU-bound threads, whether
executing in usermode (which is one use case for
<tt>CONFIG_NO_HZ_FULL=y</tt>) or in the kernel.
That said, CPU-bound loops in the kernel must execute
<tt>cond_resched_rcu_qs()</tt> at least once per few tens of milliseconds
in order to avoid receiving an IPI from RCU.
<p>
Finally, RCU's status as a synchronization primitive means that
any RCU failure can result in arbitrary memory corruption that can be
extremely difficult to debug.
This means that RCU must be extremely reliable, which in
practice also means that RCU must have an aggressive stress-test
suite.
This stress-test suite is called <tt>rcutorture</tt>.
<p>
Although the need for <tt>rcutorture</tt> was no surprise,
the current immense popularity of the Linux kernel is posing
interesting—and perhaps unprecedented—validation
challenges.
To see this, keep in mind that there are well over one billion
instances of the Linux kernel running today, given Android
smartphones, Linux-powered televisions, and servers.
This number can be expected to increase sharply with the advent of
the celebrated Internet of Things.
<p>
Suppose that RCU contains a race condition that manifests on average
once per million years of runtime.
This bug will be occurring about three times per <i>day</i> across
the installed base.
RCU could simply hide behind hardware error rates, given that no one
should really expect their smartphone to last for a million years.
However, anyone taking too much comfort from this thought should
consider the fact that in most jurisdictions, a successful multi-year
test of a given mechanism, which might include a Linux kernel,
suffices for a number of types of safety-critical certifications.
In fact, rumor has it that the Linux kernel is already being used
in production for safety-critical applications.
I don't know about you, but I would feel quite bad if a bug in RCU
killed someone.
Which might explain my recent focus on validation and verification.
<h2><a name="Other RCU Flavors">Other RCU Flavors</a></h2>
<p>
One of the more surprising things about RCU is that there are now
no fewer than five <i>flavors</i>, or API families.
In addition, the primary flavor that has been the sole focus up to
this point has two different implementations, non-preemptible and
preemptible.
The other four flavors are listed below, with requirements for each
described in a separate section.
<ol>
<li> <a href="#Bottom-Half Flavor">Bottom-Half Flavor</a>
<li> <a href="#Sched Flavor">Sched Flavor</a>
<li> <a href="#Sleepable RCU">Sleepable RCU</a>
<li> <a href="#Tasks RCU">Tasks RCU</a>
<li> <a href="#Waiting for Multiple Grace Periods">
Waiting for Multiple Grace Periods</a>
</ol>
<h3><a name="Bottom-Half Flavor">Bottom-Half Flavor</a></h3>
<p>
The softirq-disable (AKA “bottom-half”,
hence the “_bh” abbreviations)
flavor of RCU, or <i>RCU-bh</i>, was developed by
Dipankar Sarma to provide a flavor of RCU that could withstand the
network-based denial-of-service attacks researched by Robert
Olsson.
These attacks placed so much networking load on the system
that some of the CPUs never exited softirq execution,
which in turn prevented those CPUs from ever executing a context switch,
which, in the RCU implementation of that time, prevented grace periods
from ever ending.
The result was an out-of-memory condition and a system hang.
<p>
The solution was the creation of RCU-bh, which does
<tt>local_bh_disable()</tt>
across its read-side critical sections, and which uses the transition
from one type of softirq processing to another as a quiescent state
in addition to context switch, idle, user mode, and offline.
This means that RCU-bh grace periods can complete even when some of
the CPUs execute in softirq indefinitely, thus allowing algorithms
based on RCU-bh to withstand network-based denial-of-service attacks.
<p>
Because
<tt>rcu_read_lock_bh()</tt> and <tt>rcu_read_unlock_bh()</tt>
disable and re-enable softirq handlers, any attempt to start a softirq
handlers during the
RCU-bh read-side critical section will be deferred.
In this case, <tt>rcu_read_unlock_bh()</tt>
will invoke softirq processing, which can take considerable time.
One can of course argue that this softirq overhead should be associated
with the code following the RCU-bh read-side critical section rather
than <tt>rcu_read_unlock_bh()</tt>, but the fact
is that most profiling tools cannot be expected to make this sort
of fine distinction.
For example, suppose that a three-millisecond-long RCU-bh read-side
critical section executes during a time of heavy networking load.
There will very likely be an attempt to invoke at least one softirq
handler during that three milliseconds, but any such invocation will
be delayed until the time of the <tt>rcu_read_unlock_bh()</tt>.
This can of course make it appear at first glance as if
<tt>rcu_read_unlock_bh()</tt> was executing very slowly.
<p>
The
<a href="https://lwn.net/Articles/609973/#RCU Per-Flavor API Table">RCU-bh API</a>
includes
<tt>rcu_read_lock_bh()</tt>,
<tt>rcu_read_unlock_bh()</tt>,
<tt>rcu_dereference_bh()</tt>,
<tt>rcu_dereference_bh_check()</tt>,
<tt>synchronize_rcu_bh()</tt>,
<tt>synchronize_rcu_bh_expedited()</tt>,
<tt>call_rcu_bh()</tt>,
<tt>rcu_barrier_bh()</tt>, and
<tt>rcu_read_lock_bh_held()</tt>.
<h3><a name="Sched Flavor">Sched Flavor</a></h3>
<p>
Before preemptible RCU, waiting for an RCU grace period had the
side effect of also waiting for all pre-existing interrupt
and NMI handlers.
However, there are legitimate preemptible-RCU implementations that
do not have this property, given that any point in the code outside
of an RCU read-side critical section can be a quiescent state.
Therefore, <i>RCU-sched</i> was created, which follows “classic”
RCU in that an RCU-sched grace period waits for for pre-existing
interrupt and NMI handlers.
In kernels built with <tt>CONFIG_PREEMPT=n</tt>, the RCU and RCU-sched
APIs have identical implementations, while kernels built with
<tt>CONFIG_PREEMPT=y</tt> provide a separate implementation for each.
<p>
Note well that in <tt>CONFIG_PREEMPT=y</tt> kernels,
<tt>rcu_read_lock_sched()</tt> and <tt>rcu_read_unlock_sched()</tt>
disable and re-enable preemption, respectively.
This means that if there was a preemption attempt during the
RCU-sched read-side critical section, <tt>rcu_read_unlock_sched()</tt>
will enter the scheduler, with all the latency and overhead entailed.
Just as with <tt>rcu_read_unlock_bh()</tt>, this can make it look
as if <tt>rcu_read_unlock_sched()</tt> was executing very slowly.
However, the highest-priority task won't be preempted, so that task
will enjoy low-overhead <tt>rcu_read_unlock_sched()</tt> invocations.
<p>
The
<a href="https://lwn.net/Articles/609973/#RCU Per-Flavor API Table">RCU-sched API</a>
includes
<tt>rcu_read_lock_sched()</tt>,
<tt>rcu_read_unlock_sched()</tt>,
<tt>rcu_read_lock_sched_notrace()</tt>,
<tt>rcu_read_unlock_sched_notrace()</tt>,
<tt>rcu_dereference_sched()</tt>,
<tt>rcu_dereference_sched_check()</tt>,
<tt>synchronize_sched()</tt>,
<tt>synchronize_rcu_sched_expedited()</tt>,
<tt>call_rcu_sched()</tt>,
<tt>rcu_barrier_sched()</tt>, and
<tt>rcu_read_lock_sched_held()</tt>.
However, anything that disables preemption also marks an RCU-sched
read-side critical section, including
<tt>preempt_disable()</tt> and <tt>preempt_enable()</tt>,
<tt>local_irq_save()</tt> and <tt>local_irq_restore()</tt>,
and so on.
<h3><a name="Sleepable RCU">Sleepable RCU</a></h3>
<p>
For well over a decade, someone saying “I need to block within
an RCU read-side critical section” was a reliable indication
that this someone did not understand RCU.
After all, if you are always blocking in an RCU read-side critical
section, you can probably afford to use a higher-overhead synchronization
mechanism.
However, that changed with the advent of the Linux kernel's notifiers,
whose RCU read-side critical
sections almost never sleep, but sometimes need to.
This resulted in the introduction of
<a href="https://lwn.net/Articles/202847/">sleepable RCU</a>,
or <i>SRCU</i>.
<p>
SRCU allows different domains to be defined, with each such domain
defined by an instance of an <tt>srcu_struct</tt> structure.
A pointer to this structure must be passed in to each SRCU function,
for example, <tt>synchronize_srcu(&ss)</tt>, where
<tt>ss</tt> is the <tt>srcu_struct</tt> structure.
The key benefit of these domains is that a slow SRCU reader in one
domain does not delay an SRCU grace period in some other domain.
That said, one consequence of these domains is that read-side code
must pass a “cookie” from <tt>srcu_read_lock()</tt>
to <tt>srcu_read_unlock()</tt>, for example, as follows:
<blockquote>
<pre>
1 int idx;
2
3 idx = srcu_read_lock(&ss);
4 do_something();
5 srcu_read_unlock(&ss, idx);
</pre>
</blockquote>
<p>
As noted above, it is legal to block within SRCU read-side critical sections,
however, with great power comes great responsibility.
If you block forever in one of a given domain's SRCU read-side critical
sections, then that domain's grace periods will also be blocked forever.
Of course, one good way to block forever is to deadlock, which can
happen if any operation in a given domain's SRCU read-side critical
section can block waiting, either directly or indirectly, for that domain's
grace period to elapse.
For example, this results in a self-deadlock:
<blockquote>
<pre>
1 int idx;
2
3 idx = srcu_read_lock(&ss);
4 do_something();
5 synchronize_srcu(&ss);
6 srcu_read_unlock(&ss, idx);
</pre>
</blockquote>
<p>
However, if line 5 acquired a mutex that was held across
a <tt>synchronize_srcu()</tt> for domain <tt>ss</tt>,
deadlock would still be possible.
Furthermore, if line 5 acquired a mutex that was held across
a <tt>synchronize_srcu()</tt> for some other domain <tt>ss1</tt>,
and if an <tt>ss1</tt>-domain SRCU read-side critical section
acquired another mutex that was held across as <tt>ss</tt>-domain
<tt>synchronize_srcu()</tt>,
deadlock would again be possible.
Such a deadlock cycle could extend across an arbitrarily large number
of different SRCU domains.
Again, with great power comes great responsibility.
<p>
Unlike the other RCU flavors, SRCU read-side critical sections can
run on idle and even offline CPUs.
This ability requires that <tt>srcu_read_lock()</tt> and
<tt>srcu_read_unlock()</tt> contain memory barriers, which means
that SRCU readers will run a bit slower than would RCU readers.
It also motivates the <tt>smp_mb__after_srcu_read_unlock()</tt>
API, which, in combination with <tt>srcu_read_unlock()</tt>,
guarantees a full memory barrier.
<p>
The
<a href="https://lwn.net/Articles/609973/#RCU Per-Flavor API Table">SRCU API</a>
includes
<tt>srcu_read_lock()</tt>,
<tt>srcu_read_unlock()</tt>,
<tt>srcu_dereference()</tt>,
<tt>srcu_dereference_check()</tt>,
<tt>synchronize_srcu()</tt>,
<tt>synchronize_srcu_expedited()</tt>,
<tt>call_srcu()</tt>,
<tt>srcu_barrier()</tt>, and
<tt>srcu_read_lock_held()</tt>.
It also includes
<tt>DEFINE_SRCU()</tt>,
<tt>DEFINE_STATIC_SRCU()</tt>, and
<tt>init_srcu_struct()</tt>
APIs for defining and initializing <tt>srcu_struct</tt> structures.
<h3><a name="Tasks RCU">Tasks RCU</a></h3>
<p>
Some forms of tracing use “trampolines” to handle the
binary rewriting required to install different types of probes.
It would be good to be able to free old trampolines, which sounds
like a job for some form of RCU.
However, because it is necessary to be able to install a trace
anywhere in the code, it is not possible to use read-side markers
such as <tt>rcu_read_lock()</tt> and <tt>rcu_read_unlock()</tt>.
In addition, it does not work to have these markers in the trampoline
itself, because there would need to be instructions following
<tt>rcu_read_unlock()</tt>.
Although <tt>synchronize_rcu()</tt> would guarantee that execution
reached the <tt>rcu_read_unlock()</tt>, it would not be able to
guarantee that execution had completely left the trampoline.
<p>
The solution, in the form of
<a href="https://lwn.net/Articles/607117/"><i>Tasks RCU</i></a>,
is to have implicit
read-side critical sections that are delimited by voluntary context
switches, that is, calls to <tt>schedule()</tt>,
<tt>cond_resched_rcu_qs()</tt>, and
<tt>synchronize_rcu_tasks()</tt>.
In addition, transitions to and from userspace execution also delimit
tasks-RCU read-side critical sections.
<p>
The tasks-RCU API is quite compact, consisting only of
<tt>call_rcu_tasks()</tt>,
<tt>synchronize_rcu_tasks()</tt>, and
<tt>rcu_barrier_tasks()</tt>.
<h3><a name="Waiting for Multiple Grace Periods">
Waiting for Multiple Grace Periods</a></h3>
<p>
Perhaps you have an RCU protected data structure that is accessed from
RCU read-side critical sections, from softirq handlers, and from
hardware interrupt handlers.
That is three flavors of RCU, the normal flavor, the bottom-half flavor,
and the sched flavor.
How to wait for a compound grace period?
<p>
The best approach is usually to “just say no!” and
insert <tt>rcu_read_lock()</tt> and <tt>rcu_read_unlock()</tt>
around each RCU read-side critical section, regardless of what
environment it happens to be in.
But suppose that some of the RCU read-side critical sections are
on extremely hot code paths, and that use of <tt>CONFIG_PREEMPT=n</tt>
is not a viable option, so that <tt>rcu_read_lock()</tt> and
<tt>rcu_read_unlock()</tt> are not free.
What then?
<p>
You <i>could</i> wait on all three grace periods in succession, as follows:
<blockquote>
<pre>
1 synchronize_rcu();
2 synchronize_rcu_bh();
3 synchronize_sched();
</pre>
</blockquote>
<p>
This works, but triples the update-side latency penalty.
In cases where this is not acceptable, <tt>synchronize_rcu_mult()</tt>
may be used to wait on all three flavors of grace period concurrently:
<blockquote>
<pre>
1 synchronize_rcu_mult(call_rcu, call_rcu_bh, call_rcu_sched);
</pre>
</blockquote>
<p>
But what if it is necessary to also wait on SRCU?
This can be done as follows:
<blockquote>
<pre>
1 static void call_my_srcu(struct rcu_head *head,
2 void (*func)(struct rcu_head *head))
3 {
4 call_srcu(&my_srcu, head, func);
5 }
6
7 synchronize_rcu_mult(call_rcu, call_rcu_bh, call_rcu_sched, call_my_srcu);
</pre>
</blockquote>
<p>
If you needed to wait on multiple different flavors of SRCU
(but why???), you would need to create a wrapper function resembling
<tt>call_my_srcu()</tt> for each SRCU flavor.
<table>
<tr><th> </th></tr>
<tr><th align="left">Quick Quiz:</th></tr>
<tr><td>
But what if I need to wait for multiple RCU flavors, but I also need
the grace periods to be expedited?
</td></tr>
<tr><th align="left">Answer:</th></tr>
<tr><td bgcolor="#ffffff"><font color="ffffff">
If you are using expedited grace periods, there should be less penalty
for waiting on them in succession.
But if that is nevertheless a problem, you can use workqueues
or multiple kthreads to wait on the various expedited grace
periods concurrently.
</font></td></tr>
<tr><td> </td></tr>
</table>
<p>
Again, it is usually better to adjust the RCU read-side critical sections
to use a single flavor of RCU, but when this is not feasible, you can use
<tt>synchronize_rcu_mult()</tt>.
<h2><a name="Possible Future Changes">Possible Future Changes</a></h2>
<p>
One of the tricks that RCU uses to attain update-side scalability is
to increase grace-period latency with increasing numbers of CPUs.
If this becomes a serious problem, it will be necessary to rework the
grace-period state machine so as to avoid the need for the additional
latency.
<p>
Expedited grace periods scan the CPUs, so their latency and overhead
increases with increasing numbers of CPUs.
If this becomes a serious problem on large systems, it will be necessary
to do some redesign to avoid this scalability problem.
<p>
RCU disables CPU hotplug in a few places, perhaps most notably in the
expedited grace-period and <tt>rcu_barrier()</tt> operations.
If there is a strong reason to use expedited grace periods in CPU-hotplug
notifiers, it will be necessary to avoid disabling CPU hotplug.
This would introduce some complexity, so there had better be a <i>very</i>
good reason.
<p>
The tradeoff between grace-period latency on the one hand and interruptions
of other CPUs on the other hand may need to be re-examined.
The desire is of course for zero grace-period latency as well as zero
interprocessor interrupts undertaken during an expedited grace period
operation.
While this ideal is unlikely to be achievable, it is quite possible that
further improvements can be made.
<p>
The multiprocessor implementations of RCU use a combining tree that
groups CPUs so as to reduce lock contention and increase cache locality.
However, this combining tree does not spread its memory across NUMA
nodes nor does it align the CPU groups with hardware features such
as sockets or cores.
Such spreading and alignment is currently believed to be unnecessary
because the hotpath read-side primitives do not access the combining
tree, nor does <tt>call_rcu()</tt> in the common case.
If you believe that your architecture needs such spreading and alignment,
then your architecture should also benefit from the
<tt>rcutree.rcu_fanout_leaf</tt> boot parameter, which can be set
to the number of CPUs in a socket, NUMA node, or whatever.
If the number of CPUs is too large, use a fraction of the number of
CPUs.
If the number of CPUs is a large prime number, well, that certainly
is an “interesting” architectural choice!
More flexible arrangements might be considered, but only if
<tt>rcutree.rcu_fanout_leaf</tt> has proven inadequate, and only
if the inadequacy has been demonstrated by a carefully run and
realistic system-level workload.
<p>
Please note that arrangements that require RCU to remap CPU numbers will
require extremely good demonstration of need and full exploration of
alternatives.
<p>
There is an embarrassingly large number of flavors of RCU, and this
number has been increasing over time.
Perhaps it will be possible to combine some at some future date.
<p>
RCU's various kthreads are reasonably recent additions.
It is quite likely that adjustments will be required to more gracefully
handle extreme loads.
It might also be necessary to be able to relate CPU utilization by
RCU's kthreads and softirq handlers to the code that instigated this
CPU utilization.
For example, RCU callback overhead might be charged back to the
originating <tt>call_rcu()</tt> instance, though probably not
in production kernels.
<h2><a name="Summary">Summary</a></h2>
<p>
This document has presented more than two decade's worth of RCU
requirements.
Given that the requirements keep changing, this will not be the last
word on this subject, but at least it serves to get an important
subset of the requirements set forth.
<h2><a name="Acknowledgments">Acknowledgments</a></h2>
I am grateful to Steven Rostedt, Lai Jiangshan, Ingo Molnar,
Oleg Nesterov, Borislav Petkov, Peter Zijlstra, Boqun Feng, and
Andy Lutomirski for their help in rendering
this article human readable, and to Michelle Rankin for her support
of this effort.
Other contributions are acknowledged in the Linux kernel's git archive.
The cartoon is copyright (c) 2013 by Melissa Broussard,
and is provided
under the terms of the Creative Commons Attribution-Share Alike 3.0
United States license.
</body></html>
|
